NFSv4 S. Shepler
Internet-Draft Editor
Intended status: Standards Track March 6, 2006
Expires: September 7, 2006
NFSv4 Minor Version 1
draft-ietf-nfsv4-minorversion1-02.txt
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Copyright Notice
Copyright (C) The Internet Society (2006).
Abstract
This Internet-Draft describes the NFSv4 minor version 1 protocol
extensions. These most significant of these extensions are commonly
called: Sessions, Directory Delegations, and parallel NFS or pNFS
Requirements Language
The key words "MUST", "MUST NOT", "REQUIRED", "SHALL", "SHALL NOT",
"SHOULD", "SHOULD NOT", "RECOMMENDED", "MAY", and "OPTIONAL" in this
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document are to be interpreted as described in RFC 2119 [1].
Table of Contents
1. Protocol Data Types . . . . . . . . . . . . . . . . . . . . . 9
1.1. Basic Data Types . . . . . . . . . . . . . . . . . . . . 9
1.2. Structured Data Types . . . . . . . . . . . . . . . . . 10
2. Filehandles . . . . . . . . . . . . . . . . . . . . . . . . . 19
2.1. Obtaining the First Filehandle . . . . . . . . . . . . . 19
2.1.1. Root Filehandle . . . . . . . . . . . . . . . . . . 20
2.1.2. Public Filehandle . . . . . . . . . . . . . . . . . 20
2.2. Filehandle Types . . . . . . . . . . . . . . . . . . . . 20
2.2.1. General Properties of a Filehandle . . . . . . . . . 21
2.2.2. Persistent Filehandle . . . . . . . . . . . . . . . 22
2.2.3. Volatile Filehandle . . . . . . . . . . . . . . . . 22
2.3. One Method of Constructing a Volatile Filehandle . . . . 23
2.4. Client Recovery from Filehandle Expiration . . . . . . . 24
3. File Attributes . . . . . . . . . . . . . . . . . . . . . . . 25
3.1. Mandatory Attributes . . . . . . . . . . . . . . . . . . 26
3.2. Recommended Attributes . . . . . . . . . . . . . . . . . 26
3.3. Named Attributes . . . . . . . . . . . . . . . . . . . . 27
3.4. Classification of Attributes . . . . . . . . . . . . . . 27
3.5. Mandatory Attributes - Definitions . . . . . . . . . . . 28
3.6. Recommended Attributes - Definitions . . . . . . . . . . 30
3.7. Time Access . . . . . . . . . . . . . . . . . . . . . . 38
3.8. Interpreting owner and owner_group . . . . . . . . . . . 38
3.9. Character Case Attributes . . . . . . . . . . . . . . . 40
3.10. Quota Attributes . . . . . . . . . . . . . . . . . . . . 40
3.11. mounted_on_fileid . . . . . . . . . . . . . . . . . . . 41
3.12. send_impl_id and recv_impl_id . . . . . . . . . . . . . 42
3.13. fs_layouttype . . . . . . . . . . . . . . . . . . . . . 43
3.14. layouttype . . . . . . . . . . . . . . . . . . . . . . . 43
3.15. layouthint . . . . . . . . . . . . . . . . . . . . . . . 43
3.16. Access Control Lists . . . . . . . . . . . . . . . . . . 43
3.16.1. ACE type . . . . . . . . . . . . . . . . . . . . . . 45
3.16.2. ACE Access Mask . . . . . . . . . . . . . . . . . . 46
3.16.3. ACE flag . . . . . . . . . . . . . . . . . . . . . . 51
3.16.4. ACE who . . . . . . . . . . . . . . . . . . . . . . 53
3.16.5. Mode Attribute . . . . . . . . . . . . . . . . . . . 54
3.16.6. Interaction Between Mode and ACL Attributes . . . . 55
4. Filesystem Migration and Replication . . . . . . . . . . . . 69
4.1. Replication . . . . . . . . . . . . . . . . . . . . . . 69
4.2. Migration . . . . . . . . . . . . . . . . . . . . . . . 70
4.3. Interpretation of the fs_locations Attribute . . . . . . 70
4.4. Filehandle Recovery for Migration or Replication . . . . 72
5. NFS Server Name Space . . . . . . . . . . . . . . . . . . . . 72
5.1. Server Exports . . . . . . . . . . . . . . . . . . . . . 72
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5.2. Browsing Exports . . . . . . . . . . . . . . . . . . . . 72
5.3. Server Pseudo Filesystem . . . . . . . . . . . . . . . . 73
5.4. Multiple Roots . . . . . . . . . . . . . . . . . . . . . 73
5.5. Filehandle Volatility . . . . . . . . . . . . . . . . . 74
5.6. Exported Root . . . . . . . . . . . . . . . . . . . . . 74
5.7. Mount Point Crossing . . . . . . . . . . . . . . . . . . 74
5.8. Security Policy and Name Space Presentation . . . . . . 75
6. File Locking and Share Reservations . . . . . . . . . . . . . 76
6.1. Locking . . . . . . . . . . . . . . . . . . . . . . . . 76
6.1.1. Client ID . . . . . . . . . . . . . . . . . . . . . 77
6.1.2. Server Release of Clientid . . . . . . . . . . . . . 79
6.1.3. lock_owner and stateid Definition . . . . . . . . . 80
6.1.4. Use of the stateid and Locking . . . . . . . . . . . 82
6.1.5. Sequencing of Lock Requests . . . . . . . . . . . . 84
6.1.6. Recovery from Replayed Requests . . . . . . . . . . 85
6.1.7. Releasing lock_owner State . . . . . . . . . . . . . 85
6.1.8. Use of Open Confirmation . . . . . . . . . . . . . . 85
6.2. Lock Ranges . . . . . . . . . . . . . . . . . . . . . . 87
6.3. Upgrading and Downgrading Locks . . . . . . . . . . . . 87
6.4. Blocking Locks . . . . . . . . . . . . . . . . . . . . . 87
6.5. Lease Renewal . . . . . . . . . . . . . . . . . . . . . 88
6.6. Crash Recovery . . . . . . . . . . . . . . . . . . . . . 89
6.6.1. Client Failure and Recovery . . . . . . . . . . . . 89
6.6.2. Server Failure and Recovery . . . . . . . . . . . . 90
6.6.3. Network Partitions and Recovery . . . . . . . . . . 92
6.7. Recovery from a Lock Request Timeout or Abort . . . . . 95
6.8. Server Revocation of Locks . . . . . . . . . . . . . . . 96
6.9. Share Reservations . . . . . . . . . . . . . . . . . . . 97
6.10. OPEN/CLOSE Operations . . . . . . . . . . . . . . . . . 97
6.10.1. Close and Retention of State Information . . . . . . 98
6.11. Open Upgrade and Downgrade . . . . . . . . . . . . . . . 99
6.12. Short and Long Leases . . . . . . . . . . . . . . . . . 99
6.13. Clocks, Propagation Delay, and Calculating Lease
Expiration . . . . . . . . . . . . . . . . . . . . . . . 100
6.14. Migration, Replication and State . . . . . . . . . . . . 100
6.14.1. Migration and State . . . . . . . . . . . . . . . . 101
6.14.2. Replication and State . . . . . . . . . . . . . . . 102
6.14.3. Notification of Migrated Lease . . . . . . . . . . . 102
6.14.4. Migration and the Lease_time Attribute . . . . . . . 103
7. Client-Side Caching . . . . . . . . . . . . . . . . . . . . . 103
7.1. Performance Challenges for Client-Side Caching . . . . . 104
7.2. Delegation and Callbacks . . . . . . . . . . . . . . . . 105
7.2.1. Delegation Recovery . . . . . . . . . . . . . . . . 106
7.3. Data Caching . . . . . . . . . . . . . . . . . . . . . . 108
7.3.1. Data Caching and OPENs . . . . . . . . . . . . . . . 108
7.3.2. Data Caching and File Locking . . . . . . . . . . . 109
7.3.3. Data Caching and Mandatory File Locking . . . . . . 111
7.3.4. Data Caching and File Identity . . . . . . . . . . . 111
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7.4. Open Delegation . . . . . . . . . . . . . . . . . . . . 112
7.4.1. Open Delegation and Data Caching . . . . . . . . . . 115
7.4.2. Open Delegation and File Locks . . . . . . . . . . . 116
7.4.3. Handling of CB_GETATTR . . . . . . . . . . . . . . . 116
7.4.4. Recall of Open Delegation . . . . . . . . . . . . . 119
7.4.5. Clients that Fail to Honor Delegation Recalls . . . 121
7.4.6. Delegation Revocation . . . . . . . . . . . . . . . 122
7.5. Data Caching and Revocation . . . . . . . . . . . . . . 122
7.5.1. Revocation Recovery for Write Open Delegation . . . 123
7.6. Attribute Caching . . . . . . . . . . . . . . . . . . . 124
7.7. Data and Metadata Caching and Memory Mapped Files . . . 126
7.8. Name Caching . . . . . . . . . . . . . . . . . . . . . . 128
7.9. Directory Caching . . . . . . . . . . . . . . . . . . . 129
8. Security Negotiation . . . . . . . . . . . . . . . . . . . . 130
9. Clarification of Security Negotiation in NFSv4.1 . . . . . . 130
9.1. PUTFH + LOOKUP . . . . . . . . . . . . . . . . . . . . . 130
9.2. PUTFH + LOOKUPP . . . . . . . . . . . . . . . . . . . . 131
9.3. PUTFH + SECINFO . . . . . . . . . . . . . . . . . . . . 131
9.4. PUTFH + Anything Else . . . . . . . . . . . . . . . . . 131
10. NFSv4.1 Sessions . . . . . . . . . . . . . . . . . . . . . . 132
10.1. Sessions Background . . . . . . . . . . . . . . . . . . 132
10.1.1. Introduction to Sessions . . . . . . . . . . . . . . 132
10.1.2. Motivation . . . . . . . . . . . . . . . . . . . . . 133
10.1.3. Problem Statement . . . . . . . . . . . . . . . . . 134
10.1.4. NFSv4 Session Extension Characteristics . . . . . . 136
10.2. Transport Issues . . . . . . . . . . . . . . . . . . . . 136
10.2.1. Session Model . . . . . . . . . . . . . . . . . . . 136
10.2.2. Connection State . . . . . . . . . . . . . . . . . . 137
10.2.3. NFSv4 Channels, Sessions and Connections . . . . . . 138
10.2.4. Reconnection, Trunking and Failover . . . . . . . . 140
10.2.5. Server Duplicate Request Cache . . . . . . . . . . . 141
10.3. Session Initialization and Transfer Models . . . . . . . 142
10.3.1. Session Negotiation . . . . . . . . . . . . . . . . 142
10.3.2. RDMA Requirements . . . . . . . . . . . . . . . . . 144
10.3.3. RDMA Connection Resources . . . . . . . . . . . . . 144
10.3.4. TCP and RDMA Inline Transfer Model . . . . . . . . . 145
10.3.5. RDMA Direct Transfer Model . . . . . . . . . . . . . 148
10.4. Connection Models . . . . . . . . . . . . . . . . . . . 151
10.4.1. TCP Connection Model . . . . . . . . . . . . . . . . 152
10.4.2. Negotiated RDMA Connection Model . . . . . . . . . . 153
10.4.3. Automatic RDMA Connection Model . . . . . . . . . . 154
10.5. Buffer Management, Transfer, Flow Control . . . . . . . 154
10.6. Retry and Replay . . . . . . . . . . . . . . . . . . . . 157
10.7. The Back Channel . . . . . . . . . . . . . . . . . . . . 158
10.8. COMPOUND Sizing Issues . . . . . . . . . . . . . . . . . 159
10.9. Data Alignment . . . . . . . . . . . . . . . . . . . . . 159
10.10. NFSv4 Integration . . . . . . . . . . . . . . . . . . . 161
10.10.1. Minor Versioning . . . . . . . . . . . . . . . . . . 161
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10.10.2. Slot Identifiers and Server Duplicate Request
Cache . . . . . . . . . . . . . . . . . . . . . . . 161
10.10.3. Resolving server callback races with sessions . . . 165
10.10.4. COMPOUND and CB_COMPOUND . . . . . . . . . . . . . . 166
10.10.5. eXternal Data Representation Efficiency . . . . . . 167
10.10.6. Effect of Sessions on Existing Operations . . . . . 167
10.10.7. Authentication Efficiencies . . . . . . . . . . . . 168
10.11. Sessions Security Considerations . . . . . . . . . . . . 169
10.11.1. Authentication . . . . . . . . . . . . . . . . . . . 171
11. Directory Delegations . . . . . . . . . . . . . . . . . . . . 172
11.1. Introduction to Directory Delegations . . . . . . . . . 172
11.2. Directory Delegation Design (in brief) . . . . . . . . . 173
11.3. Recommended Attributes in support of Directory
Delegations . . . . . . . . . . . . . . . . . . . . . . 174
11.4. Delegation Recall . . . . . . . . . . . . . . . . . . . 175
11.5. Delegation Recovery . . . . . . . . . . . . . . . . . . 175
12. Introduction . . . . . . . . . . . . . . . . . . . . . . . . 175
13. General Definitions . . . . . . . . . . . . . . . . . . . . . 178
13.1. Metadata Server . . . . . . . . . . . . . . . . . . . . 178
13.2. Client . . . . . . . . . . . . . . . . . . . . . . . . . 178
13.3. Storage Device . . . . . . . . . . . . . . . . . . . . . 178
13.4. Storage Protocol . . . . . . . . . . . . . . . . . . . . 179
13.5. Control Protocol . . . . . . . . . . . . . . . . . . . . 179
13.6. Metadata . . . . . . . . . . . . . . . . . . . . . . . . 179
13.7. Layout . . . . . . . . . . . . . . . . . . . . . . . . . 180
14. pNFS protocol semantics . . . . . . . . . . . . . . . . . . . 180
14.1. Definitions . . . . . . . . . . . . . . . . . . . . . . 180
14.1.1. Layout Types . . . . . . . . . . . . . . . . . . . . 180
14.1.2. Layout Iomode . . . . . . . . . . . . . . . . . . . 181
14.1.3. Layout Segments . . . . . . . . . . . . . . . . . . 181
14.1.4. Device IDs . . . . . . . . . . . . . . . . . . . . . 182
14.1.5. Aggregation Schemes . . . . . . . . . . . . . . . . 183
14.2. Guarantees Provided by Layouts . . . . . . . . . . . . . 183
14.3. Getting a Layout . . . . . . . . . . . . . . . . . . . . 184
14.4. Committing a Layout . . . . . . . . . . . . . . . . . . 185
14.4.1. LAYOUTCOMMIT and mtime/atime/change . . . . . . . . 186
14.4.2. LAYOUTCOMMIT and size . . . . . . . . . . . . . . . 186
14.4.3. LAYOUTCOMMIT and layoutupdate . . . . . . . . . . . 187
14.5. Recalling a Layout . . . . . . . . . . . . . . . . . . . 187
14.5.1. Basic Operation . . . . . . . . . . . . . . . . . . 188
14.5.2. Recall Callback Robustness . . . . . . . . . . . . . 189
14.5.3. Recall/Return Sequencing . . . . . . . . . . . . . . 190
14.6. Metadata Server Write Propagation . . . . . . . . . . . 192
14.7. Crash Recovery . . . . . . . . . . . . . . . . . . . . . 193
14.7.1. Leases . . . . . . . . . . . . . . . . . . . . . . . 193
14.7.2. Client Recovery . . . . . . . . . . . . . . . . . . 194
14.7.3. Metadata Server Recovery . . . . . . . . . . . . . . 195
14.7.4. Storage Device Recovery . . . . . . . . . . . . . . 197
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15. Security Considerations . . . . . . . . . . . . . . . . . . . 198
15.1. File Layout Security . . . . . . . . . . . . . . . . . . 199
15.2. Object Layout Security . . . . . . . . . . . . . . . . . 199
15.3. Block/Volume Layout Security . . . . . . . . . . . . . . 201
16. The NFSv4 File Layout Type . . . . . . . . . . . . . . . . . 201
16.1. File Striping and Data Access . . . . . . . . . . . . . 202
16.1.1. Sparse and Dense Storage Device Data Layouts . . . . 203
16.1.2. Metadata and Storage Device Roles . . . . . . . . . 205
16.1.3. Device Multipathing . . . . . . . . . . . . . . . . 206
16.1.4. Operations Issued to Storage Devices . . . . . . . . 206
16.2. Global Stateid Requirements . . . . . . . . . . . . . . 207
16.3. The Layout Iomode . . . . . . . . . . . . . . . . . . . 207
16.4. Storage Device State Propagation . . . . . . . . . . . . 208
16.4.1. Lock State Propagation . . . . . . . . . . . . . . . 208
16.4.2. Open-mode Validation . . . . . . . . . . . . . . . . 209
16.4.3. File Attributes . . . . . . . . . . . . . . . . . . 209
16.5. Storage Device Component File Size . . . . . . . . . . . 210
16.6. Crash Recovery Considerations . . . . . . . . . . . . . 211
16.7. Security Considerations . . . . . . . . . . . . . . . . 211
16.8. Alternate Approaches . . . . . . . . . . . . . . . . . . 211
17. Layouts and Aggregation . . . . . . . . . . . . . . . . . . . 212
17.1. Simple Map . . . . . . . . . . . . . . . . . . . . . . . 213
17.2. Block Extent Map . . . . . . . . . . . . . . . . . . . . 213
17.3. Striped Map (RAID 0) . . . . . . . . . . . . . . . . . . 213
17.4. Replicated Map . . . . . . . . . . . . . . . . . . . . . 213
17.5. Concatenated Map . . . . . . . . . . . . . . . . . . . . 214
17.6. Nested Map . . . . . . . . . . . . . . . . . . . . . . . 214
18. Minor Versioning . . . . . . . . . . . . . . . . . . . . . . 214
19. Internationalization . . . . . . . . . . . . . . . . . . . . 216
19.1. Stringprep profile for the utf8str_cs type . . . . . . . 218
19.2. Stringprep profile for the utf8str_cis type . . . . . . 219
19.3. Stringprep profile for the utf8str_mixed type . . . . . 221
19.4. UTF-8 Related Errors . . . . . . . . . . . . . . . . . . 222
20. Error Definitions . . . . . . . . . . . . . . . . . . . . . . 222
21. NFS version 4.1 Procedures . . . . . . . . . . . . . . . . . 231
21.1. Procedure 0: NULL - No Operation . . . . . . . . . . . . 231
21.2. Procedure 1: COMPOUND - Compound Operations . . . . . . 232
22. NFS version 4.1 Operations . . . . . . . . . . . . . . . . . 234
22.1. Operation 3: ACCESS - Check Access Rights . . . . . . . 235
22.2. Operation 4: CLOSE - Close File . . . . . . . . . . . . 237
22.3. Operation 5: COMMIT - Commit Cached Data . . . . . . . . 238
22.4. Operation 6: CREATE - Create a Non-Regular File Object . 241
22.5. Operation 7: DELEGPURGE - Purge Delegations Awaiting
Recovery . . . . . . . . . . . . . . . . . . . . . . . . 244
22.6. Operation 8: DELEGRETURN - Return Delegation . . . . . . 245
22.7. Operation 9: GETATTR - Get Attributes . . . . . . . . . 245
22.8. Operation 10: GETFH - Get Current Filehandle . . . . . . 247
22.9. Operation 11: LINK - Create Link to a File . . . . . . . 248
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22.10. Operation 12: LOCK - Create Lock . . . . . . . . . . . . 249
22.11. Operation 13: LOCKT - Test For Lock . . . . . . . . . . 253
22.12. Operation 14: LOCKU - Unlock File . . . . . . . . . . . 255
22.13. Operation 15: LOOKUP - Lookup Filename . . . . . . . . . 256
22.14. Operation 16: LOOKUPP - Lookup Parent Directory . . . . 258
22.15. Operation 17: NVERIFY - Verify Difference in
Attributes . . . . . . . . . . . . . . . . . . . . . . . 259
22.16. Operation 18: OPEN - Open a Regular File . . . . . . . . 260
22.17. Operation 19: OPENATTR - Open Named Attribute
Directory . . . . . . . . . . . . . . . . . . . . . . . 269
22.18. Operation 20: OPEN_CONFIRM - Confirm Open . . . . . . . 271
22.19. Operation 21: OPEN_DOWNGRADE - Reduce Open File Access . 273
22.20. Operation 22: PUTFH - Set Current Filehandle . . . . . . 274
22.21. Operation 24: PUTROOTFH - Set Root Filehandle . . . . . 275
22.22. Operation 25: READ - Read from File . . . . . . . . . . 276
22.23. Operation 26: READDIR - Read Directory . . . . . . . . . 278
22.24. Operation 27: READLINK - Read Symbolic Link . . . . . . 282
22.25. Operation 28: REMOVE - Remove Filesystem Object . . . . 283
22.26. Operation 29: RENAME - Rename Directory Entry . . . . . 285
22.27. Operation 30: RENEW - Renew a Lease . . . . . . . . . . 287
22.28. Operation 31: RESTOREFH - Restore Saved Filehandle . . . 288
22.29. Operation 32: SAVEFH - Save Current Filehandle . . . . . 289
22.30. Operation 33: SECINFO - Obtain Available Security . . . 290
22.31. Operation 34: SETATTR - Set Attributes . . . . . . . . . 293
22.32. Operation 35: SETCLIENTID - Negotiate Clientid . . . . . 296
22.33. Operation 36: SETCLIENTID_CONFIRM - Confirm Clientid . . 300
22.34. Operation 37: VERIFY - Verify Same Attributes . . . . . 303
22.35. Operation 38: WRITE - Write to File . . . . . . . . . . 304
22.36. Operation 39: RELEASE_LOCKOWNER - Release Lockowner
State . . . . . . . . . . . . . . . . . . . . . . . . . 309
22.37. Operation 10044: ILLEGAL - Illegal operation . . . . . . 310
22.38. SECINFO_NO_NAME - Get Security on Unnamed Object . . . . 310
22.39. CREATECLIENTID - Instantiate Clientid . . . . . . . . . 312
22.40. CREATESESSION - Create New Session and Confirm
Clientid . . . . . . . . . . . . . . . . . . . . . . . . 317
22.41. BIND_BACKCHANNEL - Create a callback channel binding . . 322
22.42. DESTROYSESSION - Destroy existing session . . . . . . . 324
22.43. SEQUENCE - Supply per-procedure sequencing and control . 325
22.44. GET_DIR_DELEGATION - Get a directory delegation . . . . 326
22.45. LAYOUTGET - Get Layout Information . . . . . . . . . . . 330
22.46. LAYOUTCOMMIT - Commit writes made using a layout . . . . 332
22.47. LAYOUTRETURN - Release Layout Information . . . . . . . 336
22.48. GETDEVICEINFO - Get Device Information . . . . . . . . . 337
22.49. GETDEVICELIST . . . . . . . . . . . . . . . . . . . . . 338
23. NFS version 4.1 Callback Procedures . . . . . . . . . . . . . 340
23.1. Procedure 0: CB_NULL - No Operation . . . . . . . . . . 340
23.2. Procedure 1: CB_COMPOUND - Compound Operations . . . . . 340
24. NFS version 4.1 Callback Operations . . . . . . . . . . . . . 342
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24.1. Operation 3: CB_GETATTR - Get Attributes . . . . . . . . 342
24.2. Operation 4: CB_RECALL - Recall an Open Delegation . . . 343
24.3. Operation 10044: CB_ILLEGAL - Illegal Callback
Operation . . . . . . . . . . . . . . . . . . . . . . . 344
24.4. CB_RECALLCREDIT - change flow control limits . . . . . . 345
24.5. CB_SEQUENCE - Supply callback channel sequencing and
control . . . . . . . . . . . . . . . . . . . . . . . . 346
24.6. CB_NOTIFY - Notify directory changes . . . . . . . . . . 348
24.7. CB_RECALL_ANY - Keep any N delegations . . . . . . . . . 351
24.8. CB_SIZECHANGED . . . . . . . . . . . . . . . . . . . . . 354
24.9. CB_LAYOUTRECALL . . . . . . . . . . . . . . . . . . . . 355
25. References . . . . . . . . . . . . . . . . . . . . . . . . . 357
25.1. Normative References . . . . . . . . . . . . . . . . . . 357
25.2. Informative References . . . . . . . . . . . . . . . . . 357
Appendix A. Acknowledgments . . . . . . . . . . . . . . . . . . 358
Author's Address . . . . . . . . . . . . . . . . . . . . . . . . 358
Intellectual Property and Copyright Statements . . . . . . . . . 360
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1. Protocol Data Types
The syntax and semantics to describe the data types of the NFS
version 4 protocol are defined in the XDR RFC1832 [2] and RPC RFC1831
[3] documents. The next sections build upon the XDR data types to
define types and structures specific to this protocol.
1.1. Basic Data Types
These are the base NFSv4 data types.
+---------------+---------------------------------------------------+
| Data Type | Definition |
+---------------+---------------------------------------------------+
| int32_t | typedef int int32_t; |
| uint32_t | typedef unsigned int uint32_t; |
| int64_t | typedef hyper int64_t; |
| uint64_t | typedef unsigned hyper uint64_t; |
| attrlist4 | typedef opaque attrlist4<> |
| | Used for file/directory attributes |
| bitmap4 | typedef uint32_t bitmap4<> |
| | Used in attribute array encoding. |
| changeid4 | typedef uint64_t changeid4; |
| | Used in definition of change_info |
| clientid4 | typedef uint64_t clientid4; |
| | Shorthand reference to client identification |
| component4 | typedef utf8str_cs component4; |
| | Represents path name components |
| count4 | typedef uint32_t count4; |
| | Various count parameters (READ, WRITE, COMMIT) |
| length4 | typedef uint64_t length4; |
| | Describes LOCK lengths |
| linktext4 | typedef utf8str_cs linktext4; |
| | Symbolic link contents |
| mode4 | typedef uint32_t mode4; |
| | Mode attribute data type |
| nfs_cookie4 | typedef uint64_t nfs_cookie4; |
| | Opaque cookie value for READDIR |
| nfs_fh4 | typedef opaque nfs_fh4<NFS4_FHSIZE> |
| | Filehandle definition; NFS4_FHSIZE is defined as |
| | 128 |
| nfs_ftype4 | enum nfs_ftype4; |
| | Various defined file types |
| nfsstat4 | enum nfsstat4; |
| | Return value for operations |
| offset4 | typedef uint64_t offset4; |
| | Various offset designations (READ, WRITE, LOCK, |
| | COMMIT) |
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| pathname4 | typedef component4 pathname4<> |
| | Represents path name for fs_locations |
| qop4 | typedef uint32_t qop4; |
| | Quality of protection designation in SECINFO |
| sec_oid4 | typedef opaque sec_oid4<> |
| | Security Object Identifier The sec_oid4 data type |
| | is not really opaque. Instead contains an ASN.1 |
| | OBJECT IDENTIFIER as used by GSS-API in the |
| | mech_type argument to GSS_Init_sec_context. See |
| | RFC2743 [4] for details. |
| seqid4 | typedef uint32_t seqid4; |
| | Sequence identifier used for file locking |
| utf8string | typedef opaque utf8string<> |
| | UTF-8 encoding for strings |
| utf8str_cis | typedef opaque utf8str_cis; |
| | Case-insensitive UTF-8 string |
| utf8str_cs | typedef opaque utf8str_cs; |
| | Case-sensitive UTF-8 string |
| utf8str_mixed | typedef opaque utf8str_mixed; |
| | UTF-8 strings with a case sensitive prefix and a |
| | case insensitive suffix. |
| verifier4 | typedef opaque verifier4[NFS4_VERIFIER_SIZE]; |
| | Verifier used for various operations (COMMIT, |
| | CREATE, OPEN, READDIR, SETCLIENTID, |
| | SETCLIENTID_CONFIRM, WRITE) NFS4_VERIFIER_SIZE is |
| | defined as 8. |
+---------------+---------------------------------------------------+
End of Base Data Types
Table 1
1.2. Structured Data Types
1.2.1. nfstime4
struct nfstime4 {
int64_t seconds;
uint32_t nseconds;
}
The nfstime4 structure gives the number of seconds and nanoseconds
since midnight or 0 hour January 1, 1970 Coordinated Universal Time
(UTC). Values greater than zero for the seconds field denote dates
after the 0 hour January 1, 1970. Values less than zero for the
seconds field denote dates before the 0 hour January 1, 1970. In
both cases, the nseconds field is to be added to the seconds field
for the final time representation. For example, if the time to be
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represented is one-half second before 0 hour January 1, 1970, the
seconds field would have a value of negative one (-1) and the
nseconds fields would have a value of one-half second (500000000).
Values greater than 999,999,999 for nseconds are considered invalid.
This data type is used to pass time and date information. A server
converts to and from its local representation of time when processing
time values, preserving as much accuracy as possible. If the
precision of timestamps stored for a filesystem object is less than
defined, loss of precision can occur. An adjunct time maintenance
protocol is recommended to reduce client and server time skew.
1.2.2. time_how4
enum time_how4 {
SET_TO_SERVER_TIME4 = 0,
SET_TO_CLIENT_TIME4 = 1
};
1.2.3. settime4
union settime4 switch (time_how4 set_it) {
case SET_TO_CLIENT_TIME4:
nfstime4 time;
default:
void;
};
The above definitions are used as the attribute definitions to set
time values. If set_it is SET_TO_SERVER_TIME4, then the server uses
its local representation of time for the time value.
1.2.4. specdata4
struct specdata4 {
uint32_t specdata1; /* major device number */
uint32_t specdata2; /* minor device number */
};
This data type represents additional information for the device file
types NF4CHR and NF4BLK.
1.2.5. fsid4
struct fsid4 {
uint64_t major;
uint64_t minor;
};
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1.2.6. fs_location4
struct fs_location4 {
utf8str_cis server<>
pathname4 rootpath;
};
1.2.7. fs_locations4
struct fs_locations4 {
pathname4 fs_root;
fs_location4 locations<>
};
The fs_location4 and fs_locations4 data types are used for the
fs_locations recommended attribute which is used for migration and
replication support.
1.2.8. fattr4
struct fattr4 {
bitmap4 attrmask;
attrlist4 attr_vals;
};
The fattr4 structure is used to represent file and directory
attributes.
The bitmap is a counted array of 32 bit integers used to contain bit
values. The position of the integer in the array that contains bit n
can be computed from the expression (n / 32) and its bit within that
integer is (n mod 32).
0 1
+-----------+-----------+-----------+--
| count | 31 .. 0 | 63 .. 32 |
+-----------+-----------+-----------+--
1.2.9. change_info4
struct change_info4 {
bool atomic;
changeid4 before;
changeid4 after;
};
This structure is used with the CREATE, LINK, REMOVE, RENAME
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operations to let the client know the value of the change attribute
for the directory in which the target filesystem object resides.
1.2.10. clientaddr4
struct clientaddr4 {
/* see struct rpcb in RFC1833 */
string r_netid<> /* network id */
string r_addr<> /* universal address */
};
The clientaddr4 structure is used as part of the SETCLIENTID
operation to either specify the address of the client that is using a
clientid or as part of the callback registration. The r_netid and
r_addr fields are specified in RFC1833 [9], but they are
underspecified in RFC1833 [9] as far as what they should look like
for specific protocols.
For TCP over IPv4 and for UDP over IPv4, the format of r_addr is the
US-ASCII string:
h1.h2.h3.h4.p1.p2
The prefix, "h1.h2.h3.h4", is the standard textual form for
representing an IPv4 address, which is always four octets long.
Assuming big-endian ordering, h1, h2, h3, and h4, are respectively,
the first through fourth octets each converted to ASCII-decimal.
Assuming big-endian ordering, p1 and p2 are, respectively, the first
and second octets each converted to ASCII-decimal. For example, if a
host, in big-endian order, has an address of 0x0A010307 and there is
a service listening on, in big endian order, port 0x020F (decimal
527), then complete universal address is "10.1.3.7.2.15".
For TCP over IPv4 the value of r_netid is the string "tcp". For UDP
over IPv4 the value of r_netid is the string "udp".
For TCP over IPv6 and for UDP over IPv6, the format of r_addr is the
US-ASCII string:
x1:x2:x3:x4:x5:x6:x7:x8.p1.p2
The suffix "p1.p2" is the service port, and is computed the same way
as with universal addresses for TCP and UDP over IPv4. The prefix,
"x1:x2:x3:x4:x5:x6:x7:x8", is the standard textual form for
representing an IPv6 address as defined in Section 2.2 of RFC1884
[5]. Additionally, the two alternative forms specified in Section
2.2 of RFC1884 [5] are also acceptable.
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For TCP over IPv6 the value of r_netid is the string "tcp6". For UDP
over IPv6 the value of r_netid is the string "udp6".
1.2.11. cb_client4
struct cb_client4 {
unsigned int cb_program;
clientaddr4 cb_location;
};
This structure is used by the client to inform the server of its call
back address; includes the program number and client address.
1.2.12. nfs_client_id4
struct nfs_client_id4 {
verifier4 verifier;
opaque id<NFS4_OPAQUE_LIMIT>
};
This structure is part of the arguments to the SETCLIENTID operation.
NFS4_OPAQUE_LIMIT is defined as 1024.
1.2.13. open_owner4
struct open_owner4 {
clientid4 clientid;
opaque owner<NFS4_OPAQUE_LIMIT>
};
This structure is used to identify the owner of open state.
NFS4_OPAQUE_LIMIT is defined as 1024.
1.2.14. lock_owner4
struct lock_owner4 {
clientid4 clientid;
opaque owner<NFS4_OPAQUE_LIMIT>
};
This structure is used to identify the owner of file locking state.
NFS4_OPAQUE_LIMIT is defined as 1024.
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1.2.15. open_to_lock_owner4
struct open_to_lock_owner4 {
seqid4 open_seqid;
stateid4 open_stateid;
seqid4 lock_seqid;
lock_owner4 lock_owner;
};
This structure is used for the first LOCK operation done for an
open_owner4. It provides both the open_stateid and lock_owner such
that the transition is made from a valid open_stateid sequence to
that of the new lock_stateid sequence. Using this mechanism avoids
the confirmation of the lock_owner/lock_seqid pair since it is tied
to established state in the form of the open_stateid/open_seqid.
1.2.16. stateid4
struct stateid4 {
uint32_t seqid;
opaque other[12];
};
This structure is used for the various state sharing mechanisms
between the client and server. For the client, this data structure
is read-only. The starting value of the seqid field is undefined.
The server is required to increment the seqid field monotonically at
each transition of the stateid. This is important since the client
will inspect the seqid in OPEN stateids to determine the order of
OPEN processing done by the server.
1.2.17. layouttype4
enum layouttype4 {
LAYOUT_NFSV4_FILES = 1,
LAYOUT_OSD2_OBJECTS = 2,
LAYOUT_BLOCK_VOLUME = 3
};
A layout type specifies the layout being used. The implication is
that clients have "layout drivers" that support one or more layout
types. The file server advertises the layout types it supports
through the LAYOUT_TYPES file system attribute. A client asks for
layouts of a particular type in LAYOUTGET, and passes those layouts
to its layout driver. The set of well known layout types must be
defined. As well, a private range of layout types is to be defined
by this document. This would allow custom installations to introduce
new layout types.
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[[Comment.1: Determine private range of layout types]]
New layout types must be specified in RFCs approved by the IESG
before becoming part of the pNFS specification.
The LAYOUT_NFSV4_FILES enumeration specifies that the NFSv4 file
layout type is to be used. The LAYOUT_OSD2_OBJECTS enumeration
specifies that the object layout, as defined in [10], is to be used.
Similarly, the LAYOUT_BLOCK_VOLUME enumeration that the block/volume
layout, as defined in [11], is to be used.
1.2.18. pnfs_deviceid4
typedef uint32_t pnfs_deviceid4; /* 32-bit device ID */
Layout information includes device IDs that specify a storage device
through a compact handle. Addressing and type information is
obtained with the GETDEVICEINFO operation. A client must not assume
that device IDs are valid across metadata server reboots. The device
ID is qualified by the layout type and are unique per file system
(FSID). This allows different layout drivers to generate device IDs
without the need for co-ordination. See Section 14.1.4 for more
details.
1.2.19. pnfs_deviceaddr4
struct pnfs_netaddr4 {
string r_netid<> /* network ID */
string r_addr<> /* universal address */
};
struct pnfs_deviceaddr4 {
pnfs_layouttype4 type;
opaque device_addr<>
};
The device address is used to set up a communication channel with the
storage device. Different layout types will require different types
of structures to define how they communicate with storage devices.
The opaque device_addr field must be interpreted based on the
specified layout type.
Currently, the only defined device address is that for the NFSv4 file
layout (struct pnfs_netaddr4), which identifies a storage device by
network IP address and port number. This is sufficient for the
clients to communicate with the NFSv4 storage devices, and may also
be sufficient for object-based storage drivers to communicate with
OSDs. The other device address we expect to support is a SCSI volume
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identifier. The final protocol specification will detail the allowed
values for device_type and the format of their associated location
information.
[NOTE: other device addresses will be added as the respective
specifications mature. It has been suggested that a separate
device_type enumeration is used as a switch to the pnfs_deviceaddr4
structure (e.g., if multiple types of addresses exist for the same
layout type). Until such a time as a real case is made and the
respective layout types have matured, the device address structure
will be left as is.]
1.2.20. pnfs_devlist_item4
struct pnfs_devlist_item4 {
pnfs_deviceid4 id;
pnfs_deviceaddr4 addr;
};
An array of these values is returned by the GETDEVICELIST operation.
They define the set of devices associated with a file system.
1.2.21. pnfs_layout4
struct pnfs_layout4 {
offset4 offset;
length4 length;
pnfs_layoutiomode4 iomode;
pnfs_layouttype4 type;
opaque layout<>;
};
The pnfs_layout4 structure defines a layout for a file. The layout
type specific data is opaque within this structure and must be
interepreted based on the layout type. Currently, only the NFSv4
file layout type is defined; see Section 16.1 for its definition.
Since layouts are sub-dividable, the offset and length together with
the file's filehandle, the clientid, iomode, and layout type,
identifies the layout.
[[Comment.2: there is a discussion of moving the striping
information, or more generally the "aggregation scheme", up to the
generic layout level. This creates a two-layer system where the top
level is a switch on different data placement layouts, and the next
level down is a switch on different data storage types. This lets
different layouts (e.g., striping or mirroring or redundant servers)
to be layered over different storage devices. This would move
geometry information out of nfsv4_file_layouttype4 and up into a
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generic pnfs_striped_layout type that would specify a set of
pnfs_deviceid4 and pnfs_devicetype4 to use for storage. Instead of
nfsv4_file_layouttype4, there would be pnfs_nfsv4_devicetype4.]]
1.2.22. pnfs_layoutupdate4
struct pnfs_layoutupdate4 {
pnfs_layouttype4 type;
opaque layoutupdate_data<>;
};
The pnfs_layoutupdate4 structure is used by the client to return
'updated' layout information to the metadata server at LAYOUTCOMMIT
time. This structure provides a channel to pass layout type specific
information back to the metadata server. E.g., for block/volume
layout types this could include the list of reserved blocks that were
written. The contents of the opaque layoutupdate_data argument are
determined by the layout type and are defined in their context. The
NFSv4 file-based layout does not use this structure, thus the
update_data field should have a zero length.
1.2.23. layouthint4
struct pnfs_layouthint4 {
pnfs_layouttype4 type;
opaque layouthint_data<>
};
The layouthint4 structure is used by the client to pass in a hint
about the type of layout it would like created for a particular file.
It is the structure specified by the FILE_LAYOUT_HINT attribute
described below. The metadata server may ignore the hint, or may
selectively ignore fields within the hint. This hint should be
provided at create time as part of the initial attributes within
OPEN. The NFSv4 file-based layout uses the "nfsv4_file_layouthint"
structure as defined in Section 16.1.
1.2.24. pnfs_layoutiomode4
enum pnfs_layoutiomode4 {
LAYOUTIOMODE_READ = 1,
LAYOUTIOMODE_RW = 2,
LAYOUTIOMODE_ANY = 3
};
The iomode specifies whether the client intends to read or write
(with the possibility of reading) the data represented by the layout.
The ANY iomode MUST NOT be used for LAYOUTGET, however, it can be
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used for LAYOUTRETURN and LAYOUTRECALL. The ANY iomode specifies
that layouts pertaining to both READ and RW iomodes are being
returned or recalled, respectively. The metadata server's use of the
iomode may depend on the layout type being used. The storage devices
may validate I/O accesses against the iomode and reject invalid
accesses.
1.2.25. nfs_impl_id4
struct nfs_impl_id4 {
utf8str_cis nii_domain;
utf8str_cs nii_name;
nfstime4 nii_date;
};
This structure is used to identify client and server implementation
detail. The nii_domain field is the DNS domain name that the
implementer is associated with. The nii_name field is the product
name of the implementation and is completely free form. It is
encouraged that the nii_name be used to distinguish machine
architecture, machine platforms, revisions, versions, and patch
levels. The nii_date field is the timestamp of when the software
instance was published or built.
1.2.26. impl_ident4
struct impl_ident4 {
clientid4 ii_clientid;
struct nfs_impl_id4 ii_impl_id;
};
This is used for exchanging implementation identification between
client and server.
2. Filehandles
The filehandle in the NFS protocol is a per server unique identifier
for a filesystem object. The contents of the filehandle are opaque
to the client. Therefore, the server is responsible for translating
the filehandle to an internal representation of the filesystem
object.
2.1. Obtaining the First Filehandle
The operations of the NFS protocol are defined in terms of one or
more filehandles. Therefore, the client needs a filehandle to
initiate communication with the server. With the NFS version 2
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protocol [RFC1094] and the NFS version 3 protocol [RFC1813], there
exists an ancillary protocol to obtain this first filehandle. The
MOUNT protocol, RPC program number 100005, provides the mechanism of
translating a string based filesystem path name to a filehandle which
can then be used by the NFS protocols.
The MOUNT protocol has deficiencies in the area of security and use
via firewalls. This is one reason that the use of the public
filehandle was introduced in [RFC2054] and [RFC2055]. With the use
of the public filehandle in combination with the LOOKUP operation in
the NFS version 2 and 3 protocols, it has been demonstrated that the
MOUNT protocol is unnecessary for viable interaction between NFS
client and server.
Therefore, the NFS version 4 protocol will not use an ancillary
protocol for translation from string based path names to a
filehandle. Two special filehandles will be used as starting points
for the NFS client.
2.1.1. Root Filehandle
The first of the special filehandles is the ROOT filehandle. The
ROOT filehandle is the "conceptual" root of the filesystem name space
at the NFS server. The client uses or starts with the ROOT
filehandle by employing the PUTROOTFH operation. The PUTROOTFH
operation instructs the server to set the "current" filehandle to the
ROOT of the server's file tree. Once this PUTROOTFH operation is
used, the client can then traverse the entirety of the server's file
tree with the LOOKUP operation. A complete discussion of the server
name space is in the section "NFS Server Name Space".
2.1.2. Public Filehandle
The second special filehandle is the PUBLIC filehandle. Unlike the
ROOT filehandle, the PUBLIC filehandle may be bound or represent an
arbitrary filesystem object at the server. The server is responsible
for this binding. It may be that the PUBLIC filehandle and the ROOT
filehandle refer to the same filesystem object. However, it is up to
the administrative software at the server and the policies of the
server administrator to define the binding of the PUBLIC filehandle
and server filesystem object. The client may not make any
assumptions about this binding. The client uses the PUBLIC
filehandle via the PUTPUBFH operation.
2.2. Filehandle Types
In the NFS version 2 and 3 protocols, there was one type of
filehandle with a single set of semantics. This type of filehandle
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is termed "persistent" in NFS Version 4. The semantics of a
persistent filehandle remain the same as before. A new type of
filehandle introduced in NFS Version 4 is the "volatile" filehandle,
which attempts to accommodate certain server environments.
The volatile filehandle type was introduced to address server
functionality or implementation issues which make correct
implementation of a persistent filehandle infeasible. Some server
environments do not provide a filesystem level invariant that can be
used to construct a persistent filehandle. The underlying server
filesystem may not provide the invariant or the server's filesystem
programming interfaces may not provide access to the needed
invariant. Volatile filehandles may ease the implementation of
server functionality such as hierarchical storage management or
filesystem reorganization or migration. However, the volatile
filehandle increases the implementation burden for the client.
Since the client will need to handle persistent and volatile
filehandles differently, a file attribute is defined which may be
used by the client to determine the filehandle types being returned
by the server.
2.2.1. General Properties of a Filehandle
The filehandle contains all the information the server needs to
distinguish an individual file. To the client, the filehandle is
opaque. The client stores filehandles for use in a later request and
can compare two filehandles from the same server for equality by
doing a byte-by-byte comparison. However, the client MUST NOT
otherwise interpret the contents of filehandles. If two filehandles
from the same server are equal, they MUST refer to the same file.
Servers SHOULD try to maintain a one-to-one correspondence between
filehandles and files but this is not required. Clients MUST use
filehandle comparisons only to improve performance, not for correct
behavior. All clients need to be prepared for situations in which it
cannot be determined whether two filehandles denote the same object
and in such cases, avoid making invalid assumptions which might cause
incorrect behavior. Further discussion of filehandle and attribute
comparison in the context of data caching is presented in the section
"Data Caching and File Identity".
As an example, in the case that two different path names when
traversed at the server terminate at the same filesystem object, the
server SHOULD return the same filehandle for each path. This can
occur if a hard link is used to create two file names which refer to
the same underlying file object and associated data. For example, if
paths /a/b/c and /a/d/c refer to the same file, the server SHOULD
return the same filehandle for both path names traversals.
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2.2.2. Persistent Filehandle
A persistent filehandle is defined as having a fixed value for the
lifetime of the filesystem object to which it refers. Once the
server creates the filehandle for a filesystem object, the server
MUST accept the same filehandle for the object for the lifetime of
the object. If the server restarts or reboots the NFS server must
honor the same filehandle value as it did in the server's previous
instantiation. Similarly, if the filesystem is migrated, the new NFS
server must honor the same filehandle as the old NFS server.
The persistent filehandle will be become stale or invalid when the
filesystem object is removed. When the server is presented with a
persistent filehandle that refers to a deleted object, it MUST return
an error of NFS4ERR_STALE. A filehandle may become stale when the
filesystem containing the object is no longer available. The file
system may become unavailable if it exists on removable media and the
media is no longer available at the server or the filesystem in whole
has been destroyed or the filesystem has simply been removed from the
server's name space (i.e. unmounted in a UNIX environment).
2.2.3. Volatile Filehandle
A volatile filehandle does not share the same longevity
characteristics of a persistent filehandle. The server may determine
that a volatile filehandle is no longer valid at many different
points in time. If the server can definitively determine that a
volatile filehandle refers to an object that has been removed, the
server should return NFS4ERR_STALE to the client (as is the case for
persistent filehandles). In all other cases where the server
determines that a volatile filehandle can no longer be used, it
should return an error of NFS4ERR_FHEXPIRED.
The mandatory attribute "fh_expire_type" is used by the client to
determine what type of filehandle the server is providing for a
particular filesystem. This attribute is a bitmask with the
following values:
FH4_PERSISTENT The value of FH4_PERSISTENT is used to indicate a
persistent filehandle, which is valid until the object is removed
from the filesystem. The server will not return NFS4ERR_FHEXPIRED
for this filehandle. FH4_PERSISTENT is defined as a value in
which none of the bits specified below are set.
FH4_VOLATILE_ANY The filehandle may expire at any time, except as
specifically excluded (i.e. FH4_NO_EXPIRE_WITH_OPEN).
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FH4_NOEXPIRE_WITH_OPEN May only be set when FH4_VOLATILE_ANY is set.
If this bit is set, then the meaning of FH4_VOLATILE_ANY is
qualified to exclude any expiration of the filehandle when it is
open.
FH4_VOL_MIGRATION The filehandle will expire as a result of
migration. If FH4_VOL_ANY is set, FH4_VOL_MIGRATION is redundant.
FH4_VOL_RENAME The filehandle will expire during rename. This
includes a rename by the requesting client or a rename by any
other client. If FH4_VOL_ANY is set, FH4_VOL_RENAME is redundant.
Servers which provide volatile filehandles that may expire while open
(i.e. if FH4_VOL_MIGRATION or FH4_VOL_RENAME is set or if
FH4_VOLATILE_ANY is set and FH4_NOEXPIRE_WITH_OPEN not set), should
deny a RENAME or REMOVE that would affect an OPEN file of any of the
components leading to the OPEN file. In addition, the server should
deny all RENAME or REMOVE requests during the grace period upon
server restart.
Note that the bits FH4_VOL_MIGRATION and FH4_VOL_RENAME allow the
client to determine that expiration has occurred whenever a specific
event occurs, without an explicit filehandle expiration error from
the server. FH4_VOL_ANY does not provide this form of information.
In situations where the server will expire many, but not all
filehandles upon migration (e.g. all but those that are open),
FH4_VOLATILE_ANY (in this case with FH4_NOEXPIRE_WITH_OPEN) is a
better choice since the client may not assume that all filehandles
will expire when migration occurs, and it is likely that additional
expirations will occur (as a result of file CLOSE) that are separated
in time from the migration event itself.
2.3. One Method of Constructing a Volatile Filehandle
A volatile filehandle, while opaque to the client could contain:
[volatile bit = 1 | server boot time | slot | generation number]
o slot is an index in the server volatile filehandle table
o generation number is the generation number for the table entry/
slot
When the client presents a volatile filehandle, the server makes the
following checks, which assume that the check for the volatile bit
has passed. If the server boot time is less than the current server
boot time, return NFS4ERR_FHEXPIRED. If slot is out of range, return
NFS4ERR_BADHANDLE. If the generation number does not match, return
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NFS4ERR_FHEXPIRED.
When the server reboots, the table is gone (it is volatile).
If volatile bit is 0, then it is a persistent filehandle with a
different structure following it.
2.4. Client Recovery from Filehandle Expiration
If possible, the client SHOULD recover from the receipt of an
NFS4ERR_FHEXPIRED error. The client must take on additional
responsibility so that it may prepare itself to recover from the
expiration of a volatile filehandle. If the server returns
persistent filehandles, the client does not need these additional
steps.
For volatile filehandles, most commonly the client will need to store
the component names leading up to and including the filesystem object
in question. With these names, the client should be able to recover
by finding a filehandle in the name space that is still available or
by starting at the root of the server's filesystem name space.
If the expired filehandle refers to an object that has been removed
from the filesystem, obviously the client will not be able to recover
from the expired filehandle.
It is also possible that the expired filehandle refers to a file that
has been renamed. If the file was renamed by another client, again
it is possible that the original client will not be able to recover.
However, in the case that the client itself is renaming the file and
the file is open, it is possible that the client may be able to
recover. The client can determine the new path name based on the
processing of the rename request. The client can then regenerate the
new filehandle based on the new path name. The client could also use
the compound operation mechanism to construct a set of operations
like:
RENAME A B
LOOKUP B
GETFH
Note that the COMPOUND procedure does not provide atomicity. This
example only reduces the overhead of recovering from an expired
filehandle.
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3. File Attributes
To meet the requirements of extensibility and increased
interoperability with non-UNIX platforms, attributes must be handled
in a flexible manner. The NFS version 3 fattr3 structure contains a
fixed list of attributes that not all clients and servers are able to
support or care about. The fattr3 structure can not be extended as
new needs arise and it provides no way to indicate non-support. With
the NFS version 4 protocol, the client is able query what attributes
the server supports and construct requests with only those supported
attributes (or a subset thereof).
To this end, attributes are divided into three groups: mandatory,
recommended, and named. Both mandatory and recommended attributes
are supported in the NFS version 4 protocol by a specific and well-
defined encoding and are identified by number. They are requested by
setting a bit in the bit vector sent in the GETATTR request; the
server response includes a bit vector to list what attributes were
returned in the response. New mandatory or recommended attributes
may be added to the NFS protocol between major revisions by
publishing a standards-track RFC which allocates a new attribute
number value and defines the encoding for the attribute. See the
section "Minor Versioning" for further discussion.
Named attributes are accessed by the new OPENATTR operation, which
accesses a hidden directory of attributes associated with a file
system object. OPENATTR takes a filehandle for the object and
returns the filehandle for the attribute hierarchy. The filehandle
for the named attributes is a directory object accessible by LOOKUP
or READDIR and contains files whose names represent the named
attributes and whose data bytes are the value of the attribute. For
example:
+----------+-----------+---------------------------------+
| LOOKUP | "foo" | ; look up file |
| GETATTR | attrbits | |
| OPENATTR | | ; access foo's named attributes |
| LOOKUP | "x11icon" | ; look up specific attribute |
| READ | 0,4096 | ; read stream of bytes |
+----------+-----------+---------------------------------+
Named attributes are intended for data needed by applications rather
than by an NFS client implementation. NFS implementors are strongly
encouraged to define their new attributes as recommended attributes
by bringing them to the IETF standards-track process.
The set of attributes which are classified as mandatory is
deliberately small since servers must do whatever it takes to support
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them. A server should support as many of the recommended attributes
as possible but by their definition, the server is not required to
support all of them. Attributes are deemed mandatory if the data is
both needed by a large number of clients and is not otherwise
reasonably computable by the client when support is not provided on
the server.
Note that the hidden directory returned by OPENATTR is a convenience
for protocol processing. The client should not make any assumptions
about the server's implementation of named attributes and whether the
underlying filesystem at the server has a named attribute directory
or not. Therefore, operations such as SETATTR and GETATTR on the
named attribute directory are undefined.
3.1. Mandatory Attributes
These MUST be supported by every NFS version 4 client and server in
order to ensure a minimum level of interoperability. The server must
store and return these attributes and the client must be able to
function with an attribute set limited to these attributes. With
just the mandatory attributes some client functionality may be
impaired or limited in some ways. A client may ask for any of these
attributes to be returned by setting a bit in the GETATTR request and
the server must return their value.
3.2. Recommended Attributes
These attributes are understood well enough to warrant support in the
NFS version 4 protocol. However, they may not be supported on all
clients and servers. A client may ask for any of these attributes to
be returned by setting a bit in the GETATTR request but must handle
the case where the server does not return them. A client may ask for
the set of attributes the server supports and should not request
attributes the server does not support. A server should be tolerant
of requests for unsupported attributes and simply not return them
rather than considering the request an error. It is expected that
servers will support all attributes they comfortably can and only
fail to support attributes which are difficult to support in their
operating environments. A server should provide attributes whenever
they don't have to "tell lies" to the client. For example, a file
modification time should be either an accurate time or should not be
supported by the server. This will not always be comfortable to
clients but the client is better positioned decide whether and how to
fabricate or construct an attribute or whether to do without the
attribute.
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3.3. Named Attributes
These attributes are not supported by direct encoding in the NFS
Version 4 protocol but are accessed by string names rather than
numbers and correspond to an uninterpreted stream of bytes which are
stored with the filesystem object. The name space for these
attributes may be accessed by using the OPENATTR operation. The
OPENATTR operation returns a filehandle for a virtual "attribute
directory" and further perusal of the name space may be done using
READDIR and LOOKUP operations on this filehandle. Named attributes
may then be examined or changed by normal READ and WRITE and CREATE
operations on the filehandles returned from READDIR and LOOKUP.
Named attributes may have attributes.
It is recommended that servers support arbitrary named attributes. A
client should not depend on the ability to store any named attributes
in the server's filesystem. If a server does support named
attributes, a client which is also able to handle them should be able
to copy a file's data and meta-data with complete transparency from
one location to another; this would imply that names allowed for
regular directory entries are valid for named attribute names as
well.
Names of attributes will not be controlled by this document or other
IETF standards track documents. See the section "IANA
Considerations" for further discussion.
3.4. Classification of Attributes
Each of the Mandatory and Recommended attributes can be classified in
one of three categories: per server, per filesystem, or per
filesystem object. Note that it is possible that some per filesystem
attributes may vary within the filesystem. See the "homogeneous"
attribute for its definition. Note that the attributes
time_access_set and time_modify_set are not listed in this section
because they are write-only attributes corresponding to time_access
and time_modify, and are used in a special instance of SETATTR.
o The per server attribute is:
lease_time
o The per filesystem attributes are:
supp_attr, fh_expire_type, link_support, symlink_support,
unique_handles, aclsupport, cansettime, case_insensitive,
case_preserving, chown_restricted, files_avail, files_free,
files_total, fs_locations, homogeneous, maxfilesize, maxname,
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maxread, maxwrite, no_trunc, space_avail, space_free,
space_total, time_delta, fs_layouttype, send_impl_id,
recv_impl_id
o The per filesystem object attributes are:
type, change, size, named_attr, fsid, rdattr_error, filehandle,
ACL, archive, fileid, hidden, maxlink, mimetype, mode,
numlinks, owner, owner_group, rawdev, space_used, system,
time_access, time_backup, time_create, time_metadata,
time_modify, mounted_on_fileid, layouttype, layouthint,
layout_blksize, layout_alignment
For quota_avail_hard, quota_avail_soft, and quota_used see their
definitions below for the appropriate classification.
3.5. Mandatory Attributes - Definitions
+-----------------+----+------------+--------+----------------------+
| name | # | Data Type | Access | Description |
+-----------------+----+------------+--------+----------------------+
| supp_attr | 0 | bitmap | READ | The bit vector which |
| | | | | would retrieve all |
| | | | | mandatory and |
| | | | | recommended |
| | | | | attributes that are |
| | | | | supported for this |
| | | | | object. The scope of |
| | | | | this attribute |
| | | | | applies to all |
| | | | | objects with a |
| | | | | matching fsid. |
| type | 1 | nfs4_ftype | READ | The type of the |
| | | | | object (file, |
| | | | | directory, symlink, |
| | | | | etc.) |
| fh_expire_type | 2 | uint32 | READ | Server uses this to |
| | | | | specify filehandle |
| | | | | expiration behavior |
| | | | | to the client. See |
| | | | | the section |
| | | | | "Filehandles" for |
| | | | | additional |
| | | | | description. |
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| change | 3 | uint64 | READ | A value created by |
| | | | | the server that the |
| | | | | client can use to |
| | | | | determine if file |
| | | | | data, directory |
| | | | | contents or |
| | | | | attributes of the |
| | | | | object have been |
| | | | | modified. The server |
| | | | | may return the |
| | | | | object's |
| | | | | time_metadata |
| | | | | attribute for this |
| | | | | attribute's value |
| | | | | but only if the |
| | | | | filesystem object |
| | | | | can not be updated |
| | | | | more frequently than |
| | | | | the resolution of |
| | | | | time_metadata. |
| size | 4 | uint64 | R/W | The size of the |
| | | | | object in bytes. |
| link_support | 5 | bool | READ | True, if the |
| | | | | object's filesystem |
| | | | | supports hard links. |
| symlink_support | 6 | bool | READ | True, if the |
| | | | | object's filesystem |
| | | | | supports symbolic |
| | | | | links. |
| named_attr | 7 | bool | READ | True, if this object |
| | | | | has named |
| | | | | attributes. In other |
| | | | | words, object has a |
| | | | | non-empty named |
| | | | | attribute directory. |
| fsid | 8 | fsid4 | READ | Unique filesystem |
| | | | | identifier for the |
| | | | | filesystem holding |
| | | | | this object. fsid |
| | | | | contains major and |
| | | | | minor components |
| | | | | each of which are |
| | | | | uint64. |
| unique_handles | 9 | bool | READ | True, if two |
| | | | | distinct filehandles |
| | | | | guaranteed to refer |
| | | | | to two different |
| | | | | filesystem objects. |
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| lease_time | 10 | nfs_lease4 | READ | Duration of leases |
| | | | | at server in |
| | | | | seconds. |
| rdattr_error | 11 | enum | READ | Error returned from |
| | | | | getattr during |
| | | | | readdir. |
| filehandle | 19 | nfs_fh4 | READ | The filehandle of |
| | | | | this object |
| | | | | (primarily for |
| | | | | readdir requests). |
+-----------------+----+------------+--------+----------------------+
3.6. Recommended Attributes - Definitions
+--------------------+-----+--------------+--------+----------------+
| name | # | Data Type | Access | Description |
+--------------------+-----+--------------+--------+----------------+
| ACL | 12 | nfsace4<> | R/W | The access |
| | | | | control list |
| | | | | for the |
| | | | | object. |
| aclsupport | 13 | uint32 | READ | Indicates what |
| | | | | types of ACLs |
| | | | | are supported |
| | | | | on the current |
| | | | | filesystem. |
| archive | 14 | bool | R/W | True, if this |
| | | | | file has been |
| | | | | archived since |
| | | | | the time of |
| | | | | last |
| | | | | modification |
| | | | | (deprecated in |
| | | | | favor of |
| | | | | time_backup). |
| cansettime | 15 | bool | READ | True, if the |
| | | | | server able to |
| | | | | change the |
| | | | | times for a |
| | | | | filesystem |
| | | | | object as |
| | | | | specified in a |
| | | | | SETATTR |
| | | | | operation. |
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| case_insensitive | 16 | bool | READ | True, if |
| | | | | filename |
| | | | | comparisons on |
| | | | | this |
| | | | | filesystem are |
| | | | | case |
| | | | | insensitive. |
| case_preserving | 17 | bool | READ | True, if |
| | | | | filename case |
| | | | | on this |
| | | | | filesystem are |
| | | | | preserved. |
| chown_restricted | 18 | bool | READ | If TRUE, the |
| | | | | server will |
| | | | | reject any |
| | | | | request to |
| | | | | change either |
| | | | | the owner or |
| | | | | the group |
| | | | | associated |
| | | | | with a file if |
| | | | | the caller is |
| | | | | not a |
| | | | | privileged |
| | | | | user (for |
| | | | | example, |
| | | | | "root" in UNIX |
| | | | | operating |
| | | | | environments |
| | | | | or in Windows |
| | | | | 2000 the "Take |
| | | | | Ownership" |
| | | | | privilege). |
| fileid | 20 | uint64 | READ | A number |
| | | | | uniquely |
| | | | | identifying |
| | | | | the file |
| | | | | within the |
| | | | | filesystem. |
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| files_avail | 21 | uint64 | READ | File slots |
| | | | | available to |
| | | | | this user on |
| | | | | the filesystem |
| | | | | containing |
| | | | | this object - |
| | | | | this should be |
| | | | | the smallest |
| | | | | relevant |
| | | | | limit. |
| files_free | 22 | uint64 | READ | Free file |
| | | | | slots on the |
| | | | | filesystem |
| | | | | containing |
| | | | | this object - |
| | | | | this should be |
| | | | | the smallest |
| | | | | relevant |
| | | | | limit. |
| files_total | 23 | uint64 | READ | Total file |
| | | | | slots on the |
| | | | | filesystem |
| | | | | containing |
| | | | | this object. |
| fs_locations | 24 | fs_locations | READ | Locations |
| | | | | where this |
| | | | | filesystem may |
| | | | | be found. If |
| | | | | the server |
| | | | | returns |
| | | | | NFS4ERR_MOVED |
| | | | | as an error, |
| | | | | this attribute |
| | | | | MUST be |
| | | | | supported. |
| hidden | 25 | bool | R/W | True, if the |
| | | | | file is |
| | | | | considered |
| | | | | hidden with |
| | | | | respect to the |
| | | | | Windows API? |
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| homogeneous | 26 | bool | READ | True, if this |
| | | | | object's |
| | | | | filesystem is |
| | | | | homogeneous, |
| | | | | i.e. are per |
| | | | | filesystem |
| | | | | attributes the |
| | | | | same for all |
| | | | | filesystem's |
| | | | | objects. |
| maxfilesize | 27 | uint64 | READ | Maximum |
| | | | | supported file |
| | | | | size for the |
| | | | | filesystem of |
| | | | | this object. |
| maxlink | 28 | uint32 | READ | Maximum number |
| | | | | of links for |
| | | | | this object. |
| maxname | 29 | uint32 | READ | Maximum |
| | | | | filename size |
| | | | | supported for |
| | | | | this object. |
| maxread | 30 | uint64 | READ | Maximum read |
| | | | | size supported |
| | | | | for this |
| | | | | object. |
| maxwrite | 31 | uint64 | READ | Maximum write |
| | | | | size supported |
| | | | | for this |
| | | | | object. This |
| | | | | attribute |
| | | | | SHOULD be |
| | | | | supported if |
| | | | | the file is |
| | | | | writable. Lack |
| | | | | of this |
| | | | | attribute can |
| | | | | lead to the |
| | | | | client either |
| | | | | wasting |
| | | | | bandwidth or |
| | | | | not receiving |
| | | | | the best |
| | | | | performance. |
| mimetype | 32 | utf8<> | R/W | MIME body |
| | | | | type/subtype |
| | | | | of this |
| | | | | object. |
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| mode | 33 | mode4 | R/W | UNIX-style |
| | | | | mode and |
| | | | | permission |
| | | | | bits for this |
| | | | | object. |
| no_trunc | 34 | bool | READ | True, if a |
| | | | | name longer |
| | | | | than name_max |
| | | | | is used, an |
| | | | | error be |
| | | | | returned and |
| | | | | name is not |
| | | | | truncated. |
| numlinks | 35 | uint32 | READ | Number of hard |
| | | | | links to this |
| | | | | object. |
| owner | 36 | utf8<> | R/W | The string |
| | | | | name of the |
| | | | | owner of this |
| | | | | object. |
| owner_group | 37 | utf8<> | R/W | The string |
| | | | | name of the |
| | | | | group |
| | | | | ownership of |
| | | | | this object. |
| quota_avail_hard | 38 | uint64 | READ | For definition |
| | | | | see "Quota |
| | | | | Attributes" |
| | | | | section below. |
| quota_avail_soft | 39 | uint64 | READ | For definition |
| | | | | see "Quota |
| | | | | Attributes" |
| | | | | section below. |
| quota_used | 40 | uint64 | READ | For definition |
| | | | | see "Quota |
| | | | | Attributes" |
| | | | | section below. |
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| rawdev | 41 | specdata4 | READ | Raw device |
| | | | | identifier. |
| | | | | UNIX device |
| | | | | major/minor |
| | | | | node |
| | | | | information. |
| | | | | If the value |
| | | | | of type is not |
| | | | | NF4BLK or |
| | | | | NF4CHR, the |
| | | | | value return |
| | | | | SHOULD NOT be |
| | | | | considered |
| | | | | useful. |
| space_avail | 42 | uint64 | READ | Disk space in |
| | | | | bytes |
| | | | | available to |
| | | | | this user on |
| | | | | the filesystem |
| | | | | containing |
| | | | | this object - |
| | | | | this should be |
| | | | | the smallest |
| | | | | relevant |
| | | | | limit. |
| space_free | 43 | uint64 | READ | Free disk |
| | | | | space in bytes |
| | | | | on the |
| | | | | filesystem |
| | | | | containing |
| | | | | this object - |
| | | | | this should be |
| | | | | the smallest |
| | | | | relevant |
| | | | | limit. |
| space_total | 44 | uint64 | READ | Total disk |
| | | | | space in bytes |
| | | | | on the |
| | | | | filesystem |
| | | | | containing |
| | | | | this object. |
| space_used | 45 | uint64 | READ | Number of |
| | | | | filesystem |
| | | | | bytes |
| | | | | allocated to |
| | | | | this object. |
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| system | 46 | bool | R/W | True, if this |
| | | | | file is a |
| | | | | "system" file |
| | | | | with respect |
| | | | | to the Windows |
| | | | | API? |
| time_access | 47 | nfstime4 | READ | The time of |
| | | | | last access to |
| | | | | the object by |
| | | | | a read that |
| | | | | was satisfied |
| | | | | by the server. |
| time_access_set | 48 | settime4 | WRITE | Set the time |
| | | | | of last access |
| | | | | to the object. |
| | | | | SETATTR use |
| | | | | only. |
| time_backup | 49 | nfstime4 | R/W | The time of |
| | | | | last backup of |
| | | | | the object. |
| time_create | 50 | nfstime4 | R/W | The time of |
| | | | | creation of |
| | | | | the object. |
| | | | | This attribute |
| | | | | does not have |
| | | | | any relation |
| | | | | to the |
| | | | | traditional |
| | | | | UNIX file |
| | | | | attribute |
| | | | | "ctime" or |
| | | | | "change time". |
| time_delta | 51 | nfstime4 | READ | Smallest |
| | | | | useful server |
| | | | | time |
| | | | | granularity. |
| time_metadata | 52 | nfstime4 | READ | The time of |
| | | | | last meta-data |
| | | | | modification |
| | | | | of the object. |
| time_modify | 53 | nfstime4 | READ | The time of |
| | | | | last |
| | | | | modification |
| | | | | to the object. |
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| time_modify_set | 54 | settime4 | WRITE | Set the time |
| | | | | of last |
| | | | | modification |
| | | | | to the object. |
| | | | | SETATTR use |
| | | | | only. |
| mounted_on_fileid | 55 | uint64 | READ | Like fileid, |
| | | | | but if the |
| | | | | target |
| | | | | filehandle is |
| | | | | the root of a |
| | | | | filesystem |
| | | | | return the |
| | | | | fileid of the |
| | | | | underlying |
| | | | | directory. |
| send_impl_id | TBD | impl_ident4 | WRITE | Client |
| | | | | provides |
| | | | | server with |
| | | | | implementation |
| | | | | identity via |
| | | | | SETATTR. |
| recv_impl_id | TBD | nfs_impl_id4 | READ | Client obtains |
| | | | | server |
| | | | | implementation |
| | | | | via GETATTR. |
| dir_notif_delay | TBD | R/W | READ | notification |
| | | | | delays on |
| | | | | directory |
| | | | | attributes |
| dirent_notif_delay | TBD | R/W | READ | notification |
| | | | | delays on |
| | | | | child |
| | | | | attributes |
| fs_layouttype | TBD | layouttype4 | READ | Layout types |
| | | | | available for |
| | | | | the |
| | | | | filesystem. |
| layouttype | TBD | layouttype4 | READ | Layout types |
| | | | | available for |
| | | | | the file. |
| layouthint | TBD | layouthint4 | WRITE | Client |
| | | | | specified hint |
| | | | | for file |
| | | | | layout. |
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| layout_blksize | TBD | uint32_t | READ | Preferred |
| | | | | block size for |
| | | | | layout related |
| | | | | I/O. |
| layout_alignment | TBD | uint32_t | READ | Preferred |
| | | | | alignment for |
| | | | | layout related |
| | | | | I/O. |
| | TBD | | READ | desc |
| | TBD | | READ | desc |
+--------------------+-----+--------------+--------+----------------+
3.7. Time Access
As defined above, the time_access attribute represents the time of
last access to the object by a read that was satisfied by the server.
The notion of what is an "access" depends on server's operating
environment and/or the server's filesystem semantics. For example,
for servers obeying POSIX semantics, time_access would be updated
only by the READLINK, READ, and READDIR operations and not any of the
operations that modify the content of the object. Of course, setting
the corresponding time_access_set attribute is another way to modify
the time_access attribute.
Whenever the file object resides on a writable filesystem, the server
should make best efforts to record time_access into stable storage.
However, to mitigate the performance effects of doing so, and most
especially whenever the server is satisfying the read of the object's
content from its cache, the server MAY cache access time updates and
lazily write them to stable storage. It is also acceptable to give
administrators of the server the option to disable time_access
updates.
3.8. Interpreting owner and owner_group
The recommended attributes "owner" and "owner_group" (and also users
and groups within the "acl" attribute) are represented in terms of a
UTF-8 string. To avoid a representation that is tied to a particular
underlying implementation at the client or server, the use of the
UTF-8 string has been chosen. Note that section 6.1 of [RFC2624]
provides additional rationale. It is expected that the client and
server will have their own local representation of owner and
owner_group that is used for local storage or presentation to the end
user. Therefore, it is expected that when these attributes are
transferred between the client and server that the local
representation is translated to a syntax of the form "user@
dns_domain". This will allow for a client and server that do not use
the same local representation the ability to translate to a common
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syntax that can be interpreted by both.
Similarly, security principals may be represented in different ways
by different security mechanisms. Servers normally translate these
representations into a common format, generally that used by local
storage, to serve as a means of identifying the users corresponding
to these security principals. When these local identifiers are
translated to the form of the owner attribute, associated with files
created by such principals they identify, in a common format, the
users associated with each corresponding set of security principals.
The translation used to interpret owner and group strings is not
specified as part of the protocol. This allows various solutions to
be employed. For example, a local translation table may be consulted
that maps between a numeric id to the user@dns_domain syntax. A name
service may also be used to accomplish the translation. A server may
provide a more general service, not limited by any particular
translation (which would only translate a limited set of possible
strings) by storing the owner and owner_group attributes in local
storage without any translation or it may augment a translation
method by storing the entire string for attributes for which no
translation is available while using the local representation for
those cases in which a translation is available.
Servers that do not provide support for all possible values of the
owner and owner_group attributes, should return an error
(NFS4ERR_BADOWNER) when a string is presented that has no
translation, as the value to be set for a SETATTR of the owner,
owner_group, or acl attributes. When a server does accept an owner
or owner_group value as valid on a SETATTR (and similarly for the
owner and group strings in an acl), it is promising to return that
same string when a corresponding GETATTR is done. Configuration
changes and ill-constructed name translations (those that contain
aliasing) may make that promise impossible to honor. Servers should
make appropriate efforts to avoid a situation in which these
attributes have their values changed when no real change to ownership
has occurred.
The "dns_domain" portion of the owner string is meant to be a DNS
domain name. For example, user@ietf.org. Servers should accept as
valid a set of users for at least one domain. A server may treat
other domains as having no valid translations. A more general
service is provided when a server is capable of accepting users for
multiple domains, or for all domains, subject to security
constraints.
In the case where there is no translation available to the client or
server, the attribute value must be constructed without the "@".
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Therefore, the absence of the @ from the owner or owner_group
attribute signifies that no translation was available at the sender
and that the receiver of the attribute should not use that string as
a basis for translation into its own internal format. Even though
the attribute value can not be translated, it may still be useful.
In the case of a client, the attribute string may be used for local
display of ownership.
To provide a greater degree of compatibility with previous versions
of NFS (i.e. v2 and v3), which identified users and groups by 32-bit
unsigned uid's and gid's, owner and group strings that consist of
decimal numeric values with no leading zeros can be given a special
interpretation by clients and servers which choose to provide such
support. The receiver may treat such a user or group string as
representing the same user as would be represented by a v2/v3 uid or
gid having the corresponding numeric value. A server is not
obligated to accept such a string, but may return an NFS4ERR_BADOWNER
instead. To avoid this mechanism being used to subvert user and
group translation, so that a client might pass all of the owners and
groups in numeric form, a server SHOULD return an NFS4ERR_BADOWNER
error when there is a valid translation for the user or owner
designated in this way. In that case, the client must use the
appropriate name@domain string and not the special form for
compatibility.
The owner string "nobody" may be used to designate an anonymous user,
which will be associated with a file created by a security principal
that cannot be mapped through normal means to the owner attribute.
3.9. Character Case Attributes
With respect to the case_insensitive and case_preserving attributes,
each UCS-4 character (which UTF-8 encodes) has a "long descriptive
name" [RFC1345] which may or may not included the word "CAPITAL" or
"SMALL". The presence of SMALL or CAPITAL allows an NFS server to
implement unambiguous and efficient table driven mappings for case
insensitive comparisons, and non-case-preserving storage. For
general character handling and internationalization issues, see the
section "Internationalization".
3.10. Quota Attributes
For the attributes related to filesystem quotas, the following
definitions apply:
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quota_avail_soft The value in bytes which represents the amount of
additional disk space that can be allocated to this file or
directory before the user may reasonably be warned. It is
understood that this space may be consumed by allocations to other
files or directories though there is a rule as to which other
files or directories.
quota_avail_hard The value in bytes which represent the amount of
additional disk space beyond the current allocation that can be
allocated to this file or directory before further allocations
will be refused. It is understood that this space may be consumed
by allocations to other files or directories.
quota_used The value in bytes which represent the amount of disc
space used by this file or directory and possibly a number of
other similar files or directories, where the set of "similar"
meets at least the criterion that allocating space to any file or
directory in the set will reduce the "quota_avail_hard" of every
other file or directory in the set.
Note that there may be a number of distinct but overlapping sets
of files or directories for which a quota_used value is
maintained. E.g. "all files with a given owner", "all files with
a given group owner". etc.
The server is at liberty to choose any of those sets but should do
so in a repeatable way. The rule may be configured per-filesystem
or may be "choose the set with the smallest quota".
3.11. mounted_on_fileid
UNIX-based operating environments connect a filesystem into the
namespace by connecting (mounting) the filesystem onto the existing
file object (the mount point, usually a directory) of an existing
filesystem. When the mount point's parent directory is read via an
API like readdir(), the return results are directory entries, each
with a component name and a fileid. The fileid of the mount point's
directory entry will be different from the fileid that the stat()
system call returns. The stat() system call is returning the fileid
of the root of the mounted filesystem, whereas readdir() is returning
the fileid stat() would have returned before any filesystems were
mounted on the mount point.
Unlike NFS version 3, NFS version 4 allows a client's LOOKUP request
to cross other filesystems. The client detects the filesystem
crossing whenever the filehandle argument of LOOKUP has an fsid
attribute different from that of the filehandle returned by LOOKUP.
A UNIX-based client will consider this a "mount point crossing".
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UNIX has a legacy scheme for allowing a process to determine its
current working directory. This relies on readdir() of a mount
point's parent and stat() of the mount point returning fileids as
previously described. The mounted_on_fileid attribute corresponds to
the fileid that readdir() would have returned as described
previously.
While the NFS version 4 client could simply fabricate a fileid
corresponding to what mounted_on_fileid provides (and if the server
does not support mounted_on_fileid, the client has no choice), there
is a risk that the client will generate a fileid that conflicts with
one that is already assigned to another object in the filesystem.
Instead, if the server can provide the mounted_on_fileid, the
potential for client operational problems in this area is eliminated.
If the server detects that there is no mounted point at the target
file object, then the value for mounted_on_fileid that it returns is
the same as that of the fileid attribute.
The mounted_on_fileid attribute is RECOMMENDED, so the server SHOULD
provide it if possible, and for a UNIX-based server, this is
straightforward. Usually, mounted_on_fileid will be requested during
a READDIR operation, in which case it is trivial (at least for UNIX-
based servers) to return mounted_on_fileid since it is equal to the
fileid of a directory entry returned by readdir(). If
mounted_on_fileid is requested in a GETATTR operation, the server
should obey an invariant that has it returning a value that is equal
to the file object's entry in the object's parent directory, i.e.
what readdir() would have returned. Some operating environments
allow a series of two or more filesystems to be mounted onto a single
mount point. In this case, for the server to obey the aforementioned
invariant, it will need to find the base mount point, and not the
intermediate mount points.
3.12. send_impl_id and recv_impl_id
These recommended attributes are used to identify the client and
server. In the case of the send_impl_id attribute, the client sends
its clientid4 value along with the nfs_impl_id4. The use of the
clientid4 value allows the server to identify and match specific
client interaction. In the case of the recv_impl_id attribute, the
client receives the nfs_impl_id4 value.
Access to this identification information can be most useful at both
client and server. Being able to identify specific implementations
can help in planning by administrators or implementers. For example,
diagnostic software may extract this information in an attempt to
identify implementation problems, performance workload behaviors or
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general usage statistics. Since the intent of having access to this
information is for planning or general diagnosis only, the client and
server MUST NOT interpret this implementation identity information in
a way that affects interoperational behavior of the implementation.
The reason is the if clients and servers did such a thing, they might
use fewer capabilities of the protocol than the peer can support, or
the client and server might refuse to interoperate.
Because it is likely some implementations will violate the protocol
specification and interpret the identity information, implementations
MUST allow the users of the NFSv4 client and server to set the
contents of the sent nfs_impl_id structure to any value.
Even though these attributes are recommended, if the server supports
one of them it MUST support the other.
3.13. fs_layouttype
This attribute applies to a file system and indicates what layout
types are supported by the file system. We expect this attribute to
be queried when a client encounters a new fsid. This attribute is
used by the client to determine if it has applicable layout drivers.
3.14. layouttype
This attribute indicates the particular layout type(s) used for a
file. This is for informational purposes only. The client needs to
use the LAYOUTGET operation in order to get enough information (e.g.,
specific device information) in order to perform I/O.
3.15. layouthint
This attribute may be set on newly created files to influence the
metadata server's choice for the file's layout. It is suggested that
this attribute is set as one of the initial attributes within the
OPEN call. The metadata server may ignore this attribute. This
attribute is a sub-set of the layout structure returned by LAYOUTGET.
For example, instead of specifying particular devices, this would be
used to suggest the stripe width of a file. It is up to the server
implementation to determine which fields within the layout it uses.
[[Comment.3: it has been suggested that the HINT is a well defined
type other than pnfs_layoutdata4, similar to pnfs_layoutupdate4.]]
3.16. Access Control Lists
The NFS version 4 ACL attribute is an array of access control entries
(ACE). Although, the client can read and write the ACL attribute,
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the NFSv4 model is the server does all access control based on the
server's interpretation of the ACL. If at any point the client wants
to check access without issuing an operation that modifies or reads
data or metadata, the client can use the OPEN and ACCESS operations
to do so. There are various access control entry types, as defined
in Section 3.16.1. The server is able to communicate which ACE types
are supported by returning the appropriate value within the
aclsupport attribute. Each ACE covers one or more operations on a
file or directory as described in Section 3.16.2. It may also
contain one or more flags that modify the semantics of the ACE as
defined in Section 3.16.3.
The NFS ACE attribute is defined as follows:
typedef uint32_t acetype4;
typedef uint32_t aceflag4;
typedef uint32_t acemask4;
struct nfsace4 {
acetype4 type;
aceflag4 flag;
acemask4 access_mask;
utf8str_mixed who;
};
To determine if a request succeeds, each nfsace4 entry is processed
in order by the server. Only ACEs which have a "who" that matches
the requester are considered. Each ACE is processed until all of the
bits of the requester's access have been ALLOWED. Once a bit (see
below) has been ALLOWED by an ACCESS_ALLOWED_ACE, it is no longer
considered in the processing of later ACEs. If an ACCESS_DENIED_ACE
is encountered where the requester's access still has unALLOWED bits
in common with the "access_mask" of the ACE, the request is denied.
However, unlike the ALLOWED and DENIED ACE types, the ALARM and AUDIT
ACE types do not affect a requester's access, and instead are for
triggering events as a result of a requester's access attempt.
Therefore, all AUDIT and ALARM ACEs are processed until end of the
ACL. When the ACL is fully processed, if there are bits in
requester's mask that have not been considered whether the server
allows or denies, the access is denied. Even though a request is
denied, servers may choose to have other restrictions or
implementation defined security policies in place. In those cases,
access may be decided outside of what is in the ACL. Examples of
such security policies or restrictions are:
o The owner of the file will always be able granted ACE4_WRITE_ACL
and ACE4_READ_ACL permissions. This would prevent the user from
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getting into the situation where they can't ever modify the ACL.
o The ACL may say that an entity is to be granted ACE4_WRITE_DATA
permission, but the file system is mounted read only, therefore
write access is denied.
As mentioned before, this is one of the reasons that client
implementations are not recommended to do their own access checking.
The NFS version 4 ACL model is quite rich. Some server platforms may
provide access control functionality that goes beyond the UNIX-style
mode attribute, but which is not as rich as the NFS ACL model. So
that users can take advantage of this more limited functionality, the
server may indicate that it supports ACLs as long as it follows the
guidelines for mapping between its ACL model and the NFS version 4
ACL model.
The situation is complicated by the fact that a server may have
multiple modules that enforce ACLs. For example, the enforcement for
NFS version 4 access may be different from the enforcement for local
access, and both may be different from the enforcement for access
through other protocols such as SMB. So it may be useful for a
server to accept an ACL even if not all of its modules are able to
support it.
The guiding principle in all cases is that the server must not accept
ACLs that appear to make the file more secure than it really is.
3.16.1. ACE type
Type Description
_____________________________________________________
ALLOW Explicitly grants the access defined in
acemask4 to the file or directory.
DENY Explicitly denies the access defined in
acemask4 to the file or directory.
AUDIT LOG (system dependent) any access
attempt to a file or directory which
uses any of the access methods specified
in acemask4.
ALARM Generate a system ALARM (system
dependent) when any access attempt is
made to a file or directory for the
access methods specified in acemask4.
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A server need not support all of the above ACE types. The bitmask
constants used to represent the above definitions within the
aclsupport attribute are as follows:
const ACL4_SUPPORT_ALLOW_ACL = 0x00000001;
const ACL4_SUPPORT_DENY_ACL = 0x00000002;
const ACL4_SUPPORT_AUDIT_ACL = 0x00000004;
const ACL4_SUPPORT_ALARM_ACL = 0x00000008;
The semantics of the "type" field follow the descriptions provided
above.
The constants used for the type field (acetype4) are as follows:
const ACE4_ACCESS_ALLOWED_ACE_TYPE = 0x00000000;
const ACE4_ACCESS_DENIED_ACE_TYPE = 0x00000001;
const ACE4_SYSTEM_AUDIT_ACE_TYPE = 0x00000002;
const ACE4_SYSTEM_ALARM_ACE_TYPE = 0x00000003;
Clients should not attempt to set an ACE unless the server claims
support for that ACE type. If the server receives a request to set
an ACE that it cannot store, it MUST reject the request with
NFS4ERR_ATTRNOTSUPP. If the server receives a request to set an ACE
that it can store but cannot enforce, the server SHOULD reject the
request with NFS4ERR_ATTRNOTSUPP.
Example: suppose a server can enforce NFS ACLs for NFS access but
cannot enforce ACLs for local access. If arbitrary processes can run
on the server, then the server SHOULD NOT indicate ACL support. On
the other hand, if only trusted administrative programs run locally,
then the server may indicate ACL support.
3.16.2. ACE Access Mask
The access_mask field contains values based on the following:
ACE4_READ_DATA
Operation(s) affected:
READ
OPEN
Discussion:
Permission to read the data of the file.
ACE4_LIST_DIRECTORY
Operation(s) affected:
READDIR
Discussion:
Permission to list the contents of a directory.
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ACE4_WRITE_DATA
Operation(s) affected:
WRITE
OPEN
Discussion:
Permission to modify a file's data anywhere in the file's
offset range. This includes the ability to write to any
arbitrary offset and as a result to grow the file.
ACE4_ADD_FILE
Operation(s) affected:
CREATE
OPEN
Discussion:
Permission to add a new file in a directory. The CREATE
operation is affected when nfs_ftype4 is NF4LNK, NF4BLK,
NF4CHR, NF4SOCK, or NF4FIFO. (NF4DIR is not listed because
it is covered by ACE4_ADD_SUBDIRECTORY.) OPEN is affected
when used to create a regular file.
ACE4_APPEND_DATA
Operation(s) affected:
WRITE
OPEN
Discussion:
The ability to modify a file's data, but only starting at
EOF. This allows for the notion of append-only files, by
allowing ACE4_APPEND_DATA and denying ACE4_WRITE_DATA to
the same user or group. If a file has an ACL such as the
one described above and a WRITE request is made for
somewhere other than EOF, the server SHOULD return
NFS4ERR_ACCESS.
ACE4_ADD_SUBDIRECTORY
Operation(s) affected:
CREATE
Discussion:
Permission to create a subdirectory in a directory. The
CREATE operation is affected when nfs_ftype4 is NF4DIR.
ACE4_READ_NAMED_ATTRS
Operation(s) affected:
OPENATTR
Discussion:
Permission to read the named attributes of a file or to
lookup the named attributes directory. OPENATTR is
affected when it is not used to create a named attribute
directory. This is when 1.) createdir is TRUE, but a
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named attribute directory already exists, or 2.) createdir
is FALSE.
ACE4_WRITE_NAMED_ATTRS
Operation(s) affected:
OPENATTR
Discussion:
Permission to write the named attributes of a file or
to create a named attribute directory. OPENATTR is
affected when it is used to create a named attribute
directory. This is when createdir is TRUE and no named
attribute directory exists. The ability to check whether
or not a named attribute directory exists depends on the
ability to look it up, therefore, users also need the
ACE4_READ_NAMED_ATTRS permission in order to create a
named attribute directory.
ACE4_EXECUTE
Operation(s) affected:
LOOKUP
Discussion:
Permission to execute a file or traverse/search a
directory.
ACE4_DELETE_CHILD
Operation(s) affected:
REMOVE
Discussion:
Permission to delete a file or directory within a
directory. See section "ACE4_DELETE vs.
ACE4_DELETE_CHILD" for information on how these two access
mask bits interact.
ACE4_READ_ATTRIBUTES
Operation(s) affected:
GETATTR of file system object attributes
Discussion:
The ability to read basic attributes (non-ACLs) of a file.
On a UNIX system, basic attributes can be thought of as
the stat level attributes. Allowing this access mask bit
would mean the entity can execute "ls -l" and stat.
ACE4_WRITE_ATTRIBUTES
Operation(s) affected:
SETATTR of time_access_set, time_backup,
time_create, time_modify_set
Discussion:
Permission to change the times associated with a file
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or directory to an arbitrary value. A user having
ACE4_WRITE_DATA permission, but lacking
ACE4_WRITE_ATTRIBUTES must be allowed to implicitly set
the times associated with a file.
ACE4_DELETE
Operation(s) affected:
REMOVE
Discussion:
Permission to delete the file or directory. See section
"ACE4_DELETE vs. ACE4_DELETE_CHILD" for information on how
these two access mask bits interact.
ACE4_READ_ACL
Operation(s) affected:
GETATTR of acl
Discussion:
Permission to read the ACL.
ACE4_WRITE_ACL
Operation(s) affected:
SETATTR of acl and mode
Discussion:
Permission to write the acl and mode attributes.
ACE4_WRITE_OWNER
Operation(s) affected:
SETATTR of owner and owner_group
Discussions:
Permission to write the owner and owner_group attributes.
On UNIX systems, this is the ability to execute chown or
chgrp.
ACE4_SYNCHRONIZE
Operation(s) affected:
NONE
Discussion:
Permission to access file locally at the server with
synchronized reads and writes.
The bitmask constants used for the access mask field are as follows:
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const ACE4_READ_DATA = 0x00000001;
const ACE4_LIST_DIRECTORY = 0x00000001;
const ACE4_WRITE_DATA = 0x00000002;
const ACE4_ADD_FILE = 0x00000002;
const ACE4_APPEND_DATA = 0x00000004;
const ACE4_ADD_SUBDIRECTORY = 0x00000004;
const ACE4_READ_NAMED_ATTRS = 0x00000008;
const ACE4_WRITE_NAMED_ATTRS = 0x00000010;
const ACE4_EXECUTE = 0x00000020;
const ACE4_DELETE_CHILD = 0x00000040;
const ACE4_READ_ATTRIBUTES = 0x00000080;
const ACE4_WRITE_ATTRIBUTES = 0x00000100;
const ACE4_DELETE = 0x00010000;
const ACE4_READ_ACL = 0x00020000;
const ACE4_WRITE_ACL = 0x00040000;
const ACE4_WRITE_OWNER = 0x00080000;
const ACE4_SYNCHRONIZE = 0x00100000;
Server implementations need not provide the granularity of control
that is implied by this list of masks. For example, POSIX-based
systems might not distinguish APPEND_DATA (the ability to append to a
file) from WRITE_DATA (the ability to modify existing contents); both
masks would be tied to a single "write" permission. When such a
server returns attributes to the client, it would show both
APPEND_DATA and WRITE_DATA if and only if the write permission is
enabled.
If a server receives a SETATTR request that it cannot accurately
implement, it should error in the direction of more restricted
access. For example, suppose a server cannot distinguish overwriting
data from appending new data, as described in the previous paragraph.
If a client submits an ACE where APPEND_DATA is set but WRITE_DATA is
not (or vice versa), the server should reject the request with
NFS4ERR_ATTRNOTSUPP. Nonetheless, if the ACE has type DENY, the
server may silently turn on the other bit, so that both APPEND_DATA
and WRITE_DATA are denied.
3.16.2.1. ACE4_DELETE vs. ACE4_DELETE_CHILD
There are two separate access mask bits that govern the ability to
delete a file: ACE4_DELETE and ACE4_DELETE_CHILD. ACE4_DELETE is
intended to be specified by the ACL for the object to be deleted, and
ACE4_DELETE_CHILD is intended to be specified by the ACL of the
parent directory.
In addition to ACE4_DELETE and ACE4_DELETE_CHILD, many systems also
consider the "sticky bit" (MODE4_SVTX) and the appropriate "write"
mode bit when determining whether to allow a file to be deleted. The
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mode bit for write corresponds to ACE4_WRITE_DATA, which is the same
physical bit as ACE4_ADD_FILE. Therefore, ACE4_ADD_FILE can come
into play when determining permission to delete.
In the algorithm below, the strategy is that ACE4_DELETE and
ACE4_DELETE_CHILD take precedence over the sticky bit, and the sticky
bit takes precedence over the "write" mode bits (reflected in
ACE4_ADD_FILE).
Server implementations SHOULD grant or deny permission to delete
based on the following algorithm.
if ACE4_EXECUTE is denied by the parent directory ACL:
deny delete
else if ACE4_EXECUTE is unspecified by the parent
directory ACL:
deny delete
else if ACE4_DELETE is allowed by the target object ACL:
allow delete
else if ACE4_DELETE_CHILD is allowed by the parent
directory ACL:
allow delete
else if ACE4_DELETE_CHILD is denied by the
parent directory ACL:
deny delete
else if ACE4_ADD_FILE is allowed by the parent directory ACL:
if MODE4_SVTX is set for the parent directory:
if the principal owns the parent directory OR
the principal owns the target object OR
ACE4_WRITE_DATA is allowed by the target
object ACL:
allow delete
else:
deny delete
else:
allow delete
else:
deny delete
3.16.3. ACE flag
The "flag" field contains values based on the following descriptions.
ACE4_FILE_INHERIT_ACE
Can be placed on a directory and indicates that this ACE should be
added to each new non-directory file created.
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ACE4_DIRECTORY_INHERIT_ACE
Can be placed on a directory and indicates that this ACE should be
added to each new directory created.
ACE4_INHERIT_ONLY_ACE
Can be placed on a directory but does not apply to the directory,
only to newly created files/directories as specified by the above
two flags.
ACE4_NO_PROPAGATE_INHERIT_ACE
Can be placed on a directory. Normally when a new directory is
created and an ACE exists on the parent directory which is marked
ACE4_DIRECTORY_INHERIT_ACE, two ACEs are placed on the new
directory. One for the directory itself and one which is an
inheritable ACE for newly created directories. This flag tells
the server to not place an ACE on the newly created directory
which is inheritable by subdirectories of the created directory.
ACE4_SUCCESSFUL_ACCESS_ACE_FLAG
ACE4_FAILED_ACCESS_ACE_FLAG
The ACE4_SUCCESSFUL_ACCESS_ACE_FLAG (SUCCESS) and
ACE4_FAILED_ACCESS_ACE_FLAG (FAILED) flag bits relate only to
ACE4_SYSTEM_AUDIT_ACE_TYPE (AUDIT) and ACE4_SYSTEM_ALARM_ACE_TYPE
(ALARM) ACE types. If during the processing of the file's ACL,
the server encounters an AUDIT or ALARM ACE that matches the
principal attempting the OPEN, the server notes that fact, and the
presence, if any, of the SUCCESS and FAILED flags encountered in
the AUDIT or ALARM ACE. Once the server completes the ACL
processing, and the share reservation processing, and the OPEN
call, it then notes if the OPEN succeeded or failed. If the OPEN
succeeded, and if the SUCCESS flag was set for a matching AUDIT or
ALARM, then the appropriate AUDIT or ALARM event occurs. If the
OPEN failed, and if the FAILED flag was set for the matching AUDIT
or ALARM, then the appropriate AUDIT or ALARM event occurs.
Clearly either or both of the SUCCESS or FAILED can be set, but if
neither is set, the AUDIT or ALARM ACE is not useful.
The previously described processing applies to that of the ACCESS
operation as well. The difference being that "success" or
"failure" does not mean whether ACCESS returns NFS4_OK or not.
Success means whether ACCESS returns all requested and supported
bits. Failure means whether ACCESS failed to return a bit that
was requested and supported.
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ACE4_IDENTIFIER_GROUP
Indicates that the "who" refers to a GROUP as defined under UNIX
or a GROUP ACCOUNT as defined under Windows. Clients and servers
may ignore the ACE4_IDENTIFIER_GROUP flag on ACEs with a who value
equal to one of the special identifiers outlined in section "ACE
who".
The bitmask constants used for the flag field are as follows:
const ACE4_FILE_INHERIT_ACE = 0x00000001;
const ACE4_DIRECTORY_INHERIT_ACE = 0x00000002;
const ACE4_NO_PROPAGATE_INHERIT_ACE = 0x00000004;
const ACE4_INHERIT_ONLY_ACE = 0x00000008;
const ACE4_SUCCESSFUL_ACCESS_ACE_FLAG = 0x00000010;
const ACE4_FAILED_ACCESS_ACE_FLAG = 0x00000020;
const ACE4_IDENTIFIER_GROUP = 0x00000040;
A server need not support any of these flags. If the server supports
flags that are similar to, but not exactly the same as, these flags,
the implementation may define a mapping between the protocol-defined
flags and the implementation-defined flags. Again, the guiding
principle is that the file not appear to be more secure than it
really is.
For example, suppose a client tries to set an ACE with
ACE4_FILE_INHERIT_ACE set but not ACE4_DIRECTORY_INHERIT_ACE. If the
server does not support any form of ACL inheritance, the server
should reject the request with NFS4ERR_ATTRNOTSUPP. If the server
supports a single "inherit ACE" flag that applies to both files and
directories, the server may reject the request (i.e., requiring the
client to set both the file and directory inheritance flags). The
server may also accept the request and silently turn on the
ACE4_DIRECTORY_INHERIT_ACE flag.
3.16.4. ACE who
There are several special identifiers ("who") which need to be
understood universally, rather than in the context of a particular
DNS domain. Some of these identifiers cannot be understood when an
NFS client accesses the server, but have meaning when a local process
accesses the file. The ability to display and modify these
permissions is permitted over NFS, even if none of the access methods
on the server understands the identifiers.
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Who Description
_______________________________________________________________
"OWNER" The owner of the file.
"GROUP" The group associated with the file.
"EVERYONE" The world, including the owner and
owning group.
"INTERACTIVE" Accessed from an interactive terminal.
"NETWORK" Accessed via the network.
"DIALUP" Accessed as a dialup user to the server.
"BATCH" Accessed from a batch job.
"ANONYMOUS" Accessed without any authentication.
"AUTHENTICATED" Any authenticated user (opposite of
ANONYMOUS)
"SERVICE" Access from a system service.
To avoid conflict, these special identifiers are distinguish by an
appended "@" and should appear in the form "xxxx@" (note: no domain
name after the "@"). For example: ANONYMOUS@.
3.16.4.1. Discussion on EVERYONE@
It is important to note that "EVERYONE@" is not equivalent to the
UNIX "other" entity. This is because, by definition, UNIX "other"
does not include the owner or owning group of a file. "EVERYONE@"
means literally everyone, including the owner or owning group.
3.16.4.2. Discussion on OWNER@ and GROUP@
Due to the use of the special identifiers "OWNER@" and "GROUP@" to
indicate that an ACE applies to the the owner and owning group,
respectively, associated with a file, the ACL cannot be used to
determine the owner and owning group of a file. This information
should be indicated by the values of the owner and owner_group file
attributes returned by the server.
3.16.5. Mode Attribute
The NFS version 4 mode attribute is based on the UNIX mode bits. The
following bits are defined:
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const MODE4_SUID = 0x800; /* set user id on execution */
const MODE4_SGID = 0x400; /* set group id on execution */
const MODE4_SVTX = 0x200; /* save text even after use */
const MODE4_RUSR = 0x100; /* read permission: owner */
const MODE4_WUSR = 0x080; /* write permission: owner */
const MODE4_XUSR = 0x040; /* execute permission: owner */
const MODE4_RGRP = 0x020; /* read permission: group */
const MODE4_WGRP = 0x010; /* write permission: group */
const MODE4_XGRP = 0x008; /* execute permission: group */
const MODE4_ROTH = 0x004; /* read permission: other */
const MODE4_WOTH = 0x002; /* write permission: other */
const MODE4_XOTH = 0x001; /* execute permission: other */
Bits MODE4_RUSR, MODE4_WUSR, and MODE4_XUSR apply to the principal
identified in the owner attribute. Bits MODE4_RGRP, MODE4_WGRP, and
MODE4_XGRP apply to the principals identified in the owner_group
attribute. Bits MODE4_ROTH, MODE4_WOTH, MODE4_XOTH apply to any
principal that does not match that in the owner group, and does not
have a group matching that of the owner_group attribute.
The remaining bits are not defined by this protocol and MUST NOT be
used. The minor version mechanism must be used to define further bit
usage.
Note that in UNIX, if a file has the MODE4_SGID bit set and no
MODE4_XGRP bit set, then READ and WRITE must use mandatory file
locking.
3.16.6. Interaction Between Mode and ACL Attributes
As defined, there is a certain amount of overlap between ACL and mode
file attributes. Even though there is overlap, ACLs don't contain
all the information specified by a mode and modes can't possibly
contain all the information specified by an ACL.
For servers that support both mode and ACL, the mode's MODE4_R*,
MODE4_W* and MODE4_X* values should be computed from the ACL and
should be recomputed upon each SETATTR of ACL. Similarly, upon
SETATTR of mode, the ACL should be modified in order to allow the
mode computed from the ACL to be the same as the mode given to
SETATTR. The mode computed from any given ACL should be
deterministic. This means that given an ACL, the same mode will
always be computed.
For servers that support ACL and not mode, clients may handle
applications which set and get the mode by creating the correct ACL
to send to the server and by computing the mode from the ACL,
respectively. In this case, the methods used by the server to keep
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the mode in sync with the ACL can also be used by the client. These
methods are explained in sections Section 3.16.6.3 Section 3.16.6.1
and Section 3.16.6.2.
Since the mode can't possibly represent all of the information that
is defined by an ACL, there are some descrepencies to be aware of.
As explained in the section "Deficiencies in a Mode Representation of
an ACL", the mode bits computed from the ACL could potentially convey
more restrictive permissions than what would be granted via the ACL.
Because of this clients are not recommended to do their own access
checks based on the mode of a file.
Because the mode attribute includes bits (i.e. MODE4_SUID,
MODE4_SGID, MODE4_SVTX) that have nothing to do with ACL semantics,
it is permitted for clients to specify both the ACL attribute and
mode in the same SETATTR operation. However, because there is no
prescribed order for processing the attributes in a SETATTR, clients
may see differing results. For recommendations on how to achieve
consistent behavior, see Section 3.16.6.4 for recommendations.
3.16.6.1. Recomputing mode upon SETATTR of ACL
Keeping the mode and ACL attributes synchronized is important, but as
mentioned previously, the mode cannot possibly represent all of the
information in the ACL. Still, the mode should be modified to
represent the access as accurately as possible.
The general algorithm to assign a new mode attribute to an object
based on a new ACL being set is:
1. Walk through the ACEs in order, looking for ACEs with a "who"
value of OWNER@, GROUP@, or EVERYONE@.
2. It is understood that ACEs with a "who" value of OWNER@ affect
the *USR bits of the mode, GROUP@ affect *GRP bits, and EVERYONE@
affect *USR, *GRP, and *OTH bits.
3. If such an ACE specifies ALLOW or DENY for ACE4_READ_DATA,
ACE4_WRITE_DATA, or ACE4_EXECUTE, and the mode bits affected have
not been determined yet, set them to one (if ALLOW) or zero (if
DENY).
4. Upon completion, any mode bits as yet undetermined have a value
of zero.
This pseudocode more precisely describes the algorithm:
/* octal constants for the mode bits */
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RUSR = 0400
WUSR = 0200
XUSR = 0100
RGRP = 0040
WGRP = 0020
XGRP = 0010
ROTH = 0004
WOTH = 0002
XOTH = 0001
/*
* old_mode represents the previous value
* of the mode of the object.
*/
mode_t mode = 0, seen = 0;
for each ACE a {
if a.type is ALLOW or DENY and
ACE4_INHERIT_ONLY_ACE is not set in a.flags {
if a.who is OWNER@ {
if ((a.mask & ACE4_READ_DATA) &&
(! (seen & RUSR))) {
seen |= RUSR;
if a.type is ALLOW {
mode |= RUSR;
}
}
if ((a.mask & ACE4_WRITE_DATA) &&
(! (seen & WUSR))) {
seen |= WUSR;
if a.type is ALLOW {
mode |= WUSR;
}
}
if ((a.mask & ACE4_EXECUTE) &&
(! (seen & XUSR))) {
seen |= XUSR;
if a.type is ALLOW {
mode |= XUSR;
}
}
} else if a.who is GROUP@ {
if ((a.mask & ACE4_READ_DATA) &&
(! (seen & RGRP))) {
seen |= RGRP;
if a.type is ALLOW {
mode |= RGRP;
}
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}
if ((a.mask & ACE4_WRITE_DATA) &&
(! (seen & WGRP))) {
seen |= WGRP;
if a.type is ALLOW {
mode |= WGRP;
}
}
if ((a.mask & ACE4_EXECUTE) &&
(! (seen & XGRP))) {
seen |= XGRP;
if a.type is ALLOW {
mode |= XGRP;
}
}
} else if a.who is EVERYONE@ {
if (a.mask & ACE4_READ_DATA) {
if ! (seen & RUSR) {
seen |= RUSR;
if a.type is ALLOW {
mode |= RUSR;
}
}
if ! (seen & RGRP) {
seen |= RGRP;
if a.type is ALLOW {
mode |= RGRP;
}
}
if ! (seen & ROTH) {
seen |= ROTH;
if a.type is ALLOW {
mode |= ROTH;
}
}
}
if (a.mask & ACE4_WRITE_DATA) {
if ! (seen & WUSR) {
seen |= WUSR;
if a.type is ALLOW {
mode |= WUSR;
}
}
if ! (seen & WGRP) {
seen |= WGRP;
if a.type is ALLOW {
mode |= WGRP;
}
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}
if ! (seen & WOTH) {
seen |= WOTH;
if a.type is ALLOW {
mode |= WOTH;
}
}
}
if (a.mask & ACE4_EXECUTE) {
if ! (seen & XUSR) {
seen |= XUSR;
if a.type is ALLOW {
mode |= XUSR;
}
}
if ! (seen & XGRP) {
seen |= XGRP;
if a.type is ALLOW {
mode |= XGRP;
}
}
if ! (seen & XOTH) {
seen |= XOTH;
if a.type is ALLOW {
mode |= XOTH;
}
}
}
}
}
}
return mode | (old_mode & (SUID | SGID | SVTX))
3.16.6.2. Applying the mode given to CREATE or OPEN to an inherited ACL
The goal of implementing ACL inheritance is for newly created objects
to inherit the ACLs they were intended to inherit, but without
disregarding the mode that is given with the arguments to the CREATE
or OPEN operations. The general algorithm is as follows:
1. Form an ACL on the newly created object that is the concatenation
of all inheritable ACEs from its parent directory. Note that
there may be zero inheritable ACEs; thus, an object may start
with an empty ACL.
2. For each ACE in the new ACL, adjust its flags if necessary, and
possibly create two ACEs in place of one. This is necessary to
honor the intent of the inheritance- related flags and to
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preserve information about the original inheritable ACEs in the
case that they will be modified by other steps. The algorithm is
as follows:
A. If the ACE4_NO_PROPAGATE_INHERIT_ACE is set, or if the object
being created is not a directory, then clear the following
flags:
ACE4_NO_PROPAGATE_INHERIT_ACE
ACE4_FILE_INHERIT_ACE
ACE4_DIRECTORY_INHERIT_ACE
ACE4_INHERIT_ONLY_ACE
Continue on to the next ACE.
B. If the object being created is a directory and
ACE4_FILE_INHERIT_ACE is set, but ACE4_DIRECTORY_INHERIT_ACE
is NOT set, then we ensure that ACE4_INHERIT_ONLY_ACE is set.
Continue on to the next ACE. Otherwise:
C. If the type of the ACE is neither ALLOW nor DENY, then
continue on to the next ACE.
D. Copy the original ACE into a second, adjacent ACE.
E. On the first ACE, ensure that ACE4_INHERIT_ONLY_ACE is set.
F. On the second ACE, clear the following flags:
ACE4_NO_PROPAGATE_INHERIT_ACE
ACE4_FILE_INHERIT_ACE
ACE4_DIRECTORY_INHERIT_ACE
ACE4_INHERIT_ONLY_ACE
G. On the second ACE, if the type field is ALLOW, an
implementation MAY clear the following mask bits:
ACE4_WRITE_ACL
ACE4_WRITE_OWNER
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3. To ensure that the mode is honored, apply the algorithm for
applying a mode to a file/directory with an existing ACL on the
new object as described in Section 3.16.6.3, using the mode that
is to be used for file creation.
3.16.6.3. Applying a Mode to an Existing ACL
An existing ACL can mean two things in this context. One, that a
file/directory already exists and it has an ACL. Two, that a
directory has inheritable ACEs that will make up the ACL for any new
files or directories created therein.
The high-level goal of the behavior when a mode is set on a file with
an existing ACL is to take the new mode into account, without needing
to delete a pre-existing ACL.
When a mode is applied to an object, e.g. via SETATTR or CREATE/OPEN,
the ACL must be modified to accommodate the mode.
1. The ACL is traversed, one ACE at a time. For each ACE:
1. If the type of the ACE is neither ALLOW nor DENY, the ACE is
left unchanged. Continue to the next ACE.
2. If the ACE4_INHERIT_ONLY_ACE flag is set on the ACE, it is
left unchanged. Continue to the next ACE.
3. If either or both of ACE4_FILE_INHERIT_ACE or
ACE4_DIRECTORY_INHERIT_ACE are set:
1. A copy of the ACE is made, and placed in the ACL
immediately following the current ACE.
2. In the first ACE, the flag ACE4_INHERIT_ONLY_ACE is set.
3. In the second ACE, the following flags are cleared:
ACE4_FILE_INHERIT_ACE
ACE4_DIRECTORY_INHERIT_ACE
ACE4_NO_PROPAGATE_INHERIT_ACE
The algorithm continues on with the second ACE.
4. If the "who" field is one of the following:
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OWNER@
GROUP@
EVERYONE@
then the following mask bits are cleared:
ACE4_READ_DATA / ACE4_LIST_DIRECTORY
ACE4_WRITE_DATA / ACE4_ADD_FILE
ACE4_APPEND_DATA / ACE4_ADD_SUBDIRECTORY
ACE4_EXECUTE
At this point, we proceed to the next ACE.
5. Otherwise, if the "who" field did not match one of OWNER@,
GROUP@, or EVERYONE@, the following steps SHOULD be
performed.
1. If the type of the ACE is ALLOW, we check the preceding
ACE (if any). If it does not meet all of the following
criteria:
1. The type field is DENY.
2. The who field is the same as the current ACE.
3. The flag bit ACE4_IDENTIFIER_GROUP is the same as it
is in the current ACE, and no other flag bits are
set.
4. The mask bits are a subset of the mask bits of the
current ACE, and are also a subset of the following:
ACE4_READ_DATA / ACE4_LIST_DIRECTORY
ACE4_WRITE_DATA / ACE4_ADD_FILE
ACE4_APPEND_DATA / ACE4_ADD_SUBDIRECTORY
ACE4_EXECUTE
then an ACE of type DENY, with a who equal to the current
ACE, flag bits equal to (<current-ACE-flags> &
ACE4_IDENTIFIER_GROUP), and no mask bits, is prepended.
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2. The following modifications are made to the prepended
ACE. The intent is to mask the following ACE to disallow
ACE4_READ_DATA, ACE4_WRITE_DATA, ACE4_APPEND_DATA, or
ACE4_EXECUTE, based upon the group permissions of the new
mode. As a special case, if the ACE matches the current
owner of the file, the owner bits are used, rather than
the group bits. This is reflected in the algorithm
below.
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Let there be three bits defined:
#define READ 04
#define WRITE 02
#define EXEC 01
Let "amode" be the new mode, right-shifted three
bits, in order to have the group permission bits
placed in the three low order bits of amode,
i.e. amode = mode >> 3
If ACE4_IDENTIFIER_GROUP is not set in the flags,
and the "who" field of the ACE matches the owner
of the file, we shift amode three more bits, in
order to have the owner permission bits placed in
the three low order bits of amode:
amode = amode >> 3
amode is now used as follows:
If ACE4_READ_DATA is set on the current ACE:
If READ is set on amode:
ACE4_READ_DATA is cleared on the prepended ACE
else:
ACE4_READ_DATA is set on the prepended ACE
If ACE4_WRITE_DATA is set on the current ACE:
If WRITE is set on amode:
ACE4_WRITE_DATA is cleared on the prepended ACE
else:
ACE4_WRITE_DATA is set on the prepended ACE
If ACE4_APPEND_DATA is set on the current ACE:
If WRITE is set on amode:
ACE4_APPEND_DATA is cleared on the
prepended ACE
else:
ACE4_APPEND_DATA is set on the prepended ACE
If ACE4_EXECUTE is set on the current ACE:
If EXEC is set on amode:
ACE4_EXECUTE is cleared on the prepended ACE
else:
ACE4_EXECUTE is set on the prepended ACE
3. To conform with POSIX, and prevent cases where the owner
of the file is given permissions via an explicit group,
we implement the following step.
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If ACE4_IDENTIFIER_GROUP is set in the flags field of
the ALLOW ACE:
Let "mode" be the mode that we are chmoding to:
extramode = (mode >> 3) & 07
ownermode = mode >> 6
extramode &= ~ownermode
If extramode is not zero:
If extramode & READ:
Clear ACE4_READ_DATA in both the
prepended DENY ACE and the ALLOW ACE
If extramode & WRITE:
Clear ACE4_WRITE_DATA and ACE_APPEND_DATA
in both the prepended DENY ACE and the
ALLOW ACE
If extramode & EXEC:
Clear ACE4_EXECUTE in both the prepended
DENY ACE and the ALLOW ACE
2. If there are at least six ACEs, the final six ACEs are examined.
If they are not equal to the following ACEs:
A1) OWNER@:::DENY
A2) OWNER@:ACE4_WRITE_ACL/ACE4_WRITE_OWNER/
ACE4_WRITE_ATTRIBUTES/ACE4_WRITE_NAMED_ATTRIBUTES::ALLOW
A3) GROUP@::ACE4_IDENTIFIER_GROUP:DENY
A4) GROUP@::ACE4_IDENTIFIER_GROUP:ALLOW
A5) EVERYONE@:ACE4_WRITE_ACL/ACE4_WRITE_OWNER/
ACE4_WRITE_ATTRIBUTES/ACE4_WRITE_NAMED_ATTRIBUTES::DENY
A6) EVERYONE@:ACE4_READ_ACL/ACE4_READ_ATTRIBUTES/
ACE4_READ_NAMED_ATTRIBUTES/ACE4_SYNCHRONIZE::ALLOW
Then six ACEs matching the above are appended.
3. The final six ACEs are adjusted according to the incoming mode.
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/* octal constants for the mode bits */
RUSR = 0400
WUSR = 0200
XUSR = 0100
RGRP = 0040
WGRP = 0020
XGRP = 0010
ROTH = 0004
WOTH = 0002
XOTH = 0001
If RUSR is set: set ACE4_READ_DATA in A2
else: set ACE4_READ_DATA in A1
If WUSR is set: set ACE4_WRITE_DATA and ACE4_APPEND_DATA in A2
else: set ACE4_WRITE_DATA and ACE4_APPEND_DATA in A1
If XUSR is set: set ACE4_EXECUTE in A2
else: set ACE4_EXECUTE in A1
If RGRP is set: set ACE4_READ_DATA in A4
else: set ACE4_READ_DATA in A3
If WGRP is set: set ACE4_WRITE_DATA and ACE4_APPEND_DATA in A4
else: set ACE4_WRITE_DATA and ACE4_APPEND_DATA in A3
If XGRP is set: set ACE4_EXECUTE in A4
else: set ACE4_EXECUTE in A3
If ROTH is set: set ACE4_READ_DATA in A6
else: set ACE4_READ_DATA in A5
If WOTH is set: set ACE4_WRITE_DATA and ACE4_APPEND_DATA in A6
else: set ACE4_WRITE_DATA and ACE4_APPEND_DATA in A5
If XOTH is set: set ACE4_EXECUTE in A6
else: set ACE4_EXECUTE in A5
3.16.6.4. ACL and mode in the same SETATTR
The only reason that a mode and ACL should be set in the same SETATTR
is if the user wants to set the SUID, SGID and SVTX bits along with
setting the permissions by means of an ACL. There is still no way to
enforce which order the attributes will be set in, and it is likely
that different orders of operations will produce different results.
3.16.6.4.1. Client Side Recommendations
If an application needs to enforce a certain behavior, it is
recommended that the client implementations set mode and ACL in
separate SETATTR requests. This will produce consistent and expected
results.
If an application wants to set SUID, SGID and SVTX bits and an ACL:
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In the first SETATTR, set the mode with SUID, SGID and SVTX bits
as desired and all other bits with a value of 0.
In a following SETATTR (preferably in the same COMPOUND) set the
ACL.
3.16.6.4.2. Server Side Recommendations
If both mode and ACL are given to SETATTR, server implementations
should verify that the mode and ACL don't conflict, i.e. the mode
computed from the given ACL must be the same as the given mode,
excluding the SUID, SGID and SVTX bits. The algorithm for assigning
a new mode based on the ACL can be used. This is described in
section Section 3.16.6.1. If a server receives a request to set both
mode and ACL, but the two conflict, the server should return
NFS4ERR_INVAL.
3.16.6.5. Inheritance and turning it off
The inheritance of access permissions may be problematic if a user
cannot prevent their file from inheriting unwanted permissions. For
example, a user, "samf", sets up a shared project directory to be
used by everyone working on Project Foo. "lisagab" is a part of
Project Foo, but is working on something that should not be seen by
anyone else. How can "lisagab" make sure that any new files that she
creates in this shared project directory do not inherit anything that
could compromise the security of her work?
More relevant to the implementors of NFS version 4 clients and
servers is the question of how to communicate the fact that user,
"lisagab", doesn't want any permissions to be inherited to her newly
created file or directory.
To do this, implementors should standardize on what the behavior of
CREATE and OPEN must be if:
1. just mode is given
In this case, inheritance will take place, but the mode will be
applied to the inherited ACL as described in Section 3.16.6.1,
thereby modifying the ACL.
2. just ACL is given
In this case, inheritance will not take place, and the ACL as
defined in the CREATE or OPEN will be set without modification.
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3. both mode and ACL are given
In this case, implementors should verify that the mode and ACL
don't conflict, i.e. the mode computed from the given ACL must be
the same as the given mode. The algorithm for assigning a new
mode based on the ACL can be used. This is described in
Section 3.16.6.1) If a server receives a request to set both mode
and ACL, but the two conflict, the server should return
NFS4ERR_INVAL. If the mode and ACL don't conflict, inheritance
will not take placeand both, the mode and ACL, will be set
without modification.
4. neither mode nor ACL are given
In this case, inheritance will take place and no modifications to
the ACL will happen. It is worth noting that if no inheritable
ACEs exist on the parent directory, the file will be created with
an empty ACL, thus granting no accesses.
3.16.6.6. Deficiencies in a Mode Representation of an ACL
In the presence of an ACL, there are certain cases when the
representation of the mode is not guaranteed to be accurate. An
example of a situation is detailed below.
As mentioned in Section 3.16.6, the representation of the mode is
deterministic, but not guaranteed to be accurate. The mode bits
potentially convey a more restrictive permission than what will
actually be granted via the ACL.
Given the following ACL of two ACEs:
GROUP@:ACE4_READ_DATA/ACE4_WRITE_DATA/ACE4_EXECUTE:
ACE4_IDENTIFIER_GROUP:ALLOW
EVERYONE@:ACE4_READ_DATA/ACE4_WRITE_DATA/ACE4_EXECUTE::DENY
we would compute a mode of 0070. However, it is possible, even
likely, that the owner might be a member of the object's owning
group, and thus, the owner would be granted read, write, and execute
access to the object. This would conflict with the mode of 0070,
where an owner would be denied this access.
The only way to overcome this deficiency would be to determine
whether the object's owner is a member of the object's owning group.
This is difficult, but worse, on a POSIX or any UNIX-like system, it
is a process' membership in a group that is important, not a user's.
Thus, any fixed mode intended to represent the above ACL can be
incorrect.
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Example: administrative databases (possibly /etc/passwd and /etc/
group) indicate that the user "bob" is a member of the group "staff".
An object has the ACL given above, is owned by "bob", and has an
owning group of "staff". User "bob" has logged into the system, and
thus processes have been created owned by "bob" and having membership
in group "staff".
A mode representation of the above ACL could thus be 0770, due to
user "bob" having membership in group "staff". Now, the
administrative databases are changed, such that user "bob" is no
longer in group "staff". User "bob" logs in to the system again, and
thus more processes are created, this time owned by "bob" but NOT in
group "staff".
A mode of 0770 is inaccurate for processes not belonging to group
"staff". But even if the mode of the file were proactively changed
to 0070 at the time the group database was edited, mode 0070 would be
inaccurate for the pre-existing processes owned by user "bob" and
having membership in group "staff".
4. Filesystem Migration and Replication
With the use of the recommended attribute "fs_locations", the NFS
version 4 server has a method of providing filesystem migration or
replication services. For the purposes of migration and replication,
a filesystem will be defined as all files that share a given fsid
(both major and minor values are the same).
The fs_locations attribute provides a list of filesystem locations.
These locations are specified by providing the server name (either
DNS domain or IP address) and the path name representing the root of
the filesystem. Depending on the type of service being provided, the
list will provide a new location or a set of alternate locations for
the filesystem. The client will use this information to redirect its
requests to the new server.
4.1. Replication
It is expected that filesystem replication will be used in the case
of read-only data. Typically, the filesystem will be replicated on
two or more servers. The fs_locations attribute will provide the
list of these locations to the client. On first access of the
filesystem, the client should obtain the value of the fs_locations
attribute. If, in the future, the client finds the server
unresponsive, the client may attempt to use another server specified
by fs_locations.
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If applicable, the client must take the appropriate steps to recover
valid filehandles from the new server. This is described in more
detail in the following sections.
4.2. Migration
Filesystem migration is used to move a filesystem from one server to
another. Migration is typically used for a filesystem that is
writable and has a single copy. The expected use of migration is for
load balancing or general resource reallocation. The protocol does
not specify how the filesystem will be moved between servers. This
server-to-server transfer mechanism is left to the server
implementor. However, the method used to communicate the migration
event between client and server is specified here.
Once the servers participating in the migration have completed the
move of the filesystem, the error NFS4ERR_MOVED will be returned for
subsequent requests received by the original server. The
NFS4ERR_MOVED error is returned for all operations except PUTFH and
GETATTR. Upon receiving the NFS4ERR_MOVED error, the client will
obtain the value of the fs_locations attribute. The client will then
use the contents of the attribute to redirect its requests to the
specified server. To facilitate the use of GETATTR, operations such
as PUTFH must also be accepted by the server for the migrated file
system's filehandles. Note that if the server returns NFS4ERR_MOVED,
the server MUST support the fs_locations attribute.
If the client requests more attributes than just fs_locations, the
server may return fs_locations only. This is to be expected since
the server has migrated the filesystem and may not have a method of
obtaining additional attribute data.
The server implementor needs to be careful in developing a migration
solution. The server must consider all of the state information
clients may have outstanding at the server. This includes but is not
limited to locking/share state, delegation state, and asynchronous
file writes which are represented by WRITE and COMMIT verifiers. The
server should strive to minimize the impact on its clients during and
after the migration process.
4.3. Interpretation of the fs_locations Attribute
The fs_location attribute is structured in the following way:
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struct fs_location {
utf8str_cis server<>;
pathname4 rootpath;
};
struct fs_locations {
pathname4 fs_root;
fs_location locations<>;
};
The fs_location struct is used to represent the location of a
filesystem by providing a server name and the path to the root of the
filesystem. For a multi-homed server or a set of servers that use
the same rootpath, an array of server names may be provided. An
entry in the server array is an UTF8 string and represents one of a
traditional DNS host name, IPv4 address, or IPv6 address. It is not
a requirement that all servers that share the same rootpath be listed
in one fs_location struct. The array of server names is provided for
convenience. Servers that share the same rootpath may also be listed
in separate fs_location entries in the fs_locations attribute.
The fs_locations struct and attribute then contains an array of
locations. Since the name space of each server may be constructed
differently, the "fs_root" field is provided. The path represented
by fs_root represents the location of the filesystem in the server's
name space. Therefore, the fs_root path is only associated with the
server from which the fs_locations attribute was obtained. The
fs_root path is meant to aid the client in locating the filesystem at
the various servers listed.
As an example, there is a replicated filesystem located at two
servers (servA and servB). At servA the filesystem is located at
path "/a/b/c". At servB the filesystem is located at path "/x/y/z".
In this example the client accesses the filesystem first at servA
with a multi-component lookup path of "/a/b/c/d". Since the client
used a multi-component lookup to obtain the filehandle at "/a/b/c/d",
it is unaware that the filesystem's root is located in servA's name
space at "/a/b/c". When the client switches to servB, it will need
to determine that the directory it first referenced at servA is now
represented by the path "/x/y/z/d" on servB. To facilitate this, the
fs_locations attribute provided by servA would have a fs_root value
of "/a/b/c" and two entries in fs_location. One entry in fs_location
will be for itself (servA) and the other will be for servB with a
path of "/x/y/z". With this information, the client is able to
substitute "/x/y/z" for the "/a/b/c" at the beginning of its access
path and construct "/x/y/z/d" to use for the new server.
See the section "Security Considerations" for a discussion on the
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recommendations for the security flavor to be used by any GETATTR
operation that requests the "fs_locations" attribute.
4.4. Filehandle Recovery for Migration or Replication
Filehandles for filesystems that are replicated or migrated generally
have the same semantics as for filesystems that are not replicated or
migrated. For example, if a filesystem has persistent filehandles
and it is migrated to another server, the filehandle values for the
filesystem will be valid at the new server.
For volatile filehandles, the servers involved likely do not have a
mechanism to transfer filehandle format and content between
themselves. Therefore, a server may have difficulty in determining
if a volatile filehandle from an old server should return an error of
NFS4ERR_FHEXPIRED. Therefore, the client is informed, with the use
of the fh_expire_type attribute, whether volatile filehandles will
expire at the migration or replication event. If the bit
FH4_VOL_MIGRATION is set in the fh_expire_type attribute, the client
must treat the volatile filehandle as if the server had returned the
NFS4ERR_FHEXPIRED error. At the migration or replication event in
the presence of the FH4_VOL_MIGRATION bit, the client will not
present the original or old volatile filehandle to the new server.
The client will start its communication with the new server by
recovering its filehandles using the saved file names.
5. NFS Server Name Space
5.1. Server Exports
On a UNIX server the name space describes all the files reachable by
pathnames under the root directory or "/". On a Windows NT server
the name space constitutes all the files on disks named by mapped
disk letters. NFS server administrators rarely make the entire
server's filesystem name space available to NFS clients. More often
portions of the name space are made available via an "export"
feature. In previous versions of the NFS protocol, the root
filehandle for each export is obtained through the MOUNT protocol;
the client sends a string that identifies the export of name space
and the server returns the root filehandle for it. The MOUNT
protocol supports an EXPORTS procedure that will enumerate the
server's exports.
5.2. Browsing Exports
The NFS version 4 protocol provides a root filehandle that clients
can use to obtain filehandles for these exports via a multi-component
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LOOKUP. A common user experience is to use a graphical user
interface (perhaps a file "Open" dialog window) to find a file via
progressive browsing through a directory tree. The client must be
able to move from one export to another export via single-component,
progressive LOOKUP operations.
This style of browsing is not well supported by the NFS version 2 and
3 protocols. The client expects all LOOKUP operations to remain
within a single server filesystem. For example, the device attribute
will not change. This prevents a client from taking name space paths
that span exports.
An automounter on the client can obtain a snapshot of the server's
name space using the EXPORTS procedure of the MOUNT protocol. If it
understands the server's pathname syntax, it can create an image of
the server's name space on the client. The parts of the name space
that are not exported by the server are filled in with a "pseudo
filesystem" that allows the user to browse from one mounted
filesystem to another. There is a drawback to this representation of
the server's name space on the client: it is static. If the server
administrator adds a new export the client will be unaware of it.
5.3. Server Pseudo Filesystem
NFS version 4 servers avoid this name space inconsistency by
presenting all the exports within the framework of a single server
name space. An NFS version 4 client uses LOOKUP and READDIR
operations to browse seamlessly from one export to another. Portions
of the server name space that are not exported are bridged via a
"pseudo filesystem" that provides a view of exported directories
only. A pseudo filesystem has a unique fsid and behaves like a
normal, read only filesystem.
Based on the construction of the server's name space, it is possible
that multiple pseudo filesystems may exist. For example,
/a pseudo filesystem
/a/b real filesystem
/a/b/c pseudo filesystem
/a/b/c/d real filesystem
Each of the pseudo filesystems are considered separate entities and
therefore will have a unique fsid.
5.4. Multiple Roots
The DOS and Windows operating environments are sometimes described as
having "multiple roots". Filesystems are commonly represented as
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disk letters. MacOS represents filesystems as top level names. NFS
version 4 servers for these platforms can construct a pseudo file
system above these root names so that disk letters or volume names
are simply directory names in the pseudo root.
5.5. Filehandle Volatility
The nature of the server's pseudo filesystem is that it is a logical
representation of filesystem(s) available from the server.
Therefore, the pseudo filesystem is most likely constructed
dynamically when the server is first instantiated. It is expected
that the pseudo filesystem may not have an on disk counterpart from
which persistent filehandles could be constructed. Even though it is
preferable that the server provide persistent filehandles for the
pseudo filesystem, the NFS client should expect that pseudo file
system filehandles are volatile. This can be confirmed by checking
the associated "fh_expire_type" attribute for those filehandles in
question. If the filehandles are volatile, the NFS client must be
prepared to recover a filehandle value (e.g. with a multi-component
LOOKUP) when receiving an error of NFS4ERR_FHEXPIRED.
5.6. Exported Root
If the server's root filesystem is exported, one might conclude that
a pseudo-filesystem is not needed. This would be wrong. Assume the
following filesystems on a server:
/ disk1 (exported)
/a disk2 (not exported)
/a/b disk3 (exported)
Because disk2 is not exported, disk3 cannot be reached with simple
LOOKUPs. The server must bridge the gap with a pseudo-filesystem.
5.7. Mount Point Crossing
The server filesystem environment may be constructed in such a way
that one filesystem contains a directory which is 'covered' or
mounted upon by a second filesystem. For example:
/a/b (filesystem 1)
/a/b/c/d (filesystem 2)
The pseudo filesystem for this server may be constructed to look
like:
/ (place holder/not exported)
/a/b (filesystem 1)
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/a/b/c/d (filesystem 2)
It is the server's responsibility to present the pseudo filesystem
that is complete to the client. If the client sends a lookup request
for the path "/a/b/c/d", the server's response is the filehandle of
the filesystem "/a/b/c/d". In previous versions of the NFS protocol,
the server would respond with the filehandle of directory "/a/b/c/d"
within the filesystem "/a/b".
The NFS client will be able to determine if it crosses a server mount
point by a change in the value of the "fsid" attribute.
5.8. Security Policy and Name Space Presentation
The application of the server's security policy needs to be carefully
considered by the implementor. One may choose to limit the
viewability of portions of the pseudo filesystem based on the
server's perception of the client's ability to authenticate itself
properly. However, with the support of multiple security mechanisms
and the ability to negotiate the appropriate use of these mechanisms,
the server is unable to properly determine if a client will be able
to authenticate itself. If, based on its policies, the server
chooses to limit the contents of the pseudo filesystem, the server
may effectively hide filesystems from a client that may otherwise
have legitimate access.
As suggested practice, the server should apply the security policy of
a shared resource in the server's namespace to the components of the
resource's ancestors. For example:
/
/a/b
/a/b/c
The /a/b/c directory is a real filesystem and is the shared resource.
The security policy for /a/b/c is Kerberos with integrity. The
server should apply the same security policy to /, /a, and /a/b.
This allows for the extension of the protection of the server's
namespace to the ancestors of the real shared resource.
For the case of the use of multiple, disjoint security mechanisms in
the server's resources, the security for a particular object in the
server's namespace should be the union of all security mechanisms of
all direct descendants.
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6. File Locking and Share Reservations
Integrating locking into the NFS protocol necessarily causes it to be
stateful. With the inclusion of share reservations the protocol
becomes substantially more dependent on state than the traditional
combination of NFS and NLM [XNFS]. There are three components to
making this state manageable:
o Clear division between client and server
o Ability to reliably detect inconsistency in state between client
and server
o Simple and robust recovery mechanisms
In this model, the server owns the state information. The client
communicates its view of this state to the server as needed. The
client is also able to detect inconsistent state before modifying a
file.
To support Win32 share reservations it is necessary to atomically
OPEN or CREATE files. Having a separate share/unshare operation
would not allow correct implementation of the Win32 OpenFile API. In
order to correctly implement share semantics, the previous NFS
protocol mechanisms used when a file is opened or created (LOOKUP,
CREATE, ACCESS) need to be replaced. The NFS version 4 protocol has
an OPEN operation that subsumes the NFS version 3 methodology of
LOOKUP, CREATE, and ACCESS. However, because many operations require
a filehandle, the traditional LOOKUP is preserved to map a file name
to filehandle without establishing state on the server. The policy
of granting access or modifying files is managed by the server based
on the client's state. These mechanisms can implement policy ranging
from advisory only locking to full mandatory locking.
6.1. Locking
It is assumed that manipulating a lock is rare when compared to READ
and WRITE operations. It is also assumed that crashes and network
partitions are relatively rare. Therefore it is important that the
READ and WRITE operations have a lightweight mechanism to indicate if
they possess a held lock. A lock request contains the heavyweight
information required to establish a lock and uniquely define the lock
owner.
The following sections describe the transition from the heavy weight
information to the eventual stateid used for most client and server
locking and lease interactions.
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6.1.1. Client ID
For each LOCK request, the client must identify itself to the server.
This is done in such a way as to allow for correct lock
identification and crash recovery. A sequence of a SETCLIENTID
operation followed by a SETCLIENTID_CONFIRM operation is required to
establish the identification onto the server. Establishment of
identification by a new incarnation of the client also has the effect
of immediately breaking any leased state that a previous incarnation
of the client might have had on the server, as opposed to forcing the
new client incarnation to wait for the leases to expire. Breaking
the lease state amounts to the server removing all lock, share
reservation, and, where the server is not supporting the
CLAIM_DELEGATE_PREV claim type, all delegation state associated with
same client with the same identity. For discussion of delegation
state recovery, see the section "Delegation Recovery".
Client identification is encapsulated in the following structure:
struct nfs_client_id4 {
verifier4 verifier;
opaque id<NFS4_OPAQUE_LIMIT>;
};
The first field, verifier is a client incarnation verifier that is
used to detect client reboots. Only if the verifier is different
from that the server has previously recorded the client (as
identified by the second field of the structure, id) does the server
start the process of canceling the client's leased state.
The second field, id is a variable length string that uniquely
defines the client.
There are several considerations for how the client generates the id
string:
o The string should be unique so that multiple clients do not
present the same string. The consequences of two clients
presenting the same string range from one client getting an error
to one client having its leased state abruptly and unexpectedly
canceled.
o The string should be selected so the subsequent incarnations (e.g.
reboots) of the same client cause the client to present the same
string. The implementor is cautioned from an approach that
requires the string to be recorded in a local file because this
precludes the use of the implementation in an environment where
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there is no local disk and all file access is from an NFS version
4 server.
o The string should be different for each server network address
that the client accesses, rather than common to all server network
addresses. The reason is that it may not be possible for the
client to tell if same server is listening on multiple network
addresses. If the client issues SETCLIENTID with the same id
string to each network address of such a server, the server will
think it is the same client, and each successive SETCLIENTID will
cause the server to begin the process of removing the client's
previous leased state.
o The algorithm for generating the string should not assume that the
client's network address won't change. This includes changes
between client incarnations and even changes while the client is
stilling running in its current incarnation. This means that if
the client includes just the client's and server's network address
in the id string, there is a real risk, after the client gives up
the network address, that another client, using a similar
algorithm for generating the id string, will generate a
conflicting id string.
o Given the above considerations, an example of a well generated id
string is one that includes:
o The server's network address.
o The client's network address.
o For a user level NFS version 4 client, it should contain
additional information to distinguish the client from other user
level clients running on the same host, such as a process id or
other unique sequence.
o Additional information that tends to be unique, such as one or
more of:
* The client machine's serial number (for privacy reasons, it is
best to perform some one way function on the serial number).
* A MAC address.
* The timestamp of when the NFS version 4 software was first
installed on the client (though this is subject to the
previously mentioned caution about using information that is
stored in a file, because the file might only be accessible
over NFS version 4).
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* A true random number. However since this number ought to be
the same between client incarnations, this shares the same
problem as that of the using the timestamp of the software
installation.
As a security measure, the server MUST NOT cancel a client's leased
state if the principal established the state for a given id string is
not the same as the principal issuing the SETCLIENTID.
Note that SETCLIENTID and SETCLIENTID_CONFIRM has a secondary purpose
of establishing the information the server needs to make callbacks to
the client for purpose of supporting delegations. It is permitted to
change this information via SETCLIENTID and SETCLIENTID_CONFIRM
within the same incarnation of the client without removing the
client's leased state.
Once a SETCLIENTID and SETCLIENTID_CONFIRM sequence has successfully
completed, the client uses the short hand client identifier, of type
clientid4, instead of the longer and less compact nfs_client_id4
structure. This short hand client identifier (a clientid) is
assigned by the server and should be chosen so that it will not
conflict with a clientid previously assigned by the server. This
applies across server restarts or reboots. When a clientid is
presented to a server and that clientid is not recognized, as would
happen after a server reboot, the server will reject the request with
the error NFS4ERR_STALE_CLIENTID. When this happens, the client must
obtain a new clientid by use of the SETCLIENTID operation and then
proceed to any other necessary recovery for the server reboot case
(See the section "Server Failure and Recovery").
The client must also employ the SETCLIENTID operation when it
receives a NFS4ERR_STALE_STATEID error using a stateid derived from
its current clientid, since this also indicates a server reboot which
has invalidated the existing clientid (see the next section
"lock_owner and stateid Definition" for details).
See the detailed descriptions of SETCLIENTID and SETCLIENTID_CONFIRM
for a complete specification of the operations.
6.1.2. Server Release of Clientid
If the server determines that the client holds no associated state
for its clientid, the server may choose to release the clientid. The
server may make this choice for an inactive client so that resources
are not consumed by those intermittently active clients. If the
client contacts the server after this release, the server must ensure
the client receives the appropriate error so that it will use the
SETCLIENTID/SETCLIENTID_CONFIRM sequence to establish a new identity.
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It should be clear that the server must be very hesitant to release a
clientid since the resulting work on the client to recover from such
an event will be the same burden as if the server had failed and
restarted. Typically a server would not release a clientid unless
there had been no activity from that client for many minutes.
Note that if the id string in a SETCLIENTID request is properly
constructed, and if the client takes care to use the same principal
for each successive use of SETCLIENTID, then, barring an active
denial of service attack, NFS4ERR_CLID_INUSE should never be
returned.
However, client bugs, server bugs, or perhaps a deliberate change of
the principal owner of the id string (such as the case of a client
that changes security flavors, and under the new flavor, there is no
mapping to the previous owner) will in rare cases result in
NFS4ERR_CLID_INUSE.
In that event, when the server gets a SETCLIENTID for a client id
that currently has no state, or it has state, but the lease has
expired, rather than returning NFS4ERR_CLID_INUSE, the server MUST
allow the SETCLIENTID, and confirm the new clientid if followed by
the appropriate SETCLIENTID_CONFIRM.
6.1.3. lock_owner and stateid Definition
When requesting a lock, the client must present to the server the
clientid and an identifier for the owner of the requested lock.
These two fields are referred to as the lock_owner and the definition
of those fields are:
o A clientid returned by the server as part of the client's use of
the SETCLIENTID operation.
o A variable length opaque array used to uniquely define the owner
of a lock managed by the client.
This may be a thread id, process id, or other unique value.
When the server grants the lock, it responds with a unique stateid.
The stateid is used as a shorthand reference to the lock_owner, since
the server will be maintaining the correspondence between them.
The server is free to form the stateid in any manner that it chooses
as long as it is able to recognize invalid and out-of-date stateids.
This requirement includes those stateids generated by earlier
instances of the server. From this, the client can be properly
notified of a server restart. This notification will occur when the
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client presents a stateid to the server from a previous
instantiation.
The server must be able to distinguish the following situations and
return the error as specified:
o The stateid was generated by an earlier server instance (i.e.
before a server reboot). The error NFS4ERR_STALE_STATEID should
be returned.
o The stateid was generated by the current server instance but the
stateid no longer designates the current locking state for the
lockowner-file pair in question (i.e. one or more locking
operations has occurred). The error NFS4ERR_OLD_STATEID should be
returned.
This error condition will only occur when the client issues a
locking request which changes a stateid while an I/O request that
uses that stateid is outstanding.
o The stateid was generated by the current server instance but the
stateid does not designate a locking state for any active
lockowner-file pair. The error NFS4ERR_BAD_STATEID should be
returned.
This error condition will occur when there has been a logic error
on the part of the client or server. This should not happen.
One mechanism that may be used to satisfy these requirements is for
the server to,
o divide the "other" field of each stateid into two fields:
* A server verifier which uniquely designates a particular server
instantiation.
* An index into a table of locking-state structures.
o utilize the "seqid" field of each stateid, such that seqid is
monotonically incremented for each stateid that is associated with
the same index into the locking-state table.
By matching the incoming stateid and its field values with the state
held at the server, the server is able to easily determine if a
stateid is valid for its current instantiation and state. If the
stateid is not valid, the appropriate error can be supplied to the
client.
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6.1.4. Use of the stateid and Locking
All READ, WRITE and SETATTR operations contain a stateid. For the
purposes of this section, SETATTR operations which change the size
attribute of a file are treated as if they are writing the area
between the old and new size (i.e. the range truncated or added to
the file by means of the SETATTR), even where SETATTR is not
explicitly mentioned in the text.
If the lock_owner performs a READ or WRITE in a situation in which it
has established a lock or share reservation on the server (any OPEN
constitutes a share reservation) the stateid (previously returned by
the server) must be used to indicate what locks, including both
record locks and share reservations, are held by the lockowner. If
no state is established by the client, either record lock or share
reservation, a stateid of all bits 0 is used. Regardless whether a
stateid of all bits 0, or a stateid returned by the server is used,
if there is a conflicting share reservation or mandatory record lock
held on the file, the server MUST refuse to service the READ or WRITE
operation.
Share reservations are established by OPEN operations and by their
nature are mandatory in that when the OPEN denies READ or WRITE
operations, that denial results in such operations being rejected
with error NFS4ERR_LOCKED. Record locks may be implemented by the
server as either mandatory or advisory, or the choice of mandatory or
advisory behavior may be determined by the server on the basis of the
file being accessed (for example, some UNIX-based servers support a
"mandatory lock bit" on the mode attribute such that if set, record
locks are required on the file before I/O is possible). When record
locks are advisory, they only prevent the granting of conflicting
lock requests and have no effect on READs or WRITEs. Mandatory
record locks, however, prevent conflicting I/O operations. When they
are attempted, they are rejected with NFS4ERR_LOCKED. When the
client gets NFS4ERR_LOCKED on a file it knows it has the proper share
reservation for, it will need to issue a LOCK request on the region
of the file that includes the region the I/O was to be performed on,
with an appropriate locktype (i.e. READ*_LT for a READ operation,
WRITE*_LT for a WRITE operation).
With NFS version 3, there was no notion of a stateid so there was no
way to tell if the application process of the client sending the READ
or WRITE operation had also acquired the appropriate record lock on
the file. Thus there was no way to implement mandatory locking.
With the stateid construct, this barrier has been removed.
Note that for UNIX environments that support mandatory file locking,
the distinction between advisory and mandatory locking is subtle. In
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fact, advisory and mandatory record locks are exactly the same in so
far as the APIs and requirements on implementation. If the mandatory
lock attribute is set on the file, the server checks to see if the
lockowner has an appropriate shared (read) or exclusive (write)
record lock on the region it wishes to read or write to. If there is
no appropriate lock, the server checks if there is a conflicting lock
(which can be done by attempting to acquire the conflicting lock on
the behalf of the lockowner, and if successful, release the lock
after the READ or WRITE is done), and if there is, the server returns
NFS4ERR_LOCKED.
For Windows environments, there are no advisory record locks, so the
server always checks for record locks during I/O requests.
Thus, the NFS version 4 LOCK operation does not need to distinguish
between advisory and mandatory record locks. It is the NFS version 4
server's processing of the READ and WRITE operations that introduces
the distinction.
Every stateid other than the special stateid values noted in this
section, whether returned by an OPEN-type operation (i.e. OPEN,
OPEN_DOWNGRADE), or by a LOCK-type operation (i.e. LOCK or LOCKU),
defines an access mode for the file (i.e. READ, WRITE, or READ-
WRITE) as established by the original OPEN which began the stateid
sequence, and as modified by subsequent OPENs and OPEN_DOWNGRADEs
within that stateid sequence. When a READ, WRITE, or SETATTR which
specifies the size attribute, is done, the operation is subject to
checking against the access mode to verify that the operation is
appropriate given the OPEN with which the operation is associated.
In the case of WRITE-type operations (i.e. WRITEs and SETATTRs which
set size), the server must verify that the access mode allows writing
and return an NFS4ERR_OPENMODE error if it does not. In the case, of
READ, the server may perform the corresponding check on the access
mode, or it may choose to allow READ on opens for WRITE only, to
accommodate clients whose write implementation may unavoidably do
reads (e.g. due to buffer cache constraints). However, even if READs
are allowed in these circumstances, the server MUST still check for
locks that conflict with the READ (e.g. another open specify denial
of READs). Note that a server which does enforce the access mode
check on READs need not explicitly check for conflicting share
reservations since the existence of OPEN for read access guarantees
that no conflicting share reservation can exist.
A stateid of all bits 1 (one) MAY allow READ operations to bypass
locking checks at the server. However, WRITE operations with a
stateid with bits all 1 (one) MUST NOT bypass locking checks and are
treated exactly the same as if a stateid of all bits 0 were used.
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A lock may not be granted while a READ or WRITE operation using one
of the special stateids is being performed and the range of the lock
request conflicts with the range of the READ or WRITE operation. For
the purposes of this paragraph, a conflict occurs when a shared lock
is requested and a WRITE operation is being performed, or an
exclusive lock is requested and either a READ or a WRITE operation is
being performed. A SETATTR that sets size is treated similarly to a
WRITE as discussed above.
6.1.5. Sequencing of Lock Requests
Locking is different than most NFS operations as it requires "at-
most-one" semantics that are not provided by ONCRPC. ONCRPC over a
reliable transport is not sufficient because a sequence of locking
requests may span multiple TCP connections. In the face of
retransmission or reordering, lock or unlock requests must have a
well defined and consistent behavior. To accomplish this, each lock
request contains a sequence number that is a consecutively increasing
integer. Different lock_owners have different sequences. The server
maintains the last sequence number (L) received and the response that
was returned. The first request issued for any given lock_owner is
issued with a sequence number of zero.
Note that for requests that contain a sequence number, for each
lock_owner, there should be no more than one outstanding request.
If a request (r) with a previous sequence number (r < L) is received,
it is rejected with the return of error NFS4ERR_BAD_SEQID. Given a
properly-functioning client, the response to (r) must have been
received before the last request (L) was sent. If a duplicate of
last request (r == L) is received, the stored response is returned.
If a request beyond the next sequence (r == L + 2) is received, it is
rejected with the return of error NFS4ERR_BAD_SEQID. Sequence
history is reinitialized whenever the SETCLIENTID/SETCLIENTID_CONFIRM
sequence changes the client verifier.
Since the sequence number is represented with an unsigned 32-bit
integer, the arithmetic involved with the sequence number is mod
2^32. For an example of modulo arithetic involving sequence numbers
see [RFC793].
It is critical the server maintain the last response sent to the
client to provide a more reliable cache of duplicate non-idempotent
requests than that of the traditional cache described in [Juszczak].
The traditional duplicate request cache uses a least recently used
algorithm for removing unneeded requests. However, the last lock
request and response on a given lock_owner must be cached as long as
the lock state exists on the server.
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The client MUST monotonically increment the sequence number for the
CLOSE, LOCK, LOCKU, OPEN, OPEN_CONFIRM, and OPEN_DOWNGRADE
operations. This is true even in the event that the previous
operation that used the sequence number received an error. The only
exception to this rule is if the previous operation received one of
the following errors: NFS4ERR_STALE_CLIENTID, NFS4ERR_STALE_STATEID,
NFS4ERR_BAD_STATEID, NFS4ERR_BAD_SEQID, NFS4ERR_BADXDR,
NFS4ERR_RESOURCE, NFS4ERR_NOFILEHANDLE.
6.1.6. Recovery from Replayed Requests
As described above, the sequence number is per lock_owner. As long
as the server maintains the last sequence number received and follows
the methods described above, there are no risks of a Byzantine router
re-sending old requests. The server need only maintain the
(lock_owner, sequence number) state as long as there are open files
or closed files with locks outstanding.
LOCK, LOCKU, OPEN, OPEN_DOWNGRADE, and CLOSE each contain a sequence
number and therefore the risk of the replay of these operations
resulting in undesired effects is non-existent while the server
maintains the lock_owner state.
6.1.7. Releasing lock_owner State
When a particular lock_owner no longer holds open or file locking
state at the server, the server may choose to release the sequence
number state associated with the lock_owner. The server may make
this choice based on lease expiration, for the reclamation of server
memory, or other implementation specific details. In any event, the
server is able to do this safely only when the lock_owner no longer
is being utilized by the client. The server may choose to hold the
lock_owner state in the event that retransmitted requests are
received. However, the period to hold this state is implementation
specific.
In the case that a LOCK, LOCKU, OPEN_DOWNGRADE, or CLOSE is
retransmitted after the server has previously released the lock_owner
state, the server will find that the lock_owner has no files open and
an error will be returned to the client. If the lock_owner does have
a file open, the stateid will not match and again an error is
returned to the client.
6.1.8. Use of Open Confirmation
In the case that an OPEN is retransmitted and the lock_owner is being
used for the first time or the lock_owner state has been previously
released by the server, the use of the OPEN_CONFIRM operation will
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prevent incorrect behavior. When the server observes the use of the
lock_owner for the first time, it will direct the client to perform
the OPEN_CONFIRM for the corresponding OPEN. This sequence
establishes the use of an lock_owner and associated sequence number.
Since the OPEN_CONFIRM sequence connects a new open_owner on the
server with an existing open_owner on a client, the sequence number
may have any value. The OPEN_CONFIRM step assures the server that
the value received is the correct one. See the section "OPEN_CONFIRM
- Confirm Open" for further details.
There are a number of situations in which the requirement to confirm
an OPEN would pose difficulties for the client and server, in that
they would be prevented from acting in a timely fashion on
information received, because that information would be provisional,
subject to deletion upon non-confirmation. Fortunately, these are
situations in which the server can avoid the need for confirmation
when responding to open requests. The two constraints are:
o The server must not bestow a delegation for any open which would
require confirmation.
o The server MUST NOT require confirmation on a reclaim-type open
(i.e. one specifying claim type CLAIM_PREVIOUS or
CLAIM_DELEGATE_PREV).
These constraints are related in that reclaim-type opens are the only
ones in which the server may be required to send a delegation. For
CLAIM_NULL, sending the delegation is optional while for
CLAIM_DELEGATE_CUR, no delegation is sent.
Delegations being sent with an open requiring confirmation are
troublesome because recovering from non-confirmation adds undue
complexity to the protocol while requiring confirmation on reclaim-
type opens poses difficulties in that the inability to resolve the
status of the reclaim until lease expiration may make it difficult to
have timely determination of the set of locks being reclaimed (since
the grace period may expire).
Requiring open confirmation on reclaim-type opens is avoidable
because of the nature of the environments in which such opens are
done. For CLAIM_PREVIOUS opens, this is immediately after server
reboot, so there should be no time for lockowners to be created,
found to be unused, and recycled. For CLAIM_DELEGATE_PREV opens, we
are dealing with a client reboot situation. A server which supports
delegation can be sure that no lockowners for that client have been
recycled since client initialization and thus can ensure that
confirmation will not be required.
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6.2. Lock Ranges
The protocol allows a lock owner to request a lock with a byte range
and then either upgrade or unlock a sub-range of the initial lock.
It is expected that this will be an uncommon type of request. In any
case, servers or server filesystems may not be able to support sub-
range lock semantics. In the event that a server receives a locking
request that represents a sub-range of current locking state for the
lock owner, the server is allowed to return the error
NFS4ERR_LOCK_RANGE to signify that it does not support sub-range lock
operations. Therefore, the client should be prepared to receive this
error and, if appropriate, report the error to the requesting
application.
The client is discouraged from combining multiple independent locking
ranges that happen to be adjacent into a single request since the
server may not support sub-range requests and for reasons related to
the recovery of file locking state in the event of server failure.
As discussed in the section "Server Failure and Recovery" below, the
server may employ certain optimizations during recovery that work
effectively only when the client's behavior during lock recovery is
similar to the client's locking behavior prior to server failure.
6.3. Upgrading and Downgrading Locks
If a client has a write lock on a record, it can request an atomic
downgrade of the lock to a read lock via the LOCK request, by setting
the type to READ_LT. If the server supports atomic downgrade, the
request will succeed. If not, it will return NFS4ERR_LOCK_NOTSUPP.
The client should be prepared to receive this error, and if
appropriate, report the error to the requesting application.
If a client has a read lock on a record, it can request an atomic
upgrade of the lock to a write lock via the LOCK request by setting
the type to WRITE_LT or WRITEW_LT. If the server does not support
atomic upgrade, it will return NFS4ERR_LOCK_NOTSUPP. If the upgrade
can be achieved without an existing conflict, the request will
succeed. Otherwise, the server will return either NFS4ERR_DENIED or
NFS4ERR_DEADLOCK. The error NFS4ERR_DEADLOCK is returned if the
client issued the LOCK request with the type set to WRITEW_LT and the
server has detected a deadlock. The client should be prepared to
receive such errors and if appropriate, report the error to the
requesting application.
6.4. Blocking Locks
Some clients require the support of blocking locks. The NFS version
4 protocol must not rely on a callback mechanism and therefore is
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unable to notify a client when a previously denied lock has been
granted. Clients have no choice but to continually poll for the
lock. This presents a fairness problem. Two new lock types are
added, READW and WRITEW, and are used to indicate to the server that
the client is requesting a blocking lock. The server should maintain
an ordered list of pending blocking locks. When the conflicting lock
is released, the server may wait the lease period for the first
waiting client to re-request the lock. After the lease period
expires the next waiting client request is allowed the lock. Clients
are required to poll at an interval sufficiently small that it is
likely to acquire the lock in a timely manner. The server is not
required to maintain a list of pending blocked locks as it is used to
increase fairness and not correct operation. Because of the
unordered nature of crash recovery, storing of lock state to stable
storage would be required to guarantee ordered granting of blocking
locks.
Servers may also note the lock types and delay returning denial of
the request to allow extra time for a conflicting lock to be
released, allowing a successful return. In this way, clients can
avoid the burden of needlessly frequent polling for blocking locks.
The server should take care in the length of delay in the event the
client retransmits the request.
6.5. Lease Renewal
The purpose of a lease is to allow a server to remove stale locks
that are held by a client that has crashed or is otherwise
unreachable. It is not a mechanism for cache consistency and lease
renewals may not be denied if the lease interval has not expired.
The following events cause implicit renewal of all of the leases for
a given client (i.e. all those sharing a given clientid). Each of
these is a positive indication that the client is still active and
that the associated state held at the server, for the client, is
still valid.
o An OPEN with a valid clientid.
o Any operation made with a valid stateid (CLOSE, DELEGRETURN, LOCK,
LOCKU, OPEN, OPEN_CONFIRM, OPEN_DOWNGRADE, READ, SETATTR, WRITE).
This does not include the special stateids of all bits 0 or all
bits 1.
Note that if the client had restarted or rebooted, the client
would not be making these requests without issuing the
SETCLIENTID/SETCLIENTID_CONFIRM sequence. The use of the
SETCLIENTID/SETCLIENTID_CONFIRM sequence (one that changes the
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client verifier) notifies the server to drop the locking state
associated with the client. SETCLIENTID/SETCLIENTID_CONFIRM never
renews a lease.
If the server has rebooted, the stateids (NFS4ERR_STALE_STATEID
error) or the clientid (NFS4ERR_STALE_CLIENTID error) will not be
valid hence preventing spurious renewals.
This approach allows for low overhead lease renewal which scales
well. In the typical case no extra RPC calls are required for lease
renewal and in the worst case one RPC is required every lease period
(i.e. a RENEW operation). The number of locks held by the client is
not a factor since all state for the client is involved with the
lease renewal action.
Since all operations that create a new lease also renew existing
leases, the server must maintain a common lease expiration time for
all valid leases for a given client. This lease time can then be
easily updated upon implicit lease renewal actions.
6.6. Crash Recovery
The important requirement in crash recovery is that both the client
and the server know when the other has failed. Additionally, it is
required that a client sees a consistent view of data across server
restarts or reboots. All READ and WRITE operations that may have
been queued within the client or network buffers must wait until the
client has successfully recovered the locks protecting the READ and
WRITE operations.
6.6.1. Client Failure and Recovery
In the event that a client fails, the server may recover the client's
locks when the associated leases have expired. Conflicting locks
from another client may only be granted after this lease expiration.
If the client is able to restart or reinitialize within the lease
period the client may be forced to wait the remainder of the lease
period before obtaining new locks.
To minimize client delay upon restart, lock requests are associated
with an instance of the client by a client supplied verifier. This
verifier is part of the initial SETCLIENTID call made by the client.
The server returns a clientid as a result of the SETCLIENTID
operation. The client then confirms the use of the clientid with
SETCLIENTID_CONFIRM. The clientid in combination with an opaque
owner field is then used by the client to identify the lock owner for
OPEN. This chain of associations is then used to identify all locks
for a particular client.
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Since the verifier will be changed by the client upon each
initialization, the server can compare a new verifier to the verifier
associated with currently held locks and determine that they do not
match. This signifies the client's new instantiation and subsequent
loss of locking state. As a result, the server is free to release
all locks held which are associated with the old clientid which was
derived from the old verifier.
Note that the verifier must have the same uniqueness properties of
the verifier for the COMMIT operation.
6.6.2. Server Failure and Recovery
If the server loses locking state (usually as a result of a restart
or reboot), it must allow clients time to discover this fact and re-
establish the lost locking state. The client must be able to re-
establish the locking state without having the server deny valid
requests because the server has granted conflicting access to another
client. Likewise, if there is the possibility that clients have not
yet re-established their locking state for a file, the server must
disallow READ and WRITE operations for that file. The duration of
this recovery period is equal to the duration of the lease period.
A client can determine that server failure (and thus loss of locking
state) has occurred, when it receives one of two errors. The
NFS4ERR_STALE_STATEID error indicates a stateid invalidated by a
reboot or restart. The NFS4ERR_STALE_CLIENTID error indicates a
clientid invalidated by reboot or restart. When either of these are
received, the client must establish a new clientid (See the section
"Client ID") and re-establish the locking state as discussed below.
The period of special handling of locking and READs and WRITEs, equal
in duration to the lease period, is referred to as the "grace
period". During the grace period, clients recover locks and the
associated state by reclaim-type locking requests (i.e. LOCK
requests with reclaim set to true and OPEN operations with a claim
type of CLAIM_PREVIOUS). During the grace period, the server must
reject READ and WRITE operations and non-reclaim locking requests
(i.e. other LOCK and OPEN operations) with an error of NFS4ERR_GRACE.
If the server can reliably determine that granting a non-reclaim
request will not conflict with reclamation of locks by other clients,
the NFS4ERR_GRACE error does not have to be returned and the non-
reclaim client request can be serviced. For the server to be able to
service READ and WRITE operations during the grace period, it must
again be able to guarantee that no possible conflict could arise
between an impending reclaim locking request and the READ or WRITE
operation. If the server is unable to offer that guarantee, the
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NFS4ERR_GRACE error must be returned to the client.
For a server to provide simple, valid handling during the grace
period, the easiest method is to simply reject all non-reclaim
locking requests and READ and WRITE operations by returning the
NFS4ERR_GRACE error. However, a server may keep information about
granted locks in stable storage. With this information, the server
could determine if a regular lock or READ or WRITE operation can be
safely processed.
For example, if a count of locks on a given file is available in
stable storage, the server can track reclaimed locks for the file and
when all reclaims have been processed, non-reclaim locking requests
may be processed. This way the server can ensure that non-reclaim
locking requests will not conflict with potential reclaim requests.
With respect to I/O requests, if the server is able to determine that
there are no outstanding reclaim requests for a file by information
from stable storage or another similar mechanism, the processing of
I/O requests could proceed normally for the file.
To reiterate, for a server that allows non-reclaim lock and I/O
requests to be processed during the grace period, it MUST determine
that no lock subsequently reclaimed will be rejected and that no lock
subsequently reclaimed would have prevented any I/O operation
processed during the grace period.
Clients should be prepared for the return of NFS4ERR_GRACE errors for
non-reclaim lock and I/O requests. In this case the client should
employ a retry mechanism for the request. A delay (on the order of
several seconds) between retries should be used to avoid overwhelming
the server. Further discussion of the general issue is included in
[Floyd]. The client must account for the server that is able to
perform I/O and non-reclaim locking requests within the grace period
as well as those that can not do so.
A reclaim-type locking request outside the server's grace period can
only succeed if the server can guarantee that no conflicting lock or
I/O request has been granted since reboot or restart.
A server may, upon restart, establish a new value for the lease
period. Therefore, clients should, once a new clientid is
established, refetch the lease_time attribute and use it as the basis
for lease renewal for the lease associated with that server.
However, the server must establish, for this restart event, a grace
period at least as long as the lease period for the previous server
instantiation. This allows the client state obtained during the
previous server instance to be reliably re-established.
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6.6.3. Network Partitions and Recovery
If the duration of a network partition is greater than the lease
period provided by the server, the server will have not received a
lease renewal from the client. If this occurs, the server may free
all locks held for the client. As a result, all stateids held by the
client will become invalid or stale. Once the client is able to
reach the server after such a network partition, all I/O submitted by
the client with the now invalid stateids will fail with the server
returning the error NFS4ERR_EXPIRED. Once this error is received,
the client will suitably notify the application that held the lock.
As a courtesy to the client or as an optimization, the server may
continue to hold locks on behalf of a client for which recent
communication has extended beyond the lease period. If the server
receives a lock or I/O request that conflicts with one of these
courtesy locks, the server must free the courtesy lock and grant the
new request.
When a network partition is combined with a server reboot, there are
edge conditions that place requirements on the server in order to
avoid silent data corruption following the server reboot. Two of
these edge conditions are known, and are discussed below.
The first edge condition has the following scenario:
1. Client A acquires a lock.
2. Client A and server experience mutual network partition, such
that client A is unable to renew its lease.
3. Client A's lease expires, so server releases lock.
4. Client B acquires a lock that would have conflicted with that of
Client A.
5. Client B releases the lock
6. Server reboots
7. Network partition between client A and server heals.
8. Client A issues a RENEW operation, and gets back a
NFS4ERR_STALE_CLIENTID.
9. Client A reclaims its lock within the server's grace period.
Thus, at the final step, the server has erroneously granted client
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A's lock reclaim. If client B modified the object the lock was
protecting, client A will experience object corruption.
The second known edge condition follows:
1. Client A acquires a lock.
2. Server reboots.
3. Client A and server experience mutual network partition, such
that client A is unable to reclaim its lock within the grace
period.
4. Server's reclaim grace period ends. Client A has no locks
recorded on server.
5. Client B acquires a lock that would have conflicted with that of
Client A.
6. Client B releases the lock
7. Server reboots a second time
8. Network partition between client A and server heals.
9. Client A issues a RENEW operation, and gets back a
NFS4ERR_STALE_CLIENTID.
10. Client A reclaims its lock within the server's grace period.
As with the first edge condition, the final step of the scenario of
the second edge condition has the server erroneously granting client
A's lock reclaim.
Solving the first and second edge conditions requires that the server
either assume after it reboots that edge condition occurs, and thus
return NFS4ERR_NO_GRACE for all reclaim attempts, or that the server
record some information stable storage. The amount of information
the server records in stable storage is in inverse proportion to how
harsh the server wants to be whenever the edge conditions occur. The
server that is completely tolerant of all edge conditions will record
in stable storage every lock that is acquired, removing the lock
record from stable storage only when the lock is unlocked by the
client and the lock's lockowner advances the sequence number such
that the lock release is not the last stateful event for the
lockowner's sequence. For the two aforementioned edge conditions,
the harshest a server can be, and still support a grace period for
reclaims, requires that the server record in stable storage
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information some minimal information. For example, a server
implementation could, for each client, save in stable storage a
record containing:
o the client's id string
o a boolean that indicates if the client's lease expired or if there
was administrative intervention (see the section, Server
Revocation of Locks) to revoke a record lock, share reservation,
or delegation
o a timestamp that is updated the first time after a server boot or
reboot the client acquires record locking, share reservation, or
delegation state on the server. The timestamp need not be updated
on subsequent lock requests until the server reboots.
The server implementation would also record in the stable storage the
timestamps from the two most recent server reboots.
Assuming the above record keeping, for the first edge condition,
after the server reboots, the record that client A's lease expired
means that another client could have acquired a conflicting record
lock, share reservation, or delegation. Hence the server must reject
a reclaim from client A with the error NFS4ERR_NO_GRACE.
For the second edge condition, after the server reboots for a second
time, the record that the client had an unexpired record lock, share
reservation, or delegation established before the server's previous
incarnation means that the server must reject a reclaim from client A
with the error NFS4ERR_NO_GRACE.
Regardless of the level and approach to record keeping, the server
MUST implement one of the following strategies (which apply to
reclaims of share reservations, record locks, and delegations):
1. Reject all reclaims with NFS4ERR_NO_GRACE. This is superharsh,
but necessary if the server does not want to record lock state in
stable storage.
2. Record sufficient state in stable storage such that all known
edge conditions involving server reboot, including the two noted
in this section, are detected. False positives are acceptable.
Note that at this time, it is not known if there are other edge
conditions.
In the event, after a server reboot, the server determines that
there is unrecoverable damage or corruption to the the stable
storage, then for all clients and/or locks affected, the server
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MUST return NFS4ERR_NO_GRACE.
A mandate for the client's handling of the NFS4ERR_NO_GRACE error is
outside the scope of this specification, since the strategies for
such handling are very dependent on the client's operating
environment. However, one potential approach is described below.
When the client receives NFS4ERR_NO_GRACE, it could examine the
change attribute of the objects the client is trying to reclaim state
for, and use that to determine whether to re-establish the state via
normal OPEN or LOCK requests. This is acceptable provided the
client's operating environment allows it. In otherwords, the client
implementor is advised to document for his users the behavior. The
client could also inform the application that its record lock or
share reservations (whether they were delegated or not) have been
lost, such as via a UNIX signal, a GUI pop-up window, etc. See the
section, "Data Caching and Revocation" for a discussion of what the
client should do for dealing with unreclaimed delegations on client
state.
For further discussion of revocation of locks see the section "Server
Revocation of Locks".
6.7. Recovery from a Lock Request Timeout or Abort
In the event a lock request times out, a client may decide to not
retry the request. The client may also abort the request when the
process for which it was issued is terminated (e.g. in UNIX due to a
signal). It is possible though that the server received the request
and acted upon it. This would change the state on the server without
the client being aware of the change. It is paramount that the
client re-synchronize state with server before it attempts any other
operation that takes a seqid and/or a stateid with the same
lock_owner. This is straightforward to do without a special re-
synchronize operation.
Since the server maintains the last lock request and response
received on the lock_owner, for each lock_owner, the client should
cache the last lock request it sent such that the lock request did
not receive a response. From this, the next time the client does a
lock operation for the lock_owner, it can send the cached request, if
there is one, and if the request was one that established state (e.g.
a LOCK or OPEN operation), the server will return the cached result
or if never saw the request, perform it. The client can follow up
with a request to remove the state (e.g. a LOCKU or CLOSE operation).
With this approach, the sequencing and stateid information on the
client and server for the given lock_owner will re-synchronize and in
turn the lock state will re-synchronize.
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6.8. Server Revocation of Locks
At any point, the server can revoke locks held by a client and the
client must be prepared for this event. When the client detects that
its locks have been or may have been revoked, the client is
responsible for validating the state information between itself and
the server. Validating locking state for the client means that it
must verify or reclaim state for each lock currently held.
The first instance of lock revocation is upon server reboot or re-
initialization. In this instance the client will receive an error
(NFS4ERR_STALE_STATEID or NFS4ERR_STALE_CLIENTID) and the client will
proceed with normal crash recovery as described in the previous
section.
The second lock revocation event is the inability to renew the lease
before expiration. While this is considered a rare or unusual event,
the client must be prepared to recover. Both the server and client
will be able to detect the failure to renew the lease and are capable
of recovering without data corruption. For the server, it tracks the
last renewal event serviced for the client and knows when the lease
will expire. Similarly, the client must track operations which will
renew the lease period. Using the time that each such request was
sent and the time that the corresponding reply was received, the
client should bound the time that the corresponding renewal could
have occurred on the server and thus determine if it is possible that
a lease period expiration could have occurred.
The third lock revocation event can occur as a result of
administrative intervention within the lease period. While this is
considered a rare event, it is possible that the server's
administrator has decided to release or revoke a particular lock held
by the client. As a result of revocation, the client will receive an
error of NFS4ERR_ADMIN_REVOKED. In this instance the client may
assume that only the lock_owner's locks have been lost. The client
notifies the lock holder appropriately. The client may not assume
the lease period has been renewed as a result of failed operation.
When the client determines the lease period may have expired, the
client must mark all locks held for the associated lease as
"unvalidated". This means the client has been unable to re-establish
or confirm the appropriate lock state with the server. As described
in the previous section on crash recovery, there are scenarios in
which the server may grant conflicting locks after the lease period
has expired for a client. When it is possible that the lease period
has expired, the client must validate each lock currently held to
ensure that a conflicting lock has not been granted. The client may
accomplish this task by issuing an I/O request, either a pending I/O
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or a zero-length read, specifying the stateid associated with the
lock in question. If the response to the request is success, the
client has validated all of the locks governed by that stateid and
re-established the appropriate state between itself and the server.
If the I/O request is not successful, then one or more of the locks
associated with the stateid was revoked by the server and the client
must notify the owner.
6.9. Share Reservations
A share reservation is a mechanism to control access to a file. It
is a separate and independent mechanism from record locking. When a
client opens a file, it issues an OPEN operation to the server
specifying the type of access required (READ, WRITE, or BOTH) and the
type of access to deny others (deny NONE, READ, WRITE, or BOTH). If
the OPEN fails the client will fail the application's open request.
Pseudo-code definition of the semantics:
if (request.access == 0)
return (NFS4ERR_INVAL)
else
if ((request.access & file_state.deny)) ||
(request.deny & file_state.access))
return (NFS4ERR_DENIED)
This checking of share reservations on OPEN is done with no exception
for an existing OPEN for the same open_owner.
The constants used for the OPEN and OPEN_DOWNGRADE operations for the
access and deny fields are as follows:
const OPEN4_SHARE_ACCESS_READ = 0x00000001;
const OPEN4_SHARE_ACCESS_WRITE = 0x00000002;
const OPEN4_SHARE_ACCESS_BOTH = 0x00000003;
const OPEN4_SHARE_DENY_NONE = 0x00000000;
const OPEN4_SHARE_DENY_READ = 0x00000001;
const OPEN4_SHARE_DENY_WRITE = 0x00000002;
const OPEN4_SHARE_DENY_BOTH = 0x00000003;
6.10. OPEN/CLOSE Operations
To provide correct share semantics, a client MUST use the OPEN
operation to obtain the initial filehandle and indicate the desired
access and what if any access to deny. Even if the client intends to
use a stateid of all 0's or all 1's, it must still obtain the
filehandle for the regular file with the OPEN operation so the
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appropriate share semantics can be applied. For clients that do not
have a deny mode built into their open programming interfaces, deny
equal to NONE should be used.
The OPEN operation with the CREATE flag, also subsumes the CREATE
operation for regular files as used in previous versions of the NFS
protocol. This allows a create with a share to be done atomically.
The CLOSE operation removes all share reservations held by the
lock_owner on that file. If record locks are held, the client SHOULD
release all locks before issuing a CLOSE. The server MAY free all
outstanding locks on CLOSE but some servers may not support the CLOSE
of a file that still has record locks held. The server MUST return
failure, NFS4ERR_LOCKS_HELD, if any locks would exist after the
CLOSE.
The LOOKUP operation will return a filehandle without establishing
any lock state on the server. Without a valid stateid, the server
will assume the client has the least access. For example, a file
opened with deny READ/WRITE cannot be accessed using a filehandle
obtained through LOOKUP because it would not have a valid stateid
(i.e. using a stateid of all bits 0 or all bits 1).
6.10.1. Close and Retention of State Information
Since a CLOSE operation requests deallocation of a stateid, dealing
with retransmission of the CLOSE, may pose special difficulties,
since the state information, which normally would be used to
determine the state of the open file being designated, might be
deallocated, resulting in an NFS4ERR_BAD_STATEID error.
Servers may deal with this problem in a number of ways. To provide
the greatest degree assurance that the protocol is being used
properly, a server should, rather than deallocate the stateid, mark
it as close-pending, and retain the stateid with this status, until
later deallocation. In this way, a retransmitted CLOSE can be
recognized since the stateid points to state information with this
distinctive status, so that it can be handled without error.
When adopting this strategy, a server should retain the state
information until the earliest of:
o Another validly sequenced request for the same lockowner, that is
not a retransmission.
o The time that a lockowner is freed by the server due to period
with no activity.
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o All locks for the client are freed as a result of a SETCLIENTID.
Servers may avoid this complexity, at the cost of less complete
protocol error checking, by simply responding NFS4_OK in the event of
a CLOSE for a deallocated stateid, on the assumption that this case
must be caused by a retransmitted close. When adopting this
approach, it is desirable to at least log an error when returning a
no-error indication in this situation. If the server maintains a
reply-cache mechanism, it can verify the CLOSE is indeed a
retransmission and avoid error logging in most cases.
6.11. Open Upgrade and Downgrade
When an OPEN is done for a file and the lockowner for which the open
is being done already has the file open, the result is to upgrade the
open file status maintained on the server to include the access and
deny bits specified by the new OPEN as well as those for the existing
OPEN. The result is that there is one open file, as far as the
protocol is concerned, and it includes the union of the access and
deny bits for all of the OPEN requests completed. Only a single
CLOSE will be done to reset the effects of both OPENs. Note that the
client, when issuing the OPEN, may not know that the same file is in
fact being opened. The above only applies if both OPENs result in
the OPENed object being designated by the same filehandle.
When the server chooses to export multiple filehandles corresponding
to the same file object and returns different filehandles on two
different OPENs of the same file object, the server MUST NOT "OR"
together the access and deny bits and coalesce the two open files.
Instead the server must maintain separate OPENs with separate
stateids and will require separate CLOSEs to free them.
When multiple open files on the client are merged into a single open
file object on the server, the close of one of the open files (on the
client) may necessitate change of the access and deny status of the
open file on the server. This is because the union of the access and
deny bits for the remaining opens may be smaller (i.e. a proper
subset) than previously. The OPEN_DOWNGRADE operation is used to
make the necessary change and the client should use it to update the
server so that share reservation requests by other clients are
handled properly.
6.12. Short and Long Leases
When determining the time period for the server lease, the usual
lease tradeoffs apply. Short leases are good for fast server
recovery at a cost of increased RENEW or READ (with zero length)
requests. Longer leases are certainly kinder and gentler to servers
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trying to handle very large numbers of clients. The number of RENEW
requests drop in proportion to the lease time. The disadvantages of
long leases are slower recovery after server failure (the server must
wait for the leases to expire and the grace period to elapse before
granting new lock requests) and increased file contention (if client
fails to transmit an unlock request then server must wait for lease
expiration before granting new locks).
Long leases are usable if the server is able to store lease state in
non-volatile memory. Upon recovery, the server can reconstruct the
lease state from its non-volatile memory and continue operation with
its clients and therefore long leases would not be an issue.
6.13. Clocks, Propagation Delay, and Calculating Lease Expiration
To avoid the need for synchronized clocks, lease times are granted by
the server as a time delta. However, there is a requirement that the
client and server clocks do not drift excessively over the duration
of the lock. There is also the issue of propagation delay across the
network which could easily be several hundred milliseconds as well as
the possibility that requests will be lost and need to be
retransmitted.
To take propagation delay into account, the client should subtract it
from lease times (e.g. if the client estimates the one-way
propagation delay as 200 msec, then it can assume that the lease is
already 200 msec old when it gets it). In addition, it will take
another 200 msec to get a response back to the server. So the client
must send a lock renewal or write data back to the server 400 msec
before the lease would expire.
The server's lease period configuration should take into account the
network distance of the clients that will be accessing the server's
resources. It is expected that the lease period will take into
account the network propagation delays and other network delay
factors for the client population. Since the protocol does not allow
for an automatic method to determine an appropriate lease period, the
server's administrator may have to tune the lease period.
6.14. Migration, Replication and State
When responsibility for handling a given file system is transferred
to a new server (migration) or the client chooses to use an alternate
server (e.g. in response to server unresponsiveness) in the context
of file system replication, the appropriate handling of state shared
between the client and server (i.e. locks, leases, stateids, and
clientids) is as described below. The handling differs between
migration and replication. For related discussion of file server
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state and recover of such see the sections under "File Locking and
Share Reservations"
If server replica or a server immigrating a filesystem agrees to, or
is expected to, accept opaque values from the client that originated
from another server, then it is a wise implementation practice for
the servers to encode the "opaque" values in network byte order.
This way, servers acting as replicas or immigrating filesystems will
be able to parse values like stateids, directory cookies,
filehandles, etc. even if their native byte order is different from
other servers cooperating in the replication and migration of the
filesystem.
6.14.1. Migration and State
In the case of migration, the servers involved in the migration of a
filesystem SHOULD transfer all server state from the original to the
new server. This must be done in a way that is transparent to the
client. This state transfer will ease the client's transition when a
filesystem migration occurs. If the servers are successful in
transferring all state, the client will continue to use stateids
assigned by the original server. Therefore the new server must
recognize these stateids as valid. This holds true for the clientid
as well. Since responsibility for an entire filesystem is
transferred with a migration event, there is no possibility that
conflicts will arise on the new server as a result of the transfer of
locks.
As part of the transfer of information between servers, leases would
be transferred as well. The leases being transferred to the new
server will typically have a different expiration time from those for
the same client, previously on the old server. To maintain the
property that all leases on a given server for a given client expire
at the same time, the server should advance the expiration time to
the later of the leases being transferred or the leases already
present. This allows the client to maintain lease renewal of both
classes without special effort.
The servers may choose not to transfer the state information upon
migration. However, this choice is discouraged. In this case, when
the client presents state information from the original server, the
client must be prepared to receive either NFS4ERR_STALE_CLIENTID or
NFS4ERR_STALE_STATEID from the new server. The client should then
recover its state information as it normally would in response to a
server failure. The new server must take care to allow for the
recovery of state information as it would in the event of server
restart.
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6.14.2. Replication and State
Since client switch-over in the case of replication is not under
server control, the handling of state is different. In this case,
leases, stateids and clientids do not have validity across a
transition from one server to another. The client must re-establish
its locks on the new server. This can be compared to the re-
establishment of locks by means of reclaim-type requests after a
server reboot. The difference is that the server has no provision to
distinguish requests reclaiming locks from those obtaining new locks
or to defer the latter. Thus, a client re-establishing a lock on the
new server (by means of a LOCK or OPEN request), may have the
requests denied due to a conflicting lock. Since replication is
intended for read-only use of filesystems, such denial of locks
should not pose large difficulties in practice. When an attempt to
re-establish a lock on a new server is denied, the client should
treat the situation as if his original lock had been revoked.
6.14.3. Notification of Migrated Lease
In the case of lease renewal, the client may not be submitting
requests for a filesystem that has been migrated to another server.
This can occur because of the implicit lease renewal mechanism. The
client renews leases for all filesystems when submitting a request to
any one filesystem at the server.
In order for the client to schedule renewal of leases that may have
been relocated to the new server, the client must find out about
lease relocation before those leases expire. To accomplish this, all
operations which implicitly renew leases for a client (i.e. OPEN,
CLOSE, READ, WRITE, RENEW, LOCK, LOCKT, LOCKU), will return the error
NFS4ERR_LEASE_MOVED if responsibility for any of the leases to be
renewed has been transferred to a new server. This condition will
continue until the client receives an NFS4ERR_MOVED error and the
server receives the subsequent GETATTR(fs_locations) for an access to
each filesystem for which a lease has been moved to a new server.
When a client receives an NFS4ERR_LEASE_MOVED error, it should
perform an operation on each filesystem associated with the server in
question. When the client receives an NFS4ERR_MOVED error, the
client can follow the normal process to obtain the new server
information (through the fs_locations attribute) and perform renewal
of those leases on the new server. If the server has not had state
transferred to it transparently, the client will receive either
NFS4ERR_STALE_CLIENTID or NFS4ERR_STALE_STATEID from the new server,
as described above, and the client can then recover state information
as it does in the event of server failure.
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6.14.4. Migration and the Lease_time Attribute
In order that the client may appropriately manage its leases in the
case of migration, the destination server must establish proper
values for the lease_time attribute.
When state is transferred transparently, that state should include
the correct value of the lease_time attribute. The lease_time
attribute on the destination server must never be less than that on
the source since this would result in premature expiration of leases
granted by the source server. Upon migration in which state is
transferred transparently, the client is under no obligation to re-
fetch the lease_time attribute and may continue to use the value
previously fetched (on the source server).
If state has not been transferred transparently (i.e. the client sees
a real or simulated server reboot), the client should fetch the value
of lease_time on the new (i.e. destination) server, and use it for
subsequent locking requests. However the server must respect a grace
period at least as long as the lease_time on the source server, in
order to ensure that clients have ample time to reclaim their locks
before potentially conflicting non-reclaimed locks are granted. The
means by which the new server obtains the value of lease_time on the
old server is left to the server implementations. It is not
specified by the NFS version 4 protocol.
7. Client-Side Caching
Client-side caching of data, of file attributes, and of file names is
essential to providing good performance with the NFS protocol.
Providing distributed cache coherence is a difficult problem and
previous versions of the NFS protocol have not attempted it.
Instead, several NFS client implementation techniques have been used
to reduce the problems that a lack of coherence poses for users.
These techniques have not been clearly defined by earlier protocol
specifications and it is often unclear what is valid or invalid
client behavior.
The NFS version 4 protocol uses many techniques similar to those that
have been used in previous protocol versions. The NFS version 4
protocol does not provide distributed cache coherence. However, it
defines a more limited set of caching guarantees to allow locks and
share reservations to be used without destructive interference from
client side caching.
In addition, the NFS version 4 protocol introduces a delegation
mechanism which allows many decisions normally made by the server to
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be made locally by clients. This mechanism provides efficient
support of the common cases where sharing is infrequent or where
sharing is read-only.
7.1. Performance Challenges for Client-Side Caching
Caching techniques used in previous versions of the NFS protocol have
been successful in providing good performance. However, several
scalability challenges can arise when those techniques are used with
very large numbers of clients. This is particularly true when
clients are geographically distributed which classically increases
the latency for cache revalidation requests.
The previous versions of the NFS protocol repeat their file data
cache validation requests at the time the file is opened. This
behavior can have serious performance drawbacks. A common case is
one in which a file is only accessed by a single client. Therefore,
sharing is infrequent.
In this case, repeated reference to the server to find that no
conflicts exist is expensive. A better option with regards to
performance is to allow a client that repeatedly opens a file to do
so without reference to the server. This is done until potentially
conflicting operations from another client actually occur.
A similar situation arises in connection with file locking. Sending
file lock and unlock requests to the server as well as the read and
write requests necessary to make data caching consistent with the
locking semantics (see the section "Data Caching and File Locking")
can severely limit performance. When locking is used to provide
protection against infrequent conflicts, a large penalty is incurred.
This penalty may discourage the use of file locking by applications.
The NFS version 4 protocol provides more aggressive caching
strategies with the following design goals:
.IP o Compatibility with a large range of server semantics. .IP o
Provide the same caching benefits as previous versions of the NFS
protocol when unable to provide the more aggressive model. .IP o
Requirements for aggressive caching are organized so that a large
portion of the benefit can be obtained even when not all of the
requirements can be met. .LP The appropriate requirements for the
server are discussed in later sections in which specific forms of
caching are covered. (see the section "Open Delegation").
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7.2. Delegation and Callbacks
Recallable delegation of server responsibilities for a file to a
client improves performance by avoiding repeated requests to the
server in the absence of inter-client conflict. With the use of a
"callback" RPC from server to client, a server recalls delegated
responsibilities when another client engages in sharing of a
delegated file.
A delegation is passed from the server to the client, specifying the
object of the delegation and the type of delegation. There are
different types of delegations but each type contains a stateid to be
used to represent the delegation when performing operations that
depend on the delegation. This stateid is similar to those
associated with locks and share reservations but differs in that the
stateid for a delegation is associated with a clientid and may be
used on behalf of all the open_owners for the given client. A
delegation is made to the client as a whole and not to any specific
process or thread of control within it.
Because callback RPCs may not work in all environments (due to
firewalls, for example), correct protocol operation does not depend
on them. Preliminary testing of callback functionality by means of a
CB_NULL procedure determines whether callbacks can be supported. The
CB_NULL procedure checks the continuity of the callback path. A
server makes a preliminary assessment of callback availability to a
given client and avoids delegating responsibilities until it has
determined that callbacks are supported. Because the granting of a
delegation is always conditional upon the absence of conflicting
access, clients must not assume that a delegation will be granted and
they must always be prepared for OPENs to be processed without any
delegations being granted.
Once granted, a delegation behaves in most ways like a lock. There
is an associated lease that is subject to renewal together with all
of the other leases held by that client.
Unlike locks, an operation by a second client to a delegated file
will cause the server to recall a delegation through a callback.
On recall, the client holding the delegation must flush modified
state (such as modified data) to the server and return the
delegation. The conflicting request will not receive a response
until the recall is complete. The recall is considered complete when
the client returns the delegation or the server times out on the
recall and revokes the delegation as a result of the timeout.
Following the resolution of the recall, the server has the
information necessary to grant or deny the second client's request.
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At the time the client receives a delegation recall, it may have
substantial state that needs to be flushed to the server. Therefore,
the server should allow sufficient time for the delegation to be
returned since it may involve numerous RPCs to the server. If the
server is able to determine that the client is diligently flushing
state to the server as a result of the recall, the server may extend
the usual time allowed for a recall. However, the time allowed for
recall completion should not be unbounded.
An example of this is when responsibility to mediate opens on a given
file is delegated to a client (see the section "Open Delegation").
The server will not know what opens are in effect on the client.
Without this knowledge the server will be unable to determine if the
access and deny state for the file allows any particular open until
the delegation for the file has been returned.
A client failure or a network partition can result in failure to
respond to a recall callback. In this case, the server will revoke
the delegation which in turn will render useless any modified state
still on the client.
7.2.1. Delegation Recovery
There are three situations that delegation recovery must deal with:
o Client reboot or restart
o Server reboot or restart
o Network partition (full or callback-only)
In the event the client reboots or restarts, the failure to renew
leases will result in the revocation of record locks and share
reservations. Delegations, however, may be treated a bit
differently.
There will be situations in which delegations will need to be
reestablished after a client reboots or restarts. The reason for
this is the client may have file data stored locally and this data
was associated with the previously held delegations. The client will
need to reestablish the appropriate file state on the server.
To allow for this type of client recovery, the server MAY extend the
period for delegation recovery beyond the typical lease expiration
period. This implies that requests from other clients that conflict
with these delegations will need to wait. Because the normal recall
process may require significant time for the client to flush changed
state to the server, other clients need be prepared for delays that
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occur because of a conflicting delegation. This longer interval
would increase the window for clients to reboot and consult stable
storage so that the delegations can be reclaimed. For open
delegations, such delegations are reclaimed using OPEN with a claim
type of CLAIM_DELEGATE_PREV. (See the sections on "Data Caching and
Revocation" and "Operation 18: OPEN" for discussion of open
delegation and the details of OPEN respectively).
A server MAY support a claim type of CLAIM_DELEGATE_PREV, but if it
does, it MUST NOT remove delegations upon SETCLIENTID_CONFIRM, and
instead MUST, for a period of time no less than that of the value of
the lease_time attribute, maintain the client's delegations to allow
time for the client to issue CLAIM_DELEGATE_PREV requests. The
server that supports CLAIM_DELEGATE_PREV MUST support the DELEGPURGE
operation.
When the server reboots or restarts, delegations are reclaimed (using
the OPEN operation with CLAIM_PREVIOUS) in a similar fashion to
record locks and share reservations. However, there is a slight
semantic difference. In the normal case if the server decides that a
delegation should not be granted, it performs the requested action
(e.g. OPEN) without granting any delegation. For reclaim, the
server grants the delegation but a special designation is applied so
that the client treats the delegation as having been granted but
recalled by the server. Because of this, the client has the duty to
write all modified state to the server and then return the
delegation. This process of handling delegation reclaim reconciles
three principles of the NFS version 4 protocol:
o Upon reclaim, a client reporting resources assigned to it by an
earlier server instance must be granted those resources.
o The server has unquestionable authority to determine whether
delegations are to be granted and, once granted, whether they are
to be continued.
o The use of callbacks is not to be depended upon until the client
has proven its ability to receive them.
When a network partition occurs, delegations are subject to freeing
by the server when the lease renewal period expires. This is similar
to the behavior for locks and share reservations. For delegations,
however, the server may extend the period in which conflicting
requests are held off. Eventually the occurrence of a conflicting
request from another client will cause revocation of the delegation.
A loss of the callback path (e.g. by later network configuration
change) will have the same effect. A recall request will fail and
revocation of the delegation will result.
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A client normally finds out about revocation of a delegation when it
uses a stateid associated with a delegation and receives the error
NFS4ERR_EXPIRED. It also may find out about delegation revocation
after a client reboot when it attempts to reclaim a delegation and
receives that same error. Note that in the case of a revoked write
open delegation, there are issues because data may have been modified
by the client whose delegation is revoked and separately by other
clients. See the section "Revocation Recovery for Write Open
Delegation" for a discussion of such issues. Note also that when
delegations are revoked, information about the revoked delegation
will be written by the server to stable storage (as described in the
section "Crash Recovery"). This is done to deal with the case in
which a server reboots after revoking a delegation but before the
client holding the revoked delegation is notified about the
revocation.
7.3. Data Caching
When applications share access to a set of files, they need to be
implemented so as to take account of the possibility of conflicting
access by another application. This is true whether the applications
in question execute on different clients or reside on the same
client.
Share reservations and record locks are the facilities the NFS
version 4 protocol provides to allow applications to coordinate
access by providing mutual exclusion facilities. The NFS version 4
protocol's data caching must be implemented such that it does not
invalidate the assumptions that those using these facilities depend
upon.
7.3.1. Data Caching and OPENs
In order to avoid invalidating the sharing assumptions that
applications rely on, NFS version 4 clients should not provide cached
data to applications or modify it on behalf of an application when it
would not be valid to obtain or modify that same data via a READ or
WRITE operation.
Furthermore, in the absence of open delegation (see the section "Open
Delegation") two additional rules apply. Note that these rules are
obeyed in practice by many NFS version 2 and version 3 clients.
o First, cached data present on a client must be revalidated after
doing an OPEN. Revalidating means that the client fetches the
change attribute from the server, compares it with the cached
change attribute, and if different, declares the cached data (as
well as the cached attributes) as invalid. This is to ensure that
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the data for the OPENed file is still correctly reflected in the
client's cache. This validation must be done at least when the
client's OPEN operation includes DENY=WRITE or BOTH thus
terminating a period in which other clients may have had the
opportunity to open the file with WRITE access. Clients may
choose to do the revalidation more often (i.e. at OPENs specifying
DENY=NONE) to parallel the NFS version 3 protocol's practice for
the benefit of users assuming this degree of cache revalidation.
Since the change attribute is updated for data and metadata
modifications, some client implementors may be tempted to use the
time_modify attribute and not change to validate cached data, so
that metadata changes do not spuriously invalidate clean data.
The implementor is cautioned in this approach. The change
attribute is guaranteed to change for each update to the file,
whereas time_modify is guaranteed to change only at the
granularity of the time_delta attribute. Use by the client's data
cache validation logic of time_modify and not change runs the risk
of the client incorrectly marking stale data as valid.
o Second, modified data must be flushed to the server before closing
a file OPENed for write. This is complementary to the first rule.
If the data is not flushed at CLOSE, the revalidation done after
client OPENs as file is unable to achieve its purpose. The other
aspect to flushing the data before close is that the data must be
committed to stable storage, at the server, before the CLOSE
operation is requested by the client. In the case of a server
reboot or restart and a CLOSEd file, it may not be possible to
retransmit the data to be written to the file. Hence, this
requirement.
7.3.2. Data Caching and File Locking
For those applications that choose to use file locking instead of
share reservations to exclude inconsistent file access, there is an
analogous set of constraints that apply to client side data caching.
These rules are effective only if the file locking is used in a way
that matches in an equivalent way the actual READ and WRITE
operations executed. This is as opposed to file locking that is
based on pure convention. For example, it is possible to manipulate
a two-megabyte file by dividing the file into two one-megabyte
regions and protecting access to the two regions by file locks on
bytes zero and one. A lock for write on byte zero of the file would
represent the right to do READ and WRITE operations on the first
region. A lock for write on byte one of the file would represent the
right to do READ and WRITE operations on the second region. As long
as all applications manipulating the file obey this convention, they
will work on a local filesystem. However, they may not work with the
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NFS version 4 protocol unless clients refrain from data caching.
The rules for data caching in the file locking environment are:
o First, when a client obtains a file lock for a particular region,
the data cache corresponding to that region (if any cache data
exists) must be revalidated. If the change attribute indicates
that the file may have been updated since the cached data was
obtained, the client must flush or invalidate the cached data for
the newly locked region. A client might choose to invalidate all
of non-modified cached data that it has for the file but the only
requirement for correct operation is to invalidate all of the data
in the newly locked region.
o Second, before releasing a write lock for a region, all modified
data for that region must be flushed to the server. The modified
data must also be written to stable storage.
Note that flushing data to the server and the invalidation of cached
data must reflect the actual byte ranges locked or unlocked.
Rounding these up or down to reflect client cache block boundaries
will cause problems if not carefully done. For example, writing a
modified block when only half of that block is within an area being
unlocked may cause invalid modification to the region outside the
unlocked area. This, in turn, may be part of a region locked by
another client. Clients can avoid this situation by synchronously
performing portions of write operations that overlap that portion
(initial or final) that is not a full block. Similarly, invalidating
a locked area which is not an integral number of full buffer blocks
would require the client to read one or two partial blocks from the
server if the revalidation procedure shows that the data which the
client possesses may not be valid.
The data that is written to the server as a prerequisite to the
unlocking of a region must be written, at the server, to stable
storage. The client may accomplish this either with synchronous
writes or by following asynchronous writes with a COMMIT operation.
This is required because retransmission of the modified data after a
server reboot might conflict with a lock held by another client.
A client implementation may choose to accommodate applications which
use record locking in non-standard ways (e.g. using a record lock as
a global semaphore) by flushing to the server more data upon an LOCKU
than is covered by the locked range. This may include modified data
within files other than the one for which the unlocks are being done.
In such cases, the client must not interfere with applications whose
READs and WRITEs are being done only within the bounds of record
locks which the application holds. For example, an application locks
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a single byte of a file and proceeds to write that single byte. A
client that chose to handle a LOCKU by flushing all modified data to
the server could validly write that single byte in response to an
unrelated unlock. However, it would not be valid to write the entire
block in which that single written byte was located since it includes
an area that is not locked and might be locked by another client.
Client implementations can avoid this problem by dividing files with
modified data into those for which all modifications are done to
areas covered by an appropriate record lock and those for which there
are modifications not covered by a record lock. Any writes done for
the former class of files must not include areas not locked and thus
not modified on the client.
7.3.3. Data Caching and Mandatory File Locking
Client side data caching needs to respect mandatory file locking when
it is in effect. The presence of mandatory file locking for a given
file is indicated when the client gets back NFS4ERR_LOCKED from a
READ or WRITE on a file it has an appropriate share reservation for.
When mandatory locking is in effect for a file, the client must check
for an appropriate file lock for data being read or written. If a
lock exists for the range being read or written, the client may
satisfy the request using the client's validated cache. If an
appropriate file lock is not held for the range of the read or write,
the read or write request must not be satisfied by the client's cache
and the request must be sent to the server for processing. When a
read or write request partially overlaps a locked region, the request
should be subdivided into multiple pieces with each region (locked or
not) treated appropriately.
7.3.4. Data Caching and File Identity
When clients cache data, the file data needs to be organized
according to the filesystem object to which the data belongs. For
NFS version 3 clients, the typical practice has been to assume for
the purpose of caching that distinct filehandles represent distinct
filesystem objects. The client then has the choice to organize and
maintain the data cache on this basis.
In the NFS version 4 protocol, there is now the possibility to have
significant deviations from a "one filehandle per object" model
because a filehandle may be constructed on the basis of the object's
pathname. Therefore, clients need a reliable method to determine if
two filehandles designate the same filesystem object. If clients
were simply to assume that all distinct filehandles denote distinct
objects and proceed to do data caching on this basis, caching
inconsistencies would arise between the distinct client side objects
which mapped to the same server side object.
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By providing a method to differentiate filehandles, the NFS version 4
protocol alleviates a potential functional regression in comparison
with the NFS version 3 protocol. Without this method, caching
inconsistencies within the same client could occur and this has not
been present in previous versions of the NFS protocol. Note that it
is possible to have such inconsistencies with applications executing
on multiple clients but that is not the issue being addressed here.
For the purposes of data caching, the following steps allow an NFS
version 4 client to determine whether two distinct filehandles denote
the same server side object:
o If GETATTR directed to two filehandles returns different values of
the fsid attribute, then the filehandles represent distinct
objects.
o If GETATTR for any file with an fsid that matches the fsid of the
two filehandles in question returns a unique_handles attribute
with a value of TRUE, then the two objects are distinct.
o If GETATTR directed to the two filehandles does not return the
fileid attribute for both of the handles, then it cannot be
determined whether the two objects are the same. Therefore,
operations which depend on that knowledge (e.g. client side data
caching) cannot be done reliably.
o If GETATTR directed to the two filehandles returns different
values for the fileid attribute, then they are distinct objects.
o Otherwise they are the same object.
7.4. Open Delegation
When a file is being OPENed, the server may delegate further handling
of opens and closes for that file to the opening client. Any such
delegation is recallable, since the circumstances that allowed for
the delegation are subject to change. In particular, the server may
receive a conflicting OPEN from another client, the server must
recall the delegation before deciding whether the OPEN from the other
client may be granted. Making a delegation is up to the server and
clients should not assume that any particular OPEN either will or
will not result in an open delegation. The following is a typical
set of conditions that servers might use in deciding whether OPEN
should be delegated:
o The client must be able to respond to the server's callback
requests. The server will use the CB_NULL procedure for a test of
callback ability.
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o The client must have responded properly to previous recalls.
o There must be no current open conflicting with the requested
delegation.
o There should be no current delegation that conflicts with the
delegation being requested.
o The probability of future conflicting open requests should be low
based on the recent history of the file.
o The existence of any server-specific semantics of OPEN/CLOSE that
would make the required handling incompatible with the prescribed
handling that the delegated client would apply (see below).
There are two types of open delegations, read and write. A read open
delegation allows a client to handle, on its own, requests to open a
file for reading that do not deny read access to others. Multiple
read open delegations may be outstanding simultaneously and do not
conflict. A write open delegation allows the client to handle, on
its own, all opens. Only one write open delegation may exist for a
given file at a given time and it is inconsistent with any read open
delegations.
When a client has a read open delegation, it may not make any changes
to the contents or attributes of the file but it is assured that no
other client may do so. When a client has a write open delegation,
it may modify the file data since no other client will be accessing
the file's data. The client holding a write delegation may only
affect file attributes which are intimately connected with the file
data: size, time_modify, change.
When a client has an open delegation, it does not send OPENs or
CLOSEs to the server but updates the appropriate status internally.
For a read open delegation, opens that cannot be handled locally
(opens for write or that deny read access) must be sent to the
server.
When an open delegation is made, the response to the OPEN contains an
open delegation structure which specifies the following:
o the type of delegation (read or write)
o space limitation information to control flushing of data on close
(write open delegation only, see the section "Open Delegation and
Data Caching")
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o an nfsace4 specifying read and write permissions
o a stateid to represent the delegation for READ and WRITE
The delegation stateid is separate and distinct from the stateid for
the OPEN proper. The standard stateid, unlike the delegation
stateid, is associated with a particular lock_owner and will continue
to be valid after the delegation is recalled and the file remains
open.
When a request internal to the client is made to open a file and open
delegation is in effect, it will be accepted or rejected solely on
the basis of the following conditions. Any requirement for other
checks to be made by the delegate should result in open delegation
being denied so that the checks can be made by the server itself.
o The access and deny bits for the request and the file as described
in the section "Share Reservations".
o The read and write permissions as determined below.
The nfsace4 passed with delegation can be used to avoid frequent
ACCESS calls. The permission check should be as follows:
o If the nfsace4 indicates that the open may be done, then it should
be granted without reference to the server.
o If the nfsace4 indicates that the open may not be done, then an
ACCESS request must be sent to the server to obtain the definitive
answer.
The server may return an nfsace4 that is more restrictive than the
actual ACL of the file. This includes an nfsace4 that specifies
denial of all access. Note that some common practices such as
mapping the traditional user "root" to the user "nobody" may make it
incorrect to return the actual ACL of the file in the delegation
response.
The use of delegation together with various other forms of caching
creates the possibility that no server authentication will ever be
performed for a given user since all of the user's requests might be
satisfied locally. Where the client is depending on the server for
authentication, the client should be sure authentication occurs for
each user by use of the ACCESS operation. This should be the case
even if an ACCESS operation would not be required otherwise. As
mentioned before, the server may enforce frequent authentication by
returning an nfsace4 denying all access with every open delegation.
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7.4.1. Open Delegation and Data Caching
OPEN delegation allows much of the message overhead associated with
the opening and closing files to be eliminated. An open when an open
delegation is in effect does not require that a validation message be
sent to the server. The continued endurance of the "read open
delegation" provides a guarantee that no OPEN for write and thus no
write has occurred. Similarly, when closing a file opened for write
and if write open delegation is in effect, the data written does not
have to be flushed to the server until the open delegation is
recalled. The continued endurance of the open delegation provides a
guarantee that no open and thus no read or write has been done by
another client.
For the purposes of open delegation, READs and WRITEs done without an
OPEN are treated as the functional equivalents of a corresponding
type of OPEN. This refers to the READs and WRITEs that use the
special stateids consisting of all zero bits or all one bits.
Therefore, READs or WRITEs with a special stateid done by another
client will force the server to recall a write open delegation. A
WRITE with a special stateid done by another client will force a
recall of read open delegations.
With delegations, a client is able to avoid writing data to the
server when the CLOSE of a file is serviced. The file close system
call is the usual point at which the client is notified of a lack of
stable storage for the modified file data generated by the
application. At the close, file data is written to the server and
through normal accounting the server is able to determine if the
available filesystem space for the data has been exceeded (i.e.
server returns NFS4ERR_NOSPC or NFS4ERR_DQUOT). This accounting
includes quotas. The introduction of delegations requires that a
alternative method be in place for the same type of communication to
occur between client and server.
In the delegation response, the server provides either the limit of
the size of the file or the number of modified blocks and associated
block size. The server must ensure that the client will be able to
flush data to the server of a size equal to that provided in the
original delegation. The server must make this assurance for all
outstanding delegations. Therefore, the server must be careful in
its management of available space for new or modified data taking
into account available filesystem space and any applicable quotas.
The server can recall delegations as a result of managing the
available filesystem space. The client should abide by the server's
state space limits for delegations. If the client exceeds the stated
limits for the delegation, the server's behavior is undefined.
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Based on server conditions, quotas or available filesystem space, the
server may grant write open delegations with very restrictive space
limitations. The limitations may be defined in a way that will
always force modified data to be flushed to the server on close.
With respect to authentication, flushing modified data to the server
after a CLOSE has occurred may be problematic. For example, the user
of the application may have logged off the client and unexpired
authentication credentials may not be present. In this case, the
client may need to take special care to ensure that local unexpired
credentials will in fact be available. This may be accomplished by
tracking the expiration time of credentials and flushing data well in
advance of their expiration or by making private copies of
credentials to assure their availability when needed.
7.4.2. Open Delegation and File Locks
When a client holds a write open delegation, lock operations are
performed locally. This includes those required for mandatory file
locking. This can be done since the delegation implies that there
can be no conflicting locks. Similarly, all of the revalidations
that would normally be associated with obtaining locks and the
flushing of data associated with the releasing of locks need not be
done.
When a client holds a read open delegation, lock operations are not
performed locally. All lock operations, including those requesting
non-exclusive locks, are sent to the server for resolution.
7.4.3. Handling of CB_GETATTR
The server needs to employ special handling for a GETATTR where the
target is a file that has a write open delegation in effect. The
reason for this is that the client holding the write delegation may
have modified the data and the server needs to reflect this change to
the second client that submitted the GETATTR. Therefore, the client
holding the write delegation needs to be interrogated. The server
will use the CB_GETATTR operation. The only attributes that the
server can reliably query via CB_GETATTR are size and change.
Since CB_GETATTR is being used to satisfy another client's GETATTR
request, the server only needs to know if the client holding the
delegation has a modified version of the file. If the client's copy
of the delegated file is not modified (data or size), the server can
satisfy the second client's GETATTR request from the attributes
stored locally at the server. If the file is modified, the server
only needs to know about this modified state. If the server
determines that the file is currently modified, it will respond to
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the second client's GETATTR as if the file had been modified locally
at the server.
Since the form of the change attribute is determined by the server
and is opaque to the client, the client and server need to agree on a
method of communicating the modified state of the file. For the size
attribute, the client will report its current view of the file size.
For the change attribute, the handling is more involved.
For the client, the following steps will be taken when receiving a
write delegation:
o The value of the change attribute will be obtained from the server
and cached. Let this value be represented by c.
o The client will create a value greater than c that will be used
for communicating modified data is held at the client. Let this
value be represented by d.
o When the client is queried via CB_GETATTR for the change
attribute, it checks to see if it holds modified data. If the
file is modified, the value d is returned for the change attribute
value. If this file is not currently modified, the client returns
the value c for the change attribute.
For simplicity of implementation, the client MAY for each CB_GETATTR
return the same value d. This is true even if, between successive
CB_GETATTR operations, the client again modifies in the file's data
or metadata in its cache. The client can return the same value
because the only requirement is that the client be able to indicate
to the server that the client holds modified data. Therefore, the
value of d may always be c + 1.
While the change attribute is opaque to the client in the sense that
it has no idea what units of time, if any, the server is counting
change with, it is not opaque in that the client has to treat it as
an unsigned integer, and the server has to be able to see the results
of the client's changes to that integer. Therefore, the server MUST
encode the change attribute in network order when sending it to the
client. The client MUST decode it from network order to its native
order when receiving it and the client MUST encode it network order
when sending it to the server. For this reason, change is defined as
an unsigned integer rather than an opaque array of octets.
For the server, the following steps will be taken when providing a
write delegation:
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o Upon providing a write delegation, the server will cache a copy of
the change attribute in the data structure it uses to record the
delegation. Let this value be represented by sc.
o When a second client sends a GETATTR operation on the same file to
the server, the server obtains the change attribute from the first
client. Let this value be cc.
o If the value cc is equal to sc, the file is not modified and the
server returns the current values for change, time_metadata, and
time_modify (for example) to the second client.
o If the value cc is NOT equal to sc, the file is currently modified
at the first client and most likely will be modified at the server
at a future time. The server then uses its current time to
construct attribute values for time_metadata and time_modify. A
new value of sc, which we will call nsc, is computed by the
server, such that nsc >= sc + 1. The server then returns the
constructed time_metadata, time_modify, and nsc values to the
requester. The server replaces sc in the delegation record with
nsc. To prevent the possibility of time_modify, time_metadata,
and change from appearing to go backward (which would happen if
the client holding the delegation fails to write its modified data
to the server before the delegation is revoked or returned), the
server SHOULD update the file's metadata record with the
constructed attribute values. For reasons of reasonable
performance, committing the constructed attribute values to stable
storage is OPTIONAL.
As discussed earlier in this section, the client MAY return the same
cc value on subsequent CB_GETATTR calls, even if the file was
modified in the client's cache yet again between successive
CB_GETATTR calls. Therefore, the server must assume that the file
has been modified yet again, and MUST take care to ensure that the
new nsc it constructs and returns is greater than the previous nsc it
returned. An example implementation's delegation record would
satisfy this mandate by including a boolean field (let us call it
"modified") that is set to false when the delegation is granted, and
an sc value set at the time of grant to the change attribute value.
The modified field would be set to true the first time cc != sc, and
would stay true until the delegation is returned or revoked. The
processing for constructing nsc, time_modify, and time_metadata would
use this pseudo code:
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if (!modified) {
do CB_GETATTR for change and size;
if (cc != sc)
modified = TRUE;
} else {
do CB_GETATTR for size;
}
if (modified) {
sc = sc + 1;
time_modify = time_metadata = current_time;
update sc, time_modify, time_metadata into file's metadata;
}
return to client (that sent GETATTR) the attributes
it requested, but make sure size comes from what
CB_GETATTR returned. Do not update the file's metadata
with the client's modified size.
In the case that the file attribute size is different than the
server's current value, the server treats this as a modification
regardless of the value of the change attribute retrieved via
CB_GETATTR and responds to the second client as in the last step.
This methodology resolves issues of clock differences between client
and server and other scenarios where the use of CB_GETATTR break
down.
It should be noted that the server is under no obligation to use
CB_GETATTR and therefore the server MAY simply recall the delegation
to avoid its use.
7.4.4. Recall of Open Delegation
The following events necessitate recall of an open delegation:
o Potentially conflicting OPEN request (or READ/WRITE done with
"special" stateid)
o SETATTR issued by another client
o REMOVE request for the file
o RENAME request for the file as either source or target of the
RENAME
Whether a RENAME of a directory in the path leading to the file
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results in recall of an open delegation depends on the semantics of
the server filesystem. If that filesystem denies such RENAMEs when a
file is open, the recall must be performed to determine whether the
file in question is, in fact, open.
In addition to the situations above, the server may choose to recall
open delegations at any time if resource constraints make it
advisable to do so. Clients should always be prepared for the
possibility of recall.
When a client receives a recall for an open delegation, it needs to
update state on the server before returning the delegation. These
same updates must be done whenever a client chooses to return a
delegation voluntarily. The following items of state need to be
dealt with:
o If the file associated with the delegation is no longer open and
no previous CLOSE operation has been sent to the server, a CLOSE
operation must be sent to the server.
o If a file has other open references at the client, then OPEN
operations must be sent to the server. The appropriate stateids
will be provided by the server for subsequent use by the client
since the delegation stateid will not longer be valid. These OPEN
requests are done with the claim type of CLAIM_DELEGATE_CUR. This
will allow the presentation of the delegation stateid so that the
client can establish the appropriate rights to perform the OPEN.
(see the section "Operation 18: OPEN" for details.)
o If there are granted file locks, the corresponding LOCK operations
need to be performed. This applies to the write open delegation
case only.
o For a write open delegation, if at the time of recall the file is
not open for write, all modified data for the file must be flushed
to the server. If the delegation had not existed, the client
would have done this data flush before the CLOSE operation.
o For a write open delegation when a file is still open at the time
of recall, any modified data for the file needs to be flushed to
the server.
o With the write open delegation in place, it is possible that the
file was truncated during the duration of the delegation. For
example, the truncation could have occurred as a result of an OPEN
UNCHECKED with a size attribute value of zero. Therefore, if a
truncation of the file has occurred and this operation has not
been propagated to the server, the truncation must occur before
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any modified data is written to the server.
In the case of write open delegation, file locking imposes some
additional requirements. To precisely maintain the associated
invariant, it is required to flush any modified data in any region
for which a write lock was released while the write delegation was in
effect. However, because the write open delegation implies no other
locking by other clients, a simpler implementation is to flush all
modified data for the file (as described just above) if any write
lock has been released while the write open delegation was in effect.
An implementation need not wait until delegation recall (or deciding
to voluntarily return a delegation) to perform any of the above
actions, if implementation considerations (e.g. resource availability
constraints) make that desirable. Generally, however, the fact that
the actual open state of the file may continue to change makes it not
worthwhile to send information about opens and closes to the server,
except as part of delegation return. Only in the case of closing the
open that resulted in obtaining the delegation would clients be
likely to do this early, since, in that case, the close once done
will not be undone. Regardless of the client's choices on scheduling
these actions, all must be performed before the delegation is
returned, including (when applicable) the close that corresponds to
the open that resulted in the delegation. These actions can be
performed either in previous requests or in previous operations in
the same COMPOUND request.
7.4.5. Clients that Fail to Honor Delegation Recalls
A client may fail to respond to a recall for various reasons, such as
a failure of the callback path from server to the client. The client
may be unaware of a failure in the callback path. This lack of
awareness could result in the client finding out long after the
failure that its delegation has been revoked, and another client has
modified the data for which the client had a delegation. This is
especially a problem for the client that held a write delegation.
The server also has a dilemma in that the client that fails to
respond to the recall might also be sending other NFS requests,
including those that renew the lease before the lease expires.
Without returning an error for those lease renewing operations, the
server leads the client to believe that the delegation it has is in
force.
This difficulty is solved by the following rules:
o When the callback path is down, the server MUST NOT revoke the
delegation if one of the following occurs:
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* The client has issued a RENEW operation and the server has
returned an NFS4ERR_CB_PATH_DOWN error. The server MUST renew
the lease for any record locks and share reservations the
client has that the server has known about (as opposed to those
locks and share reservations the client has established but not
yet sent to the server, due to the delegation). The server
SHOULD give the client a reasonable time to return its
delegations to the server before revoking the client's
delegations.
* The client has not issued a RENEW operation for some period of
time after the server attempted to recall the delegation. This
period of time MUST NOT be less than the value of the
lease_time attribute.
o When the client holds a delegation, it can not rely on operations,
except for RENEW, that take a stateid, to renew delegation leases
across callback path failures. The client that wants to keep
delegations in force across callback path failures must use RENEW
to do so.
7.4.6. Delegation Revocation
At the point a delegation is revoked, if there are associated opens
on the client, the applications holding these opens need to be
notified. This notification usually occurs by returning errors for
READ/WRITE operations or when a close is attempted for the open file.
If no opens exist for the file at the point the delegation is
revoked, then notification of the revocation is unnecessary.
However, if there is modified data present at the client for the
file, the user of the application should be notified. Unfortunately,
it may not be possible to notify the user since active applications
may not be present at the client. See the section "Revocation
Recovery for Write Open Delegation" for additional details.
7.5. Data Caching and Revocation
When locks and delegations are revoked, the assumptions upon which
successful caching depend are no longer guaranteed. For any locks or
share reservations that have been revoked, the corresponding owner
needs to be notified. This notification includes applications with a
file open that has a corresponding delegation which has been revoked.
Cached data associated with the revocation must be removed from the
client. In the case of modified data existing in the client's cache,
that data must be removed from the client without it being written to
the server. As mentioned, the assumptions made by the client are no
longer valid at the point when a lock or delegation has been revoked.
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For example, another client may have been granted a conflicting lock
after the revocation of the lock at the first client. Therefore, the
data within the lock range may have been modified by the other
client. Obviously, the first client is unable to guarantee to the
application what has occurred to the file in the case of revocation.
Notification to a lock owner will in many cases consist of simply
returning an error on the next and all subsequent READs/WRITEs to the
open file or on the close. Where the methods available to a client
make such notification impossible because errors for certain
operations may not be returned, more drastic action such as signals
or process termination may be appropriate. The justification for
this is that an invariant for which an application depends on may be
violated. Depending on how errors are typically treated for the
client operating environment, further levels of notification
including logging, console messages, and GUI pop-ups may be
appropriate.
7.5.1. Revocation Recovery for Write Open Delegation
Revocation recovery for a write open delegation poses the special
issue of modified data in the client cache while the file is not
open. In this situation, any client which does not flush modified
data to the server on each close must ensure that the user receives
appropriate notification of the failure as a result of the
revocation. Since such situations may require human action to
correct problems, notification schemes in which the appropriate user
or administrator is notified may be necessary. Logging and console
messages are typical examples.
If there is modified data on the client, it must not be flushed
normally to the server. A client may attempt to provide a copy of
the file data as modified during the delegation under a different
name in the filesystem name space to ease recovery. Note that when
the client can determine that the file has not been modified by any
other client, or when the client has a complete cached copy of file
in question, such a saved copy of the client's view of the file may
be of particular value for recovery. In other case, recovery using a
copy of the file based partially on the client's cached data and
partially on the server copy as modified by other clients, will be
anything but straightforward, so clients may avoid saving file
contents in these situations or mark the results specially to warn
users of possible problems.
Saving of such modified data in delegation revocation situations may
be limited to files of a certain size or might be used only when
sufficient disk space is available within the target filesystem.
Such saving may also be restricted to situations when the client has
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sufficient buffering resources to keep the cached copy available
until it is properly stored to the target filesystem.
7.6. Attribute Caching
The attributes discussed in this section do not include named
attributes. Individual named attributes are analogous to files and
caching of the data for these needs to be handled just as data
caching is for ordinary files. Similarly, LOOKUP results from an
OPENATTR directory are to be cached on the same basis as any other
pathnames and similarly for directory contents.
Clients may cache file attributes obtained from the server and use
them to avoid subsequent GETATTR requests. Such caching is write
through in that modification to file attributes is always done by
means of requests to the server and should not be done locally and
cached. The exception to this are modifications to attributes that
are intimately connected with data caching. Therefore, extending a
file by writing data to the local data cache is reflected immediately
in the size as seen on the client without this change being
immediately reflected on the server. Normally such changes are not
propagated directly to the server but when the modified data is
flushed to the server, analogous attribute changes are made on the
server. When open delegation is in effect, the modified attributes
may be returned to the server in the response to a CB_RECALL call.
The result of local caching of attributes is that the attribute
caches maintained on individual clients will not be coherent.
Changes made in one order on the server may be seen in a different
order on one client and in a third order on a different client.
The typical filesystem application programming interfaces do not
provide means to atomically modify or interrogate attributes for
multiple files at the same time. The following rules provide an
environment where the potential incoherences mentioned above can be
reasonably managed. These rules are derived from the practice of
previous NFS protocols.
o All attributes for a given file (per-fsid attributes excepted) are
cached as a unit at the client so that no non-serializability can
arise within the context of a single file.
o An upper time boundary is maintained on how long a client cache
entry can be kept without being refreshed from the server.
o When operations are performed that change attributes at the
server, the updated attribute set is requested as part of the
containing RPC. This includes directory operations that update
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attributes indirectly. This is accomplished by following the
modifying operation with a GETATTR operation and then using the
results of the GETATTR to update the client's cached attributes.
Note that if the full set of attributes to be cached is requested by
READDIR, the results can be cached by the client on the same basis as
attributes obtained via GETATTR.
A client may validate its cached version of attributes for a file by
fetching just both the change and time_access attributes and assuming
that if the change attribute has the same value as it did when the
attributes were cached, then no attributes other than time_access
have changed. The reason why time_access is also fetched is because
many servers operate in environments where the operation that updates
change does not update time_access. For example, POSIX file
semantics do not update access time when a file is modified by the
write system call. Therefore, the client that wants a current
time_access value should fetch it with change during the attribute
cache validation processing and update its cached time_access.
The client may maintain a cache of modified attributes for those
attributes intimately connected with data of modified regular files
(size, time_modify, and change). Other than those three attributes,
the client MUST NOT maintain a cache of modified attributes.
Instead, attribute changes are immediately sent to the server.
In some operating environments, the equivalent to time_access is
expected to be implicitly updated by each read of the content of the
file object. If an NFS client is caching the content of a file
object, whether it is a regular file, directory, or symbolic link,
the client SHOULD NOT update the time_access attribute (via SETATTR
or a small READ or READDIR request) on the server with each read that
is satisfied from cache. The reason is that this can defeat the
performance benefits of caching content, especially since an explicit
SETATTR of time_access may alter the change attribute on the server.
If the change attribute changes, clients that are caching the content
will think the content has changed, and will re-read unmodified data
from the server. Nor is the client encouraged to maintain a modified
version of time_access in its cache, since this would mean that the
client will either eventually have to write the access time to the
server with bad performance effects, or it would never update the
server's time_access, thereby resulting in a situation where an
application that caches access time between a close and open of the
same file observes the access time oscillating between the past and
present. The time_access attribute always means the time of last
access to a file by a read that was satisfied by the server. This
way clients will tend to see only time_access changes that go forward
in time.
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7.7. Data and Metadata Caching and Memory Mapped Files
Some operating environments include the capability for an application
to map a file's content into the application's address space. Each
time the application accesses a memory location that corresponds to a
block that has not been loaded into the address space, a page fault
occurs and the file is read (or if the block does not exist in the
file, the block is allocated and then instantiated in the
application's address space).
As long as each memory mapped access to the file requires a page
fault, the relevant attributes of the file that are used to detect
access and modification (time_access, time_metadata, time_modify, and
change) will be updated. However, in many operating environments,
when page faults are not required these attributes will not be
updated on reads or updates to the file via memory access (regardless
whether the file is local file or is being access remotely). A
client or server MAY fail to update attributes of a file that is
being accessed via memory mapped I/O. This has several implications:
o If there is an application on the server that has memory mapped a
file that a client is also accessing, the client may not be able
to get a consistent value of the change attribute to determine
whether its cache is stale or not. A server that knows that the
file is memory mapped could always pessimistically return updated
values for change so as to force the application to always get the
most up to date data and metadata for the file. However, due to
the negative performance implications of this, such behavior is
OPTIONAL.
o If the memory mapped file is not being modified on the server, and
instead is just being read by an application via the memory mapped
interface, the client will not see an updated time_access
attribute. However, in many operating environments, neither will
any process running on the server. Thus NFS clients are at no
disadvantage with respect to local processes.
o If there is another client that is memory mapping the file, and if
that client is holding a write delegation, the same set of issues
as discussed in the previous two bullet items apply. So, when a
server does a CB_GETATTR to a file that the client has modified in
its cache, the response from CB_GETATTR will not necessarily be
accurate. As discussed earlier, the client's obligation is to
report that the file has been modified since the delegation was
granted, not whether it has been modified again between successive
CB_GETATTR calls, and the server MUST assume that any file the
client has modified in cache has been modified again between
successive CB_GETATTR calls. Depending on the nature of the
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client's memory management system, this weak obligation may not be
possible. A client MAY return stale information in CB_GETATTR
whenever the file is memory mapped.
o The mixture of memory mapping and file locking on the same file is
problematic. Consider the following scenario, where a page size
on each client is 8192 bytes.
* Client A memory maps first page (8192 bytes) of file X
* Client B memory maps first page (8192 bytes) of file X
* Client A write locks first 4096 bytes
* Client B write locks second 4096 bytes
* Client A, via a STORE instruction modifies part of its locked
region.
* Simultaneous to client A, client B issues a STORE on part of
its locked region.
Here the challenge is for each client to resynchronize to get a
correct view of the first page. In many operating environments, the
virtual memory management systems on each client only know a page is
modified, not that a subset of the page corresponding to the
respective lock regions has been modified. So it is not possible for
each client to do the right thing, which is to only write to the
server that portion of the page that is locked. For example, if
client A simply writes out the page, and then client B writes out the
page, client A's data is lost.
Moreover, if mandatory locking is enabled on the file, then we have a
different problem. When clients A and B issue the STORE
instructions, the resulting page faults require a record lock on the
entire page. Each client then tries to extend their locked range to
the entire page, which results in a deadlock. Communicating the
NFS4ERR_DEADLOCK error to a STORE instruction is difficult at best.
If a client is locking the entire memory mapped file, there is no
problem with advisory or mandatory record locking, at least until the
client unlocks a region in the middle of the file.
Given the above issues the following are permitted:
o Clients and servers MAY deny memory mapping a file they know there
are record locks for.
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o Clients and servers MAY deny a record lock on a file they know is
memory mapped.
o A client MAY deny memory mapping a file that it knows requires
mandatory locking for I/O. If mandatory locking is enabled after
the file is opened and mapped, the client MAY deny the application
further access to its mapped file.
7.8. Name Caching
The results of LOOKUP and READDIR operations may be cached to avoid
the cost of subsequent LOOKUP operations. Just as in the case of
attribute caching, inconsistencies may arise among the various client
caches. To mitigate the effects of these inconsistencies and given
the context of typical filesystem APIs, an upper time boundary is
maintained on how long a client name cache entry can be kept without
verifying that the entry has not been made invalid by a directory
change operation performed by another client. .LP When a client is
not making changes to a directory for which there exist name cache
entries, the client needs to periodically fetch attributes for that
directory to ensure that it is not being modified. After determining
that no modification has occurred, the expiration time for the
associated name cache entries may be updated to be the current time
plus the name cache staleness bound.
When a client is making changes to a given directory, it needs to
determine whether there have been changes made to the directory by
other clients. It does this by using the change attribute as
reported before and after the directory operation in the associated
change_info4 value returned for the operation. The server is able to
communicate to the client whether the change_info4 data is provided
atomically with respect to the directory operation. If the change
values are provided atomically, the client is then able to compare
the pre-operation change value with the change value in the client's
name cache. If the comparison indicates that the directory was
updated by another client, the name cache associated with the
modified directory is purged from the client. If the comparison
indicates no modification, the name cache can be updated on the
client to reflect the directory operation and the associated timeout
extended. The post-operation change value needs to be saved as the
basis for future change_info4 comparisons.
As demonstrated by the scenario above, name caching requires that the
client revalidate name cache data by inspecting the change attribute
of a directory at the point when the name cache item was cached.
This requires that the server update the change attribute for
directories when the contents of the corresponding directory is
modified. For a client to use the change_info4 information
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appropriately and correctly, the server must report the pre and post
operation change attribute values atomically. When the server is
unable to report the before and after values atomically with respect
to the directory operation, the server must indicate that fact in the
change_info4 return value. When the information is not atomically
reported, the client should not assume that other clients have not
changed the directory.
7.9. Directory Caching
The results of READDIR operations may be used to avoid subsequent
READDIR operations. Just as in the cases of attribute and name
caching, inconsistencies may arise among the various client caches.
To mitigate the effects of these inconsistencies, and given the
context of typical filesystem APIs, the following rules should be
followed:
o Cached READDIR information for a directory which is not obtained
in a single READDIR operation must always be a consistent snapshot
of directory contents. This is determined by using a GETATTR
before the first READDIR and after the last of READDIR that
contributes to the cache.
o An upper time boundary is maintained to indicate the length of
time a directory cache entry is considered valid before the client
must revalidate the cached information.
The revalidation technique parallels that discussed in the case of
name caching. When the client is not changing the directory in
question, checking the change attribute of the directory with GETATTR
is adequate. The lifetime of the cache entry can be extended at
these checkpoints. When a client is modifying the directory, the
client needs to use the change_info4 data to determine whether there
are other clients modifying the directory. If it is determined that
no other client modifications are occurring, the client may update
its directory cache to reflect its own changes.
As demonstrated previously, directory caching requires that the
client revalidate directory cache data by inspecting the change
attribute of a directory at the point when the directory was cached.
This requires that the server update the change attribute for
directories when the contents of the corresponding directory is
modified. For a client to use the change_info4 information
appropriately and correctly, the server must report the pre and post
operation change attribute values atomically. When the server is
unable to report the before and after values atomically with respect
to the directory operation, the server must indicate that fact in the
change_info4 return value. When the information is not atomically
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reported, the client should not assume that other clients have not
changed the directory.
8. Security Negotiation
The NFSv4.0 specification contains three oversights and ambiguities
with respect to the SECINFO operation.
First, it is impossible for the client to use the SECINFO operation
to determine the correct security triple for accessing a parent
directory. This is because SECINFO takes as arguments the current
file handle and a component name. However, NFSv4.0 uses the LOOKUPP
operation to get the parent directory of the current file handle. If
the client uses the wrong security when issuing the LOOKUPP, and gets
back an NFS4ERR_WRONGSEC error, SECINFO is useless to the client.
The client is left with guessing which security the server will
accept. This defeats the purpose of SECINFO, which was to provide an
efficient method of negotiating security.
Second, there is ambiguity as to what the server should do when it is
passed a LOOKUP operation such that the server restricts access to
the current file handle with one security triple, and access to the
component with a different triple, and remote procedure call uses one
of the two security triples. Should the server allow the LOOKUP?
Third, there is a problem as to what the client must do (or can do),
whenever the server returns NFS4ERR_WRONGSEC in response to a PUTFH
operation. The NFSv4.0 specification says that client should issue a
SECINFO using the parent filehandle and the component name of the
filehandle that PUTFH was issued with. This may not be convenient
for the client.
This document resolves the above three issues in the context of
NFSv4.1.
9. Clarification of Security Negotiation in NFSv4.1
This section attempts to clarify NFSv4.1 security negotiation issues.
Unless noted otherwise, for any mention of PUTFH in this section, the
reader should interpret it as applying to PUTROOTFH and PUTPUBFH in
addition to PUTFH.
9.1. PUTFH + LOOKUP
The server implementation may decide whether to impose any
restrictions on export security administration. There are at least
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three approaches (Sc is the flavor set of the child export, Sp that
of the parent),
a) Sc <= Sp (<= for subset)
b) Sc ^ Sp != {} (^ for intersection, {} for the empty set)
c) free form
To support b (when client chooses a flavor that is not a member of
Sp) and c, PUTFH must NOT return NFS4ERR_WRONGSEC in case of security
mismatch. Instead, it should be returned from the LOOKUP that
follows.
Since the above guideline does not contradict a, it should be
followed in general.
9.2. PUTFH + LOOKUPP
Since SECINFO only works its way down, there is no way LOOKUPP can
return NFS4ERR_WRONGSEC without the server implementing
SECINFO_NO_NAME. SECINFO_NO_NAME solves this issue because via style
"parent", it works in the opposite direction as SECINFO (component
name is implicit in this case).
9.3. PUTFH + SECINFO
This case should be treated specially.
A security sensitive client should be allowed to choose a strong
flavor when querying a server to determine a file object's permitted
security flavors. The security flavor chosen by the client does not
have to be included in the flavor list of the export. Of course the
server has to be configured for whatever flavor the client selects,
otherwise the request will fail at RPC authentication.
In theory, there is no connection between the security flavor used by
SECINFO and those supported by the export. But in practice, the
client may start looking for strong flavors from those supported by
the export, followed by those in the mandatory set.
9.4. PUTFH + Anything Else
PUTFH must return NFS4ERR_WRONGSEC in case of security mismatch.
This is the most straightforward approach without having to add
NFS4ERR_WRONGSEC to every other operations.
PUTFH + SECINFO_NO_NAME (style "current_fh") is needed for the client
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to recover from NFS4ERR_WRONGSEC.
10. NFSv4.1 Sessions
10.1. Sessions Background
10.1.1. Introduction to Sessions
This draft proposes extensions to NFS version 4 [RFC3530] enabling it
to support sessions and endpoint management, and to support operation
atop RDMA-capable RPC over transports such as iWARP. [RDMAP, DDP]
These extensions enable support for exactly-once semantics by NFSv4
servers, multipathing and trunking of transport connections, and
enhanced security. The ability to operate over RDMA enables greatly
enhanced performance. Operation over existing TCP is enhanced as
well.
While discussed here with respect to IETF-chartered transports, the
proposed protocol is intended to function over other standards, such
as Infiniband. [IB]
The following are the major aspects of this proposal:
o Changes are proposed within the framework of NFSv4 minor
versioning. RPC, XDR, and the NFSv4 procedures and operations are
preserved. The proposed extension functions equally well over
existing transports and RDMA, and interoperates transparently with
existing implementations, both at the local programmatic interface
and over the wire.
o An explicit session is introduced to NFSv4, and new operations are
added to support it. The session allows for enhanced trunking,
failover and recovery, and authentication efficiency, along with
necessary support for RDMA. The session is implemented as
operations within NFSv4 COMPOUND and does not impact layering or
interoperability with existing NFSv4 implementations. The NFSv4
callback channel is dynamically associated and is connected by the
client and not the server, enhancing security and operation
through firewalls. In fact, the callback channel will be enabled
to share the same connection as the operations channel.
o An enhanced RPC layer enables NFSv4 operation atop RDMA. The
session assists RDMA-mode connection, and additional facilities
are provided for managing RDMA resources at both NFSv4 server and
client. Existing NFSv4 operations continue to function as before,
though certain size limits are negotiated. A companion draft to
this document, "RDMA Transport for ONC RPC" [RPCRDMA] is to be
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referenced for details of RPC RDMA support.
o Support for exactly-once semantics ("EOS") is enabled by the new
session facilities, by providing to the server a way to bound the
size of the duplicate request cache for a single client, and to
manage its persistent storage.
Block Diagram
+-----------------+-------------------------------------+
| NFSv4 | NFSv4 + session extensions |
+-----------------+------+----------------+-------------+
| Operations | Session | |
+------------------------+----------------+ |
| RPC/XDR | |
+-------------------------------+---------+ |
| Stream Transport | RDMA Transport |
+-------------------------------+-----------------------+
10.1.2. Motivation
NFS version 4 [RFC3530] has been granted "Proposed Standard" status.
The NFSv4 protocol was developed along several design points,
important among them: effective operation over wide-area networks,
including the Internet itself; strong security integrated into the
protocol; extensive cross-platform interoperability including
integrated locking semantics compatible with multiple operating
systems; and protocol extensibility.
The NFS version 4 protocol, however, does not provide support for
certain important transport aspects. For example, the protocol does
not address response caching, which is required to provide
correctness for retried client requests across a network partition,
nor does it provide an interoperable way to support trunking and
multipathing of connections. This leads to inefficiencies,
especially where trunking and multipathing are concerned, and
presents additional difficulties in supporting RDMA fabrics, in which
endpoints may require dedicated or specialized resources. Sessions
can be employed to unify NFS-level constructs such as the clientid,
with transport-level constructs such as transport endpoints. Each
transport endpoint draws on resources via its membership in a
session. Resource management can be more strictly maintained,
leading to greater server efficiency in implementing the protocol.
The enhanced operation over a session affords an opportunity to the
server to implement a highly reliable duplicate request cache, and
thereby export exactly-once semantics.
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NFSv4 advances the state of high-performance local sharing, by virtue
of its integrated security, locking, and delegation, and its
excellent coverage of the sharing semantics of multiple operating
systems. It is precisely this environment where exactly-once
semantics become a fundamental requirement.
Additionally, efforts to standardize a set of protocols for Remote
Direct Memory Access, RDMA, over the Internet Protocol Suite have
made significant progress. RDMA is a general solution to the problem
of CPU overhead incurred due to data copies, primarily at the
receiver. Substantial research has addressed this and has borne out
the efficacy of the approach. An overview of this is the RDDP
Problem Statement document, [RDDPPS].
Numerous upper layer protocols achieve extremely high bandwidth and
low overhead through the use of RDMA. Products from a wide variety
of vendors employ RDMA to advantage, and prototypes have demonstrated
the effectiveness of many more. Here, we are concerned specifically
with NFS and NFS-style upper layer protocols; examples from Network
Appliance [DAFS, DCK+03], Fujitsu Prime Software Technologies [FJNFS,
FJDAFS] and Harvard University [KM02] are all relevant.
By layering a session binding for NFS version 4 directly atop a
standard RDMA transport, a greatly enhanced level of performance and
transparency can be supported on a wide variety of operating system
platforms. These combined capabilities alter the landscape between
local filesystems and network attached storage, enable a new level of
performance, and lead new classes of application to take advantage of
NFS.
10.1.3. Problem Statement
Two issues drive the current proposal: correctness, and performance.
Both are instances of "raising the bar" for NFS, whereby the desire
to use NFS in new classes applications can be accommodated by
providing the basic features to make such use feasible. Such
applications include tightly coupled sharing environments such as
cluster computing, high performance computing (HPC) and information
processing such as databases. These trends are explored in depth in
[NFSPS].
The first issue, correctness, exemplified among the attributes of
local filesystems, is support for exactly-once semantics. Such
semantics have not been reliably available with NFS. Server-based
duplicate request caches [CJ89] help, but do not reliably provide
strict correctness. For the type of application which is expected to
make extensive use of the high-performance RDMA-enabled environment,
the reliable provision of such semantics is a fundamental
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requirement.
Introduction of a session to NFSv4 will address these issues. With
higher performance and enhanced semantics comes the problem of
enabling advanced endpoint management, for example high-speed
trunking, multipathing and failover. These characteristics enable
availability and performance. RFC3530 presents some issues in
permitting a single clientid to access a server over multiple
connections.
A second issue encountered in common by NFS implementations is the
CPU overhead required to implement the protocol. Primary among the
sources of this overhead is the movement of data from NFS protocol
messages to its eventual destination in user buffers or aligned
kernel buffers. The data copies consume system bus bandwidth and CPU
time, reducing the available system capacity for applications.
[RDDPPS] Achieving zero-copy with NFS has to date required
sophisticated, "header cracking" hardware and/or extensive platform-
specific virtual memory mapping tricks.
Combined in this way, NFSv4, RDMA and the emerging high-speed network
fabrics will enable delivery of performance which matches that of the
fastest local filesystems, preserving the key existing local
filesystem semantics, while enhancing them by providing network
filesystem sharing semantics.
RDMA implementations generally have other interesting properties,
such as hardware assisted protocol access, and support for user space
access to I/O. RDMA is compelling here for another reason; hardware
offloaded networking support in itself does not avoid data copies,
without resorting to implementing part of the NFS protocol in the
NIC. Support of RDMA by NFS enables the highest performance at the
architecture level rather than by implementation; this enables
ubiquitous and interoperable solutions.
By providing file access performance equivalent to that of local file
systems, NFSv4 over RDMA will enable applications running on a set of
client machines to interact through an NFSv4 file system, just as
applications running on a single machine might interact through a
local file system.
This raises the issue of whether additional protocol enhancements to
enable such interaction would be desirable and what such enhancements
would be. This is a complicated issue which the working group needs
to address and will not be further discussed in this document.
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10.1.4. NFSv4 Session Extension Characteristics
This draft will present a solution based upon minor versioning of
NFSv4. It will introduce a session to collect transport endpoints
and resources such as reply caching, which in turn enables
enhancements such as trunking, failover and recovery. It will
describe use of RDMA by employing support within an underlying RPC
layer [RPCRDMA]. Most importantly, it will focus on making the best
possible use of an RDMA transport.
These extensions are proposed as elements of a new minor revision of
NFS version 4. In this draft, NFS version 4 will be referred to
generically as "NFSv4", when describing properties common to all
minor versions. When referring specifically to properties of the
original, minor version 0 protocol, "NFSv4.0" will be used, and
changes proposed here for minor version 1 will be referred to as
"NFSv4.1".
This draft proposes only changes which are strictly upward-
compatible with existing RPC and NFS Application Programming
Interfaces (APIs).
10.2. Transport Issues
The Transport Issues section of the document explores the details of
utilizing the various supported transports.
10.2.1. Session Model
The first and most evident issue in supporting diverse transports is
how to provide for their differences. This draft proposes
introducing an explicit session.
A session introduces minimal protocol requirements, and provides for
a highly useful and convenient way to manage numerous endpoint-
related issues. The session is a local construct; it represents a
named, higher-layer object to which connections can refer, and
encapsulates properties important to each associated client.
A session is a dynamically created, long-lived server object created
by a client, used over time from one or more transport connections.
Its function is to maintain the server's state relative to the
connection(s) belonging to a client instance. This state is entirely
independent of the connection itself. The session in effect becomes
the object representing an active client on a connection or set of
connections.
Clients may create multiple sessions for a single clientid, and may
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wish to do so for optimization of transport resources, buffers, or
server behavior. A session could be created by the client to
represent a single mount point, for separate read and write
"channels", or for any number of other client-selected parameters.
The session enables several things immediately. Clients may
disconnect and reconnect (voluntarily or not) without loss of context
at the server. (Of course, locks, delegations and related
associations require special handling, and generally expire in the
extended absence of an open connection.) Clients may connect
multiple transport endpoints to this common state. The endpoints may
have all the same attributes, for instance when trunked on multiple
physical network links for bandwidth aggregation or path failover.
Or, the endpoints can have specific, special purpose attributes such
as callback channels.
The NFSv4 specification does not provide for any form of flow
control; instead it relies on the windowing provided by TCP to
throttle requests. This unfortunately does not work with RDMA, which
in general provides no operation flow control and will terminate a
connection in error when limits are exceeded. Limits are therefore
exchanged when a session is created; These limits then provide maxima
within which each session's connections must operate, they are
managed within these limits as described in [RPCRDMA]. The limits
may also be modified dynamically at the server's choosing by
manipulating certain parameters present in each NFSv4.1 request.
The presence of a maximum request limit on the session bounds the
requirements of the duplicate request cache. This can be used to
advantage by a server, which can accurately determine any storage
needs and enable it to maintain duplicate request cache persistence
and to provide reliable exactly-once semantics.
Finally, given adequate connection-oriented transport security
semantics, authentication and authorization may be cached on a per-
session basis, enabling greater efficiency in the issuing and
processing of requests on both client and server. A proposal for
transparent, server-driven implementation of this in NFSv4 has been
made. [CCM] The existence of the session greatly facilitates the
implementation of this approach. This is discussed in detail in the
Authentication Efficiencies section later in this draft.
10.2.2. Connection State
In RFC3530, the combination of a connected transport endpoint and a
clientid forms the basis of connection state. While has been made to
be workable with certain limitations, there are difficulties in
correct and robust implementation. The NFSv4.0 protocol must provide
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a server-initiated connection for the callback channel, and must
carefully specify the persistence of client state at the server in
the face of transport interruptions. The server has only the
client's transport address binding (the IP 4-tuple) to identify the
client RPC transaction stream and to use as a lookup tag on the
duplicate request cache. (A useful overview of this is in [RW96].)
If the server listens on multiple adddresses, and the client connects
to more than one, it must employ different clientid's on each,
negating its ability to aggregate bandwidth and redundancy. In
effect, each transport connection is used as the server's
representation of client state. But, transport connections are
potentially fragile and transitory.
In this proposal, a session identifier is assigned by the server upon
initial session negotiation on each connection. This identifier is
used to associate additional connections, to renegotiate after a
reconnect, to provide an abstraction for the various session
properties, and to address the duplicate request cache. No
transport-specific information is used in the duplicate request cache
implementation of an NFSv4.1 server, nor in fact the RPC XID itself.
The session identifier is unique within the server's scope and may be
subject to certain server policies such as being bounded in time.
It is envisioned that the primary transport model will be connection
oriented. Connection orientation brings with it certain potential
optimizations, such as caching of per-connection properties, which
are easily leveraged through the generality of the session. However,
it is possible that in future, other transport models could be
accommodated below the session abstraction.
10.2.3. NFSv4 Channels, Sessions and Connections
There are at least two types of NFSv4 channels: the "operations"
channel used for ordinary requests from client to server, and the
"back" channel, used for callback requests from server to client.
As mentioned above, different NFSv4 operations on these channels can
lead to different resource needs. For example, server callback
operations (CB_RECALL) are specific, small messages which flow from
server to client at arbitrary times, while data transfers such as
read and write have very different sizes and asymmetric behaviors.
It is sometimes impractical for the RDMA peers (NFSv4 client and
NFSv4 server) to post buffers for these various operations on a
single connection. Commingling of requests with responses at the
client receive queue is particularly troublesome, due both to the
need to manage both solicited and unsolicited completions, and to
provision buffers for both purposes. Due to the lack of any ordering
of callback requests versus response arrivals, without any other
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mechanisms, the client would be forced to allocate all buffers sized
to the worst case.
The callback requests are likely to be handled by a different task
context from that handling the responses. Significant demultiplexing
and thread management may be required if both are received on the
same queue. However, if callbacks are relatively rare (perhaps due
to client access patterns), many of these difficulties can be
minimized.
Also, the client may wish to perform trunking of operations channel
requests for performance reasons, or multipathing for availability.
This proposal permits both, as well as many other session and
connection possibilities, by permitting each operation to carry
session membership information and to share session (and clientid)
state in order to draw upon the appropriate resources. For example,
reads and writes may be assigned to specific, optimized connections,
or sorted and separated by any or all of size, idempotency, etc.
To address the problems described above, this proposal allows
multiple sessions to share a clientid, as well as for multiple
connections to share a session.
Single Connection model:
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NFSv4.1 Session
/ \
Operations_Channel [Back_Channel]
\ /
Connection
|
Multi-connection trunked model (2 operations channels shown):
NFSv4.1 Session
/ \
Operations_Channels [Back_Channel]
| | |
Connection Connection [Connection]
| | |
Multi-connection split-use model (2 mounts shown):
NFSv4.1 Session
/ \
(/home) (/usr/local - readonly)
/ \ |
Operations_Channel [Back_Channel] |
| | Operations_Channel
Connection [Connection] |
| | Connection
|
In this way, implementation as well as resource management may be
optimized. Each session will have its own response caching and
buffering, and each connection or channel will have its own transport
resources, as appropriate. Clients which do not require certain
behaviors may optimize such resources away completely, by using
specific sessions and not even creating the additional channels and
connections.
10.2.4. Reconnection, Trunking and Failover
Reconnection after failure references stored state on the server
associated with lease recovery during the grace period. The session
provides a convenient handle for storing and managing information
regarding the client's previous state on a per- connection basis,
e.g. to be used upon reconnection. Reconnection to a previously
existing session, and its stored resources, are covered in the
"Connection Models" section below.
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One important aspect of reconnection is that of RPC library support.
Traditionally, an Upper Layer RPC-based Protocol such as NFS leaves
all transport knowledge to the RPC layer implementation below it.
This allows NFS to operate over a wide variety of transports and has
proven to be a highly successful approach. The session, however,
introduces an abstraction which is, in a way, "between" RPC and
NFSv4.1. It is important that the session abstraction not have
ramifications within the RPC layer.
One such issue arises within the reconnection logic of RPC.
Previously, an explicit session binding operation, which established
session context for each new connection, was explored. This however
required that the session binding also be performed during reconnect,
which in turn required an RPC request. This additional request
requires new RPC semantics, both in implementation and the fact that
a new request is inserted into the RPC stream. Also, the binding of
a connection to a session required the upper layer to become "aware"
of connections, something the RPC layer abstraction architecturally
abstracts away. Therefore the session binding is not handled in
connection scope but instead explicitly carried in each request.
For Reliability Availability and Serviceability (RAS) issues such as
bandwidth aggregation and multipathing, clients frequently seek to
make multiple connections through multiple logical or physical
channels. The session is a convenient point to aggregate and manage
these resources.
10.2.5. Server Duplicate Request Cache
Server duplicate request caches, while not a part of an NFS protocol,
have become a standard, even required, part of any NFS
implementation. First described in [CJ89], the duplicate request
cache was initially found to reduce work at the server by avoiding
duplicate processing for retransmitted requests. A second, and in
the long run more important benefit, was improved correctness, as the
cache avoided certain destructive non-idempotent requests from being
reinvoked.
However, such caches do not provide correctness guarantees; they
cannot be managed in a reliable, persistent fashion. The reason is
understandable - their storage requirement is unbounded due to the
lack of any such bound in the NFS protocol, and they are dependent on
transport addresses for request matching.
As proposed in this draft, the presence of maximum request count
limits and negotiated maximum sizes allows the size and duration of
the cache to be bounded, and coupled with a long-lived session
identifier, enables its persistent storage on a per-session basis.
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This provides a single unified mechanism which provides the following
guarantees required in the NFSv4 specification, while extending them
to all requests, rather than limiting them only to a subset of state-
related requests:
"It is critical the server maintain the last response sent to the
client to provide a more reliable cache of duplicate non- idempotent
requests than that of the traditional cache described in [CJ89]..."
[RFC3530]
The maximum request count limit is the count of active operations,
which bounds the number of entries in the cache. Constraining the
size of operations additionally serves to limit the required storage
to the product of the current maximum request count and the maximum
response size. This storage requirement enables server- side
efficiencies.
Session negotiation allows the server to maintain other state. An
NFSv4.1 client invoking the session destroy operation will cause the
server to denegotiate (close) the session, allowing the server to
deallocate cache entries. Clients can potentially specify that such
caches not be kept for appropriate types of sessions (for example,
read-only sessions). This can enable more efficient server operation
resulting in improved response times, and more efficient sizing of
buffers and response caches.
Similarly, it is important for the client to explicitly learn whether
the server is able to implement reliable semantics. Knowledge of
whether these semantics are in force is critical for a highly
reliable client, one which must provide transactional integrity
guarantees. When clients request that the semantics be enabled for a
given session, the session reply must inform the client if the mode
is in fact enabled. In this way the client can confidently proceed
with operations without having to implement consistency facilities of
its own.
10.3. Session Initialization and Transfer Models
Session initialization issues, and data transfer models relevant to
both TCP and RDMA are discussed in this section.
10.3.1. Session Negotiation
The following parameters are exchanged between client and server at
session creation time. Their values allow the server to properly
size resources allocated in order to service the client's requests,
and to provide the server with a way to communicate limits to the
client for proper and optimal operation. They are exchanged prior to
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all session-related activity, over any transport type. Discussion of
their use is found in their descriptions as well as throughout this
section.
Maximum Requests
The client's desired maximum number of concurrent requests is
passed, in order to allow the server to size its reply cache
storage. The server may modify the client's requested limit
downward (or upward) to match its local policy and/or resources.
Over RDMA-capable RPC transports, the per-request management of
low-level transport message credits is handled within the RPC
layer. [RPCRDMA]
Maximum Request/Response Sizes
The maximum request and response sizes are exchanged in order to
permit allocation of appropriately sized buffers and request cache
entries. The size must allow for certain protocol minima,
allowing the receipt of maximally sized operations (e.g. RENAME
requests which contains two name strings). Note the maximum
request/response sizes cover the entire request/response message
and not simply the data payload as traditional NFS maximum read or
write size. Also note the server implementation may not, in fact
probably does not, require the reply cache entries to be sized as
large as the maximum response. The server may reduce the client's
requested sizes.
Inline Padding/Alignment
The server can inform the client of any padding which can be used
to deliver NFSv4 inline WRITE payloads into aligned buffers. Such
alignment can be used to avoid data copy operations at the server
for both TCP and inline RDMA transfers. For RDMA, the client
informs the server in each operation when padding has been
applied. [RPCRDMA]
Transport Attributes
A placeholder for transport-specific attributes is provided, with
a format to be determined. Possible examples of information to be
passed in this parameter include transport security attributes to
be used on the connection, RDMA- specific attributes, legacy
"private data" as used on existing RDMA fabrics, transport Quality
of Service attributes, etc. This information is to be passed to
the peer's transport layer by local means which is currently
outside the scope of this draft, however one attribute is provided
in the RDMA case:
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RDMA Read Resources
RDMA implementations must explicitly provision resources to
support RDMA Read requests from connected peers. These values
must be explicitly specified, to provide adequate resources for
matching the peer's expected needs and the connection's delay-
bandwidth parameters. The client provides its chosen value to the
server in the initial session creation, the value must be provided
in each client RDMA endpoint. The values are asymmetric and
should be set to zero at the server in order to conserve RDMA
resources, since clients do not issue RDMA Read operations in this
proposal. The result is communicated in the session response, to
permit matching of values across the connection. The value may
not be changed in the duration of the session, although a new
value may be requested as part of a new session.
10.3.2. RDMA Requirements
A complete discussion of the operation of RPC-based protocols atop
RDMA transports is in [RPCRDMA]. Where RDMA is considered, this
proposal assumes the use of such a layering; it addresses only the
upper layer issues relevant to making best use of RPC/RDMA.
A connection oriented (reliable sequenced) RDMA transport will be
required. There are several reasons for this. First, this model
most closely reflects the general NFSv4 requirement of long-lived and
congestion-controlled transports. Second, to operate correctly over
either an unreliable or unsequenced RDMA transport, or both, would
require significant complexity in the implementation and protocol not
appropriate for a strict minor version. For example, retransmission
on connected endpoints is explicitly disallowed in the current NFSv4
draft; it would again be required with these alternate transport
characteristics. Third, the proposal assumes a specific RDMA
ordering semantic, which presents the same set of ordering and
reliability issues to the RDMA layer over such transports.
The RDMA implementation provides for making connections to other
RDMA-capable peers. In the case of the current proposals before the
RDDP working group, these RDMA connections are preceded by a
"streaming" phase, where ordinary TCP (or NFS) traffic might flow.
However, this is not assumed here and sizes and other parameters are
explicitly exchanged upon a session entering RDMA mode.
10.3.3. RDMA Connection Resources
On transport endpoints which support automatic RDMA mode, that is,
endpoints which are created in the RDMA-enabled state, a single,
preposted buffer must initially be provided by both peers, and the
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client session negotiation must be the first exchange.
On transport endpoints supporting dynamic negotiation, a more
sophisticated negotiation is possible, but is not discussed in the
current draft.
RDMA imposes several requirements on upper layer consumers.
Registration of memory and the need to post buffers of a specific
size and number for receive operations are a primary consideration.
Registration of memory can be a relatively high-overhead operation,
since it requires pinning of buffers, assignment of attributes (e.g.
readable/writable), and initialization of hardware translation.
Preregistration is desirable to reduce overhead. These registrations
are specific to hardware interfaces and even to RDMA connection
endpoints, therefore negotiation of their limits is desirable to
manage resources effectively.
Following the basic registration, these buffers must be posted by the
RPC layer to handle receives. These buffers remain in use by the
RPC/NFSv4 implementation; the size and number of them must be known
to the remote peer in order to avoid RDMA errors which would cause a
fatal error on the RDMA connection.
The session provides a natural way for the server to manage resource
allocation to each client rather than to each transport connection
itself. This enables considerable flexibility in the administration
of transport endpoints.
10.3.4. TCP and RDMA Inline Transfer Model
The basic transfer model for both TCP and RDMA is referred to as
"inline". For TCP, this is the only transfer model supported, since
TCP carries both the RPC header and data together in the data stream.
For RDMA, the RDMA Send transfer model is used for all NFS requests
and replies, but data is optionally carried by RDMA Writes or RDMA
Reads. Use of Sends is required to ensure consistency of data and to
deliver completion notifications. The pure-Send method is typically
used where the data payload is small, or where for whatever reason
target memory for RDMA is not available.
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Inline message exchange
Client Server
: Request :
Send : ------------------------------> : untagged
: : buffer
: Response :
untagged : <------------------------------ : Send
buffer : :
Client Server
: Read request :
Send : ------------------------------> : untagged
: : buffer
: Read response with data :
untagged : <------------------------------ : Send
buffer : :
Client Server
: Write request with data :
Send : ------------------------------> : untagged
: : buffer
: Write response :
untagged : <------------------------------ : Send
buffer : :
Responses must be sent to the client on the same connection that the
request was sent. It is important that the server does not assume
any specific client implementation, in particular whether connections
within a session share any state at the client. This is also
important to preserve ordering of RDMA operations, and especially
RMDA consistency. Additionally, it ensures that the RPC RDMA layer
makes no requirement of the RDMA provider to open its memory
registration handles (Steering Tags) beyond the scope of a single
RDMA connection. This is an important security consideration.
Two values must be known to each peer prior to issuing Sends: the
maximum number of sends which may be posted, and their maximum size.
These values are referred to, respectively, as the message credits
and the maximum message size. While the message credits might vary
dynamically over the duration of the session, the maximum message
size does not. The server must commit to preserving this number of
duplicate request cache entires, and preparing a number of receive
buffers equal to or greater than its currently advertised credit
value, each of the advertised size. These ensure that transport
resources are allocated sufficient to receive the full advertised
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limits.
Note that the server must post the maximum number of session requests
to each client operations channel. The client is not required to
spread its requests in any particular fashion across connections
within a session. If the client wishes, it may create multiple
sessions, each with a single or small number of operations channels
to provide the server with this resource advantage. Or, over RDMA
the server may employ a "shared receive queue". The server can in
any case protect its resources by restricting the client's request
credits.
While tempting to consider, it is not possible to use the TCP window
as an RDMA operation flow control mechanism. First, to do so would
violate layering, requiring both senders to be aware of the existing
TCP outbound window at all times. Second, since requests are of
variable size, the TCP window can hold a widely variable number of
them, and since it cannot be reduced without actually receiving data,
the receiver cannot limit the sender. Third, any middlebox
interposing on the connection would wreck any possible scheme.
[MIDTAX] In this proposal, maximum request count limits are exchanged
at the session level to allow correct provisioning of receive buffers
by transports.
When operating over TCP or other similar transport, request limits
and sizes are still employed in NFSv4.1, but instead of being
required for correctness, they provide the basis for efficient server
implementation of the duplicate request cache. The limits are chosen
based upon the expected needs and capabilities of the client and
server, and are in fact arbitrary. Sizes may be specified by the
client as zero (requesting the server's preferred or optimal value),
and request limits may be chosen in proportion to the client's
capabilities. For example, a limit of 1000 allows 1000 requests to
be in progress, which may generally be far more than adequate to keep
local networks and servers fully utilized.
Both client and server have independent sizes and buffering, but over
RDMA fabrics client credits are easily managed by posting a receive
buffer prior to sending each request. Each such buffer may not be
completed with the corresponding reply, since responses from NFSv4
servers arrive in arbitrary order. When an operations channel is
also used for callbacks, the client must account for callback
requests by posting additional buffers. Note that implementation-
specific facilities such as a shared receive queue may also allow
optimization of these allocations.
When a session is created, the client requests a preferred buffer
size, and the server provides its answer. The server posts all
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buffers of at least this size. The client must comply by not sending
requests greater than this size. It is recommended that server
implementations do all they can to accommodate a useful range of
possible client requests. There is a provision in [RPCRDMA] to allow
the sending of client requests which exceed the server's receive
buffer size, but it requires the server to "pull" the client's
request as a "read chunk" via RDMA Read. This introduces at least
one additional network roundtrip, plus other overhead such as
registering memory for RDMA Read at the client and additional RDMA
operations at the server, and is to be avoided.
An issue therefore arises when considering the NFSv4 COMPOUND
procedures. Since an arbitrary number (total size) of operations can
be specified in a single COMPOUND procedure, its size is effectively
unbounded. This cannot be supported by RDMA Sends, and therefore
this size negotiation places a restriction on the construction and
maximum size of both COMPOUND requests and responses. If a COMPOUND
results in a reply at the server that is larger than can be sent in
an RDMA Send to the client, then the COMPOUND must terminate and the
operation which causes the overflow will provide a TOOSMALL error
status result.
10.3.5. RDMA Direct Transfer Model
Placement of data by explicitly tagged RDMA operations is referred to
as "direct" transfer. This method is typically used where the data
payload is relatively large, that is, when RDMA setup has been
performed prior to the operation, or when any overhead for setting up
and performing the transfer is regained by avoiding the overhead of
processing an ordinary receive.
The client advertises RDMA buffers in this proposed model, and not
the server. This means the "XDR Decoding with Read Chunks" described
in [RPCRDMA] is not employed by NFSv4.1 replies, and instead all
results transferred via RDMA to the client employ "XDR Decoding with
Write Chunks". There are several reasons for this.
First, it allows for a correct and secure mode of transfer. The
client may advertise specific memory buffers only during specific
times, and may revoke access when it pleases. The server is not
required to expose copies of local file buffers for individual
clients, or to lock or copy them for each client access.
Second, client credits based on fixed-size request buffers are easily
managed on the server, but for the server additional management of
buffers for client RDMA Reads is not well-bounded. For example, the
client may not perform these RDMA Read operations in a timely
fashion, therefore the server would have to protect itself against
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denial-of-service on these resources.
Third, it reduces network traffic, since buffer exposure outside the
scope and duration of a single request/response exchange necessitates
additional memory management exchanges.
There are costs associated with this decision. Primary among them is
the need for the server to employ RDMA Read for operations such as
large WRITE. The RDMA Read operation is a two-way exchange at the
RDMA layer, which incurs additional overhead relative to RDMA Write.
Additionally, RDMA Read requires resources at the data source (the
client in this proposal) to maintain state and to generate replies.
These costs are overcome through use of pipelining with credits, with
sufficient RDMA Read resources negotiated at session initiation, and
appropriate use of RDMA for writes by the client - for example only
for transfers above a certain size.
A description of which NFSv4 operation results are eligible for data
transfer via RDMA Write is in [NFSDDP]. There are only two such
operations: READ and READLINK. When XDR encoding these requests on
an RDMA transport, the NFSv4.1 client must insert the appropriate
xdr_write_list entries to indicate to the server whether the results
should be transferred via RDMA or inline with a Send. As described
in [NFSDDP], a zero-length write chunk is used to indicate an inline
result. In this way, it is unnecessary to create new operations for
RDMA-mode versions of READ and READLINK.
Another tool to avoid creation of new, RDMA-mode operations is the
Reply Chunk [RPCRDMA], which is used by RPC in RDMA mode to return
large replies via RDMA as if they were inline. Reply chunks are used
for operations such as READDIR, which returns large amounts of
information, but in many small XDR segments. Reply chunks are
offered by the client and the server can use them in preference to
inline. Reply chunks are transparent to upper layers such as NFSv4.
In any very rare cases where another NFSv4.1 operation requires
larger buffers than were negotiated when the session was created (for
example extraordinarily large RENAMEs), the underlying RPC layer may
support the use of "Message as an RDMA Read Chunk" and "RDMA Write of
Long Replies" as described in [RPCRDMA]. No additional support is
required in the NFSv4.1 client for this. The client should be
certain that its requested buffer sizes are not so small as to make
this a frequent occurrence, however.
All operations are initiated by a Send, and are completed with a
Send. This is exactly as in conventional NFSv4, but under RDMA has a
significant purpose: RDMA operations are not complete, that is,
guaranteed consistent, at the data sink until followed by a
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successful Send completion (i.e. a receive). These events provide a
natural opportunity for the initiator (client) to enable and later
disable RDMA access to the memory which is the target of each
operation, in order to provide for consistent and secure operation.
The RDMAP Send with Invalidate operation may be worth employing in
this respect, as it relieves the client of certain overhead in this
case.
A "onetime" boolean advisory to each RDMA region might become a hint
to the server that the client will use the three-tuple for only one
NFSv4 operation. For a transport such as iWARP, the server can
assist the client in invalidating the three-tuple by performing a
Send with Solicited Event and Invalidate. The server may ignore this
hint, in which case the client must perform a local invalidate after
receiving the indication from the server that the NFSv4 operation is
complete. This may be considered in a future version of this draft
and [NFSDDP].
In a trusted environment, it may be desirable for the client to
persistently enable RDMA access by the server. Such a model is
desirable for the highest level of efficiency and lowest overhead.
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RDMA message exchanges
Client Server
: Direct Read Request :
Send : ------------------------------> : untagged
: : buffer
: Segment :
tagged : <------------------------------ : RDMA Write
buffer : : :
: [Segment] :
tagged : <------------------------------ : [RDMA Write]
buffer : :
: Direct Read Response :
untagged : <------------------------------ : Send (w/Inv.)
buffer : :
Client Server
: Direct Write Request :
Send : ------------------------------> : untagged
: : buffer
: Segment :
tagged : v------------------------------ : RDMA Read
buffer : +-----------------------------> :
: : :
: [Segment] :
tagged : v------------------------------ : [RDMA Read]
buffer : +-----------------------------> :
: :
: Direct Write Response :
untagged : <------------------------------ : Send (w/Inv.)
buffer : :
10.4. Connection Models
There are three scenarios in which to discuss the connection model.
Each will be discussed individually, after describing the common case
encountered at initial connection establishment.
After a successful connection, the first request proceeds, in the
case of a new client association, to initial session creation, and
then optionally to session callback channel binding, prior to regular
operation.
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Commonly, each new client "mount" will be the action which drives
creation of a new session. However there are any number of other
approaches. Clients may choose to share a single connection and
session among all their mount points. Or, clients may support
trunking, where additional connections are created but all within a
single session. Alternatively, the client may choose to create
multiple sessions, each tuned to the buffering and reliability needs
of the mount point. For example, a readonly mount can sharply reduce
its write buffering and also makes no requirement for the server to
support reliable duplicate request caching.
Similarly, the client can choose among several strategies for
clientid usage. Sessions can share a single clientid, or create new
clientids as the client deems appropriate. For kernel-based clients
which service multiple authenticated users, a single clientid shared
across all mount points is generally the most appropriate and
flexible approach. For example, all the client's file operations may
wish to share locking state and the local client kernel takes the
responsibility for arbitrating access locally. For clients choosing
to support other authentication models, perhaps example userspace
implementations, a new clientid is indicated. Through use of session
create options, both models are supported at the client's choice.
Since the session is explicitly created and destroyed by the client,
and each client is uniquely identified, the server may be
specifically instructed to discard unneeded presistent state. For
this reason, it is possible that a server will retain any previous
state indefinitely, and place its destruction under administrative
control. Or, a server may choose to retain state for some
configurable period, provided that the period meets other NFSv4
requirements such as lease reclamation time, etc. However, since
discarding this state at the server may affect the correctness of the
server as seen by the client across network partitioning, such
discarding of state should be done only in a conservative manner.
Each client request to the server carries a new SEQUENCE operation
within each COMPOUND, which provides the session context. This
session context then governs the request control, duplicate request
caching, and other persistent parameters managed by the server for a
session.
10.4.1. TCP Connection Model
The following is a schematic diagram of the NFSv4.1 protocol
exchanges leading up to normal operation on a TCP stream.
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Client Server
TCPmode : Create Clientid(nfs_client_id4) : TCPmode
: ------------------------------> :
: :
: Clientid reply(clientid, ...) :
: <------------------------------ :
: :
: Create Session(clientid, size S, :
: maxreq N, STREAM, ...) :
: ------------------------------> :
: :
: Session reply(sessionid, size S', :
: maxreq N') :
: <------------------------------ :
: :
: <normal operation> :
: ------------------------------> :
: <------------------------------ :
: : :
No net additional exchange is added to the initial negotiation by
this proposal. In the NFSv4.1 exchange, the CREATECLIENTID replaces
SETCLIENTID (eliding the callback "clientaddr4" addressing) and
CREATESESSION subsumes the function of SETCLIENTID_CONFIRM, as
described elsewhere in this document. Callback channel binding is
optional, as in NFSv4.0. Note that the STREAM transport type is
shown above, but since the transport mode remains unchanged and
transport attributes are not necessarily exchanged, DEFAULT could
also be passed.
10.4.2. Negotiated RDMA Connection Model
One possible design which has been considered is to have a
"negotiated" RDMA connection model, supported via use of a session
bind operation as a required first step. However due to issues
mentioned earlier, this proved problematic. This section remains as
a reminder of that fact, and it is possible such a mode can be
supported.
It is not considered critical that this be supported for two reasons.
One, the session persistence provides a way for the server to
remember important session parameters, such as sizes and maximum
request counts. These values can be used to restore the endpoint
prior to making the first reply. Two, there are currently no
critical RDMA parameters to set in the endpoint at the server side of
the connection. RDMA Read resources, which are in general not
settable after entering RDMA mode, are set only at the client - the
originator of the connection. Therefore as long as the RDMA provider
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supports an automatic RDMA connection mode, no further support is
required from the NFSv4.1 protocol for reconnection.
Note, the client must provide at least as many RDMA Read resources to
its local queue for the benefit of the server when reconnecting, as
it used when negotiating the session. If this value is no longer
appropriate, the client should resynchronize its session state,
destroy the existing session, and start over with the more
appropriate values.
10.4.3. Automatic RDMA Connection Model
The following is a schematic diagram of the NFSv4.1 protocol
exchanges performed on an RDMA connection.
Client Server
RDMAmode : : : RDMAmode
: : :
Prepost : : : Prepost
receive : : : receive
: :
: Create Clientid(nfs_client_id4) :
: ------------------------------> :
: : Prepost
: Clientid reply(clientid, ...) : receive
: <------------------------------ :
Prepost : :
receive : Create Session(clientid, size S, :
: maxreq N, RDMA ...) :
: ------------------------------> :
: : Prepost <=N'
: Session reply(sessionid, size S', : receives of
: maxreq N') : size S'
: <------------------------------ :
: :
: <normal operation> :
: ------------------------------> :
: <------------------------------ :
: : :
10.5. Buffer Management, Transfer, Flow Control
Inline operations in NFSv4.1 behave effectively the same as TCP
sends. Procedure results are passed in a single message, and its
completion at the client signal the receiving process to inspect the
message.
RDMA operations are performed solely by the server in this proposal,
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as described in the previous "RDMA Direct Model" section. Since
server RDMA operations do not result in a completion at the client,
and due to ordering rules in RDMA transports, after all required RDMA
operations are complete, a Send (Send with Solicited Event for iWARP)
containing the procedure results is performed from server to client.
This Send operation will result in a completion which will signal the
client to inspect the message.
In the case of client read-type NFSv4 operations, the server will
have issued RDMA Writes to transfer the resulting data into client-
advertised buffers. The subsequent Send operation performs two
necessary functions: finalizing any active or pending DMA at the
client, and signaling the client to inspect the message.
In the case of client write-type NFSv4 operations, the server will
have issued RDMA Reads to fetch the data from the client-advertised
buffers. No data consistency issues arise at the client, but the
completion of the transfer must be acknowledged, again by a Send from
server to client.
In either case, the client advertises buffers for direct (RDMA style)
operations. The client may desire certain advertisement limits, and
may wish the server to perform remote invalidation on its behalf when
the server has completed its RDMA. This may be considered in a
future version of this draft.
In the absence of remote invalidation, the client may perform its
own, local invalidation after the operation completes. This
invalidation should occur prior to any RPCSEC GSS integrity checking,
since a validly remotely accessible buffer can possibly be modified
by the peer. However, after invalidation and the contents integrity
checked, the contents are locally secure.
Credit updates over RDMA transports are supported at the RPC layer as
described in [RPCRDMA]. In each request, the client requests a
desired number of credits to be made available to the connection on
which it sends the request. The client must not send more requests
than the number which the server has previously advertised, or in the
case of the first request, only one. If the client exceeds its
credit limit, the connection may close with a fatal RDMA error.
The server then executes the request, and replies with an updated
credit count accompanying its results. Since replies are sequenced
by their RDMA Send order, the most recent results always reflect the
server's limit. In this way the client will always know the maximum
number of requests it may safely post.
Because the client requests an arbitrary credit count in each
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request, it is relatively easy for the client to request more, or
fewer, credits to match its expected need. A client that discovered
itself frequently queuing outgoing requests due to lack of server
credits might increase its requested credits proportionately in
response. Or, a client might have a simple, configurable number.
The protocol also provides a per-operation "maxslot" exchange to
assist in dynamic adjustment at the session level, described in a
later section.
Occasionally, a server may wish to reduce the total number of credits
it offers a certain client on a connection. This could be
encountered if a client were found to be consuming its credits
slowly, or not at all. A client might notice this itself, and reduce
its requested credits in advance, for instance requesting only the
count of operations it currently has queued, plus a few as a base for
starting up again. Such mechanisms can, however, be potentially
complicated and are implementation-defined. The protocol does not
require them.
Because of the way in which RDMA fabrics function, it is not possible
for the server (or client back channel) to cancel outstanding receive
operations. Therefore, effectively only one credit can be withdrawn
per receive completion. The server (or client back channel) would
simply not replenish a receive operation when replying. The server
can still reduce the available credit advertisement in its replies to
the target value it desires, as a hint to the client that its credit
target is lower and it should expect it to be reduced accordingly.
Of course, even if the server could cancel outstanding receives, it
cannot do so, since the client may have already sent requests in
expectation of the previous limit.
This brings out an interesting scenario similar to the client
reconnect discussed earlier in "Connection Models". How does the
server reduce the credits of an inactive client?
One approach is for the server to simply close such a connection and
require the client to reconnect at a new credit limit. This is
acceptable, if inefficient, when the connection setup time is short
and where the server supports persistent session semantics.
A better approach is to provide a back channel request to return the
operations channel credits. The server may request the client to
return some number of credits, the client must comply by performing
operations on the operations channel, provided of course that the
request does not drop the client's credit count to zero (in which
case the connection would deadlock). If the client finds that it has
no requests with which to consume the credits it was previously
granted, it must send zero-length Send RDMA operations, or NULL NFSv4
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operations in order to return the resources to the server. If the
client fails to comply in a timely fashion, the server can recover
the resources by breaking the connection.
While in principle, the back channel credits could be subject to a
similar resource adjustment, in practice this is not an issue, since
the back channel is used purely for control and is expected to be
statically provisioned.
It is important to note that in addition to maximum request counts,
the sizes of buffers are negotiated per-session. This permits the
most efficient allocation of resources on both peers. There is an
important requirement on reconnection: the sizes posted by the server
at reconnect must be at least as large as previously used, to allow
recovery. Any replies that are replayed from the server's duplicate
request cache must be able to be received into client buffers. In
the case where a client has received replies to all its retried
requests (and therefore received all its expected responses), then
the client may disconnect and reconnect with different buffers at
will, since no cache replay will be required.
10.6. Retry and Replay
NFSv4.0 forbids retransmission on active connections over reliable
transports; this includes connected-mode RDMA. This restriction must
be maintained in NFSv4.1.
If one peer were to retransmit a request (or reply), it would consume
an additional credit on the other. If the server retransmitted a
reply, it would certainly result in an RDMA connection loss, since
the client would typically only post a single receive buffer for each
request. If the client retransmitted a request, the additional
credit consumed on the server might lead to RDMA connection failure
unless the client accounted for it and decreased its available
credit, leading to wasted resources.
RDMA credits present a new issue to the duplicate request cache in
NFSv4.1. The request cache may be used when a connection within a
session is lost, such as after the client reconnects. Credit
information is a dynamic property of the connection, and stale values
must not be replayed from the cache. This implies that the request
cache contents must not be blindly used when replies are issued from
it, and credit information appropriate to the channel must be
refreshed by the RPC layer.
Finally, RDMA fabrics do not guarantee that the memory handles
(Steering Tags) within each rdma three-tuple are valid on a scope
outside that of a single connection. Therefore, handles used by the
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direct operations become invalid after connection loss. The server
must ensure that any RDMA operations which must be replayed from the
request cache use the newly provided handle(s) from the most recent
request.
10.7. The Back Channel
The NFSv4 callback operations present a significant resource problem
for the RDMA enabled client. Clearly, callbacks must be negotiated
in the way credits are for the ordinary operations channel for
requests flowing from client to server. But, for callbacks to arrive
on the same RDMA endpoint as operation replies would require
dedicating additional resources, and specialized demultiplexing and
event handling. Or, callbacks may not require RDMA sevice at all
(they do not normally carry substantial data payloads). It is highly
desirable to streamline this critical path via a second
communications channel.
The session callback channel binding facility is designed for exactly
such a situation, by dynamically associating a new connected endpoint
with the session, and separately negotiating sizes and counts for
active callback channel operations. The binding operation is
firewall-friendly since it does not require the server to initiate
the connection.
This same method serves as well for ordinary TCP connection mode. It
is expected that all NFSv4.1 clients may make use of the session
facility to streamline their design.
The back channel functions exactly the same as the operations channel
except that no RDMA operations are required to perform transfers,
instead the sizes are required to be sufficiently large to carry all
data inline, and of course the client and server reverse their roles
with respect to which is in control of credit management. The same
rules apply for all transfers, with the server being required to flow
control its callback requests.
The back channel is optional. If not bound on a given session, the
server must not issue callback operations to the client. This in
turn implies that such a client must never put itself in the
situation where the server will need to do so, lest the client lose
its connection by force, or its operation be incorrect. For the same
reason, if a back channel is bound, the client is subject to
revocation of its delegations if the back channel is lost. Any
connection loss should be corrected by the client as soon as
possible.
This can be convenient for the NFSv4.1 client; if the client expects
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to make no use of back channel facilities such as delegations, then
there is no need to create it. This may save significant resources
and complexity at the client.
For these reasons, if the client wishes to use the back channel, that
channel must be bound first, before using the operations channel. In
this way, the server will not find itself in a position where it will
send callbacks on the operations channel when the client is not
prepared for them.
There is one special case, that where the back channel is bound in
fact to the operations channel's connection. This configuration
would be used normally over a TCP stream connection to exactly
implement the NFSv4.0 behavior, but over RDMA would require complex
resource and event management at both sides of the connection. The
server is not required to accept such a bind request on an RDMA
connection for this reason, though it is recommended.
10.8. COMPOUND Sizing Issues
Very large responses may pose duplicate request cache issues. Since
servers will want to bound the storage required for such a cache, the
unlimited size of response data in COMPOUND may be troublesome. If
COMPOUND is used in all its generality, then the inclusion of certain
non-idempotent operations within a single COMPOUND request may render
the entire request non-idempotent. (For example, a single COMPOUND
request which read a file or symbolic link, then removed it, would be
obliged to cache the data in order to allow identical replay).
Therefore, many requests might include operations that return any
amount of data.
It is not satisfactory for the server to reject COMPOUNDs at will
with NFS4ERR_RESOURCE when they pose such difficulties for the
server, as this results in serious interoperability problems.
Instead, any such limits must be explicitly exposed as attributes of
the session, ensuring that the server can explicitly support any
duplicate request cache needs at all times.
10.9. Data Alignment
A negotiated data alignment enables certain scatter/gather
optimizations. A facility for this is supported by [RPCRDMA]. Where
NFS file data is the payload, specific optimizations become highly
attractive.
Header padding is requested by each peer at session initiation, and
may be zero (no padding). Padding leverages the useful property that
RDMA receives preserve alignment of data, even when they are placed
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into anonymous (untagged) buffers. If requested, client inline
writes will insert appropriate pad bytes within the request header to
align the data payload on the specified boundary. The client is
encouraged to be optimistic and simply pad all WRITEs within the RPC
layer to the negotiated size, in the expectation that the server can
use them efficiently.
It is highly recommended that clients offer to pad headers to an
appropriate size. Most servers can make good use of such padding,
which allows them to chain receive buffers in such a way that any
data carried by client requests will be placed into appropriate
buffers at the server, ready for filesystem processing. The
receiver's RPC layer encounters no overhead from skipping over pad
bytes, and the RDMA layer's high performance makes the insertion and
transmission of padding on the sender a significant optimization. In
this way, the need for servers to perform RDMA Read to satisfy all
but the largest client writes is obviated. An added benefit is the
reduction of message roundtrips on the network - a potentially good
trade, where latency is present.
The value to choose for padding is subject to a number of criteria.
A primary source of variable-length data in the RPC header is the
authentication information, the form of which is client-determined,
possibly in response to server specification. The contents of
COMPOUNDs, sizes of strings such as those passed to RENAME, etc. all
go into the determination of a maximal NFSv4 request size and
therefore minimal buffer size. The client must select its offered
value carefully, so as not to overburden the server, and vice- versa.
The payoff of an appropriate padding value is higher performance.
Sender gather:
|RPC Request|Pad bytes|Length| -> |User data...|
\------+---------------------/ \
\ \
\ Receiver scatter: \-----------+- ...
/-----+----------------\ \ \
|RPC Request|Pad|Length| -> |FS buffer|->|FS buffer|->...
In the above case, the server may recycle unused buffers to the next
posted receive if unused by the actual received request, or may pass
the now-complete buffers by reference for normal write processing.
For a server which can make use of it, this removes any need for data
copies of incoming data, without resorting to complicated end-to-end
buffer advertisement and management. This includes most kernel-based
and integrated server designs, among many others. The client may
perform similar optimizations, if desired.
Padding is negotiated by the session creation operation, and
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subsequently used by the RPC RDMA layer, as described in [RPCRDMA].
10.10. NFSv4 Integration
The following section discusses the integration of the proposed RDMA
extensions with NFSv4.0.
10.10.1. Minor Versioning
Minor versioning is the existing facility to extend the NFSv4
protocol, and this proposal takes that approach.
Minor versioning of NFSv4 is relatively restrictive, and allows for
tightly limited changes only. In particular, it does not permit
adding new "procedures" (it permits adding only new "operations").
Interoperability concerns make it impossible to consider additional
layering to be a minor revision. This somewhat limits the changes
that can be proposed when considering extensions.
To support the duplicate request cache integrated with sessions and
request control, it is desirable to tag each request with an
identifier to be called a Slotid. This identifier must be passed by
NFSv4 when running atop any transport, including traditional TCP.
Therefore it is not desirable to add the Slotid to a new RPC
transport, even though such a transport is indicated for support of
RDMA. This draft and [RPCRDMA] do not propose such an approach.
Instead, this proposal conforms to the requirements of NFSv4 minor
versioning, through the use of a new operation within NFSv4 COMPOUND
procedures as detailed below.
If sessions are in use for a given clientid, this same clientid
cannot be used for non-session NFSv4 operation, including NFSv4.0.
Because the server will have allocated session-specific state to the
active clientid, it would be an unnecessary burden on the server
implementor to support and account for additional, non- session
traffic, in addition to being of no benefit. Therefore this proposal
prohibits a single clientid from doing this. Nevertheless, employing
a new clientid for such traffic is supported.
10.10.2. Slot Identifiers and Server Duplicate Request Cache
The presence of deterministic maximum request limits on a session
enables in-progress requests to be assigned unique values with useful
properties.
The RPC layer provides a transaction ID (xid), which, while required
to be unique, is not especially convenient for tracking requests.
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The transaction ID is only meaningful to the issuer (client), it
cannot be interpreted at the server except to test for equality with
previously issued requests. Because RPC operations may be completed
by the server in any order, many transaction IDs may be outstanding
at any time. The client may therefore perform a computationally
expensive lookup operation in the process of demultiplexing each
reply.
In the proposal, there is a limit to the number of active requests.
This immediately enables a convenient, computationally efficient
index for each request which is designated as a Slot Identifier, or
slotid.
When the client issues a new request, it selects a slotid in the
range 0..N-1, where N is the server's current "totalrequests" limit
granted the client on the session over which the request is to be
issued. The slotid must be unused by any of the requests which the
client has already active on the session. "Unused" here means the
client has no outstanding request for that slotid. Because the slot
id is always an integer in the range 0..N-1, client implementations
can use the slotid from a server response to efficiently match
responses with outstanding requests, such as, for example, by using
the slotid to index into a outstanding request array. This can be
used to avoid expensive hashing and lookup functions in the
performace-critical receive path.
The sequenceid, which accompanies the slotid in each request, is
important for a second, important check at the server: it must be
able to be determined efficiently whether a request using a certain
slotid is a retransmit or a new, never-before-seen request. It is
not feasible for the client to assert that it is retransmitting to
implement this, because for any given request the client cannot know
the server has seen it unless the server actually replies. Of
course, if the client has seen the server's reply, the client would
not retransmit!
The sequenceid must increase monotonically for each new transmit of a
given slotid, and must remain unchanged for any retransmission. The
server must in turn compare each newly received request's sequenceid
with the last one previously received for that slotid, to see if the
new request is:
o A new request, in which the sequenceid is greater than that
previously seen in the slot (accounting for sequence wraparound).
The server proceeds to execute the new request.
o A retransmitted request, in which the sequenceid is equal to that
last seen in the slot. Note that this request may be either
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complete, or in progress. The server performs replay processing
in these cases.
o A misordered duplicate, in which the sequenceid is less than that
previously seen in the slot. The server must drop the incoming
request, which may imply dropping the connection if the transport
is reliable, as dictated by section 3.1.1 of [RFC3530].
This last condition is possible on any connection, not just
unreliable, unordered transports. Delayed behavior on abandoned TCP
connections which are not yet closed at the server, or pathological
client implementations can cause it, among other causes. Therefore,
the server may wish to harden itself against certain repeated
occurrences of this, as it would for retransmissions in [RFC3530].
It is recommended, though not necessary for protocol correctness,
that the client simply increment the sequenceid by one for each new
request on each slotid. This reduces the wraparound window to a
minimum, and is useful for tracing and avoidance of possible
implementation errors.
The client may however, for implementation-specific reasons, choose a
different algorithm. For example it might maintain a single sequence
space for all slots in the session - e.g. employing the RPC XID
itself. The sequenceid, in any case, is never interpreted by the
server for anything but to test by comparison with previously seen
values.
The server may thereby use the slotid, in conjunction with the
sessionid and sequenceid, within the SEQUENCE portion of the request
to maintain its duplicate request cache (DRC) for the session, as
opposed to the traditional approach of ONC RPC applications that use
the XID along with certain transport information [RW96].
Unlike the XID, the slotid is always within a specific range; this
has two implications. The first implication is that for a given
session, the server need only cache the results of a limited number
of COMPOUND requests. The second implication derives from the first,
which is unlike XID-indexed DRCs, the slotid DRC by its nature cannot
be overflowed. Through use of the sequenceid to identify
retransmitted requests, it is notable that the server does not need
to actually cache the request itself, reducing the storage
requirements of the DRC further. These new facilities makes it
practical to maintain all the required entries for an effective DRC.
The slotid and sequenceid therefore take over the traditional role of
the port number in the server DRC implementation, and the session
replaces the IP address. This approach is considerably more portable
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and completely robust - it is not subject to the frequent
reassignment of ports as clients reconnect over IP networks. In
addition, the RPC XID is not used in the reply cache, enhancing
robustness of the cache in the face of any rapid reuse of XIDs by the
client.
It is required to encode the slotid information into each request in
a way that does not violate the minor versioning rules of the NFSv4.0
specification. This is accomplished here by encoding it in a control
operation within each NFSv4.1 COMPOUND and CB_COMPOUND procedure.
The operation easily piggybacks within existing messages. The
implementation section of this document describes the specific
proposal.
In general, the receipt of a new sequenced request arriving on any
valid slot is an indication that the previous DRC contents of that
slot may be discarded. In order to further assist the server in slot
management, the client is required to use the lowest available slot
when issuing a new request. In this way, the server may be able to
retire additional entries.
However, in the case where the server is actively adjusting its
granted maximum request count to the client, it may not be able to
use receipt of the slotid to retire cache entries. The slotid used
in an incoming request may not reflect the server's current idea of
the client's session limit, because the request may have been sent
from the client before the update was received. Therefore, in the
downward adjustment case, the server may have to retain a number of
duplicate request cache entries at least as large as the old value,
until operation sequencing rules allow it to infer that the client
has seen its reply.
The SEQUENCE (and CB_SEQUENCE) operation also carries a "maxslot"
value which carries additional client slot usage information. The
client must always provide its highest-numbered outstanding slot
value in the maxslot argument, and the server may reply with a new
recognized value. The client should in all cases provide the most
conservative value possible, although it can be increased somewhat
above the actual instantaneous usage to maintain some minimum or
optimal level. This provides a way for the client to yield unused
request slots back to the server, which in turn can use the
information to reallocate resources. Obviously, maxslot can never be
zero, or the session would deadlock.
The server also provides a target maxslot value to the client, which
is an indication to the client of the maxslot the server wishes the
client to be using. This permits the server to withdraw (or add)
resources from a client that has been found to not be using them, in
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order to more fairly share resources among a varying level of demand
from other clients. The client must always comply with the server's
value updates, since they indicate newly established hard limits on
the client's access to session resources. However, because of
request pipelining, the client may have active requests in flight
reflecting prior values, therefore the server must not immediately
require the client to comply.
It is worthwhile to note that Sprite RPC [BW87] defined a "channel"
which in some ways is similar to the slotid proposed here. Sprite
RPC used channels to implement parallel request processing and
request/response cache retirement.
10.10.3. Resolving server callback races with sessions
It is possible for server callbacks to arrive at the client before
the reply from related forward channel operations. For example, a
client may have been granted a delegation to a file it has opened,
but the reply to the OPEN (informing the client of the granting of
the delegation) may be delayed in the network. If a conflicting
operation arrives at the server, it will recall the delegation using
the callback channel, which may be on a different TCP connection,
perhaps even a different network. If the callback request arrives
before the related reply, the client may reply to the server with an
error.
The presence of a session between client and server can be used to
alleviate this issue. When a session is in place, each client
request is uniquely identified by its { slotid, sequenceid } pair.
By the rules under which slot entries (duplicate request cache
entries) are retired, the server has knowledge whether the client has
"seen" each of the server's replies. The server can therefore
provide sufficient information to the client to allow it to
disambiguate between an erroneous or conflicting callback and a race
condition.
To implement this, the CB_SEQUENCE operation which begins each server
callback may optionally carry a related { slotid, sequenceid }
identifier. If the client finds this identifier to be currently
outstanding (the server's reply has not been seen by the client), it
can determine that the callback has raced the reply, and act
accordingly.
The client must not simply wait forever for the expected server reply
to arrive any of the session's operations channels, because it is
possible that they will be delayed indefinitely. However, it should
endeavor to wait for a period of time, and if the time expires it can
provide a more meaningful error such as NFS4ERR_DELAY.
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[[Comment.4: We need to consider the clients' options here, and
describe them... NFS4ERR_DELAY has been discussed as a legal reply
to CB_RECALL?]]
There are other scenarios under which callbacks may race replies,
among them pnfs layout recalls, described in Section 14.5.3
[[Comment.5: fill in the blanks w/others, etc...]]
Therefore, for each client operation which might result in some sort
of server callback, the server should "remember" the { slotid,
sequenceid } pair of the client request until the slotid retirement
rules allow the server to determine that the client has, in fact,
seen the server's reply. During this time, any recalls of the
associated object should carry these identifiers, for the benefit of
the client. After this time, it is not necessary for the server to
provide this information in related callbacks, since it is certain
that a race condition can no longer occur.
10.10.4. COMPOUND and CB_COMPOUND
Support for per-operation control can be piggybacked onto NFSv4
COMPOUNDs with full transparency, by placing such facilities into
their own, new operation, and placing this operation first in each
COMPOUND under the new NFSv4 minor protocol revision. The contents
of the operation would then apply to the entire COMPOUND.
Recall that the NFSv4 minor revision is contained within the COMPOUND
header, encoded prior to the COMPOUNDed operations. By simply
requiring that the new operation always be contained in NFSv4 minor
COMPOUNDs, the control protocol can piggyback perfectly with each
request and response.
In this way, the NFSv4 RDMA Extensions may stay in compliance with
the minor versioning requirements specified in section 10 of
[RFC3530].
Referring to section 13.1 of the same document, the proposed session-
enabled COMPOUND and CB_COMPOUND have the form:
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+-----+--------------+-----------+------------+-----------+----
| tag | minorversion | numops | control op | op + args | ...
| | (== 1) | (limited) | + args | |
+-----+--------------+-----------+------------+-----------+----
and the reply's structure is:
+------------+-----+--------+-------------------------------+--//
|last status | tag | numres | status + control op + results | //
+------------+-----+--------+-------------------------------+--//
//-----------------------+----
// status + op + results | ...
//-----------------------+----
The single control operation within each NFSv4.1 COMPOUND defines the
context and operational session parameters which govern that COMPOUND
request and reply. Placing it first in the COMPOUND encoding is
required in order to allow its processing before other operations in
the COMPOUND.
10.10.5. eXternal Data Representation Efficiency
RDMA is a copy avoidance technology, and it is important to maintain
this efficiency when decoding received messages. Traditional XDR
implementations frequently use generated unmarshaling code to convert
objects to local form, incurring a data copy in the process (in
addition to subjecting the caller to recursive calls, etc). Often,
such conversions are carried out even when no size or byte order
conversion is necessary.
It is recommended that implementations pay close attention to the
details of memory referencing in such code. It is far more efficient
to inspect data in place, using native facilities to deal with word
size and byte order conversion into registers or local variables,
rather than formally (and blindly) performing the operation via
fetch, reallocate and store.
Of particular concern is the result of the READDIR operation, in
which such encoding abounds.
10.10.6. Effect of Sessions on Existing Operations
The use of a session replaces the use of the SETCLIENTID and
SETCLIENTID_CONFIRM operations, and allows certain simplification of
the RENEW and callback addressing mechanisms in the base protocol.
The cb_program and cb_location which are obtained by the server in
SETCLIENTID_CONFIRM must not be used by the server, because the
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NFSv4.1 client performs callback channel designation with
BIND_BACKCHANNEL. Therefore the SETCLIENTID and SETCLIENTID_CONFIRM
operations becomes obsolete when sessions are in use, and a server
should return an error to NFSv4.1 clients which might issue either
operation.
Another favorable result of the session is that the server is able to
avoid requiring the client to perform OPEN_CONFIRM operations. The
existence of a reliable and effective DRC means that the server will
be able to determine whether an OPEN request carrying a previously
known open_owner from a client is or is not a retransmission.
Because of this, the server no longer requires OPEN_CONFIRM to verify
whether the client is retransmitting an open request. This in turn
eliminates the server's reason for requesting OPEN_CONFIRM - the
server can simply replace any previous information on this
open_owner. Client OPEN operations are therefore streamlined,
reducing overhead and latency through avoiding the additional
OPEN_CONFIRM exchange.
Since the session carries the client liveness indication with it
implicitly, any request on a session associated with a given client
will renew that client's leases. Therefore the RENEW operation is
made unnecessary when a session is present, as any request (including
a SEQUENCE operation with or without additional NFSv4 operations)
performs its function. It is possible (though this proposal does not
make any recommendation) that the RENEW operation could be made
obsolete.
An interesting issue arises however if an error occurs on such a
SEQUENCE operation. If the SEQUENCE operation fails, perhaps due to
an invalid slotid or other non-renewal-based issue, the server may or
may not have performed the RENEW. In this case, the state of any
renewal is undefined, and the client should make no assumption that
it has been performed. In practice, this should not occur but even
if it did, it is expected the client would perform some sort of
recovery which would result in a new, successful, SEQUENCE operation
being run and the client assured that the renewal took place.
10.10.7. Authentication Efficiencies
NFSv4 requires the use of the RPCSEC_GSS ONC RPC security flavor
[RFC2203] to provide authentication, integrity, and privacy via
cryptography. The server dictates to the client the use of
RPCSEC_GSS, the service (authentication, integrity, or privacy), and
the specific GSS-API security mechanism that each remote procedure
call and result will use.
If the connection's integrity is protected by an additional means
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than RPCSEC_GSS, such as via IPsec, then the use of RPCSEC_GSS's
integrity service is nearly redundant (See the Security
Considerations section for more explanation of why it is "nearly" and
not completely redundant). Likewise, if the connection's privacy is
protected by additional means, then the use of both RPCSEC_GSS's
integrity and privacy services is nearly redundant.
Connection protection schemes, such as IPsec, are more likely to be
implemented in hardware than upper layer protocols like RPCSEC_GSS.
Hardware-based cryptography at the IPsec layer will be more efficient
than software-based cryptography at the RPCSEC_GSS layer.
When transport integrity can be obtained, it is possible for server
and client to downgrade their per-operation authentication, after an
appropriate exchange. This downgrade can in fact be as complete as
to establish security mechanisms that have zero cryptographic
overhead, effectively using the underlying integrity and privacy
services provided by transport.
Based on the above observations, a new GSS-API mechanism, called the
Channel Conjunction Mechanism [CCM], is being defined. The CCM works
by creating a GSS-API security context using as input a cookie that
the initiator and target have previously agreed to be a handle for
GSS-API context created previously over another GSS-API mechanism.
NFSv4.1 clients and servers should support CCM and they must use as
the cookie the handle from a successful RPCSEC_GSS context creation
over a non-CCM mechanism (such as Kerberos V5). The value of the
cookie will be equal to the handle field of the rpc_gss_init_res
structure from the RPCSEC_GSS specification.
The [CCM] Draft provides further discussion and examples.
10.11. Sessions Security Considerations
The NFSv4 minor version 1 retains all of existing NFSv4 security; all
security considerations present in NFSv4.0 apply to it equally.
Security considerations of any underlying RDMA transport are
additionally important, all the more so due to the emerging nature of
such transports. Examining these issues is outside the scope of this
draft.
When protecting a connection with RPCSEC_GSS, all data in each
request and response (whether transferred inline or via RDMA)
continues to receive this protection over RDMA fabrics [RPCRDMA].
However when performing data transfers via RDMA, RPCSEC_GSS
protection of the data transfer portion works against the efficiency
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which RDMA is typically employed to achieve. This is because such
data is normally managed solely by the RDMA fabric, and intentionally
is not touched by software. Therefore when employing RPCSEC_GSS
under CCM, and where integrity protection has been "downgraded", the
cooperation of the RDMA transport provider is critical to maintain
any integrity and privacy otherwise in place for the session. The
means by which the local RPCSEC_GSS implementation is integrated with
the RDMA data protection facilities are outside the scope of this
draft.
It is logical to use the same GSS context on a session's callback
channel as that used on its operations channel(s), particularly when
the connection is shared by both. The client must indicate to the
server:
- what security flavor(s) to use in the call back. A special
callback flavor might be defined for this.
- if the flavor is RPCSEC_GSS, then the client must have previously
created an RPCSEC_GSS session with the server. The client offers to
the server the the opaque handle<> value from the rpc_gss_init_res
structure, the window size of RPCSEC_GSS sequence numbers, and an
opaque gss_cb_handle.
This exchange can be performed as part of session and clientid
creation, and the issue warrants careful analysis before being
specified.
If the NFS client wishes to maintain full control over RPCSEC_GSS
protection, it may still perform its transfer operations using either
the inline or RDMA transfer model, or of course employ traditional
TCP stream operation. In the RDMA inline case, header padding is
recommended to optimize behavior at the server. At the client, close
attention should be paid to the implementation of RPCSEC_GSS
processing to minimize memory referencing and especially copying.
These are well-advised in any case!
The proposed session callback channel binding improves security over
that provided by NFSv4 for the callback channel. The connection is
client-initiated, and subject to the same firewall and routing checks
as the operations channel. The connection cannot be hijacked by an
attacker who connects to the client port prior to the intended
server. The connection is set up by the client with its desired
attributes, such as optionally securing with IPsec or similar. The
binding is fully authenticated before being activated.
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10.11.1. Authentication
Proper authentication of the principal which issues any session and
clientid in the proposed NFSv4.1 operations exactly follows the
similar requirement on client identifiers in NFSv4.0. It must not be
possible for a client to impersonate another by guessing its session
identifiers for NFSv4.1 operations, nor to bind a callback channel to
an existing session. To protect against this, NFSv4.0 requires
appropriate authentication and matching of the principal used. This
is discussed in Section 16, Security Considerations of [RFC3530].
The same requirement when using a session identifier applies to
NFSv4.1 here.
Going beyond NFSv4.0, the presence of a session associated with any
clientid may also be used to enhance NFSv4.1 security with respect to
client impersonation. In NFSv4.0, there are many operations which
carry no clientid, including in particular those which employ a
stateid argument. A rogue client which wished to carry out a denial
of service attack on another client could perform CLOSE, DELEGRETURN,
etc operations with that client's current filehandle, sequenceid and
stateid, after having obtained them from eavesdropping or other
approach. Locking and open downgrade operations could be similarly
attacked.
When an NFSv4.1 session is in place for any clientid, countermeasures
are easily applied through use of authentication by the server.
Because the sessionid is present in each request within a session,
the server may verify that the clientid is in fact originating from a
principal with the appropriate authenticated credentials, that the
sessionid belongs to the clientid, and that the stateid is valid in
these contexts. This is in general not possible with the affected
operations in NFSv4.0 due to the fact that the clientid is not
present in the requests.
In the event that authentication information is not available in the
incoming request, for example after a reconnection when the security
was previously downgraded using CCM, the server must require the
client re-establish the authentication in order that the server may
validate the other client-provided context, prior to executing any
operation. The sessionid, present in the newly retransmitted
request, combined with the retransmission detection enabled by the
NFSv4.1 duplicate request cache, are a convenient and reliable
context for the server to use for this contingency.
The server should take care to protect itself against denial of
service attacks in the creation of sessions and clientids. Clients
who connect and create sessions, only to disconnect and never use
them may leave significant state behind. (The same issue applies to
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NFSv4.0 with clients who may perform SETCLIENTID, then never perform
SETCLIENTID_CONFIRM.) Careful authentication coupled with resource
checks is highly recommended.
11. Directory Delegations
11.1. Introduction to Directory Delegations
The major addition to NFS version 4 in the area of caching is the
ability of the server to delegate certain responsibilities to the
client. When the server grants a delegation for a file to a client,
the client receives certain semantics with respect to the sharing of
that file with other clients. At OPEN, the server may provide the
client either a read or write delegation for the file. If the client
is granted a read delegation, it is assured that no other client has
the ability to write to the file for the duration of the delegation.
If the client is granted a write delegation, the client is assured
that no other client has read or write access to the file. This
reduces network traffic and server load by allowing the client to
perform certain operations on local file data and can also provide
stronger consistency for the local data.
Directory caching for the NFS version 4 protocol is similar to
previous versions. Clients typically cache directory information for
a duration determined by the client. At the end of a predefined
timeout, the client will query the server to see if the directory has
been updated. By caching attributes, clients reduce the number of
GETATTR calls made to the server to validate attributes.
Furthermore, frequently accessed files and directories, such as the
current working directory, have their attributes cached on the client
so that some NFS operations can be performed without having to make
an RPC call. By caching name and inode information about most
recently looked up entries in DNLC (Directory Name Lookup Cache),
clients do not need to send LOOKUP calls to the server every time
these files are accessed.
This caching approach works reasonably well at reducing network
traffic in many environments. However, it does not address
environments where there are numerous queries for files that do not
exist. In these cases of "misses", the client must make RPC calls to
the server in order to provide reasonable application semantics and
promptly detect the creation of new directory entries. Examples of
high miss activity are compilation in software development
environments. The current behavior of NFS limits its potential
scalability and wide-area sharing effectiveness in these types of
environments. Other distributed stateful filesystem architectures
such as AFS and DFS have proven that adding state around directory
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contents can greatly reduce network traffic in high miss
environments.
Delegation of directory contents is proposed as an extension for
NFSv4. Such an extension would provide similar traffic reduction
benefits as with file delegations. By allowing clients to cache
directory contents (in a read-only fashion) while being notified of
changes, the client can avoid making frequent requests to interrogate
the contents of slowly-changing directories, reducing network traffic
and improving client performance.
These extensions allow improved namespace cache consistency to be
achieved through delegations and synchronous recalls alone without
asking for notifications. In addition, if time-based consistency is
sufficient, asynchronous notifications can provide performance
benefits for the client, and possibly the server, under some common
operating conditions such as slowly-changing and/or very large
directories.
11.2. Directory Delegation Design (in brief)
A new operation GET_DIR_DELEGATION is used by the client to ask for a
directory delegation. The delegation covers directory attributes and
all entries in the directory. If either of these change the
delegation will be recalled synchronously. The operation causing the
recall will have to wait before the recall is complete. Any changes
to directory entry attributes will not cause the delegation to be
recalled.
In addition to asking for delegations, a client can also ask for
notifications for certain events. These events include changes to
directory attributes and/or its contents. If a client asks for
notification for a certain event, the server will notify the client
when that event occurs. This will not result in the delegation being
recalled for that client. The notifications are asynchronous and
provide a way of avoiding recalls in situations where a directory is
changing enough that the pure recall model may not be effective while
trying to allow the client to get substantial benefit. In the
absence of notifications, once the delegation is recalled the client
has to refresh its directory cache which might not be very efficient
for very large directories.
The delegation is read only and the client may not make changes to
the directory other than by performing NFSv4 operations that modify
the directory or the associated file attributes so that the server
has knowledge of these changes. In order to keep the client
namespace in sync with the server, the server will notify the client
holding the delegation of the changes made as a result. This is to
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avoid any subsequent GETATTR or READDIR calls to the server. If a
client holding the delegation makes any changes to the directory, the
delegation will not be recalled.
Delegations can be recalled by the server at any time. Normally, the
server will recall the delegation when the directory changes in a way
that is not covered by the notification, or when the directory
changes and notifications have not been requested.
Also if the server notices that handing out a delegation for a
directory is causing too many notifications to be sent out, it may
decide not to hand out a delegation for that directory or recall
existing delegations. If another client removes the directory for
which a delegation has been granted, the server will recall the
delegation.
Both the notification and recall operations need a callback path to
exist between the client and server. If the callback path does not
exist, then delegation can not be granted. Note that with the
session extensions [talpey] that should not be an issue. In the
absense of sessions, the server will have to establish a callback
path to the client to send callbacks.
11.3. Recommended Attributes in support of Directory Delegations
dir_notif_delay - notification delays on directory attributes
dir_entry_notif_delay - notification delays on child attributes
These attributes allow the client and server to negotiate the
frequency of notifications sent due to changes in attributes. These
attributes are returned as part of a GETATTR call on the directory.
The dir_notif_delay value covers all attribute changes to the
directory and the dir_entry_notif_delay covers all attribute changes
to any child in the directory.
These attributes are per directory. The client needs to get these
values by doing a GETATTR on the directory for which it wants
notifications. However these attributes are only required when the
client is interested in getting attribute notifications. For all
other types of notifications and delegation requests without
notifications, these attributes are not required.
When the client calls the GET_DIR_DELEGATION operation and asks for
attribute change notifications, it should request notification delays
that are no less than the values in the server-provided attributes.
If the client requests smaller delays, the server should not commit
to sending notifications for that change event.
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A value of zero for these attributes means the server will send the
notification as soon as the change occurs. It is not recommended to
set this value to zero since that can put a lot of burden on the
server.nfstime4 values that compute to negative values are illegal.
By granting a request for notifications, the server commits to
delaying notifications to that client by no more than the
notification delay which the client requested.
11.4. Delegation Recall
The server will recall the directory delegation by sending a callback
to the client. It will use the same callback procedure as used for
recalling file delegations. The server will recall the delegation
when the directory changes in a way that is not covered by the
notification. However the server will not recall the delegation if
attributes of an entry within the directory change. Also if the
server notices that handing out a delegation for a directory is
causing too many notifications to be sent out, it may decide not to
hand out a delegation for that directory. If another client tries to
remove the directory for which a delegation has been granted, the
server will recall the delegation.
The server will recall the delegation by sending a CB_RECALL callback
to the client. If the recall is done because of a directory changing
event, the request making that change will need to wait while the
client returns the delegation.
11.5. Delegation Recovery
Crash recovery has two main goals, avoiding the necessity of breaking
application guarantees with respect to locked files and delivery of
updates cached at the client. Neither of these applies to
directories protected by read delegations and notifications. Thus,
the client is required to establish a new delegation on a server or
client reboot.
12. Introduction
The NFSv4 protocol [6] specifies the interaction between a client
that accesses files and a server that provides access to files and is
responsible for coordinating access by multiple clients. As
described in the pNFS problem statement, this requires that all
access to a set of files exported by a single NFSv4 server be
performed by that server; at high data rates the server may become a
bottleneck.
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The parallel NFS (pNFS) extensions to NFSv4 allow data accesses to
bypass this bottleneck by permitting direct client access to the
storage devices containing the file data. When file data for a
single NFSv4 server is stored on multiple and/or higher throughput
storage devices (by comparison to the server's throughput
capability), the result can be significantly better file access
performance. The relationship among multiple clients, a single
server, and multiple storage devices for pNFS (server and clients
have access to all storage devices) is shown in this diagram:
+-----------+
|+-----------+ +-----------+
||+-----------+ | |
||| | NFSv4 + pNFS | |
+|| Clients |<------------------------------>| Server |
+| | | |
+-----------+ | |
||| +-----------+
||| |
||| |
||| Storage +-----------+ |
||| Protocol |+-----------+ |
||+----------------||+-----------+ Control|
|+-----------------||| | Protocol|
+------------------+|| Storage |------------+
+| Devices |
+-----------+
Figure 64
In this structure, the responsibility for coordination of file access
by multiple clients is shared among the server, clients, and storage
devices. This is in contrast to NFSv4 without pNFS extensions, in
which this is primarily the server's responsibility, some of which
can be delegated to clients under strictly specified conditions.
The pNFS extension to NFSv4 takes the form of new operations that
manage data location information called a "layout". The layout is
managed in a similar fashion as NFSv4 data delegations (e.g., they
are recallable and revocable). However, they are distinct
abstractions and are manipulated with new operations. When a client
holds a layout, it has rights to access the data directly using the
location information in the layout.
There are new attributes that describe general layout
characteristics. However, much of the required information cannot be
managed solely within the attribute framework, because it will need
to have a strictly limited term of validity, subject to invalidation
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by the server. This requires the use of new operations to obtain,
return, recall, and modify layouts, in addition to new attributes.
This document specifies both the NFSv4 extensions required to
distribute file access coordination between the server and its
clients and a NFSv4 file storage protocol that may be used to access
data stored on NFSv4 storage devices.
Storage protocols used to access a variety of other storage devices
are deliberately not specified here. These might include:
o Block/volume protocols such as iSCSI ([12]), and FCP ([13]). The
block/volume protocol support can be independent of the addressing
structure of the block/volume protocol used, allowing more than
one protocol to access the same file data and enabling
extensibility to other block/volume protocols.
o Object protocols such as OSD over iSCSI or Fibre Channel [14].
o Other storage protocols, including PVFS and other file systems
that are in use in HPC environments.
pNFS is designed to accommodate these protocols and be extensible to
new classes of storage protocols that may be of interest.
The distribution of file access coordination between the server and
its clients increases the level of responsibility placed on clients.
Clients are already responsible for ensuring that suitable access
checks are made to cached data and that attributes are suitably
propagated to the server. Generally, a misbehaving client that hosts
only a single-user can only impact files accessible to that single
user. Misbehavior by a client hosting multiple users may impact
files accessible to all of its users. NFSv4 delegations increase the
level of client responsibility as a client that carries out actions
requiring a delegation without obtaining that delegation will cause
its user(s) to see unexpected and/or incorrect behavior.
Some uses of pNFS extend the responsibility of clients beyond
delegations. In some configurations, the storage devices cannot
perform fine-grained access checks to ensure that clients are only
performing accesses within the bounds permitted to them by the pNFS
operations with the server (e.g., the checks may only be possible at
file system granularity rather than file granularity). In situations
where this added responsibility placed on clients creates
unacceptable security risks, pNFS configurations in which storage
devices cannot perform fine-grained access checks SHOULD NOT be used.
All pNFS server implementations MUST support NFSv4 access to any file
accessible via pNFS in order to provide an interoperable means of
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file access in such situations. See Section 15 on Security for
further discussion.
Finally, there are issues about how layouts interact with the
existing NFSv4 abstractions of data delegations and byte range
locking. These issues, and others, are also discussed here.
13. General Definitions
This protocol extension partitions the NFSv4 file system protocol
into two parts, the control path and the data path. The control path
is implemented by the extended (p)NFSv4 server. When the file system
being exported by (p)NFSv4 uses storage devices that are visible to
clients over the network, the data path may be implemented by direct
communication between the extended (p)NFSv4 file system client and
the storage devices. This leads to a few new terms used to describe
the protocol extension and some clarifications of existing terms.
13.1. Metadata Server
A pNFS "server" or "metadata server" is a server as defined by
RFC3530 [6], which additionally provides support of the pNFS minor
extension. When using the pNFS NFSv4 minor extension, the metadata
server may hold only the metadata associated with a file, while the
data can be stored on the storage devices. However, similar to
NFSv4, data may also be written through the metadata server. Note:
directory data is always accessed through the metadata server.
13.2. Client
A pNFS "client" is a client as defined by RFC3530 [6], with the
addition of supporting the pNFS minor extension server protocol and
with the addition of supporting at least one storage protocol for
performing I/O directly to storage devices.
13.3. Storage Device
This is a device, or server, that controls the file's data, but
leaves other metadata management up to the metadata server. A
storage device could be another NFS server, or an Object Storage
Device (OSD) or a block device accessed over a SAN (e.g., either
FiberChannel or iSCSI SAN). The goal of this extension is to allow
direct communication between clients and storage devices.
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13.4. Storage Protocol
This is the protocol between the pNFS client and the storage device
used to access the file data. Three following types have been
described: file protocols (e.g., NFSv4), object protocols (e.g.,
OSD), and block/volume protocols (e.g., based on SCSI-block
commands). These protocols are in turn realizable over a variety of
transport stacks. We anticipate there will be variations on these
storage protocols, including new protocols that are unknown at this
time or experimental in nature. The details of the storage protocols
will be described in other documents so that pNFS clients can be
written to use these storage protocols. Use of NFSv4 itself as a
file-based storage protocol is described in Section 16.
13.5. Control Protocol
This is a protocol used by the exported file system between the
server and storage devices. Specification of such protocols is
outside the scope of this draft. Such control protocols would be
used to control such activities as the allocation and deallocation of
storage and the management of state required by the storage devices
to perform client access control. The control protocol should not be
confused with protocols used to manage LUNs in a SAN and other
sysadmin kinds of tasks.
While the pNFS protocol allows for any control protocol, in practice
the control protocol is closely related to the storage protocol. For
example, if the storage devices are NFS servers, then the protocol
between the pNFS metadata server and the storage devices is likely to
involve NFS operations. Similarly, when object storage devices are
used, the pNFS metadata server will likely use iSCSI/OSD commands to
manipulate storage.
However, this document does not mandate any particular control
protocol. Instead, it just describes the requirements on the control
protocol for maintaining attributes like modify time, the change
attribute, and the end-of-file position.
13.6. Metadata
This is information about a file, like its name, owner, where it
stored, and so forth. The information is managed by the exported
file system server (metadata server). Metadata also includes lower-
level information like block addresses and indirect block pointers.
Depending the storage protocol, block-level metadata may or may not
be managed by the metadata server, but is instead managed by Object
Storage Devices or other servers acting as a storage device.
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13.7. Layout
A layout defines how a file's data is organized on one or more
storage devices. There are many possible layout types. They vary in
the storage protocol used to access the data, and in the aggregation
scheme that lays out the file data on the underlying storage devices.
Layouts are described in more detail below.
14. pNFS protocol semantics
This section describes the semantics of the pNFS protocol extension
to NFSv4; this is the protocol between the client and the metadata
server.
14.1. Definitions
This sub-section defines a number of terms necessary for describing
layouts and their semantics. In addition, it more precisely defines
how layouts are identified and how they can be composed of smaller
granularity layout segments.
14.1.1. Layout Types
A layout describes the mapping of a file's data to the storage
devices that hold the data. A layout is said to belong to a specific
"layout type" (see Section 1.2.17 for its RPC definition). The
layout type allows for variants to handle different storage protocols
(e.g., block/volume [11], object [10], and file [Section 16] layout
types). A metadata server, along with its control protocol, must
support at least one layout type. A private sub-range of the layout
type name space is also defined. Values from the private layout type
range can be used for internal testing or experimentation.
As an example, a file layout type could be an array of tuples (e.g.,
deviceID, file_handle), along with a definition of how the data is
stored across the devices (e.g., striping). A block/volume layout
might be an array of tuples that store <deviceID, block_number, block
count> along with information about block size and the file offset of
the first block. An object layout might be an array of tuples
<deviceID, objectID> and an additional structure (i.e., the
aggregation map) that defines how the logical byte sequence of the
file data is serialized into the different objects. Note, the actual
layouts are more complex than these simple expository examples.
This document defines a NFSv4 file layout type using a stripe-based
aggregation scheme (see Section 16). Adjunct specifications are
being drafted that precisely define other layout formats (e.g.,
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block/volume [11], and object [10] layouts) to allow interoperability
among clients and metadata servers.
14.1.2. Layout Iomode
The iomode indicates to the metadata server the client's intent to
perform either READs (only) or a mixture of I/O possibly containing
WRITEs as well as READs (i.e., READ/WRITE). For certain layout
types, it is useful for a client to specify this intent at LAYOUTGET
time. E.g., for block/volume based protocols, block allocation could
occur when a READ/WRITE iomode is specified. A special
LAYOUTIOMODE_ANY iomode is defined and can only be used for
LAYOUTRETURN and LAYOUTRECALL, not for LAYOUTGET. It specifies that
layouts pertaining to both READ and RW iomodes are being returned or
recalled, respectively.
A storage device may validate I/O with regards to the iomode; this is
dependent upon storage device implementation. Thus, if the client's
layout iomode differs from the I/O being performed the storage device
may reject the client's I/O with an error indicating a new layout
with the correct I/O mode should be fetched. E.g., if a client gets
a layout with a READ iomode and performs a WRITE to a storage device,
the storage device is allowed to reject that WRITE.
The iomode does not conflict with OPEN share modes or lock requests;
open mode checks and lock enforcement are always enforced, and are
logically separate from the pNFS layout level. As well, open modes
and locks are the preferred method for restricting user access to
data files. E.g., an OPEN of read, deny-write does not conflict with
a LAYOUTGET containing an iomode of READ/WRITE performed by another
client. Applications that depend on writing into the same file
concurrently may use byte range locking to serialize their accesses.
14.1.3. Layout Segments
Until this point, layouts have been defined in a fairly vague manner.
A layout is more precisely identified by the following tuple:
<ClientID, FH, layout type>; the FH refers to the FH of the file on
the metadata server. Note, layouts describe a file, not a byte-range
of a file.
Since a layout that describes an entire file may be very large, there
is a desire to manage layouts in smaller chunks that correspond to
byte-ranges of the file. For example, the entire layout need not be
returned, recalled, or committed. These chunks are called "layout
segments" and are further identified by the byte-range they
represent. Layout operations require the identification of the
layout segment (i.e., clientID, FH, layout type, and byte-range), as
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well as the iomode. This structure allows clients and metadata
servers to aggregate the results of layout operations into a singly
maintained layout.
It is important to define when layout segments overlap and/or
conflict with each other. For a layout segment to overlap another
layout segment both segments must be of the same layout type,
correspond to the same filehandle, and have the same iomode; in
addition, the byte-ranges of the segments must overlap. Layout
segments conflict, when they overlap and differ in the content of the
layout (i.e., the storage device/file mapping parameters differ).
Note, differing iomodes do not lead to conflicting layouts. It is
permissible for layout segments with different iomodes, pertaining to
the same byte range, to be held by the same client.
14.1.4. Device IDs
The "deviceID" is a short name for a storage device. In practice, a
significant amount of information may be required to fully identify a
storage device. Instead of embedding all that information in a
layout, a level of indirection is used. Layouts embed device IDs,
and a new operation (GETDEVICEINFO) is used to retrieve the complete
identity information about the storage device according to its layout
type. For example, the identity of a file server or object server
could be an IP address and port. The identity of a block device
could be a volume label. Due to multipath connectivity in a SAN
environment, agreement on a volume label is considered the reliable
way to locate a particular storage device.
The device ID is qualified by the layout type and unique per file
system (FSID). This allows different layout drivers to generate
device IDs without the need for co-ordination. In addition to
GETDEVICEINFO, another operation, GETDEVICELIST, has been added to
allow clients to fetch the mappings of multiple storage devices
attached to a metadata server.
Clients cannot expect the mapping between device ID and storage
device address to persist across server reboots, hence a client MUST
fetch new mappings on startup or upon detection of a metadata server
reboot unless it can revalidate its existing mappings. Not all
layout types support such revalidation, and the means of doing so is
layout specific. If data are reorganized from a storage device with
a given device ID to a different storage device (i.e., if the mapping
between storage device and data changes), the layout describing the
data MUST be recalled rather than assigning the new storage device to
the old device ID.
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14.1.5. Aggregation Schemes
Aggregation schemes can describe layouts like simple one-to-one
mapping, concatenation, and striping. A general aggregation scheme
allows nested maps so that more complex layouts can be compactly
described. The canonical aggregation type for this extension is
striping, which allows a client to access storage devices in
parallel. Even a one-to-one mapping is useful for a file server that
wishes to distribute its load among a set of other file servers.
14.2. Guarantees Provided by Layouts
Layouts delegate to the client the ability to access data out of
band. The layout guarantees the holder that the layout will be
recalled when the state encapsulated by the layout becomes invalid
(e.g., through some operation that directly or indirectly modifies
the layout) or, possibly, when a conflicting layout is requested, as
determined by the layout's iomode. When a layout is recalled, and
then returned by the client, the client retains the ability to access
file data with normal NFSv4 I/O operations through the metadata
server. Only the right to do I/O out-of-band is affected.
Holding a layout does not guarantee that a user of the layout has the
rights to access the data represented by the layout. All user access
rights MUST be obtained through the appropriate open, lock, and
access operations (i.e., those that would be used in the absence of
pNFS). However, if a valid layout for a file is not held by the
client, the storage device should reject all I/Os to that file's byte
range that originate from that client. In summary, layouts and
ordinary file access controls are independent. The act of modifying
a file for which a layout is held, does not necessarily conflict with
the holding of the layout that describes the file being modified.
However, with certain layout types (e.g., block/volume layouts), the
layout's iomode must agree with the type of I/O being performed.
Depending upon the layout type and storage protocol in use, storage
device access permissions may be granted by LAYOUTGET and may be
encoded within the type specific layout. If access permissions are
encoded within the layout, the metadata server must recall the layout
when those permissions become invalid for any reason; for example
when a file becomes unwritable or inaccessible to a client. Note,
clients are still required to perform the appropriate access
operations as described above (e.g., open and lock ops). The degree
to which it is possible for the client to circumvent these access
operations must be clearly addressed by the individual layout type
documents, as well as the consequences of doing so. In addition,
these documents must be clear about the requirements and non-
requirements for the checking performed by the server.
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If the pNFS metadata server supports mandatory byte range locks then
byte range locks must behave as specified by the NFSv4 protocol, as
observed by users of files. If a storage device is unable to
restrict access by a pNFS client who does not hold a required
mandatory byte range lock then the metadata server must not grant
layouts to a client, for that storage device, that permits any access
that conflicts with a mandatory byte range lock held by another
client. In this scenario, it is also necessary for the metadata
server to ensure that byte range locks are not granted to a client if
any other client holds a conflicting layout; in this case all
conflicting layouts must be recalled and returned before the lock
request can be granted. This requires the pNFS server to understand
the capabilities of its storage devices.
14.3. Getting a Layout
A client obtains a layout through a new operation, LAYOUTGET. The
metadata server will give out layouts of a particular type (e.g.,
block/volume, object, or file) and aggregation as requested by the
client. The client selects an appropriate layout type which the
server supports and the client is prepared to use. The layout
returned to the client may not line up exactly with the requested
byte range. A field within the LAYOUTGET request, "minlength",
specifies the minimum overlap that MUST exist between the requested
layout and the layout returned by the metadata server. The
"minlength" field should specify a size of at least one. A metadata
server may give-out multiple overlapping, non-conflicting layout
segments to the same client in response to a LAYOUTGET.
There is no implied ordering between getting a layout and performing
a file OPEN. For example, a layout may first be retrieved by placing
a LAYOUTGET operation in the same compound as the initial file OPEN.
Once the layout has been retrieved, it can be held across multiple
OPEN and CLOSE sequences.
The storage protocol used by the client to access the data on the
storage device is determined by the layout's type. The client needs
to select a "layout driver" that understands how to interpret and use
that layout. The API used by the client to talk to its drivers is
outside the scope of the pNFS extension. The storage protocol
between the client's layout driver and the actual storage is covered
by other protocols specifications such as iSCSI (block storage), OSD
(object storage) or NFS (file storage).
Although, the metadata server is in control of the layout for a file,
the pNFS client can provide hints to the server when a file is opened
or created about preferred layout type and aggregation scheme. The
pNFS extension introduces a LAYOUT_HINT attribute that the client can
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set at creation time to provide a hint to the server for new files.
It is suggested that this attribute be set as one of the initial
attributes to OPEN when creating a new file. Setting this attribute
separately, after the file has been created could make it difficult,
or impossible, for the server implementation to comply.
14.4. Committing a Layout
Due to the nature of the protocol, the file attributes, and data
location mapping (e.g., which offsets store data vs. store holes)
that exist on the metadata storage device may become inconsistent in
relation to the data stored on the storage devices; e.g., when WRITEs
occur before a layout has been committed (e.g., between a LAYOUTGET
and a LAYOUTCOMMIT). Thus, it is necessary to occasionally re-sync
this state and make it visible to other clients through the metadata
server.
The LAYOUTCOMMIT operation is responsible for committing a modified
layout segment to the metadata server. Note: the data should be
written and committed to the appropriate storage devices before the
LAYOUTCOMMIT occurs. Note, if the data is being written
asynchronously through the metadata server a COMMIT to the metadata
server is required to sync the data and make it visible on the
storage devices (see Section 14.6 for more details). The scope of
this operation depends on the storage protocol in use. For block/
volume-based layouts, it may require updating the block list that
comprises the file and committing this layout to stable storage.
While, for file-layouts it requires some synchronization of
attributes between the metadata and storage devices (i.e., mainly the
size attribute; EOF). It is important to note that the level of
synchronization is from the point of view of the client who issued
the LAYOUTCOMMIT. The updated state on the metadata server need only
reflect the state as of the client's last operation previous to the
LAYOUTCOMMIT, it need not reflect a globally synchronized state
(e.g., other clients may be performing, or may have performed I/O
since the client's last operation and the LAYOUTCOMMIT).
The control protocol is free to synchronize the attributes before it
receives a LAYOUTCOMMIT, however upon successful completion of a
LAYOUTCOMMIT, state that exists on the metadata server that describes
the file MUST be in sync with the state existing on the storage
devices that comprise that file as of the issuing client's last
operation. Thus, a client that queries the size of a file between a
WRITE to a storage device and the LAYOUTCOMMIT may observe a size
that does not reflects the actual data written.
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14.4.1. LAYOUTCOMMIT and mtime/atime/change
The change attribute and the modify/access times may be updated, by
the server, at LAYOUTCOMMIT time; since for some layout types, the
change attribute and atime/mtime can not be updated by the
appropriate I/O operation performed at a storage device. The
arguments to LAYOUTCOMMIT allow the client to provide suggested
access and modify time values to the server. Again, depending upon
the layout type, these client provided values may or may not be used.
The server should sanity check the client provided values before they
are used. For example, the server should ensure that time does not
flow backwards. According to the NFSv4 specification, The client
always has the option to set these attributes through an explicit
SETATTR operation.
As mentioned, for some layout protocols the change attribute and
mtime/atime may be updated at or after the time the I/O occurred
(e.g., if the storage device is able to communicate these attributes
to the metadata server). If, upon receiving a LAYOUTCOMMIT, the
server implementation is able to determine that the file did not
change since the last time the change attribute was updated (e.g., no
WRITEs or over-writes occurred), the implementation need not update
the change attribute; file-based protocols may have enough state to
make this determination or may update the change attribute upon each
file modification. This also applies for mtime and atime; if the
server implementation is able to determine that the file has not been
modified since the last mtime update, the server need not update
mtime at LAYOUTCOMMIT time. Once LAYOUTCOMMIT completes, the new
change attribute and mtime/atime should be visible if that file was
modified since the latest previous LAYOUTCOMMIT or LAYOUTGET.
14.4.2. LAYOUTCOMMIT and size
The file's size may be updated at LAYOUTCOMMIT time as well. The
LAYOUTCOMMIT operation contains an argument that indicates the last
byte offset to which the client wrote ("last_write_offset"). Note:
for this offset to be viewed as a file size it must be incremented by
one byte (e.g., a write to offset 0 would map into a file size of 1,
but the last write offset is 0). The metadata server may do one of
the following:
1. It may update the file's size based on the last write offset.
However, to the extent possible, the metadata server should
sanity check any value to which the file's size is going to be
set. E.g., it must not truncate the file based on the client
presenting a smaller last write offset than the file's current
size.
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2. If it has sufficient other knowledge of file size (e.g., by
querying the storage devices through the control protocol), it
may ignore the client provided argument and use the query-derived
value.
3. It may use the last write offset as a hint, subject to correction
when other information is available as above.
The method chosen to update the file's size will depend on the
storage device's and/or the control protocol's implementation. For
example, if the storage devices are block devices with no knowledge
of file size, the metadata server must rely on the client to set the
size appropriately. A new size flag and length are also returned in
the results of a LAYOUTCOMMIT. This union indicates whether a new
size was set, and to what length it was set. If a new size is set as
a result of LAYOUTCOMMIT, then the metadata server must reply with
the new size. As well, if the size is updated, the metadata server
in conjunction with the control protocol SHOULD ensure that the new
size is reflected by the storage devices immediately upon return of
the LAYOUTCOMMIT operation; e.g., a READ up to the new file size
should succeed on the storage devices (assuming no intervening
truncations). Again, if the client wants to explicitly zero-extend
or truncate a file, SETATTR must be used; it need not be used when
simply writing past EOF.
Since client layout holders may be unaware of changes made to the
file's size, through LAYOUTCOMMIT or SETATTR, by other clients, an
additional callback/notification has been added for pNFS.
CB_SIZECHANGED is a notification that the metadata server sends to
layout holders to notify them of a change in file size. This is
preferred over issuing CB_LAYOUTRECALL to each of the layout holders.
14.4.3. LAYOUTCOMMIT and layoutupdate
The LAYOUTCOMMIT operation contains a "layoutupdate" argument. This
argument is a layout type specific structure. The structure can be
used to pass arbitrary layout type specific information from the
client to the metadata server at LAYOUTCOMMIT time. For example, if
using a block/volume layout, the client can indicate to the metadata
server which reserved or allocated blocks it used and which it did
not. The "layoutupdate" structure need not be the same structure as
the layout returned by LAYOUTGET. The structure is defined by the
layout type and is opaque to LAYOUTCOMMIT.
14.5. Recalling a Layout
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14.5.1. Basic Operation
Since a layout protects a client's access to a file via a direct
client-storage-device path, a layout need only be recalled when it is
semantically unable to serve this function. Typically, this occurs
when the layout no longer encapsulates the true location of the file
over the byte range it represents. Any operation or action (e.g.,
server driven restriping or load balancing) that changes the layout
will result in a recall of the layout. A layout is recalled by the
CB_LAYOUTRECALL callback operation (see Section 24.9). This callback
can either recall a layout segment identified by a byte range, or all
the layouts associated with a file system (FSID). However, there is
no single operation to return all layouts associated with an FSID;
multiple layout segments may be returned in a single compound
operation. Section 14.5.3 discusses sequencing issues surrounding
the getting, returning, and recalling of layouts.
The iomode is also specified when recalling a layout or layout
segment. Generally, the iomode in the recall request must match the
layout, or segment, being returned; e.g., a recall with an iomode of
RW should cause the client to only return RW layout segments (not R
segments). However, a special LAYOUTIOMODE_ANY enumeration is
defined to enable recalling a layout of any type (i.e., the client
must return both read-only and read/write layouts).
A REMOVE operation may cause the metadata server to recall the layout
to prevent the client from accessing a non-existent file and to
reclaim state stored on the client. Since a REMOVE may be delayed
until the last close of the file has occurred, the recall may also be
delayed until this time. As well, once the file has been removed,
after the last reference, the client SHOULD no longer be able to
perform I/O using the layout (e.g., with file-based layouts an error
such as ESTALE could be returned).
Although, the pNFS extension does not alter the caching capabilities
of clients, or their semantics, it recognizes that some clients may
perform more aggressive write-behind caching to optimize the benefits
provided by pNFS. However, write-behind caching may impact the
latency in returning a layout in response to a CB_LAYOUTRECALL; just
as caching impacts DELEGRETURN with regards to data delegations.
Client implementations should limit the amount of dirty data they
have outstanding at any one time. Server implementations may fence
clients from performing direct I/O to the storage devices if they
perceive that the client is taking too long to return a layout once
recalled. A server may be able to monitor client progress by
watching client I/Os or by observing LAYOUTRETURNs of sub-portions of
the recalled layout. The server can also limit the amount of dirty
data to be flushed to storage devices by limiting the byte ranges
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covered in the layouts it gives out.
Once a layout has been returned, the client MUST NOT issue I/Os to
the storage devices for the file, byte range, and iomode represented
by the returned layout. If a client does issue an I/O to a storage
device for which it does not hold a layout, the storage device SHOULD
reject the I/O.
14.5.2. Recall Callback Robustness
For simplicity, the discussion thus far has assumed that pNFS client
state for a file exactly matches the pNFS server state for that file
and client regarding layout ranges and permissions. This assumption
leads to the implicit assumption that any callback results in a
LAYOUTRETURN or set of LAYOUTRETURNs that exactly match the range in
the callback, since both client and server agree about the state
being maintained. However, it can be useful if this assumption does
not always hold. For example:
o It may be useful for clients to be able to discard layout
information without calling LAYOUTRETURN. If conflicts that
require callbacks are very rare, and a server can use a multi-file
callback to recover per-client resources (e.g., via a FSID recall,
or a multi-file recall within a single compound), the result may
be significantly less client-server pNFS traffic.
o It may be similarly useful for servers to enhance information
about what layout ranges are held by a client beyond what a client
actually holds. In the extreme, a server could manage conflicts
on a per-file basis, only issuing whole-file callbacks even though
clients may request and be granted sub-file ranges.
o As well, the synchronized state assumption is not robust to minor
errors. A more robust design would allow for divergence between
client and server and the ability to recover. It is vital that a
client not assign itself layout permissions beyond what the server
has granted and that the server not forget layout permissions that
have been granted in order to avoid errors. On the other hand, if
a server believes that a client holds a layout segment that the
client does not know about, it's useful for the client to be able
to issue the LAYOUTRETURN that the server is expecting in response
to a recall.
Thus, in light of the above, it is useful for a server to be able to
issue callbacks for layout ranges it has not granted to a client, and
for a client to return ranges it does not hold. A pNFS client must
always return layout segments that comprise the full range specified
by the recall. Note, the full recalled layout range need not be
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returned as part of a single operation, but may be returned in
segments. This allows the client to stage the flushing of dirty
data, layout commits, and returns. Also, it indicates to the
metadata server that the client is making progress.
In order to ensure client/server convergence on the layout state, the
final LAYOUTRETURN operation in a sequence of returns for a
particular recall, SHOULD specify the entire range being recalled,
even if layout segments pertaining to partial ranges were previously
returned. In addition, if the client holds no layout segment that
overlaps the range being recalled, the client should return the
NFS4ERR_NOMATCHING_LAYOUT error code. This allows the server to
update its view of the client's layout state.
14.5.3. Recall/Return Sequencing
As with other stateful operations, pNFS requires the correct
sequencing of layout operations. This proposal assumes that sessions
will precede or accompany pNFS into NFSv4.x and thus, pNFS will
require the use of sessions. If the sessions proposal does not
precede pNFS, then this proposal needs to be modified to provide for
the correct sequencing of pNFS layout operations. Also, this
specification is reliant on the sessions protocol to provide the
correct sequencing between regular operations and callbacks. It is
the server's responsibility to avoid inconsistencies regarding the
layouts it hands out and the client's responsibility to properly
serialize its layout requests.
One critical issue with operation sequencing concerns callbacks. The
protocol must defend against races between the reply to a LAYOUTGET
operation and a subsequent CB_LAYOUTRECALL. It MUST NOT be possible
for a client to process the CB_LAYOUTRECALL for a layout that it has
not received in a reply message to a LAYOUTGET.
The callback races section (Section 10.10.3) describes the sessions
mechanism for allowing the client to detect such situations in order
to not process such a CB_LAYOUTRECALL. The LAYOUTGET operation is in
this case the dependent operation which the server should reference
in any layout recall, if it remains active in the server's slot
table.
14.5.3.1. Client Side Considerations
Consider a pNFS client that has issued a LAYOUTGET and then receives
an overlapping recall callback for the same file. There are two
possibilities, which in the absence of a session, the client cannot
distinguish when the callback arrives:
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1. The server processed the LAYOUTGET before issuing the recall, so
the LAYOUTGET response is in flight, and must be waited for
because it may be carrying layout info that will need to be
returned to deal with the recall callback.
2. The server issued the callback before receiving the LAYOUTGET.
The server will not respond to the LAYOUTGET until the recall
callback is processed.
This can cause deadlock, as the client must wait for the LAYOUTGET
response before processing the recall in the first case, but that
response will not arrive until after the recall is processed in the
second case. In the presence of a session, the server will provide
the client with the { slotid , sequenceid } of any earlier LAYOUTGET
which remains unconfirmed at the server by the session slot usage
rules. This allows the client to disambiguate between the two cases,
in case 1, the server will provide the reference, whereas in case 2
it will not (because there is no dependent client operation).
Therefore, the action at the client will only require waiting in the
case that the client has not yet seen the sever's earlier reply to
the LAYOUTGET.
Without the session, this deadlock can be avoided by adhering to the
following requirements:
o A LAYOUTGET MUST be rejected with an error (i.e.,
NFS4ERR_RECALLCONFLICT) if there's an overlapping outstanding
recall callback to the same client
o When processing a recall, the client MUST wait for a response to
all conflicting outstanding LAYOUTGETs before performing any
RETURN that could be affected by any such response.
o The client SHOULD wait for responses to all operations required to
complete a recall before sending any LAYOUTGETs that would
conflict with the recall because the server is likely to return
errors for them.
Now the client can wait for the LAYOUTGET response, as it will be
received in both cases.
14.5.3.2. Server Side Considerations
Consider a related situation from the pNFS server's point of view.
The server has issued a recall callback and receives an overlapping
LAYOUTGET for the same file before the LAYOUTRETURN(s) that respond
to the recall callback. Again, there are two cases:
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1. The client issued the LAYOUTGET before processing the recall
callback.
2. The client issued the LAYOUTGET after processing the recall
callback, but it arrived before the LAYOUTRETURN that completed
that processing.
The simplest approach is to always reject the overlapping LAYOUTGET.
The client has two ways to avoid this result - it can issue the
LAYOUTGET as a subsequent element of a COMPOUND containing the
LAYOUTRETURN that completes the recall callback, or it can wait for
the response to that LAYOUTRETURN.
There is little a session can do to disambiguate between these two
cases, because both operations are independent of one another. They
are simply asynchronous events which crossed. The situation can even
occur if the session is configured to use a single connection for
both operations and callbacks.
This leads to a more general problem; in the absence of a callback if
a client issues concurrent overlapping LAYOUTGET and LAYOUTRETURN
operations, it is possible for the server to process them in either
order. Again, a client must take the appropriate precautions in
serializing its actions.
[ASIDE: HighRoad forbids a client from doing this, as the per-file
layout stateid will cause one of the two operations to be rejected
with a stale layout stateid. This approach is simpler and produces
better results by comparison to allowing concurrent operations, at
least for this sort of conflict case, because server execution of
operations in an order not anticipated by the client may produce
results that are not useful to the client (e.g., if a LAYOUTRETURN is
followed by a concurrent overlapping LAYOUTGET, but executed in the
other order, the client will not retain layout extents for the
overlapping range).]
14.6. Metadata Server Write Propagation
Asynchronous writes written through the metadata server may be
propagated lazily to the storage devices. For data written
asynchronously through the metadata server, a client performing a
read at the appropriate storage device is not guaranteed to see the
newly written data until a COMMIT occurs at the metadata server.
While the write is pending, reads to the storage device can give out
either the old data, the new data, or a mixture thereof. After
either a synchronous write completes, or a COMMIT is received (for
asynchronously written data), the metadata server must ensure that
storage devices give out the new data and that the data has been
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written to stable storage. If the server implements its storage in
any way such that it cannot obey these constraints, then it must
recall the layouts to prevent reads being done that cannot be handled
correctly.
14.7. Crash Recovery
Crash recovery is complicated due to the distributed nature of the
pNFS protocol. In general, crash recovery for layouts is similar to
crash recovery for delegations in the base NFSv4 protocol. However,
the client's ability to perform I/O without contacting the metadata
server introduces subtleties that must be handled correctly if file
system corruption is to be avoided.
14.7.1. Leases
The layout lease period plays a critical role in crash recovery.
Depending on the capabilities of the storage protocol, it is crucial
that the client is able to maintain an accurate layout lease timer to
ensure that I/Os are not issued to storage devices after expiration
of the layout lease period. In order for the client to do so, it
must know which operations renew a lease.
14.7.1.1. Lease Renewal
The current NFSv4 specification allows for implicit lease renewals to
occur upon receiving an I/O. However, due to the distributed pNFS
architecture, implicit lease renewals are limited to operations
performed at the metadata server; this includes I/O performed through
the metadata server. So, a client must not assume that READ and
WRITE I/O to storage devices implicitly renew lease state.
If sessions are required for pNFS, as has been suggested, then the
SEQUENCE operation is to be used to explicitly renew leases. It is
proposed that the SEQUENCE operation be extended to return all the
specific information that RENEW does, but not as an error as RENEW
returns it. Since, when using session, beginning each compound with
the SEQUENCE op allows renews to be performed without an additional
operation and without an additional request. Again, the client must
not rely on any operation to the storage devices to renew a lease.
Using the SEQUENCE operation for renewals, simplifies the client's
perception of lease renewal.
14.7.1.2. Client Lease Timer
Depending on the storage protocol and layout type in use, it may be
crucial that the client not issue I/Os to storage devices if the
corresponding layout's lease has expired. Doing so may lead to file
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system corruption if the layout has been given out and used by
another client. In order to prevent this, the client must maintain
an accurate lease timer for all layouts held. RFC3530 has the
following to say regarding the maintenance of a client lease timer:
...the client must track operations which will renew the lease
period. Using the time that each such request was sent and the
time that the corresponding reply was received, the client should
bound the time that the corresponding renewal could have occurred
on the server and thus determine if it is possible that a lease
period expiration could have occurred.
To be conservative, the client should start its lease timer based on
the time that the it issued the operation to the metadata server,
rather than based on the time of the response.
It is also necessary to take propagation delay into account when
requesting a renewal of the lease:
...the client should subtract it from lease times (e.g., if the
client estimates the one-way propagation delay as 200 msec, then
it can assume that the lease is already 200 msec old when it gets
it). In addition, it will take another 200 msec to get a response
back to the server. So the client must send a lock renewal or
write data back to the server 400 msec before the lease would
expire.
Thus, the client must be aware of the one-way propagation delay and
should issue renewals well in advance of lease expiration. Clients,
to the extent possible, should try not to issue I/Os that may extend
past the lease expiration time period. However, since this is not
always possible, the storage protocol must be able to protect against
the effects of inflight I/Os, as is discussed later.
14.7.2. Client Recovery
Client recovery for layouts works in much the same way as NFSv4
client recovery works for other lock/delegation state. When an NFSv4
client reboots, it will lose all information about the layouts that
it previously owned. There are two methods by which the server can
reclaim these resources and allow otherwise conflicting layouts to be
provided to other clients.
The first is through the expiry of the client's lease. If the client
recovery time is longer than the lease period, the client's lease
will expire and the server will know that state may be released. for
layouts the server may release the state immediately upon lease
expiry or it may allow the layout to persist awaiting possible lease
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revival, as long as there are no conflicting requests.
On the other hand, the client may recover in less time than it takes
for the lease period to expire. In such a case, the client will
contact the server through the standard SETCLIENTID protocol. The
server will find that the client's id matches the id of the previous
client invocation, but that the verifier is different. The server
uses this as a signal to release all the state associated with the
client's previous invocation.
14.7.3. Metadata Server Recovery
The server recovery case is slightly more complex. In general, the
recovery process again follows the standard NFSv4 recovery model: the
client will discover that the metadata server has rebooted when it
receives an unexpected STALE_STATEID or STALE_CLIENTID reply from the
server; it will then proceed to try to reclaim its previous
delegations during the server's recovery grace period. However,
layouts have a slightly different mechanism for reclaim. The problem
is that a client which uses LAYOUTGET to reclaim a layout might not
get the same layout it had previously. The range might be different
or it might get the same range but the content of the layout might be
different. For example, if using a block/volume-based layout, the
blocks provisionally assigned by the layout might be different, in
which case the client will have to write the corresponding blocks
again.
Instead of reclaiming a layout with LAYOUTGET, a client can attempt
to commit data written before the file server crash by setting a
reclaim bit on the LAYOUTCOMMIT operation. This should only be done
for data that the client has already written using a layout obtained
before the server restart. For data still dirty in the client
memory, the client should get a new layout segment after the server's
grace period has elapsed. Alternatively, the client can write that
data through the metadata server using the standard NFSv4 WRITE. In
the case that the client has written dirty data to a provisionally
allocated region of the layout, but was unable to commit the layout
changes for this data before the server rebooted, the client may be
unable to reliably re-read the data from the data storage devices in
order to write it again via the metadata server. In this case the
client needs to inform the metadata server that the layout has
changed, before the server has completed its recovery grace period
and starts allowing updates to the file-system. For this purpose,
the LAYOUTCOMMIT operation contains a "reclaim" field. During the
metadata server's recovery grace period (and only during the recovery
grace period) the client may send a LAYOUTCOMMIT request with the
"reclaim" field set to "true". This indicates that the client is
attempting to commit changes to the file layout that occurred prior
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to the reboot of the metadata server. The "layout update" field of
the request must contain the portion of the layout that the client
held prior to the metadata server reboot which covers the outstanding
writes. The metadata server is free to apply consistency checks on
the layout update provided by the client, and reject the request if
the checks fail. If the checks do not fail, then the server MUST
commit the changes to the file layout contained in the "layoutupdate"
field of the LAYOUTCOMMIT request, ensuring that the clients
outstanding writes are not lost.
During the recovery grace period the metadata server should apply the
standard approach to handling WRITE and LAYOUTGET requests. That is,
if the server can reliably determine that servicing such a request
will not conflict with an impending LAYOUTCOMMIT reclaim request, it
may choose to service the request. If the server is unable to offer
this guarantee, it MUST reject the request with status NFS4ERR_GRACE.
For a metadata server to provide simple, valid handling during the
grace period with respect to pNFS layouts, the easiest method is to
simply reject all non-reclaim pNFS requests and WRITE operations by
returning the NFS4ERR_GRACE error. However, depending on the storage
protocol and server implementation, the server may be able to
determine that a particular request is safe. For example, a server
may save provisional allocation mappings for each file to stable
storage, and use this information during the recovery grace period to
determine that a WRITE request is safe. Under such circumstances,
the WRITE request MAY be serviced. To re-iterate, for a server to
allow non-reclaim pNFS requests and WRITE operations to be serviced
during the recovery grace period, it MUST determine that the request
will not conflict with any subsequent LAYOUTCOMMIT with reclaim
request.
There is an important safety concern associated with layouts that
does not come into play in the standard NFSv4 case. If a standard
NFSv4 client makes use of a stale delegation, while reading, the
consequence could be to deliver stale data to an application. If
writing, using a stale delegation or a stale state stateid for an
open or lock would result in the rejection of the client's write with
the appropriate stale stateid error.
However, the pNFS layout enables the client to directly access the
file system storage; if this access is not properly managed by the
NFSv4 server the client can potentially corrupt the file system data
or metadata. Thus, it is vitally important that the client discover
that the metadata server has rebooted, and that the client stops
using stale layouts before the metadata server gives them away to
other clients. To ensure this, the client must be implemented so
that layouts are never used to access the storage after the client's
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lease timer has expired. It is crucial that clients have precise
knowledge of the lease periods of their layouts. For specific
details on lease renewal and client lease timers, see Section 14.7.1.
The prohibition on using stale layouts applies to all layout related
accesses, especially the flushing of dirty data to the storage
devices. If the client's lease timer expires because the client
could not contact the server for any reason, the client MUST
immediately stop using the layout until the server can be contacted
and the layout can be officially recovered or reclaimed. However,
this is only part of the solution. It is also necessary to deal with
the consequences of I/Os already in flight.
The issue of the effects of I/Os started before lease expiration and
possibly continuing through lease expiration is the responsibility of
the data storage protocol and as such is layout type specific. There
are two approaches the data storage protocol can take. The protocol
may adopt a global solution which prevents all I/Os from being
executed after the lease expiration and thus is safe against a client
who issues I/Os after lease expiration. This is the preferred
solution and the solution used by NFSv4 file based layouts (see
Section 16.6); as well, the object storage device protocol allows
storage to fence clients after lease expiration. Alternatively, the
storage protocol may rely on proper client operation and only deal
with the effects of lingering I/Os. These solutions may impact the
client layout-driver, the metadata server layout-driver, and the
control protocol.
14.7.4. Storage Device Recovery
Storage device crash recovery is mostly dependent upon the layout
type in use. However, there are a few general techniques a client
can use if it discovers a storage device has crashed while holding
asynchronously written, non-committed, data. First and foremost, it
is important to realize that the client is the only one who has the
information necessary to recover asynchronously written data; since,
it holds the dirty data and most probably nobody else does. Second,
the best solution is for the client to err on the side or caution and
attempt to re-write the dirty data through another path.
The client, rather than hold the asynchronously written data
indefinitely, is encouraged to, and can make sure that the data is
written by using other paths to that data. The client may write the
data to the metadata server, either synchronously or asynchronously
with a subsequent COMMIT. Once it does this, there is no need to
wait for the original storage device. In the event that the data
range to be committed is transferred to a different storage device,
as indicated in a new layout, the client may write to that storage
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device. Once the data has been committed at that storage device,
either through a synchronous write or through a commit to that
storage device (e.g., through the NFSv4 COMMIT operation for the
NFSv4 file layout), the client should consider the transfer of
responsibility for the data to the new server as strong evidence that
this is the intended and most effective method for the client to get
the data written. In either case, once the write is on stable
storage (through either the storage device or metadata server), there
is no need to continue either attempting to commit or attempting to
synchronously write the data to the original storage device or wait
for that storage device to become available. That storage device may
never be visible to the client again.
This approach does have a "lingering write" problem, similar to
regular NFSv4. Suppose a WRITE is issued to a storage device for
which no response is received. The client breaks the connection,
trying to re-establish a new one, and gets a recall of the layout.
The client issues the I/O for the dirty data through an alternative
path, for example, through the metadata server and it succeeds. The
client then goes on to perform additional writes that all succeed.
If at some time later, the original write to the storage device
succeeds, data inconsistency could result. The same problem can
occur in regular NFSv4. For example, a WRITE is held in a switch for
some period of time while other writes are issued and replied to, if
the original WRITE finally succeeds, the same issues can occur.
However, this is solved by sessions in NFSv4.x.
15. Security Considerations
The pNFS extension partitions the NFSv4 file system protocol into two
parts, the control path and the data path (i.e., storage protocol).
The control path contains all the new operations described by this
extension; all existing NFSv4 security mechanisms and features apply
to the control path. The combination of components in a pNFS system
(see Figure 64) is required to preserve the security properties of
NFSv4 with respect to an entity accessing data via a client,
including security countermeasures to defend against threats that
NFSv4 provides defenses for in environments where these threats are
considered significant.
In some cases, the security countermeasures for connections to
storage devices may take the form of physical isolation or a
recommendation not to use pNFS in an environment. For example, it is
currently infeasible to provide confidentiality protection for some
storage device access protocols to protect against eavesdropping; in
environments where eavesdropping on such protocols is of sufficient
concern to require countermeasures, physical isolation of the
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communication channel (e.g., via direct connection from client(s) to
storage device(s)) and/or a decision to forego use of pNFS (e.g., and
fall back to NFSv4) may be appropriate courses of action.
In full generality where communication with storage devices is
subject to the same threats as client-server communication, the
protocols used for that communication need to provide security
mechanisms comparable to those available via RPSEC_GSS for NFSv4.
Many situations in which pNFS is likely to be used will not be
subject to the overall threat profile for which NFSv4 is required to
provide countermeasures.
pNFS implementations MUST NOT remove NFSv4's access controls. The
combination of clients, storage devices, and the server are
responsible for ensuring that all client to storage device file data
access respects NFSv4 ACLs and file open modes. This entails
performing both of these checks on every access in the client, the
storage device, or both. If a pNFS configuration performs these
checks only in the client, the risk of a misbehaving client obtaining
unauthorized access is an important consideration in determining when
it is appropriate to use such a pNFS configuration. Such
configurations SHOULD NOT be used when client- only access checks do
not provide sufficient assurance that NFSv4 access control is being
applied correctly.
The following subsections describe security considerations
specifically applicable to each of the three major storage device
protocol types supported for pNFS.
[Requiring strict equivalence to NFSv4 security mechanisms is the
wrong approach. Will need to lay down a set of statements that each
protocol has to make starting with access check location/properties.]
15.1. File Layout Security
A NFSv4 file layout type is defined in Section 16; see Section 16.7
for additional security considerations and details. In summary, the
NFSv4 file layout type requires that all I/O access checks MUST be
performed by the storage devices, as defined by the NFSv4
specification. If another file layout type is being used, additional
access checks may be required. But in all cases, the access control
performed by the storage devices must be at least as strict as that
specified by the NFSv4 protocol.
15.2. Object Layout Security
The object storage protocol MUST implement the security aspects
described in version 1 of the T10 OSD protocol definition [14]. The
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remainder of this section gives an overview of the security mechanism
described in that standard. The goal is to give the reader a basic
understanding of the object security model. Any discrepancies
between this text and the actual standard are obviously to be
resolved in favor of the OSD standard.
The object storage protocol relies on a cryptographically secure
capability to control accesses at the object storage devices.
Capabilities are generated by the metadata server, returned to the
client, and used by the client as described below to authenticate
their requests to the Object Storage Device (OSD). Capabilities
therefore achieve the required access and open mode checking. They
allow the file server to define and check a policy (e.g., open mode)
and the OSD to check and enforce that policy without knowing the
details (e.g., user IDs and ACLs). Since capabilities are tied to
layouts, and since they are used to enforce access control, the
server should recall layouts and revoke capabilities when the file
ACL or mode changes in order to signal the clients.
Each capability is specific to a particular object, an operation on
that object, a byte range w/in the object, and has an explicit
expiration time. The capabilities are signed with a secret key that
is shared by the object storage devices (OSD) and the metadata
managers. clients do not have device keys so they are unable to forge
capabilities. The the following sketch of the algorithm should help
the reader understand the basic model.
LAYOUTGET returns
{CapKey = MAC<SecretKey>(CapArgs), CapArgs}
The client uses CapKey to sign all the requests it issues for that
object using the respective CapArgs. In other words, the CapArgs
appears in the request to the storage device, and that request is
signed with the CapKey as follows:
ReqMAC = MAC<CapKey>(Req, Nonceln)
The following is sent to the OSD: {CapArgs, Req, Nonceln, ReqMAC}.
The OSD uses the SecretKey it shares with the metadata server to
compare the ReqMAC the client sent with a locally computed
MAC<MAC<SecretKey>(CapArgs)>(Req, Nonceln)
and if they match the OSD assumes that the capabilities came from an
authentic metadata server and allows access to the object, as allowed
by the CapArgs. Therefore, if the server LAYOUTGET reply, holding
CapKey and CapArgs, is snooped by another client, it can be used to
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generate valid OSD requests (within the CapArgs access restriction).
To provide the required privacy requirements for the capabilities
returned by LAYOUTGET, the GSS-API can be used, e.g. by using a
session key known to the file server and to the client to encrypt the
whole layout or parts of it. Two general ways to provide privacy in
the absence of GSS-API that are independent of NFSv4 are either an
isolated network such as a VLAN or a secure channel provided by
IPsec.
15.3. Block/Volume Layout Security
As typically used, block/volume protocols rely on clients to enforce
file access checks since the storage devices are generally unaware of
the files they are storing and in particular are unaware of which
blocks belongs to which file. In such environments, the physical
addresses of blocks are exported to pNFS clients via layouts. An
alternative method of block/volume protocol use is for the storage
devices to export virtualized block addresses, which do reflect the
files to which blocks belong. These virtual block addresses are
exported to pNFS clients via layouts. This allows the storage device
to make appropriate access checks, while mapping virtual block
addresses to physical block addresses.
In environments where access control is important and client-only
access checks provide insufficient assurance of access control
enforcement (e.g., there is concern about a malicious of
malfunctioning client skipping the access checks) and where physical
block addresses are exported to clients, the storage devices will
generally be unable to compensate for these client deficiencies.
In such threat environments, block/volume protocols SHOULD NOT be
used with pNFS, unless the storage device is able to implement the
appropriate access checks, via use of virtualized block addresses, or
other means. NFSv4 without pNFS or pNFS with a different type of
storage protocol would be a more suitable means to access files in
such environments. Storage-device/protocol-specific methods (e.g.
LUN masking/mapping) may be available to prevent malicious or high-
risk clients from directly accessing storage devices.
16. The NFSv4 File Layout Type
This section describes the semantics and format of NFSv4 file-based
layouts.
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16.1. File Striping and Data Access
The file layout type describes a method for striping data across
multiple devices. The data for each stripe unit is stored within an
NFSv4 file located on a particular storage device. The structures
used to describe the stripe layout are as follows:
enum stripetype4 {
STRIPE_SPARSE = 1,
STRIPE_DENSE = 2
};
struct nfsv4_file_layouthint {
stripetype4 stripe_type;
length4 stripe_unit;
uint32_t stripe_width;
};
struct nfsv4_file_layout { /* Per data stripe */
pnfs_deviceid4 dev_id<>;
nfs_fh4 fh;
};
struct nfsv4_file_layouttype4 { /* Per file */
stripetype4 stripe_type;
length4 stripe_unit;
length4 file_size;
nfsv4_file_layout dev_list<>;
};
The file layout specifies an ordered array of <deviceID, filehandle>
tuples, as well as the stripe size, type of stripe layout (discussed
a little later), and the file's current size as of LAYOUTGET time.
The filehandle, "fh", identifies the file on a storage device
identified by "dev_id", that holds a particular stripe of the file.
The "dev_id" array can be used for multipathing and is discussed
further in Section 16.1.3. The stripe width is determined by the
stripe unit size multiplied by the number of devices in the dev_list.
The stripe held by <dev_id, fh> is determined by that tuples position
within the device list, "dev_list". For example, consider a dev_list
consisting of the following <dev_id, fh> pairs:
<(1,0x12), (2,0x13), (1,0x15)> and stripe_unit = 32KB
The stripe width is 32KB * 3 devices = 96KB. The first entry
specifies that on device 1 in the data file with filehandle 0x12
holds the first 32KB of data (and every 32KB stripe beginning where
the file's offset % 96KB == 0).
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Devices may be repeated multiple times within the device list array;
this is shown where storage device 1 holds both the first and third
stripe of data. Filehandles can only be repeated if a sparse stripe
type is used. Data is striped across the devices in the order listed
in the device list array in increments of the stripe size. A data
file stored on a storage device MUST map to a single file as defined
by the metadata server; i.e., data from two files as viewed by the
metadata server MUST NOT be stored within the same data file on any
storage device.
The "stripe_type" field specifies how the data is laid out within the
data file on a storage device. It allows for two different data
layouts: sparse and dense or packed. The stripe type determines the
calculation that must be made to map the client visible file offset
to the offset within the data file located on the storage device.
The layout hint structure is described in more detail in
Section 3.15. It is used, by the client, as by the FILE_LAYOUT_HINT
attribute to specify the type of layout to be used for a newly
created file.
16.1.1. Sparse and Dense Storage Device Data Layouts
The stripe_type field allows for two storage device data file
representations. Example sparse and dense storage device data
layouts are illustrated below:
Sparse file-layout (stripe_unit = 4KB)
------------------
Is represented by the following file layout on the storage devices:
Offset ID:0 ID:1 ID:2
0 +--+ +--+ +--+ +--+ indicates a
|//| | | | | |//| stripe that
4KB +--+ +--+ +--+ +--+ contains data
| | |//| | |
8KB +--+ +--+ +--+
| | | | |//|
12KB +--+ +--+ +--+
|//| | | | |
16KB +--+ +--+ +--+
| | |//| | |
+--+ +--+ +--+
The sparse file-layout has holes for the byte ranges not exported by
that storage device. This allows clients to access data using the
real offset into the file, regardless of the storage device's
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position within the stripe. However, if a client writes to one of
the holes (e.g., offset 4-12KB on device 1), then an error MUST be
returned by the storage device. This requires that the storage
device have knowledge of the layout for each file.
When using a sparse layout, the offset into the storage device data
file is the same as the offset into the main file.
Dense/packed file-layout (stripe_unit = 4KB)
------------------------
Is represented by the following file layout on the storage devices:
Offset ID:0 ID:1 ID:2
0 +--+ +--+ +--+
|//| |//| |//|
4KB +--+ +--+ +--+
|//| |//| |//|
8KB +--+ +--+ +--+
|//| |//| |//|
12KB +--+ +--+ +--+
|//| |//| |//|
16KB +--+ +--+ +--+
|//| |//| |//|
+--+ +--+ +--+
The dense or packed file-layout does not leave holes on the storage
devices. Each stripe unit is spread across the storage devices. As
such, the storage devices need not know the file's layout since the
client is allowed to write to any offset.
The calculation to determine the byte offset within the data file for
dense storage device layouts is:
stripe_width = stripe_unit * N; where N = |dev_list|
dev_offset = floor(file_offset / stripe_width) * stripe_unit +
file_offset % stripe_unit
Regardless of the storage device data file layout, the calculation to
determine the index into the device array is the same:
dev_idx = floor(file_offset / stripe_unit) mod N
Section 16.5 describe the semantics for dealing with reads to holes
within the striped file. This is of particular concern, since each
individual component stripe file (i.e., the component of the striped
file that lives on a particular storage device) may be of different
length. Thus, clients may experience 'short' reads when reading off
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the end of one of these component files.
16.1.2. Metadata and Storage Device Roles
In many cases, the metadata server and the storage device will be
separate pieces of physical hardware. The specification text is
written as if that were always case. However, it can be the case
that the same physical hardware is used to implement both a metadata
and storage device and in this case, the specification text's
references to these two entities are to be understood as referring to
the same physical hardware implementing two distinct roles and it is
important that it be clearly understood on behalf of which role the
hardware is executing at any given time.
Two sub-cases can be distinguished. In the first sub-case, the same
physical hardware is used to implement both a metadata and data
server in which each role is addressed through a distinct network
interface (e.g., IP addresses for the metadata server and storage
device are distinct). As long as the storage device address is
obtained from the layout and is distinct from the metadata server's
address, using the device ID therein to obtain the appropriate
storage device address, it is always clear, for any given request, to
what role it is directed, based on the destination IP address.
However, it may also be the case that even though the metadata server
and storage device are distinct from one client's point of view, the
roles may be reversed according to another client's point of view.
For example, in the cluster file system model a metadata server to
one client, may be a storage device to another client. Thus, it is
safer to always mark the filehandle so that operations addressed to
storage devices can be distinguished.
The second sub-case is where both the metadata and storage device
have the same network address. This requires us to make the
distinction as to which role each request is directed, on a another
basis. Since the network address is the same, the request is
understood as being directed at one or the other, based on the
filehandle of the first current filehandle value for the request. If
the first current file handle is one derived from a layout (i.e., it
is specified within the layout) (and it is recommended that these be
distinguishable), then the request is to be considered as executed by
a storage device. Otherwise, the operation is to be understood as
executed by the metadata server.
If a current filehandle is set that is inconsistent with the role to
which it is directed, then the error NFS4ERR_BADHANDLE should result.
For example, if a request is directed at the storage device, because
the first current handle is from a layout, any attempt to set the
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current filehandle to be a value not from a layout should be
rejected. Similarly, if the first current file handle was for a
value not from a layout, a subsequent attempt to set the current file
handle to a value obtained from a layout should be rejected.
16.1.3. Device Multipathing
The NFSv4 file layout supports multipathing to 'equivalent' devices.
Device-level multipathing is primarily of use in the case of a data
server failure --- it allows the client to switch to another storage
device that is exporting the same data stripe, without having to
contact the metadata server for a new layout.
To support device multipathing, an array of device IDs is encoded
within the data stripe portion of the file's layout. This array
represents an ordered list of devices where the first element has the
highest priority. Each device in the list MUST be 'equivalent' to
every other device in the list and each device must be attempted in
the order specified.
Equivalent devices MUST export the same system image (e.g., the
stateids and filehandles that they use are the same) and must provide
the same consistency guarantees. Two equivalent storage devices must
also have sufficient connections to the storage, such that writing to
one storage device is equivalent to writing to another, this also
applies to reading. Also, if multiple copies of the same data exist,
reading from one must provide access to all existing copies. As
such, it is unlikely that multipathing will provide additional
benefit in the case of an I/O error.
[NOTE: the error cases in which a client is expected to attempt an
equivalent storage device should be specified.]
16.1.4. Operations Issued to Storage Devices
Clients MUST use the filehandle described within the layout when
accessing data on the storage devices. When using the layout's
filehandle, the client MUST only issue READ, WRITE, PUTFH, COMMIT,
and NULL operations to the storage device associated with that
filehandle. If a client issues an operation other than those
specified above, using the filehandle and storage device listed in
the client's layout, that storage device SHOULD return an error to
the client. The client MUST follow the instruction implied by the
layout (i.e., which filehandles to use on which devices). As
described in Section 14.2, a client MUST NOT issue I/Os to storage
devices for which it does not hold a valid layout. The storage
devices may reject such requests.
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GETATTR and SETATTR MUST be directed to the metadata server. In the
case of a SETATTR of the size attribute, the control protocol is
responsible for propagating size updates/truncations to the storage
devices. In the case of extending WRITEs to the storage devices, the
new size must be visible on the metadata server once a LAYOUTCOMMIT
has completed (see Section 14.4.2). Section 16.5, describes the
mechanism by which the client is to handle storage device file's that
do not reflect the metadata server's size.
16.2. Global Stateid Requirements
Note, there are no stateids returned embedded within the layout. The
client MUST use the stateid representing open or lock state as
returned by an earlier metadata operation (e.g., OPEN, LOCK), or a
special stateid to perform I/O on the storage devices, as in regular
NFSv4. Special stateid usage for I/O is subject to the NFSv4
protocol specification. The stateid used for I/O MUST have the same
effect and be subject to the same validation on storage device as it
would if the I/O was being performed on the metadata server itself in
the absence of pNFS. This has the implication that stateids are
globally valid on both the metadata and storage devices. This
requires the metadata server to propagate changes in lock and open
state to the storage devices, so that the storage devices can
validate I/O accesses. This is discussed further in Section 16.4.
Depending on when stateids are propagated, the existence of a valid
stateid on the storage device may act as proof of a valid layout.
[NOTE: a number of proposals have been made that have the possibility
of limiting the amount of validation performed by the storage device,
if any of these proposals are accepted or obtain consensus, the
global stateid requirement can be revisited.]
16.3. The Layout Iomode
The layout iomode need not used by the metadata server when servicing
NFSv4 file-based layouts, although in some circumstances it may be
useful to use. For example, if the server implementation supports
reading from read-only replicas or mirrors, it would be useful for
the server to return a layout enabling the client to do so. As such,
the client should set the iomode based on its intent to read or write
the data. The client may default to an iomode of READ/WRITE
(LAYOUTIOMODE_RW). The iomode need not be checked by the storage
devices when clients perform I/O. However, the storage devices SHOULD
still validate that the client holds a valid layout and return an
error if the client does not.
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16.4. Storage Device State Propagation
Since the metadata server, which handles lock and open-mode state
changes, as well as ACLs, may not be collocated with the storage
devices where I/O access are validated, as such, the server
implementation MUST take care of propagating changes of this state to
the storage devices. Once the propagation to the storage devices is
complete, the full effect of those changes must be in effect at the
storage devices. However, some state changes need not be propagated
immediately, although all changes SHOULD be propagated promptly.
These state propagations have an impact on the design of the control
protocol, even though the control protocol is outside of the scope of
this specification. Immediate propagation refers to the synchronous
propagation of state from the metadata server to the storage
device(s); the propagation must be complete before returning to the
client.
16.4.1. Lock State Propagation
Mandatory locks MUST be made effective at the storage devices before
the request that establishes them returns to the caller. Thus,
mandatory lock state MUST be synchronously propagated to the storage
devices. On the other hand, since advisory lock state is not used
for checking I/O accesses at the storage devices, there is no
semantic reason for propagating advisory lock state to the storage
devices. However, since all lock, unlock, open downgrades and
upgrades affect the sequence ID stored within the stateid, the
stateid changes which may cause difficulty if this state is not
propagated. Thus, when a client uses a stateid on a storage device
for I/O with a newer sequence number than the one the storage device
has, the storage device should query the metadata server and get any
pending updates to that stateid. This allows stateid sequence number
changes to be propagated lazily, on-demand.
[NOTE: With the reliance on the sessions protocol, there is no real
need for sequence ID portion of the stateid to be validated on I/O
accesses. It is proposed that the seq. ID checking is obsoleted.]
Since updates to advisory locks neither confer nor remove privileges,
these changes need not be propagated immediately, and may not need to
be propagated promptly. The updates to advisory locks need only be
propagated when the storage device needs to resolve a question about
a stateid. In fact, if byte-range locking is not mandatory (i.e., is
advisory) the clients are advised not to use the lock-based stateids
for I/O at all. The stateids returned by open are sufficient and
eliminate overhead for this kind of state propagation.
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16.4.2. Open-mode Validation
Open-mode validation MUST be performed against the open mode(s) held
by the storage devices. However, the server implementation may not
always require the immediate propagation of changes. Reduction in
access because of CLOSEs or DOWNGRADEs do not have to be propagated
immediately, but SHOULD be propagated promptly; whereas changes due
to revocation MUST be propagated immediately. On the other hand,
changes that expand access (e.g., new OPEN's and upgrades) don't have
to be propagated immediately but the storage device SHOULD NOT reject
a request because of mode issues without making sure that the upgrade
is not in flight.
16.4.3. File Attributes
Since the SETATTR operation has the ability to modify state that is
visible on both the metadata and storage devices (e.g., the size),
care must be taken to ensure that the resultant state across the set
of storage devices is consistent; especially when truncating or
growing the file.
As described earlier, the LAYOUTCOMMIT operation is used to ensure
that the metadata is synced with changes made to the storage devices.
For the file-based protocol, it is necessary to re-sync state such as
the size attribute, and the setting of mtime/atime. See Section 14.4
for a full description of the semantics regarding LAYOUTCOMMIT and
attribute synchronization. It should be noted, that by using a file-
based layout type, it is possible to synchronize this state before
LAYOUTCOMMIT occurs. For example, the control protocol can be used
to query the attributes present on the storage devices.
Any changes to file attributes that control authorization or access
as reflected by ACCESS calls or READs and WRITEs on the metadata
server, MUST be propagated to the storage devices for enforcement on
READ and WRITE I/O calls. If the changes made on the metadata server
result in more restrictive access permissions for any user, those
changes MUST be propagated to the storage devices synchronously.
Recall that the NFSv4 protocol [6] specifies that:
...since the NFS version 4 protocol does not impose any
requirement that READs and WRITEs issued for an open file have the
same credentials as the OPEN itself, the server still must do
appropriate access checking on the READs and WRITEs themselves.
This also includes changes to ACLs. The propagation of access right
changes due to changes in ACLs may be asynchronous only if the server
implementation is able to determine that the updated ACL is not more
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restrictive for any user specified in the old ACL. Due to the
relative infrequency of ACL updates, it is suggested that all changes
be propagated synchronously.
[NOTE: it has been suggested that the NFSv4 specification is in error
with regard to allowing principles other than those used for OPEN to
be used for file I/O. If changes within a minor version alter the
behavior of NFSv4 with regard to OPEN principals and stateids some
access control checking at the storage device can be made less
expensive. pNFS should be altered to take full advantage of these
changes.]
16.5. Storage Device Component File Size
A potential problem exists when a component data file on a particular
storage device is grown past EOF; the problem exists for both dense
and sparse layouts. Imagine the following scenario: a client creates
a new file (size == 0) and writes to byte 128KB; the client then
seeks to the beginning of the file and reads byte 100. The client
should receive 0s back as a result of the read. However, if the read
falls on a different storage device to the client's original write,
the storage device servicing the READ may still believe that the
file's size is at 0 and return no data with the EOF flag set. The
storage device can only return 0s if it knows that the file's size
has been extended. This would require the immediate propagation of
the file's size to all storage devices, which is potentially very
costly, instead, another approach as outlined below.
First, the file's size is returned within the layout by LAYOUTGET.
This size must reflect the latest size at the metadata server as set
by the most recent of either the last LAYOUTCOMMIT or SETATTR;
however, it may be more recent. Second, if a client performs a read
that is returned short (i.e., is fully within the file's size, but
the storage device indicates EOF and returns partial or no data), the
client must assume that it is a hole and substitute 0s for the data
not read up until its known local file size. If a client extends the
file, it must update its local file size. Third, if the metadata
server receives a SETATTR of the size or a LAYOUTCOMMIT that alters
the file's size, the metadata server must send out CB_SIZECHANGED
messages with the new size to clients holding layouts; it need not
send a notification to the client that performed the operation that
resulted in the size changing). Upon reception of the CB_SIZECHANGED
notification, clients must update their local size for that file. As
well, if a new file size is returned as a result to LAYOUTCOMMIT, the
client must update their local file size.
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16.6. Crash Recovery Considerations
As described in Section 14.7, the layout type specific storage
protocol is responsible for handling the effects of I/Os started
before lease expiration, extending through lease expiration. The
NFSv4 file layout type prevents all I/Os from being executed after
lease expiration, without relying on a precise client lease timer and
without requiring storage devices to maintain lease timers.
It works as follows. In the presence of sessions, each compound
begins with a SEQUENCE operation that contains the "clientID". On
the storage device, the clientID can be used to validate that the
client has a valid layout for the I/O being performed, if it does
not, the I/O is rejected. Before the metadata server takes any
action to invalidate a layout given out by a previous instance, it
must make sure that all layouts from that previous instance are
invalidated at the storage devices. Note: it is sufficient to
invalidate the stateids associated with the layout only if special
stateids are not being used for I/O at the storage devices, otherwise
the layout itself must be invalidated.
This means that a metadata server may not restripe a file until it
has contacted all of the storage devices to invalidate the layouts
from the previous instance nor may it give out locks that conflict
with locks embodied by the stateids associated with any layout from
the previous instance without either doing a specific invalidation
(as it would have to do anyway) or doing a global storage device
invalidation.
16.7. Security Considerations
The NFSv4 file layout type MUST adhere to the security considerations
outlined in Section 15. More specifically, storage devices must make
all of the required access checks on each READ or WRITE I/O as
determined by the NFSv4 protocol [6]. This impacts the control
protocol and the propagation of state from the metadata server to the
storage devices; see Section 16.4 for more details.
16.8. Alternate Approaches
Two alternate approaches exist for file-based layouts and the method
used by clients to obtain stateids used for I/O. Both approaches
embed stateids within the layout.
However, before examining these approaches it is important to
understand the distinction between clients and owners. Delegations
belong to clients, while locks (e.g., record and share reservations)
are held by owners which in turn belong to a specific client. As
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such, delegations can only protect against inter-client conflicts,
not intra-client conflicts. Layouts are held by clients and SHOULD
NOT be associated with state held by owners. Therefore, if stateids
used for data access are embedded within a layout, these stateids can
only act as delegation stateids, protecting against inter-client
conflicts; stateids pertaining to an owner can not be embedded within
the layout. This has the implication that the client MUST arbitrate
among all intra-client conflicts (e.g., arbitrating among lock
requests by different processes) before issuing pNFS operations.
Using the stateids stored within the layout, storage devices can only
arbitrate between clients (not owners).
The first alternate approach is to do away with global stateids,
stateids returned by OPEN/LOCK that are valid on the metadata server
and storage devices, and use only stateids embedded within the
layout. This approach has the drawback that the stateids used for
I/O access can not be validated against per owner state, since they
are only associated with the client holding the layout. It breaks
the semantics of tieing a stateid used for I/O to an open instance.
This has the implication that clients must delegate per owner lock
and open requests internally, rather than push the work onto the
storage devices. The storage devices can still arbitrate and enforce
inter-client lock and open state.
The second approach is a hybrid approach. This approach allows for
stateids to be embedded with the layout, but also allows for the
possibility of global stateids. If the stateid embedded within the
layout is a special stateid of all zeros, then the stateid referring
to the last successful OPEN/LOCK should be used. This approach is
recommended if it is decided that using NFSv4 as a control protocol
is required.
This proposal suggests the global stateid approach due to the cleaner
semantics it provides regarding the relationship between stateids
used for I/O and their corresponding open instance or lock state.
However, it does have a profound impact on the control protocol's
implementation and the state propagation that is required (as
described in Section 16.4).
17. Layouts and Aggregation
This section describes several aggregation schemes in a semi-formal
way to provide context for layout formats. These definitions will be
formalized in other protocols. However, the set of understood types
is part of this protocol in order to provide for basic
interoperability.
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The layout descriptions include (deviceID, objectID) tuples that
identify some storage object on some storage device. The addressing
formation associated with the deviceID is obtained with
GETDEVICEINFO. The interpretation of the objectID depends on the
storage protocol. The objectID could be a filehandle for an NFSv4
storage device. It could be a OSD object ID for an object server.
The layout for a block device generally includes additional block map
information to enumerate blocks or extents that are part of the
layout.
17.1. Simple Map
The data is located on a single storage device. In this case the
file server can act as the front end for several storage devices and
distribute files among them. Each file is limited in its size and
performance characteristics by a single storage device. The simple
map consists of (deviceID, objectID).
17.2. Block Extent Map
The data is located on a LUN in the SAN. The layout consists of an
array of (deviceID, blockID, offset, length) tuples. Each entry
describes a block extent.
17.3. Striped Map (RAID 0)
The data is striped across storage devices. The parameters of the
stripe include the number of storage devices (N) and the size of each
stripe unit (U). A full stripe of data is N * U bytes. The stripe
map consists of an ordered list of (deviceID, objectID) tuples and
the parameter value for U. The first stripe unit (the first U bytes)
are stored on the first (deviceID, objectID), the second stripe unit
on the second (deviceID, objectID) and so forth until the first
complete stripe. The data layout then wraps around so that byte
(N*U) of the file is stored on the first (deviceID, objectID) in the
list, but starting at offset U within that object. The striped
layout allows a client to read or write to the component objects in
parallel to achieve high bandwidth.
The striped map for a block device would be slightly different. The
map is an ordered list of (deviceID, blockID, blocksize), where the
deviceID is rotated among a set of devices to achieve striping.
17.4. Replicated Map
The file data is replicated on N storage devices. The map consists
of N (deviceID, objectID) tuples. When data is written using this
map, it should be written to N objects in parallel. When data is
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read, any component object can be used.
This map type is controversial because it highlights the issues with
error recovery. Those issues get interesting with any scheme that
employs redundancy. The handling of errors (e.g., only a subset of
replicas get updated) is outside the scope of this protocol
extension. Instead, it is a function of the storage protocol and the
metadata control protocol.
17.5. Concatenated Map
The map consists of an ordered set of N (deviceID, objectID, size)
tuples. Each successive tuple describes the next segment of the
file.
17.6. Nested Map
The nested map is used to compose more complex maps out of simpler
ones. The map format is an ordered set of M sub-maps, each submap
applies to a byte range within the file and has its own type such as
the ones introduced above. Any level of nesting is allowed in order
to build up complex aggregation schemes.
18. Minor Versioning
To address the requirement of an NFS protocol that can evolve as the
need arises, the NFS version 4 protocol contains the rules and
framework to allow for future minor changes or versioning.
The base assumption with respect to minor versioning is that any
future accepted minor version must follow the IETF process and be
documented in a standards track RFC. Therefore, each minor version
number will correspond to an RFC. Minor version zero of the NFS
version 4 protocol is represented by this RFC. The COMPOUND
procedure will support the encoding of the minor version being
requested by the client.
The following items represent the basic rules for the development of
minor versions. Note that a future minor version may decide to
modify or add to the following rules as part of the minor version
definition.
1. Procedures are not added or deleted
To maintain the general RPC model, NFS version 4 minor versions
will not add to or delete procedures from the NFS program.
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2. Minor versions may add operations to the COMPOUND and
CB_COMPOUND procedures.
The addition of operations to the COMPOUND and CB_COMPOUND
procedures does not affect the RPC model.
* Minor versions may append attributes to GETATTR4args,
bitmap4, and GETATTR4res.
This allows for the expansion of the attribute model to allow
for future growth or adaptation.
* Minor version X must append any new attributes after the last
documented attribute.
Since attribute results are specified as an opaque array of
per-attribute XDR encoded results, the complexity of adding
new attributes in the midst of the current definitions will
be too burdensome.
3. Minor versions must not modify the structure of an existing
operation's arguments or results.
Again the complexity of handling multiple structure definitions
for a single operation is too burdensome. New operations should
be added instead of modifying existing structures for a minor
version.
This rule does not preclude the following adaptations in a minor
version.
* adding bits to flag fields such as new attributes to
GETATTR's bitmap4 data type
* adding bits to existing attributes like ACLs that have flag
words
* extending enumerated types (including NFS4ERR_*) with new
values
4. Minor versions may not modify the structure of existing
attributes.
5. Minor versions may not delete operations.
This prevents the potential reuse of a particular operation
"slot" in a future minor version.
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6. Minor versions may not delete attributes.
7. Minor versions may not delete flag bits or enumeration values.
8. Minor versions may declare an operation as mandatory to NOT
implement.
Specifying an operation as "mandatory to not implement" is
equivalent to obsoleting an operation. For the client, it means
that the operation should not be sent to the server. For the
server, an NFS error can be returned as opposed to "dropping"
the request as an XDR decode error. This approach allows for
the obsolescence of an operation while maintaining its structure
so that a future minor version can reintroduce the operation.
1. Minor versions may declare attributes mandatory to NOT
implement.
2. Minor versions may declare flag bits or enumeration values
as mandatory to NOT implement.
9. Minor versions may downgrade features from mandatory to
recommended, or recommended to optional.
10. Minor versions may upgrade features from optional to recommended
or recommended to mandatory.
11. A client and server that support minor version X must support
minor versions 0 (zero) through X-1 as well.
12. No new features may be introduced as mandatory in a minor
version.
This rule allows for the introduction of new functionality and
forces the use of implementation experience before designating a
feature as mandatory.
13. A client MUST NOT attempt to use a stateid, filehandle, or
similar returned object from the COMPOUND procedure with minor
version X for another COMPOUND procedure with minor version Y,
where X != Y.
19. Internationalization
The primary issue in which NFS version 4 needs to deal with
internationalization, or I18N, is with respect to file names and
other strings as used within the protocol. The choice of string
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representation must allow reasonable name/string access to clients
which use various languages. The UTF-8 encoding of the UCS as
defined by ISO10646 [7] allows for this type of access and follows
the policy described in "IETF Policy on Character Sets and
Languages", RFC2277 [8].
[RFC-XXX-stringprep-XXX], otherwise know as "stringprep", documents a
framework for using Unicode/UTF-8 in networking protocols, so as "to
increase the likelihood that string input and string comparison work
in ways that make sense for typical users throughout the world." A
protocol must define a profile of stringprep "in order to fully
specify the processing options." The remainder of this
Internationalization section defines the NFS version 4 stringprep
profiles. Much of terminology used for the remainder of this section
comes from stringprep.
There are three UTF-8 string types defined for NFS version 4:
utf8str_cs, utf8str_cis, and utf8str_mixed. Separate profiles are
defined for each. Each profile defines the following, as required by
stringprep:
o The intended applicability of the profile
o The character repertoire that is the input and output to
stringprep (which is Unicode 3.2 for referenced version of
stringprep)
o The mapping tables from stringprep used (as described in section 3
of stringprep)
o Any additional mapping tables specific to the profile
o The Unicode normalization used, if any (as described in section 4
of stringprep)
o The tables from stringprep listing of characters that are
prohibited as output (as described in section 5 of stringprep)
o The bidirectional string testing used, if any (as described in
section 6 of stringprep)
o Any additional characters that are prohibited as output specific
to the profile
Stringprep discusses Unicode characters, whereas NFS version 4
renders UTF-8 characters. Since there is a one to one mapping from
UTF-8 to Unicode, where ever the remainder of this document refers to
to Unicode, the reader should assume UTF-8.
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Much of the text for the profiles comes from [RFC-XXX-nameprep-XXX].
19.1. Stringprep profile for the utf8str_cs type
Every use of the utf8str_cs type definition in the NFS version 4
protocol specification follows the profile named nfs4_cs_prep.
19.1.1. Intended applicability of the nfs4_cs_prep profile
The utf8str_cs type is a case sensitive string of UTF-8 characters.
Its primary use in NFS Version 4 is for naming components and
pathnames. Components and pathnames are stored on the server's
filesystem. Two valid distinct UTF-8 strings might be the same after
processing via the utf8str_cs profile. If the strings are two names
inside a directory, the NFS version 4 server will need to either:
o disallow the creation of a second name if it's post processed form
collides with that of an existing name, or
o allow the creation of the second name, but arrange so that after
post processing, the second name is different than the post
processed form of the first name.
19.1.2. Character repertoire of nfs4_cs_prep
The nfs4_cs_prep profile uses Unicode 3.2, as defined in stringprep's
Appendix A.1
19.1.3. Mapping used by nfs4_cs_prep
The nfs4_cs_prep profile specifies mapping using the following tables
from stringprep:
Table B.1
Table B.2 is normally not part of the nfs4_cs_prep profile as it is
primarily for dealing with case-insensitive comparisons. However, if
the NFS version 4 file server supports the case_insensitive
filesystem attribute, and if case_insensitive is true, the NFS
version 4 server MUST use Table B.2 (in addition to Table B1) when
processing utf8str_cs strings, and the NFS version 4 client MUST
assume Table B.2 (in addition to Table B.1) are being used.
If the case_preserving attribute is present and set to false, then
the NFS version 4 server MUST use table B.2 to map case when
processing utf8str_cs strings. Whether the server maps from lower to
upper case or the upper to lower case is an implementation
dependency.
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19.1.4. Normalization used by nfs4_cs_prep
The nfs4_cs_prep profile does not specify a normalization form. A
later revision of this specification may specify a particular
normalization form. Therefore, the server and client can expect that
they may receive unnormalized characters within protocol requests and
responses. If the operating environment requires normalization, then
the implementation must normalize utf8str_cs strings within the
protocol before presenting the information to an application (at the
client) or local filesystem (at the server).
19.1.5. Prohibited output for nfs4_cs_prep
The nfs4_cs_prep profile specifies prohibiting using the following
tables from stringprep:
Table C.3
Table C.4
Table C.5
Table C.6
Table C.7
Table C.8
Table C.9
19.1.6. Bidirectional output for nfs4_cs_prep
The nfs4_cs_prep profile does not specify any checking of
bidirectional strings.
19.2. Stringprep profile for the utf8str_cis type
Every use of the utf8str_cis type definition in the NFS version 4
protocol specification follows the profile named nfs4_cis_prep.
19.2.1. Intended applicability of the nfs4_cis_prep profile
The utf8str_cis type is a case insensitive string of UTF-8
characters. Its primary use in NFS Version 4 is for naming NFS
servers.
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19.2.2. Character repertoire of nfs4_cis_prep
The nfs4_cis_prep profile uses Unicode 3.2, as defined in
stringprep's Appendix A.1
19.2.3. Mapping used by nfs4_cis_prep
The nfs4_cis_prep profile specifies mapping using the following
tables from stringprep:
Table B.1
Table B.2
19.2.4. Normalization used by nfs4_cis_prep
The nfs4_cis_prep profile specifies using Unicode normalization form
KC, as described in stringprep.
19.2.5. Prohibited output for nfs4_cis_prep
The nfs4_cis_prep profile specifies prohibiting using the following
tables from stringprep:
Table C.1.2
Table C.2.2
Table C.3
Table C.4
Table C.5
Table C.6
Table C.7
Table C.8
Table C.9
19.2.6. Bidirectional output for nfs4_cis_prep
The nfs4_cis_prep profile specifies checking bidirectional strings as
described in stringprep's section 6.
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19.3. Stringprep profile for the utf8str_mixed type
Every use of the utf8str_mixed type definition in the NFS version 4
protocol specification follows the profile named nfs4_mixed_prep.
19.3.1. Intended applicability of the nfs4_mixed_prep profile
The utf8str_mixed type is a string of UTF-8 characters, with a prefix
that is case sensitive, a separator equal to '@', and a suffix that
is fully qualified domain name. Its primary use in NFS Version 4 is
for naming principals identified in an Access Control Entry.
19.3.2. Character repertoire of nfs4_mixed_prep
The nfs4_mixed_prep profile uses Unicode 3.2, as defined in
stringprep's Appendix A.1
19.3.3. Mapping used by nfs4_cis_prep
For the prefix and the separator of a utf8str_mixed string, the
nfs4_mixed_prep profile specifies mapping using the following table
from stringprep:
Table B.1
For the suffix of a utf8str_mixed string, the nfs4_mixed_prep profile
specifies mapping using the following tables from stringprep:
Table B.1
Table B.2
19.3.4. Normalization used by nfs4_mixed_prep
The nfs4_mixed_prep profile specifies using Unicode normalization
form KC, as described in stringprep.
19.3.5. Prohibited output for nfs4_mixed_prep
The nfs4_mixed_prep profile specifies prohibiting using the following
tables from stringprep:
Table C.1.2
Table C.2.2
Table C.3
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Table C.4
Table C.5
Table C.6
Table C.7
Table C.8
Table C.9
19.3.6. Bidirectional output for nfs4_mixed_prep
The nfs4_mixed_prep profile specifies checking bidirectional strings
as described in stringprep's section 6.
19.4. UTF-8 Related Errors
Where the client sends an invalid UTF-8 string, the server should
return an NFS4ERR_INVAL (Table 5) error. This includes cases in
which inappropriate prefixes are detected and where the count
includes trailing bytes that do not constitute a full UCS character.
Where the client supplied string is valid UTF-8 but contains
characters that are not supported by the server as a value for that
string (e.g. names containing characters that have more than two
octets on a filesystem that supports Unicode characters only), the
server should return an NFS4ERR_BADCHAR (Table 5) error.
Where a UTF-8 string is used as a file name, and the filesystem,
while supporting all of the characters within the name, does not
allow that particular name to be used, the server should return the
error NFS4ERR_BADNAME (Table 5). This includes situations in which
the server filesystem imposes a normalization constraint on name
strings, but will also include such situations as filesystem
prohibitions of "." and ".." as file names for certain operations,
and other such constraints.
20. Error Definitions
NFS error numbers are assigned to failed operations within a compound
request. A compound request contains a number of NFS operations that
have their results encoded in sequence in a compound reply. The
results of successful operations will consist of an NFS4_OK status
followed by the encoded results of the operation. If an NFS
operation fails, an error status will be entered in the reply and the
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compound request will be terminated.
Protocol Error Definitions
+------------------------------+--------+---------------------------+
| Error | Number | Description |
+------------------------------+--------+---------------------------+
| NFS4_OK | 0 | Indicates the operation |
| | | completed successfully. |
| NFS4ERR_ACCESS | 13 | Permission denied. The |
| | | caller does not have the |
| | | correct permission to |
| | | perform the requested |
| | | operation. Contrast this |
| | | with NFS4ERR_PERM, which |
| | | restricts itself to owner |
| | | or privileged user |
| | | permission failures. |
| NFS4ERR_ATTRNOTSUPP | 10032 | An attribute specified is |
| | | not supported by the |
| | | server. Does not apply to |
| | | the GETATTR operation. |
| NFS4ERR_ADMIN_REVOKED | 10047 | Due to administrator |
| | | intervention, the |
| | | lockowner's record locks, |
| | | share reservations, and |
| | | delegations have been |
| | | revoked by the server. |
| NFS4ERR_BADCHAR | 10040 | A UTF-8 string contains a |
| | | character which is not |
| | | supported by the server |
| | | in the context in which |
| | | it being used. |
| NFS4ERR_BAD_COOKIE | 10003 | READDIR cookie is stale. |
| NFS4ERR_BADHANDLE | 10001 | Illegal NFS filehandle. |
| | | The filehandle failed |
| | | internal consistency |
| | | checks. |
| NFS4ERR_BADIOMODE | TDB | Layout iomode is invalid. |
| NFS4ERR_BADLAYOUT | TDB | Layout specified is |
| | | invalid. |
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| NFS4ERR_BADNAME | 10041 | A name string in a |
| | | request consists of valid |
| | | UTF-8 characters |
| | | supported by the server |
| | | but the name is not |
| | | supported by the server |
| | | as a valid name for |
| | | current operation. |
| NFS4ERR_BADOWNER | 10039 | An owner, owner_group, or |
| | | ACL attribute value can |
| | | not be translated to |
| | | local representation. |
| NFS4ERR_BADTYPE | 10007 | An attempt was made to |
| | | create an object of a |
| | | type not supported by the |
| | | server. |
| NFS4ERR_BAD_RANGE | 10042 | The range for a LOCK, |
| | | LOCKT, or LOCKU operation |
| | | is not appropriate to the |
| | | allowable range of |
| | | offsets for the server. |
| NFS4ERR_BAD_SEQID | 10026 | The sequence number in a |
| | | locking request is |
| | | neither the next expected |
| | | number or the last number |
| | | processed. |
| NFS4ERR_BADSESSION | TDB | TDB |
| NFS4ERR_BADSLOT | TDB | TDB |
| NFS4ERR_BAD_STATEID | 10025 | A stateid generated by |
| | | the current server |
| | | instance, but which does |
| | | not designate any locking |
| | | state (either current or |
| | | superseded) for a current |
| | | lockowner-file pair, was |
| | | used. |
| NFS4ERR_BADXDR | 10036 | The server encountered an |
| | | XDR decoding error while |
| | | processing an operation. |
| NFS4ERR_CLID_INUSE | 10017 | The SETCLIENTID operation |
| | | has found that a client |
| | | id is already in use by |
| | | another client. |
| NFS4ERR_DEADLOCK | 10045 | The server has been able |
| | | to determine a file |
| | | locking deadlock |
| | | condition for a blocking |
| | | lock request. |
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| NFS4ERR_DELAY | 10008 | The server initiated the |
| | | request, but was not able |
| | | to complete it in a |
| | | timely fashion. The |
| | | client should wait and |
| | | then try the request with |
| | | a new RPC transaction ID. |
| | | For example, this error |
| | | should be returned from a |
| | | server that supports |
| | | hierarchical storage and |
| | | receives a request to |
| | | process a file that has |
| | | been migrated. In this |
| | | case, the server should |
| | | start the immigration |
| | | process and respond to |
| | | client with this error. |
| | | This error may also occur |
| | | when a necessary |
| | | delegation recall makes |
| | | processing a request in a |
| | | timely fashion |
| | | impossible. |
| NFS4ERR_DENIED | 10010 | An attempt to lock a file |
| | | is denied. Since this may |
| | | be a temporary condition, |
| | | the client is encouraged |
| | | to retry the lock request |
| | | until the lock is |
| | | accepted. |
| NFS4ERR_DIRDELEG_UNAVAIL | TBD | TBD |
| NFS4ERR_DQUOT | 69 | Resource (quota) hard |
| | | limit exceeded. The |
| | | user's resource limit on |
| | | the server has been |
| | | exceeded. |
| NFS4ERR_EXIST | 17 | File exists. The file |
| | | specified already exists. |
| NFS4ERR_EXPIRED | 10011 | A lease has expired that |
| | | is being used in the |
| | | current operation. |
| NFS4ERR_FBIG | 27 | File too large. The |
| | | operation would have |
| | | caused a file to grow |
| | | beyond the server's |
| | | limit. |
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| NFS4ERR_FHEXPIRED | 10014 | The filehandle provided |
| | | is volatile and has |
| | | expired at the server. |
| NFS4ERR_FILE_OPEN | 10046 | The operation can not be |
| | | successfully processed |
| | | because a file involved |
| | | in the operation is |
| | | currently open. |
| NFS4ERR_GRACE | 10013 | The server is in its |
| | | recovery or grace period |
| | | which should match the |
| | | lease period of the |
| | | server. |
| NFS4ERR_INVAL | 22 | Invalid argument or |
| | | unsupported argument for |
| | | an operation. Two |
| | | examples are attempting a |
| | | READLINK on an object |
| | | other than a symbolic |
| | | link or specifying a |
| | | value for an enum field |
| | | that is not defined in |
| | | the protocol (e.g. |
| | | nfs_ftype4). |
| NFS4ERR_IO | 5 | I/O error. A hard error |
| | | (for example, a disk |
| | | error) occurred while |
| | | processing the requested |
| | | operation. |
| NFS4ERR_ISDIR | 21 | Is a directory. The |
| | | caller specified a |
| | | directory in a |
| | | non-directory operation. |
| NFS4ERR_LAYOUTTRYLATER | TDB | Layouts are temporarily |
| | | unavailable for the file, |
| | | client should retry |
| | | later. |
| NFS4ERR_LAYOUTUNAVAILABLE | TDB | Layouts are not available |
| | | for the file or its |
| | | containing file system. |
| NFS4ERR_LEASE_MOVED | 10031 | A lease being renewed is |
| | | associated with a |
| | | filesystem that has been |
| | | migrated to a new server. |
| NFS4ERR_LOCKED | 10012 | A read or write operation |
| | | was attempted on a locked |
| | | file. |
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| NFS4ERR_LOCK_NOTSUPP | 10043 | Server does not support |
| | | atomic upgrade or |
| | | downgrade of locks. |
| NFS4ERR_LOCK_RANGE | 10028 | A lock request is |
| | | operating on a sub-range |
| | | of a current lock for the |
| | | lock owner and the server |
| | | does not support this |
| | | type of request. |
| NFS4ERR_LOCKS_HELD | 10037 | A CLOSE was attempted and |
| | | file locks would exist |
| | | after the CLOSE. |
| NFS4ERR_MINOR_VERS_MISMATCH | 10021 | The server has received a |
| | | request that specifies an |
| | | unsupported minor |
| | | version. The server must |
| | | return a COMPOUND4res |
| | | with a zero length |
| | | operations result array. |
| NFS4ERR_MLINK | 31 | Too many hard links. |
| NFS4ERR_MOVED | 10019 | The filesystem which |
| | | contains the current |
| | | filehandle object has |
| | | been relocated or |
| | | migrated to another |
| | | server. The client may |
| | | obtain the new filesystem |
| | | location by obtaining the |
| | | "fs_locations" attribute |
| | | for the current |
| | | filehandle. For further |
| | | discussion, refer to the |
| | | section "Filesystem |
| | | Migration or Relocation". |
| NFS4ERR_MOVED_DATA | TDB | |
| NFS4ERR_MOVED_DATA_AND_STATE | TDB | |
| NFS4ERR_NAMETOOLONG | 63 | The filename in an |
| | | operation was too long. |
| NFS4ERR_NOENT | 2 | No such file or |
| | | directory. The file or |
| | | directory name specified |
| | | does not exist. |
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| NFS4ERR_NOFILEHANDLE | 10020 | The logical current |
| | | filehandle value (or, in |
| | | the case of RESTOREFH, |
| | | the saved filehandle |
| | | value) has not been set |
| | | properly. This may be a |
| | | result of a malformed |
| | | COMPOUND operation (i.e. |
| | | no PUTFH or PUTROOTFH |
| | | before an operation that |
| | | requires the current |
| | | filehandle be set). |
| NFS4ERR_NO_GRACE | 10033 | A reclaim of client state |
| | | has fallen outside of the |
| | | grace period of the |
| | | server. As a result, the |
| | | server can not guarantee |
| | | that conflicting state |
| | | has not been provided to |
| | | another client. |
| NFS4ERR_NOMATCHING_LAYOUT | TBD | Client has no matching |
| | | layout (segment) to |
| | | return. |
| NFS4ERR_NOSPC | 28 | No space left on device. |
| | | The operation would have |
| | | caused the server's |
| | | filesystem to exceed its |
| | | limit. |
| NFS4ERR_NOTDIR | 20 | Not a directory. The |
| | | caller specified a |
| | | non-directory in a |
| | | directory operation. |
| NFS4ERR_NOTEMPTY | 66 | An attempt was made to |
| | | remove a directory that |
| | | was not empty. |
| NFS4ERR_NOTSUPP | 10004 | Operation is not |
| | | supported. |
| NFS4ERR_NOT_SAME | 10027 | This error is returned by |
| | | the VERIFY operation to |
| | | signify that the |
| | | attributes compared were |
| | | not the same as provided |
| | | in the client's request. |
| NFS4ERR_NXIO | 6 | I/O error. No such device |
| | | or address. |
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| NFS4ERR_OLD_STATEID | 10024 | A stateid which |
| | | designates the locking |
| | | state for a |
| | | lockowner-file at an |
| | | earlier time was used. |
| NFS4ERR_OPENMODE | 10038 | The client attempted a |
| | | READ, WRITE, LOCK or |
| | | SETATTR operation not |
| | | sanctioned by the stateid |
| | | passed (e.g. writing to a |
| | | file opened only for |
| | | read). |
| NFS4ERR_OP_ILLEGAL | 10044 | An illegal operation |
| | | value has been specified |
| | | in the argop field of a |
| | | COMPOUND or CB_COMPOUND |
| | | procedure. |
| NFS4ERR_PERM | 1 | Not owner. The operation |
| | | was not allowed because |
| | | the caller is either not |
| | | a privileged user (root) |
| | | or not the owner of the |
| | | target of the operation. |
| NFS4ERR_RECALLCONFLICT | TBD | Layout is unavailable due |
| | | to a conflicting |
| | | LAYOUTRECALL that is in |
| | | progress. |
| NFS4ERR_RECLAIM_BAD | 10034 | The reclaim provided by |
| | | the client does not match |
| | | any of the server's state |
| | | consistency checks and is |
| | | bad. |
| NFS4ERR_RECLAIM_CONFLICT | 10035 | The reclaim provided by |
| | | the client has |
| | | encountered a conflict |
| | | and can not be provided. |
| | | Potentially indicates a |
| | | misbehaving client. |
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| NFS4ERR_RESOURCE | 10018 | For the processing of the |
| | | COMPOUND procedure, the |
| | | server may exhaust |
| | | available resources and |
| | | can not continue |
| | | processing operations |
| | | within the COMPOUND |
| | | procedure. This error |
| | | will be returned from the |
| | | server in those instances |
| | | of resource exhaustion |
| | | related to the processing |
| | | of the COMPOUND |
| | | procedure. |
| NFS4ERR_RESTOREFH | 10030 | The RESTOREFH operation |
| | | does not have a saved |
| | | filehandle (identified by |
| | | SAVEFH) to operate upon. |
| NFS4ERR_ROFS | 30 | Read-only filesystem. A |
| | | modifying operation was |
| | | attempted on a read-only |
| | | filesystem. |
| NFS4ERR_SAME | 10009 | This error is returned by |
| | | the NVERIFY operation to |
| | | signify that the |
| | | attributes compared were |
| | | the same as provided in |
| | | the client's request. |
| NFS4ERR_SERVERFAULT | 10006 | An error occurred on the |
| | | server which does not map |
| | | to any of the legal NFS |
| | | version 4 protocol error |
| | | values. The client should |
| | | translate this into an |
| | | appropriate error. UNIX |
| | | clients may choose to |
| | | translate this to EIO. |
| NFS4ERR_SHARE_DENIED | 10015 | An attempt to OPEN a file |
| | | with a share reservation |
| | | has failed because of a |
| | | share conflict. |
| NFS4ERR_STALE | 70 | Invalid filehandle. The |
| | | filehandle given in the |
| | | arguments was invalid. |
| | | The file referred to by |
| | | that filehandle no longer |
| | | exists or access to it |
| | | has been revoked. |
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| NFS4ERR_STALE_CLIENTID | 10022 | A clientid not recognized |
| | | by the server was used in |
| | | a locking or |
| | | SETCLIENTID_CONFIRM |
| | | request. |
| NFS4ERR_STALE_STATEID | 10023 | A stateid generated by an |
| | | earlier server instance |
| | | was used. |
| NFS4ERR_SYMLINK | 10029 | The current filehandle |
| | | provided for a LOOKUP is |
| | | not a directory but a |
| | | symbolic link. Also used |
| | | if the final component of |
| | | the OPEN path is a |
| | | symbolic link. |
| NFS4ERR_TOOSMALL | 10005 | The encoded response to a |
| | | READDIR request exceeds |
| | | the size limit set by the |
| | | initial request. |
| NFS4ERR_UNKNOWN_LAYOUTTYPE | TBD | Layout type is unknown. |
| NFS4ERR_WRONGSEC | 10016 | The security mechanism |
| | | being used by the client |
| | | for the operation does |
| | | not match the server's |
| | | security policy. The |
| | | client should change the |
| | | security mechanism being |
| | | used and retry the |
| | | operation. |
| NFS4ERR_XDEV | 18 | Attempt to do an |
| | | operation between |
| | | different fsids. |
| NFS4ERR_ | TDB | TDB |
+------------------------------+--------+---------------------------+
Table 5
21. NFS version 4.1 Procedures
21.1. Procedure 0: NULL - No Operation
21.1.1. SYNOPSIS
21.1.2. ARGUMENTS
void;
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21.1.3. RESULTS
void;
21.1.4. DESCRIPTION
Standard NULL procedure. Void argument, void response. This
procedure has no functionality associated with it. Because of this
it is sometimes used to measure the overhead of processing a service
request. Therefore, the server should ensure that no unnecessary
work is done in servicing this procedure.
21.1.5. ERRORS
None.
21.2. Procedure 1: COMPOUND - Compound Operations
21.2.1. SYNOPSIS
compoundargs -> compoundres
21.2.2. ARGUMENTS
union nfs_argop4 switch (nfs_opnum4 argop) {
case <OPCODE>: <argument>;
...
};
struct COMPOUND4args {
utf8str_cs tag;
uint32_t minorversion;
nfs_argop4 argarray<>;
};
21.2.3. RESULTS
union nfs_resop4 switch (nfs_opnum4 resop){
case <OPCODE>: <result>;
...
};
struct COMPOUND4res {
nfsstat4 status;
utf8str_cs tag;
nfs_resop4 resarray<>;
};
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21.2.4. DESCRIPTION
The COMPOUND procedure is used to combine one or more of the NFS
operations into a single RPC request. The main NFS RPC program has
two main procedures: NULL and COMPOUND. All other operations use the
COMPOUND procedure as a wrapper.
The COMPOUND procedure is used to combine individual operations into
a single RPC request. The server interprets each of the operations
in turn. If an operation is executed by the server and the status of
that operation is NFS4_OK, then the next operation in the COMPOUND
procedure is executed. The server continues this process until there
are no more operations to be executed or one of the operations has a
status value other than NFS4_OK.
In the processing of the COMPOUND procedure, the server may find that
it does not have the available resources to execute any or all of the
operations within the COMPOUND sequence. In this case, the error
NFS4ERR_RESOURCE will be returned for the particular operation within
the COMPOUND procedure where the resource exhaustion occurred. This
assumes that all previous operations within the COMPOUND sequence
have been evaluated successfully. The results for all of the
evaluated operations must be returned to the client.
The server will generally choose between two methods of decoding the
client's request. The first would be the traditional one pass XDR
decode. If there is an XDR decoding error in this case, the RPC XDR
decode error would be returned. The second method would be to make
an initial pass to decode the basic COMPOUND request and then to XDR
decode the individual operations; the most interesting is the decode
of attributes. In this case, the server may encounter an XDR decode
error during the second pass. In this case, the server would return
the error NFS4ERR_BADXDR to signify the decode error.
The COMPOUND arguments contain a "minorversion" field. The initial
and default value for this field is 0 (zero). This field will be
used by future minor versions such that the client can communicate to
the server what minor version is being requested. If the server
receives a COMPOUND procedure with a minorversion field value that it
does not support, the server MUST return an error of
NFS4ERR_MINOR_VERS_MISMATCH and a zero length resultdata array.
Contained within the COMPOUND results is a "status" field. If the
results array length is non-zero, this status must be equivalent to
the status of the last operation that was executed within the
COMPOUND procedure. Therefore, if an operation incurred an error
then the "status" value will be the same error value as is being
returned for the operation that failed.
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Note that operations, 0 (zero) and 1 (one) are not defined for the
COMPOUND procedure. Operation 2 is not defined but reserved for
future definition and use with minor versioning. If the server
receives a operation array that contains operation 2 and the
minorversion field has a value of 0 (zero), an error of
NFS4ERR_OP_ILLEGAL, as described in the next paragraph, is returned
to the client. If an operation array contains an operation 2 and the
minorversion field is non-zero and the server does not support the
minor version, the server returns an error of
NFS4ERR_MINOR_VERS_MISMATCH. Therefore, the
NFS4ERR_MINOR_VERS_MISMATCH error takes precedence over all other
errors.
It is possible that the server receives a request that contains an
operation that is less than the first legal operation (OP_ACCESS) or
greater than the last legal operation (OP_RELEASE_LOCKOWNER). In
this case, the server's response will encode the opcode OP_ILLEGAL
rather than the illegal opcode of the request. The status field in
the ILLEGAL return results will set to NFS4ERR_OP_ILLEGAL. The
COMPOUND procedure's return results will also be NFS4ERR_OP_ILLEGAL.
The definition of the "tag" in the request is left to the
implementor. It may be used to summarize the content of the compound
request for the benefit of packet sniffers and engineers debugging
implementations. However, the value of "tag" in the response SHOULD
be the same value as provided in the request. This applies to the
tag field of the CB_COMPOUND procedure as well.
21.2.5. IMPLEMENTATION
Since an error of any type may occur after only a portion of the
operations have been evaluated, the client must be prepared to
recover from any failure. If the source of an NFS4ERR_RESOURCE error
was a complex or lengthy set of operations, it is likely that if the
number of operations were reduced the server would be able to
evaluate them successfully. Therefore, the client is responsible for
dealing with this type of complexity in recovery.
21.2.6. ERRORS
All errors defined in the protocol
22. NFS version 4.1 Operations
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22.1. Operation 3: ACCESS - Check Access Rights
22.1.1. SYNOPSIS
(cfh), accessreq -> supported, accessrights
22.1.2. ARGUMENTS
const ACCESS4_READ = 0x00000001;
const ACCESS4_LOOKUP = 0x00000002;
const ACCESS4_MODIFY = 0x00000004;
const ACCESS4_EXTEND = 0x00000008;
const ACCESS4_DELETE = 0x00000010;
const ACCESS4_EXECUTE = 0x00000020;
struct ACCESS4args {
/* CURRENT_FH: object */
uint32_t access;
};
22.1.3. RESULTS
struct ACCESS4resok {
uint32_t supported;
uint32_t access;
};
union ACCESS4res switch (nfsstat4 status) {
case NFS4_OK:
ACCESS4resok resok4;
default:
void;
};
22.1.4. DESCRIPTION
ACCESS determines the access rights that a user, as identified by the
credentials in the RPC request, has with respect to the file system
object specified by the current filehandle. The client encodes the
set of access rights that are to be checked in the bit mask "access".
The server checks the permissions encoded in the bit mask. If a
status of NFS4_OK is returned, two bit masks are included in the
response. The first, "supported", represents the access rights for
which the server can verify reliably. The second, "access",
represents the access rights available to the user for the filehandle
provided. On success, the current filehandle retains its value.
Note that the supported field will contain only as many values as was
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originally sent in the arguments. For example, if the client sends
an ACCESS operation with only the ACCESS4_READ value set and the
server supports this value, the server will return only ACCESS4_READ
even if it could have reliably checked other values.
The results of this operation are necessarily advisory in nature. A
return status of NFS4_OK and the appropriate bit set in the bit mask
does not imply that such access will be allowed to the file system
object in the future. This is because access rights can be revoked
by the server at any time.
The following access permissions may be requested:
ACCESS4_READ Read data from file or read a directory.
ACCESS4_LOOKUP Look up a name in a directory (no meaning for non-
directory objects).
ACCESS4_MODIFY Rewrite existing file data or modify existing
directory entries.
ACCESS4_EXTEND Write new data or add directory entries.
ACCESS4_DELETE Delete an existing directory entry.
ACCESS4_EXECUTE Execute file (no meaning for a directory).
On success, the current filehandle retains its value.
22.1.5. IMPLEMENTATION
In general, it is not sufficient for the client to attempt to deduce
access permissions by inspecting the uid, gid, and mode fields in the
file attributes or by attempting to interpret the contents of the ACL
attribute. This is because the server may perform uid or gid mapping
or enforce additional access control restrictions. It is also
possible that the server may not be in the same ID space as the
client. In these cases (and perhaps others), the client can not
reliably perform an access check with only current file attributes.
In the NFS version 2 protocol, the only reliable way to determine
whether an operation was allowed was to try it and see if it
succeeded or failed. Using the ACCESS operation in the NFS version 4
protocol, the client can ask the server to indicate whether or not
one or more classes of operations are permitted. The ACCESS
operation is provided to allow clients to check before doing a series
of operations which will result in an access failure. The OPEN
operation provides a point where the server can verify access to the
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file object and method to return that information to the client. The
ACCESS operation is still useful for directory operations or for use
in the case the UNIX API "access" is used on the client.
The information returned by the server in response to an ACCESS call
is not permanent. It was correct at the exact time that the server
performed the checks, but not necessarily afterwards. The server can
revoke access permission at any time.
The client should use the effective credentials of the user to build
the authentication information in the ACCESS request used to
determine access rights. It is the effective user and group
credentials that are used in subsequent read and write operations.
Many implementations do not directly support the ACCESS4_DELETE
permission. Operating systems like UNIX will ignore the
ACCESS4_DELETE bit if set on an access request on a non-directory
object. In these systems, delete permission on a file is determined
by the access permissions on the directory in which the file resides,
instead of being determined by the permissions of the file itself.
Therefore, the mask returned enumerating which access rights can be
determined will have the ACCESS4_DELETE value set to 0. This
indicates to the client that the server was unable to check that
particular access right. The ACCESS4_DELETE bit in the access mask
returned will then be ignored by the client.
22.1.6. ERRORS
NFS4ERR_ACCESS NFS4ERR_BADHANDLE NFS4ERR_BADXDR NFS4ERR_DELAY
NFS4ERR_FHEXPIRED NFS4ERR_INVAL NFS4ERR_IO NFS4ERR_MOVED
NFS4ERR_NOFILEHANDLE NFS4ERR_RESOURCE NFS4ERR_SERVERFAULT
NFS4ERR_STALE
22.2. Operation 4: CLOSE - Close File
22.2.1. SYNOPSIS
(cfh), seqid, open_stateid -> open_stateid
22.2.2. ARGUMENTS
struct CLOSE4args {
/* CURRENT_FH: object */
seqid4 seqid
stateid4 open_stateid;
};
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22.2.3. RESULTS
union CLOSE4res switch (nfsstat4 status) {
case NFS4_OK:
stateid4 open_stateid;
default:
void;
};
22.2.4. DESCRIPTION
The CLOSE operation releases share reservations for the regular or
named attribute file as specified by the current filehandle. The
share reservations and other state information released at the server
as a result of this CLOSE is only associated with the supplied
stateid. The sequence id provides for the correct ordering. State
associated with other OPENs is not affected.
If record locks are held, the client SHOULD release all locks before
issuing a CLOSE. The server MAY free all outstanding locks on CLOSE
but some servers may not support the CLOSE of a file that still has
record locks held. The server MUST return failure if any locks would
exist after the CLOSE.
On success, the current filehandle retains its value.
22.2.5. IMPLEMENTATION
Even though CLOSE returns a stateid, this stateid is not useful to
the client and should be treated as deprecated. CLOSE "shuts down"
the state associated with all OPENs for the file by a single
open_owner. As noted above, CLOSE will either release all file
locking state or return an error. Therefore, the stateid returned by
CLOSE is not useful for operations that follow.
22.2.6. ERRORS
NFS4ERR_ADMIN_REVOKED NFS4ERR_BADHANDLE NFS4ERR_BAD_SEQID
NFS4ERR_BAD_STATEID NFS4ERR_BADXDR NFS4ERR_DELAY NFS4ERR_EXPIRED
NFS4ERR_FHEXPIRED NFS4ERR_INVAL NFS4ERR_ISDIR NFS4ERR_LEASE_MOVED
NFS4ERR_LOCKS_HELD NFS4ERR_MOVED NFS4ERR_NOFILEHANDLE
NFS4ERR_OLD_STATEID NFS4ERR_RESOURCE NFS4ERR_SERVERFAULT
NFS4ERR_STALE NFS4ERR_STALE_STATEID
22.3. Operation 5: COMMIT - Commit Cached Data
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22.3.1. SYNOPSIS
(cfh), offset, count -> verifier
22.3.2. ARGUMENTS
struct COMMIT4args {
/* CURRENT_FH: file */
offset4 offset;
count4 count;
};
22.3.3. RESULTS
struct COMMIT4resok {
verifier4 writeverf;
};
union COMMIT4res switch (nfsstat4 status) {
case NFS4_OK:
COMMIT4resok resok4;
default:
void;
};
22.3.4. DESCRIPTION
The COMMIT operation forces or flushes data to stable storage for the
file specified by the current filehandle. The flushed data is that
which was previously written with a WRITE operation which had the
stable field set to UNSTABLE4.
The offset specifies the position within the file where the flush is
to begin. An offset value of 0 (zero) means to flush data starting
at the beginning of the file. The count specifies the number of
bytes of data to flush. If count is 0 (zero), a flush from offset to
the end of the file is done.
The server returns a write verifier upon successful completion of the
COMMIT. The write verifier is used by the client to determine if the
server has restarted or rebooted between the initial WRITE(s) and the
COMMIT. The client does this by comparing the write verifier
returned from the initial writes and the verifier returned by the
COMMIT operation. The server must vary the value of the write
verifier at each server event or instantiation that may lead to a
loss of uncommitted data. Most commonly this occurs when the server
is rebooted; however, other events at the server may result in
uncommitted data loss as well.
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On success, the current filehandle retains its value.
22.3.5. IMPLEMENTATION
The COMMIT operation is similar in operation and semantics to the
POSIX fsync(2) system call that synchronizes a file's state with the
disk (file data and metadata is flushed to disk or stable storage).
COMMIT performs the same operation for a client, flushing any
unsynchronized data and metadata on the server to the server's disk
or stable storage for the specified file. Like fsync(2), it may be
that there is some modified data or no modified data to synchronize.
The data may have been synchronized by the server's normal periodic
buffer synchronization activity. COMMIT should return NFS4_OK,
unless there has been an unexpected error.
COMMIT differs from fsync(2) in that it is possible for the client to
flush a range of the file (most likely triggered by a buffer-
reclamation scheme on the client before file has been completely
written).
The server implementation of COMMIT is reasonably simple. If the
server receives a full file COMMIT request, that is starting at
offset 0 and count 0, it should do the equivalent of fsync()'ing the
file. Otherwise, it should arrange to have the cached data in the
range specified by offset and count to be flushed to stable storage.
In both cases, any metadata associated with the file must be flushed
to stable storage before returning. It is not an error for there to
be nothing to flush on the server. This means that the data and
metadata that needed to be flushed have already been flushed or lost
during the last server failure.
The client implementation of COMMIT is a little more complex. There
are two reasons for wanting to commit a client buffer to stable
storage. The first is that the client wants to reuse a buffer. In
this case, the offset and count of the buffer are sent to the server
in the COMMIT request. The server then flushes any cached data based
on the offset and count, and flushes any metadata associated with the
file. It then returns the status of the flush and the write
verifier. The other reason for the client to generate a COMMIT is
for a full file flush, such as may be done at close. In this case,
the client would gather all of the buffers for this file that contain
uncommitted data, do the COMMIT operation with an offset of 0 and
count of 0, and then free all of those buffers. Any other dirty
buffers would be sent to the server in the normal fashion.
After a buffer is written by the client with the stable parameter set
to UNSTABLE4, the buffer must be considered as modified by the client
until the buffer has either been flushed via a COMMIT operation or
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written via a WRITE operation with stable parameter set to FILE_SYNC4
or DATA_SYNC4. This is done to prevent the buffer from being freed
and reused before the data can be flushed to stable storage on the
server.
When a response is returned from either a WRITE or a COMMIT operation
and it contains a write verifier that is different than previously
returned by the server, the client will need to retransmit all of the
buffers containing uncommitted cached data to the server. How this
is to be done is up to the implementor. If there is only one buffer
of interest, then it should probably be sent back over in a WRITE
request with the appropriate stable parameter. If there is more than
one buffer, it might be worthwhile retransmitting all of the buffers
in WRITE requests with the stable parameter set to UNSTABLE4 and then
retransmitting the COMMIT operation to flush all of the data on the
server to stable storage. The timing of these retransmissions is
left to the implementor.
The above description applies to page-cache-based systems as well as
buffer-cache-based systems. In those systems, the virtual memory
system will need to be modified instead of the buffer cache.
22.3.6. ERRORS
NFS4ERR_ACCESS NFS4ERR_BADHANDLE NFS4ERR_BADXDR NFS4ERR_FHEXPIRED
NFS4ERR_INVAL NFS4ERR_IO NFS4ERR_ISDIR NFS4ERR_MOVED
NFS4ERR_NOFILEHANDLE NFS4ERR_RESOURCE NFS4ERR_ROFS
NFS4ERR_SERVERFAULT NFS4ERR_STALE
22.4. Operation 6: CREATE - Create a Non-Regular File Object
22.4.1. SYNOPSIS
(cfh), name, type, attrs -> (cfh), change_info, attrs_set
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22.4.2. ARGUMENTS
union createtype4 switch (nfs_ftype4 type) {
case NF4LNK:
linktext4 linkdata;
case NF4BLK:
case NF4CHR:
specdata4 devdata;
case NF4SOCK:
case NF4FIFO:
case NF4DIR:
void;
};
struct CREATE4args {
/* CURRENT_FH: directory for creation */
createtype4 objtype;
component4 objname;
fattr4 createattrs;
};
22.4.3. RESULTS
struct CREATE4resok {
change_info4 cinfo;
bitmap4 attrset; /* attributes set */
};
union CREATE4res switch (nfsstat4 status) {
case NFS4_OK:
CREATE4resok resok4;
default:
void;
};
22.4.4. DESCRIPTION
The CREATE operation creates a non-regular file object in a directory
with a given name. The OPEN operation MUST be used to create a
regular file.
The objname specifies the name for the new object. The objtype
determines the type of object to be created: directory, symlink, etc.
If an object of the same name already exists in the directory, the
server will return the error NFS4ERR_EXIST.
For the directory where the new file object was created, the server
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returns change_info4 information in cinfo. With the atomic field of
the change_info4 struct, the server will indicate if the before and
after change attributes were obtained atomically with respect to the
file object creation.
If the objname has a length of 0 (zero), or if objname does not obey
the UTF-8 definition, the error NFS4ERR_INVAL will be returned.
The current filehandle is replaced by that of the new object.
The createattrs specifies the initial set of attributes for the
object. The set of attributes may include any writable attribute
valid for the object type. When the operation is successful, the
server will return to the client an attribute mask signifying which
attributes were successfully set for the object.
If createattrs includes neither the owner attribute nor an ACL with
an ACE for the owner, and if the server's filesystem both supports
and requires an owner attribute (or an owner ACE) then the server
MUST derive the owner (or the owner ACE). This would typically be
from the principal indicated in the RPC credentials of the call, but
the server's operating environment or filesystem semantics may
dictate other methods of derivation. Similarly, if createattrs
includes neither the group attribute nor a group ACE, and if the
server's filesystem both supports and requires the notion of a group
attribute (or group ACE), the server MUST derive the group attribute
(or the corresponding owner ACE) for the file. This could be from
the RPC call's credentials, such as the group principal if the
credentials include it (such as with AUTH_SYS), from the group
identifier associated with the principal in the credentials (for
e.g., POSIX systems have a passwd database that has the group
identifier for every user identifier), inherited from directory the
object is created in, or whatever else the server's operating
environment or filesystem semantics dictate. This applies to the
OPEN operation too.
Conversely, it is possible the client will specify in createattrs an
owner attribute or group attribute or ACL that the principal
indicated the RPC call's credentials does not have permissions to
create files for. The error to be returned in this instance is
NFS4ERR_PERM. This applies to the OPEN operation too.
22.4.5. IMPLEMENTATION
If the client desires to set attribute values after the create, a
SETATTR operation can be added to the COMPOUND request so that the
appropriate attributes will be set.
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22.4.6. ERRORS
NFS4ERR_ACCESS NFS4ERR_ATTRNOTSUPP NFS4ERR_BADCHAR NFS4ERR_BADHANDLE
NFS4ERR_BADNAME NFS4ERR_BADOWNER NFS4ERR_BADTYPE NFS4ERR_BADXDR
NFS4ERR_DELAY NFS4ERR_DQUOT NFS4ERR_EXIST NFS4ERR_FHEXPIRED
NFS4ERR_INVAL NFS4ERR_IO NFS4ERR_MOVED NFS4ERR_NAMETOOLONG
NFS4ERR_NOFILEHANDLE NFS4ERR_NOSPC NFS4ERR_NOTDIR NFS4ERR_PERM
NFS4ERR_RESOURCE NFS4ERR_ROFS NFS4ERR_SERVERFAULT NFS4ERR_STALE
22.5. Operation 7: DELEGPURGE - Purge Delegations Awaiting Recovery
22.5.1. SYNOPSIS
clientid ->
22.5.2. ARGUMENTS
struct DELEGPURGE4args {
clientid4 clientid;
};
22.5.3. RESULTS
struct DELEGPURGE4res {
nfsstat4 status;
};
22.5.4. DESCRIPTION
Purges all of the delegations awaiting recovery for a given client.
This is useful for clients which do not commit delegation information
to stable storage to indicate that conflicting requests need not be
delayed by the server awaiting recovery of delegation information.
This operation should be used by clients that record delegation
information on stable storage on the client. In this case,
DELEGPURGE should be issued immediately after doing delegation
recovery on all delegations known to the client. Doing so will
notify the server that no additional delegations for the client will
be recovered allowing it to free resources, and avoid delaying other
clients who make requests that conflict with the unrecovered
delegations. The set of delegations known to the server and the
client may be different. The reason for this is that a client may
fail after making a request which resulted in delegation but before
it received the results and committed them to the client's stable
storage.
The server MAY support DELEGPURGE, but if it does not, it MUST NOT
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support CLAIM_DELEGATE_PREV.
22.5.5. ERRORS
NFS4ERR_BADXDR NFS4ERR_NOTSUPP NFS4ERR_LEASE_MOVED NFS4ERR_MOVED
NFS4ERR_RESOURCE NFS4ERR_SERVERFAULT NFS4ERR_STALE_CLIENTID
22.6. Operation 8: DELEGRETURN - Return Delegation
22.6.1. SYNOPSIS
(cfh), stateid ->
22.6.2. ARGUMENTS
struct DELEGRETURN4args {
/* CURRENT_FH: delegated file */
stateid4 stateid;
};
22.6.3. RESULTS
struct DELEGRETURN4res {
nfsstat4 status;
};
22.6.4. DESCRIPTION
Returns the delegation represented by the current filehandle and
stateid.
Delegations may be returned when recalled or voluntarily (i.e. before
the server has recalled them). In either case the client must
properly propagate state changed under the context of the delegation
to the server before returning the delegation.
22.6.5. ERRORS
NFS4ERR_ADMIN_REVOKED NFS4ERR_BAD_STATEID NFS4ERR_BADXDR
NFS4ERR_EXPIRED NFS4ERR_INVAL NFS4ERR_LEASE_MOVED NFS4ERR_MOVED
NFS4ERR_NOFILEHANDLE NFS4ERR_NOTSUPP NFS4ERR_OLD_STATEID
NFS4ERR_RESOURCE NFS4ERR_SERVERFAULT NFS4ERR_STALE
NFS4ERR_STALE_STATEID
22.7. Operation 9: GETATTR - Get Attributes
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22.7.1. SYNOPSIS
(cfh), attrbits -> attrbits, attrvals
22.7.2. ARGUMENTS
struct GETATTR4args {
/* CURRENT_FH: directory or file */
bitmap4 attr_request;
};
22.7.3. RESULTS
struct GETATTR4resok {
fattr4 obj_attributes;
};
union GETATTR4res switch (nfsstat4 status) {
case NFS4_OK:
GETATTR4resok resok4;
default:
void;
};
22.7.4. DESCRIPTION
The GETATTR operation will obtain attributes for the filesystem
object specified by the current filehandle. The client sets a bit in
the bitmap argument for each attribute value that it would like the
server to return. The server returns an attribute bitmap that
indicates the attribute values for which it was able to return,
followed by the attribute values ordered lowest attribute number
first.
The server must return a value for each attribute that the client
requests if the attribute is supported by the server. If the server
does not support an attribute or cannot approximate a useful value
then it must not return the attribute value and must not set the
attribute bit in the result bitmap. The server must return an error
if it supports an attribute but cannot obtain its value. In that
case no attribute values will be returned.
All servers must support the mandatory attributes as specified in
File Attributes (Section 3).
On success, the current filehandle retains its value.
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22.7.5. IMPLEMENTATION
22.7.6. ERRORS
NFS4ERR_ACCESS NFS4ERR_BADHANDLE NFS4ERR_BADXDR NFS4ERR_DELAY
NFS4ERR_FHEXPIRED NFS4ERR_INVAL NFS4ERR_IO NFS4ERR_MOVED
NFS4ERR_NOFILEHANDLE NFS4ERR_RESOURCE NFS4ERR_SERVERFAULT
NFS4ERR_STALE
22.8. Operation 10: GETFH - Get Current Filehandle
22.8.1. SYNOPSIS
(cfh) -> filehandle
22.8.2. ARGUMENTS
/* CURRENT_FH: */
void;
22.8.3. RESULTS
struct GETFH4resok {
nfs_fh4 object;
};
union GETFH4res switch (nfsstat4 status) {
case NFS4_OK:
GETFH4resok resok4;
default:
void;
};
22.8.4. DESCRIPTION
This operation returns the current filehandle value.
On success, the current filehandle retains its value.
22.8.5. IMPLEMENTATION
Operations that change the current filehandle like LOOKUP or CREATE
do not automatically return the new filehandle as a result. For
instance, if a client needs to lookup a directory entry and obtain
its filehandle then the following request is needed.
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PUTFH (directory filehandle)
LOOKUP (entry name)
GETFH
22.8.6. ERRORS
NFS4ERR_BADHANDLE NFS4ERR_FHEXPIRED NFS4ERR_MOVED
NFS4ERR_NOFILEHANDLE NFS4ERR_RESOURCE NFS4ERR_SERVERFAULT
NFS4ERR_STALE
22.9. Operation 11: LINK - Create Link to a File
22.9.1. SYNOPSIS
(sfh), (cfh), newname -> (cfh), change_info
22.9.2. ARGUMENTS
struct LINK4args {
/* SAVED_FH: source object */
/* CURRENT_FH: target directory */
component4 newname;
};
22.9.3. RESULTS
struct LINK4resok {
change_info4 cinfo;
};
union LINK4res switch (nfsstat4 status) {
case NFS4_OK:
LINK4resok resok4;
default:
void;
};
22.9.4. DESCRIPTION
The LINK operation creates an additional newname for the file
represented by the saved filehandle, as set by the SAVEFH operation,
in the directory represented by the current filehandle. The existing
file and the target directory must reside within the same filesystem
on the server. On success, the current filehandle will continue to
be the target directory. If an object exists in the target directory
with the same name as newname, the server must return NFS4ERR_EXIST.
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For the target directory, the server returns change_info4 information
in cinfo. With the atomic field of the change_info4 struct, the
server will indicate if the before and after change attributes were
obtained atomically with respect to the link creation.
If the newname has a length of 0 (zero), or if newname does not obey
the UTF-8 definition, the error NFS4ERR_INVAL will be returned.
22.9.5. IMPLEMENTATION
Changes to any property of the "hard" linked files are reflected in
all of the linked files. When a link is made to a file, the
attributes for the file should have a value for numlinks that is one
greater than the value before the LINK operation.
The statement "file and the target directory must reside within the
same filesystem on the server" means that the fsid fields in the
attributes for the objects are the same. If they reside on different
filesystems, the error, NFS4ERR_XDEV, is returned. On some servers,
the filenames, "." and "..", are illegal as newname.
In the case that newname is already linked to the file represented by
the saved filehandle, the server will return NFS4ERR_EXIST.
Note that symbolic links are created with the CREATE operation.
22.9.6. ERRORS
NFS4ERR_ACCESS NFS4ERR_BADCHAR NFS4ERR_BADHANDLE NFS4ERR_BADNAME
NFS4ERR_BADXDR NFS4ERR_DELAY NFS4ERR_DQUOT NFS4ERR_EXIST
NFS4ERR_FHEXPIRED NFS4ERR_FILE_OPEN NFS4ERR_INVAL NFS4ERR_IO
NFS4ERR_ISDIR NFS4ERR_MLINK NFS4ERR_MOVED NFS4ERR_NAMETOOLONG
NFS4ERR_NOENT NFS4ERR_NOFILEHANDLE NFS4ERR_NOSPC NFS4ERR_NOTDIR
NFS4ERR_NOTSUPP NFS4ERR_RESOURCE NFS4ERR_ROFS NFS4ERR_SERVERFAULT
NFS4ERR_STALE NFS4ERR_WRONGSEC NFS4ERR_XDEV
22.10. Operation 12: LOCK - Create Lock
22.10.1. SYNOPSIS
(cfh) locktype, reclaim, offset, length, locker -> stateid
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22.10.2. ARGUMENTS
struct open_to_lock_owner4 {
seqid4 open_seqid;
stateid4 open_stateid;
seqid4 lock_seqid;
lock_owner4 lock_owner;
};
struct exist_lock_owner4 {
stateid4 lock_stateid;
seqid4 lock_seqid;
};
union locker4 switch (bool new_lock_owner) {
case TRUE:
open_to_lock_owner4 open_owner;
case FALSE:
exist_lock_owner4 lock_owner;
};
enum nfs_lock_type4 {
READ_LT = 1,
WRITE_LT = 2,
READW_LT = 3, /* blocking read */
WRITEW_LT = 4 /* blocking write */
};
struct LOCK4args {
/* CURRENT_FH: file */
nfs_lock_type4 locktype;
bool reclaim;
offset4 offset;
length4 length;
locker4 locker;
};
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22.10.3. RESULTS
struct LOCK4denied {
offset4 offset;
length4 length;
nfs_lock_type4 locktype;
lock_owner4 owner;
};
struct LOCK4resok {
stateid4 lock_stateid;
};
union LOCK4res switch (nfsstat4 status) {
case NFS4_OK:
LOCK4resok resok4;
case NFS4ERR_DENIED:
LOCK4denied denied;
default:
void;
};
22.10.4. DESCRIPTION
The LOCK operation requests a record lock for the byte range
specified by the offset and length parameters. The lock type is also
specified to be one of the nfs_lock_type4s. If this is a reclaim
request, the reclaim parameter will be TRUE;
Bytes in a file may be locked even if those bytes are not currently
allocated to the file. To lock the file from a specific offset
through the end-of-file (no matter how long the file actually is) use
a length field with all bits set to 1 (one). If the length is zero,
or if a length which is not all bits set to one is specified, and
length when added to the offset exceeds the maximum 64-bit unsigned
integer value, the error NFS4ERR_INVAL will result.
Some servers may only support locking for byte offsets that fit
within 32 bits. If the client specifies a range that includes a byte
beyond the last byte offset of the 32-bit range, but does not include
the last byte offset of the 32-bit and all of the byte offsets beyond
it, up to the end of the valid 64-bit range, such a 32-bit server
MUST return the error NFS4ERR_BAD_RANGE.
In the case that the lock is denied, the owner, offset, and length of
a conflicting lock are returned.
On success, the current filehandle retains its value.
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22.10.5. IMPLEMENTATION
If the server is unable to determine the exact offset and length of
the conflicting lock, the same offset and length that were provided
in the arguments should be returned in the denied results. The File
Locking section contains a full description of this and the other
file locking operations.
LOCK operations are subject to permission checks and to checks
against the access type of the associated file. However, the
specific right and modes required for various type of locks, reflect
the semantics of the server-exported filesystem, and are not
specified by the protocol. For example, Windows 2000 allows a write
lock of a file open for READ, while a POSIX-compliant system does
not.
When the client makes a lock request that corresponds to a range that
the lockowner has locked already (with the same or different lock
type), or to a sub-region of such a range, or to a region which
includes multiple locks already granted to that lockowner, in whole
or in part, and the server does not support such locking operations
(i.e. does not support POSIX locking semantics), the server will
return the error NFS4ERR_LOCK_RANGE. In that case, the client may
return an error, or it may emulate the required operations, using
only LOCK for ranges that do not include any bytes already locked by
that lock_owner and LOCKU of locks held by that lock_owner
(specifying an exactly-matching range and type). Similarly, when the
client makes a lock request that amounts to upgrading (changing from
a read lock to a write lock) or downgrading (changing from write lock
to a read lock) an existing record lock, and the server does not
support such a lock, the server will return NFS4ERR_LOCK_NOTSUPP.
Such operations may not perfectly reflect the required semantics in
the face of conflicting lock requests from other clients.
The locker argument specifies the lock_owner that is associated with
the LOCK request. The locker4 structure is a switched union that
indicates whether the lock_owner is known to the server or if the
lock_owner is new to the server. In the case that the lock_owner is
known to the server and has an established lock_seqid, the argument
is just the lock_owner and lock_seqid. In the case that the
lock_owner is not known to the server, the argument contains not only
the lock_owner and lock_seqid but also the open_stateid and
open_seqid. The new lock_owner case covers the very first lock done
by the lock_owner and offers a method to use the established state of
the open_stateid to transition to the use of the lock_owner.
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22.10.6. ERRORS
NFS4ERR_ACCESS NFS4ERR_ADMIN_REVOKED NFS4ERR_BADHANDLE
NFS4ERR_BAD_RANGE NFS4ERR_BAD_SEQID NFS4ERR_BAD_STATEID
NFS4ERR_BADXDR NFS4ERR_DEADLOCK NFS4ERR_DELAY NFS4ERR_DENIED
NFS4ERR_EXPIRED NFS4ERR_FHEXPIRED NFS4ERR_GRACE NFS4ERR_INVAL
NFS4ERR_ISDIR NFS4ERR_LEASE_MOVED NFS4ERR_LOCK_NOTSUPP
NFS4ERR_LOCK_RANGE NFS4ERR_MOVED NFS4ERR_NOFILEHANDLE
NFS4ERR_NO_GRACE NFS4ERR_OLD_STATEID NFS4ERR_OPENMODE
NFS4ERR_RECLAIM_BAD NFS4ERR_RECLAIM_CONFLICT NFS4ERR_RESOURCE
NFS4ERR_SERVERFAULT NFS4ERR_STALE NFS4ERR_STALE_CLIENTID
NFS4ERR_STALE_STATEID
22.11. Operation 13: LOCKT - Test For Lock
22.11.1. SYNOPSIS
(cfh) locktype, offset, length owner -> {void, NFS4ERR_DENIED ->
owner}
22.11.2. ARGUMENTS
struct LOCKT4args {
/* CURRENT_FH: file */
nfs_lock_type4 locktype;
offset4 offset;
length4 length;
lock_owner4 owner;
};
22.11.3. RESULTS
struct LOCK4denied {
offset4 offset;
length4 length;
nfs_lock_type4 locktype;
lock_owner4 owner;
};
union LOCKT4res switch (nfsstat4 status) {
case NFS4ERR_DENIED:
LOCK4denied denied;
case NFS4_OK:
void;
default:
void;
};
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22.11.4. DESCRIPTION
The LOCKT operation tests the lock as specified in the arguments. If
a conflicting lock exists, the owner, offset, length, and type of the
conflicting lock are returned; if no lock is held, nothing other than
NFS4_OK is returned. Lock types READ_LT and READW_LT are processed
in the same way in that a conflicting lock test is done without
regard to blocking or non-blocking. The same is true for WRITE_LT
and WRITEW_LT.
The ranges are specified as for LOCK. The NFS4ERR_INVAL and
NFS4ERR_BAD_RANGE errors are returned under the same circumstances as
for LOCK.
On success, the current filehandle retains its value.
22.11.5. IMPLEMENTATION
If the server is unable to determine the exact offset and length of
the conflicting lock, the same offset and length that were provided
in the arguments should be returned in the denied results. The File
Locking section contains further discussion of the file locking
mechanisms.
LOCKT uses a lock_owner4 rather a stateid4, as is used in LOCK to
identify the owner. This is because the client does not have to open
the file to test for the existence of a lock, so a stateid may not be
available.
The test for conflicting locks should exclude locks for the current
lockowner. Note that since such locks are not examined the possible
existence of overlapping ranges may not affect the results of LOCKT.
If the server does examine locks that match the lockowner for the
purpose of range checking, NFS4ERR_LOCK_RANGE may be returned.. In
the event that it returns NFS4_OK, clients may do a LOCK and receive
NFS4ERR_LOCK_RANGE on the LOCK request because of the flexibility
provided to the server.
22.11.6. ERRORS
NFS4ERR_ACCESS NFS4ERR_BADHANDLE NFS4ERR_BAD_RANGE NFS4ERR_BADXDR
NFS4ERR_DELAY NFS4ERR_DENIED NFS4ERR_FHEXPIRED NFS4ERR_GRACE
NFS4ERR_INVAL NFS4ERR_ISDIR NFS4ERR_LEASE_MOVED NFS4ERR_LOCK_RANGE
NFS4ERR_MOVED NFS4ERR_NOFILEHANDLE NFS4ERR_RESOURCE
NFS4ERR_SERVERFAULT NFS4ERR_STALE NFS4ERR_STALE_CLIENTID
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22.12. Operation 14: LOCKU - Unlock File
22.12.1. SYNOPSIS
(cfh) type, seqid, stateid, offset, length -> stateid
22.12.2. ARGUMENTS
struct LOCKU4args {
/* CURRENT_FH: file */
nfs_lock_type4 locktype;
seqid4 seqid;
stateid4 stateid;
offset4 offset;
length4 length;
};
22.12.3. RESULTS
union LOCKU4res switch (nfsstat4 status) {
case NFS4_OK:
stateid4 stateid;
default:
void;
};
22.12.4. DESCRIPTION
The LOCKU operation unlocks the record lock specified by the
parameters. The client may set the locktype field to any value that
is legal for the nfs_lock_type4 enumerated type, and the server MUST
accept any legal value for locktype. Any legal value for locktype
has no effect on the success or failure of the LOCKU operation.
The ranges are specified as for LOCK. The NFS4ERR_INVAL and
NFS4ERR_BAD_RANGE errors are returned under the same circumstances as
for LOCK.
On success, the current filehandle retains its value.
22.12.5. IMPLEMENTATION
If the area to be unlocked does not correspond exactly to a lock
actually held by the lockowner the server may return the error
NFS4ERR_LOCK_RANGE. This includes the case in which the area is not
locked, where the area is a sub-range of the area locked, where it
overlaps the area locked without matching exactly or the area
specified includes multiple locks held by the lockowner. In all of
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these cases, allowed by POSIX locking semantics, a client receiving
this error, should if it desires support for such operations,
simulate the operation using LOCKU on ranges corresponding to locks
it actually holds, possibly followed by LOCK requests for the sub-
ranges not being unlocked.
22.12.6. ERRORS
NFS4ERR_ACCESS NFS4ERR_ADMIN_REVOKED NFS4ERR_BADHANDLE
NFS4ERR_BAD_RANGE NFS4ERR_BAD_SEQID NFS4ERR_BAD_STATEID
NFS4ERR_BADXDR NFS4ERR_EXPIRED NFS4ERR_FHEXPIRED NFS4ERR_GRACE
NFS4ERR_INVAL NFS4ERR_ISDIR NFS4ERR_LEASE_MOVED NFS4ERR_LOCK_RANGE
NFS4ERR_MOVED NFS4ERR_NOFILEHANDLE NFS4ERR_OLD_STATEID
NFS4ERR_RESOURCE NFS4ERR_SERVERFAULT NFS4ERR_STALE
NFS4ERR_STALE_STATEID
22.13. Operation 15: LOOKUP - Lookup Filename
22.13.1. SYNOPSIS
(cfh), component -> (cfh)
22.13.2. ARGUMENTS
struct LOOKUP4args {
/* CURRENT_FH: directory */
component4 objname;
};
22.13.3. RESULTS
struct LOOKUP4res {
/* CURRENT_FH: object */
nfsstat4 status;
};
22.13.4. DESCRIPTION
This operation LOOKUPs or finds a filesystem object using the
directory specified by the current filehandle. LOOKUP evaluates the
component and if the object exists the current filehandle is replaced
with the component's filehandle.
If the component cannot be evaluated either because it does not exist
or because the client does not have permission to evaluate the
component, then an error will be returned and the current filehandle
will be unchanged.
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If the component is a zero length string or if any component does not
obey the UTF-8 definition, the error NFS4ERR_INVAL will be returned.
22.13.5. IMPLEMENTATION
If the client wants to achieve the effect of a multi-component
lookup, it may construct a COMPOUND request such as (and obtain each
filehandle):
PUTFH (directory filehandle)
LOOKUP "pub"
GETFH
LOOKUP "foo"
GETFH
LOOKUP "bar"
GETFH
NFS version 4 servers depart from the semantics of previous NFS
versions in allowing LOOKUP requests to cross mountpoints on the
server. The client can detect a mountpoint crossing by comparing the
fsid attribute of the directory with the fsid attribute of the
directory looked up. If the fsids are different then the new
directory is a server mountpoint. UNIX clients that detect a
mountpoint crossing will need to mount the server's filesystem. This
needs to be done to maintain the file object identity checking
mechanisms common to UNIX clients.
Servers that limit NFS access to "shares" or "exported" filesystems
should provide a pseudo-filesystem into which the exported
filesystems can be integrated, so that clients can browse the
server's name space. The clients view of a pseudo filesystem will be
limited to paths that lead to exported filesystems.
Note: previous versions of the protocol assigned special semantics to
the names "." and "..". NFS version 4 assigns no special semantics
to these names. The LOOKUPP operator must be used to lookup a parent
directory.
Note that this operation does not follow symbolic links. The client
is responsible for all parsing of filenames including filenames that
are modified by symbolic links encountered during the lookup process.
If the current filehandle supplied is not a directory but a symbolic
link, the error NFS4ERR_SYMLINK is returned as the error. For all
other non-directory file types, the error NFS4ERR_NOTDIR is returned.
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22.13.6. ERRORS
NFS4ERR_ACCESS NFS4ERR_BADCHAR NFS4ERR_BADHANDLE NFS4ERR_BADNAME
NFS4ERR_BADXDR NFS4ERR_FHEXPIRED NFS4ERR_INVAL NFS4ERR_IO
NFS4ERR_MOVED NFS4ERR_NAMETOOLONG NFS4ERR_NOENT NFS4ERR_NOFILEHANDLE
NFS4ERR_NOTDIR NFS4ERR_RESOURCE NFS4ERR_SERVERFAULT NFS4ERR_STALE
NFS4ERR_SYMLINK NFS4ERR_WRONGSEC
22.14. Operation 16: LOOKUPP - Lookup Parent Directory
22.14.1. SYNOPSIS
(cfh) -> (cfh)
22.14.2. ARGUMENTS
/* CURRENT_FH: object */
void;
22.14.3. RESULTS
struct LOOKUPP4res {
/* CURRENT_FH: directory */
nfsstat4 status;
};
22.14.4. DESCRIPTION
The current filehandle is assumed to refer to a regular directory or
a named attribute directory. LOOKUPP assigns the filehandle for its
parent directory to be the current filehandle. If there is no parent
directory an NFS4ERR_NOENT error must be returned. Therefore,
NFS4ERR_NOENT will be returned by the server when the current
filehandle is at the root or top of the server's file tree.
As for LOOKUP, LOOKUPP will also cross mountpoints.
If the current filehandle is not a directory or named attribute
directory, the error NFS4ERR_NOTDIR is returned.
If the requester's security flavor does not match that configured for
the parent directory, then the server SHOULD return NFS4ERR_WRONGSEC
(a future minor revision of NFSv4 may upgrade this to MUST) in the
LOOKUPP response. However, if the server does so, it MUST support
the new SECINFO_NO_NAME operation, so that the client can gracefully
determine the correct security flavor. See the discussion of the
SECINFO_NO_NAME operation for a description.
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22.14.5. IMPLEMENTATION
22.14.6. ERRORS
NFS4ERR_ACCESS NFS4ERR_BADHANDLE NFS4ERR_FHEXPIRED NFS4ERR_IO
NFS4ERR_MOVED NFS4ERR_NOENT NFS4ERR_NOFILEHANDLE NFS4ERR_NOTDIR
NFS4ERR_RESOURCE NFS4ERR_SERVERFAULT NFS4ERR_STALE NFS4ERR_WRONGSEC
22.15. Operation 17: NVERIFY - Verify Difference in Attributes
22.15.1. SYNOPSIS
(cfh), fattr -> -
22.15.2. ARGUMENTS
struct NVERIFY4args {
/* CURRENT_FH: object */
fattr4 obj_attributes;
};
22.15.3. RESULTS
struct NVERIFY4res {
nfsstat4 status;
};
22.15.4. DESCRIPTION
This operation is used to prefix a sequence of operations to be
performed if one or more attributes have changed on some filesystem
object. If all the attributes match then the error NFS4ERR_SAME must
be returned.
On success, the current filehandle retains its value.
22.15.5. IMPLEMENTATION
This operation is useful as a cache validation operator. If the
object to which the attributes belong has changed then the following
operations may obtain new data associated with that object. For
instance, to check if a file has been changed and obtain new data if
it has:
PUTFH (public)
LOOKUP "foobar"
NVERIFY attrbits attrs
READ 0 32767
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In the case that a recommended attribute is specified in the NVERIFY
operation and the server does not support that attribute for the
filesystem object, the error NFS4ERR_ATTRNOTSUPP is returned to the
client.
When the attribute rdattr_error or any write-only attribute (e.g.
time_modify_set) is specified, the error NFS4ERR_INVAL is returned to
the client.
22.15.6. ERRORS
NFS4ERR_ACCESS NFS4ERR_ATTRNOTSUPP NFS4ERR_BADCHAR NFS4ERR_BADHANDLE
NFS4ERR_BADXDR NFS4ERR_DELAY NFS4ERR_FHEXPIRED NFS4ERR_INVAL
NFS4ERR_IO NFS4ERR_MOVED NFS4ERR_NOFILEHANDLE NFS4ERR_RESOURCE
NFS4ERR_SAME NFS4ERR_SERVERFAULT NFS4ERR_STALE
22.16. Operation 18: OPEN - Open a Regular File
22.16.1. SYNOPSIS
(cfh), seqid, share_access, share_deny, owner, openhow, claim ->
(cfh), stateid, cinfo, rflags, open_confirm, attrset delegation
22.16.2. ARGUMENTS
struct OPEN4args {
seqid4 seqid;
uint32_t share_access;
uint32_t share_deny;
open_owner4 owner;
openflag4 openhow;
open_claim4 claim;
};
enum createmode4 {
UNCHECKED4 = 0,
GUARDED4 = 1,
EXCLUSIVE4 = 2
};
union createhow4 switch (createmode4 mode) {
case UNCHECKED4:
case GUARDED4:
fattr4 createattrs;
case EXCLUSIVE4:
verifier4 createverf;
};
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enum opentype4 {
OPEN4_NOCREATE = 0,
OPEN4_CREATE = 1
};
union openflag4 switch (opentype4 opentype) {
case OPEN4_CREATE:
createhow4 how;
default:
void;
};
/* Next definitions used for OPEN delegation */
enum limit_by4 {
NFS_LIMIT_SIZE = 1,
NFS_LIMIT_BLOCKS = 2
/* others as needed */
};
struct nfs_modified_limit4 {
uint32_t num_blocks;
uint32_t bytes_per_block;
};
union nfs_space_limit4 switch (limit_by4 limitby) {
/* limit specified as file size */
case NFS_LIMIT_SIZE:
uint64_t filesize;
/* limit specified by number of blocks */
case NFS_LIMIT_BLOCKS:
nfs_modified_limit4 mod_blocks;
} ;
enum open_delegation_type4 {
OPEN_DELEGATE_NONE = 0,
OPEN_DELEGATE_READ = 1,
OPEN_DELEGATE_WRITE = 2
};
enum open_claim_type4 {
CLAIM_NULL = 0,
CLAIM_PREVIOUS = 1,
CLAIM_DELEGATE_CUR = 2,
CLAIM_DELEGATE_PREV = 3
};
struct open_claim_delegate_cur4 {
stateid4 delegate_stateid;
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component4 file;
};
union open_claim4 switch (open_claim_type4 claim) {
/*
* No special rights to file. Ordinary OPEN of the specified file.
*/
case CLAIM_NULL:
/* CURRENT_FH: directory */
component4 file;
/*
* Right to the file established by an open previous to server
* reboot. File identified by filehandle obtained at that time
* rather than by name.
*/
case CLAIM_PREVIOUS:
/* CURRENT_FH: file being reclaimed */
open_delegation_type4 delegate_type;
/*
* Right to file based on a delegation granted by the server.
* File is specified by name.
*/
case CLAIM_DELEGATE_CUR:
/* CURRENT_FH: directory */
open_claim_delegate_cur4 delegate_cur_info;
/* Right to file based on a delegation granted to a previous boot
* instance of the client. File is specified by name.
*/
case CLAIM_DELEGATE_PREV:
/* CURRENT_FH: directory */
component4 file_delegate_prev;
};
22.16.3. RESULTS
struct open_read_delegation4 {
stateid4 stateid; /* Stateid for delegation*/
bool recall; /* Pre-recalled flag for
delegations obtained
by reclaim
(CLAIM_PREVIOUS) */
nfsace4 permissions; /* Defines users who don't
need an ACCESS call to
open for read */
};
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struct open_write_delegation4 {
stateid4 stateid; /* Stateid for delegation*/
bool recall; /* Pre-recalled flag for
delegations obtained
by reclaim
(CLAIM_PREVIOUS) */
nfs_space_limit4 space_limit; /* Defines condition that
the client must check to
determine whether the
file needs to be flushed
to the server on close.
*/
nfsace4 permissions; /* Defines users who don't
need an ACCESS call as
part of a delegated
open. */
};
union open_delegation4
switch (open_delegation_type4 delegation_type) {
case OPEN_DELEGATE_NONE:
void;
case OPEN_DELEGATE_READ:
open_read_delegation4 read;
case OPEN_DELEGATE_WRITE:
open_write_delegation4 write;
};
const OPEN4_RESULT_CONFIRM = 0x00000002;
const OPEN4_RESULT_LOCKTYPE_POSIX = 0x00000004;
struct OPEN4resok {
stateid4 stateid; /* Stateid for open */
change_info4 cinfo; /* Directory Change Info */
uint32_t rflags; /* Result flags */
bitmap4 attrset; /* attributes on create */
open_delegation4 delegation; /* Info on any open
delegation */
};
union OPEN4res switch (nfsstat4 status) {
case NFS4_OK:
/* CURRENT_FH: opened file */
OPEN4resok resok4;
default:
void;
};
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22.16.4. DESCRIPTION
The OPEN operation creates and/or opens a regular file in a directory
with the provided name. If the file does not exist at the server and
creation is desired, specification of the method of creation is
provided by the openhow parameter. The client has the choice of
three creation methods: UNCHECKED, GUARDED, or EXCLUSIVE.
If the current filehandle is a named attribute directory, OPEN will
then create or open a named attribute file. Note that exclusive
create of a named attribute is not supported. If the createmode is
EXCLUSIVE4 and the current filehandle is a named attribute directory,
the server will return EINVAL.
UNCHECKED means that the file should be created if a file of that
name does not exist and encountering an existing regular file of that
name is not an error. For this type of create, createattrs specifies
the initial set of attributes for the file. The set of attributes
may include any writable attribute valid for regular files. When an
UNCHECKED create encounters an existing file, the attributes
specified by createattrs are not used, except that when an size of
zero is specified, the existing file is truncated. If GUARDED is
specified, the server checks for the presence of a duplicate object
by name before performing the create. If a duplicate exists, an
error of NFS4ERR_EXIST is returned as the status. If the object does
not exist, the request is performed as described for UNCHECKED. For
each of these cases (UNCHECKED and GUARDED) where the operation is
successful, the server will return to the client an attribute mask
signifying which attributes were successfully set for the object.
EXCLUSIVE specifies that the server is to follow exclusive creation
semantics, using the verifier to ensure exclusive creation of the
target. The server should check for the presence of a duplicate
object by name. If the object does not exist, the server creates the
object and stores the verifier with the object. If the object does
exist and the stored verifier matches the client provided verifier,
the server uses the existing object as the newly created object. If
the stored verifier does not match, then an error of NFS4ERR_EXIST is
returned. No attributes may be provided in this case, since the
server may use an attribute of the target object to store the
verifier. If the server uses an attribute to store the exclusive
create verifier, it will signify which attribute by setting the
appropriate bit in the attribute mask that is returned in the
results.
For the target directory, the server returns change_info4 information
in cinfo. With the atomic field of the change_info4 struct, the
server will indicate if the before and after change attributes were
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obtained atomically with respect to the link creation.
Upon successful creation, the current filehandle is replaced by that
of the new object.
The OPEN operation provides for Windows share reservation capability
with the use of the share_access and share_deny fields of the OPEN
arguments. The client specifies at OPEN the required share_access
and share_deny modes. For clients that do not directly support
SHAREs (i.e. UNIX), the expected deny value is DENY_NONE. In the
case that there is a existing SHARE reservation that conflicts with
the OPEN request, the server returns the error NFS4ERR_SHARE_DENIED.
For a complete SHARE request, the client must provide values for the
owner and seqid fields for the OPEN argument. For additional
discussion of SHARE semantics see the section on 'Share
Reservations'.
In the case that the client is recovering state from a server
failure, the claim field of the OPEN argument is used to signify that
the request is meant to reclaim state previously held.
The "claim" field of the OPEN argument is used to specify the file to
be opened and the state information which the client claims to
possess. There are four basic claim types which cover the various
situations for an OPEN. They are as follows:
+---------------------+---------------------------------------------+
| open type | description |
+---------------------+---------------------------------------------+
| CLAIM_NULL | For the client, this is a new OPEN request |
| | and there is no previous state associate |
| | with the file for the client. |
| CLAIM_PREVIOUS | The client is claiming basic OPEN state for |
| | a file that was held previous to a server |
| | reboot. Generally used when a server is |
| | returning persistent filehandles; the |
| | client may not have the file name to |
| | reclaim the OPEN. |
| CLAIM_DELEGATE_CUR | The client is claiming a delegation for |
| | OPEN as granted by the server. Generally |
| | this is done as part of recalling a |
| | delegation. |
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| CLAIM_DELEGATE_PREV | The client is claiming a delegation granted |
| | to a previous client instance; used after |
| | the client reboots. The server MAY support |
| | CLAIM_DELEGATE_PREV. If it does support |
| | CLAIM_DELEGATE_PREV, SETCLIENTID_CONFIRM |
| | MUST NOT remove the client's delegation |
| | state, and the server MUST support the |
| | DELEGPURGE operation. |
+---------------------+---------------------------------------------+
For OPEN requests whose claim type is other than CLAIM_PREVIOUS (i.e.
requests other than those devoted to reclaiming opens after a server
reboot) that reach the server during its grace or lease expiration
period, the server returns an error of NFS4ERR_GRACE.
For any OPEN request, the server may return an open delegation, which
allows further opens and closes to be handled locally on the client
as described in the section Open Delegation. Note that delegation is
up to the server to decide. The client should never assume that
delegation will or will not be granted in a particular instance. It
should always be prepared for either case. A partial exception is
the reclaim (CLAIM_PREVIOUS) case, in which a delegation type is
claimed. In this case, delegation will always be granted, although
the server may specify an immediate recall in the delegation
structure.
The rflags returned by a successful OPEN allow the server to return
information governing how the open file is to be handled.
OPEN4_RESULT_CONFIRM indicates that the client MUST execute an
OPEN_CONFIRM operation before using the open file.
OPEN4_RESULT_LOCKTYPE_POSIX indicates the server's file locking
behavior supports the complete set of Posix locking techniques. From
this the client can choose to manage file locking state in a way to
handle a mis-match of file locking management.
If the component is of zero length, NFS4ERR_INVAL will be returned.
The component is also subject to the normal UTF-8, character support,
and name checks. See the section "UTF-8 Related Errors" for further
discussion.
When an OPEN is done and the specified lockowner already has the
resulting filehandle open, the result is to "OR" together the new
share and deny status together with the existing status. In this
case, only a single CLOSE need be done, even though multiple OPENs
were completed. When such an OPEN is done, checking of share
reservations for the new OPEN proceeds normally, with no exception
for the existing OPEN held by the same lockowner.
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If the underlying filesystem at the server is only accessible in a
read-only mode and the OPEN request has specified ACCESS_WRITE or
ACCESS_BOTH, the server will return NFS4ERR_ROFS to indicate a read-
only filesystem.
As with the CREATE operation, the server MUST derive the owner, owner
ACE, group, or group ACE if any of the four attributes are required
and supported by the server's filesystem. For an OPEN with the
EXCLUSIVE4 createmode, the server has no choice, since such OPEN
calls do not include the createattrs field. Conversely, if
createattrs is specified, and includes owner or group (or
corresponding ACEs) that the principal in the RPC call's credentials
does not have authorization to create files for, then the server may
return NFS4ERR_PERM.
In the case of a OPEN which specifies a size of zero (e.g.
truncation) and the file has named attributes, the named attributes
are left as is. They are not removed.
22.16.5. IMPLEMENTATION
The OPEN operation contains support for EXCLUSIVE create. The
mechanism is similar to the support in NFS version 3 [RFC1813]. As
in NFS version 3, this mechanism provides reliable exclusive
creation. Exclusive create is invoked when the how parameter is
EXCLUSIVE. In this case, the client provides a verifier that can
reasonably be expected to be unique. A combination of a client
identifier, perhaps the client network address, and a unique number
generated by the client, perhaps the RPC transaction identifier, may
be appropriate.
If the object does not exist, the server creates the object and
stores the verifier in stable storage. For filesystems that do not
provide a mechanism for the storage of arbitrary file attributes, the
server may use one or more elements of the object meta-data to store
the verifier. The verifier must be stored in stable storage to
prevent erroneous failure on retransmission of the request. It is
assumed that an exclusive create is being performed because exclusive
semantics are critical to the application. Because of the expected
usage, exclusive CREATE does not rely solely on the normally volatile
duplicate request cache for storage of the verifier. The duplicate
request cache in volatile storage does not survive a crash and may
actually flush on a long network partition, opening failure windows.
In the UNIX local filesystem environment, the expected storage
location for the verifier on creation is the meta-data (time stamps)
of the object. For this reason, an exclusive object create may not
include initial attributes because the server would have nowhere to
store the verifier.
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If the server can not support these exclusive create semantics,
possibly because of the requirement to commit the verifier to stable
storage, it should fail the OPEN request with the error,
NFS4ERR_NOTSUPP.
During an exclusive CREATE request, if the object already exists, the
server reconstructs the object's verifier and compares it with the
verifier in the request. If they match, the server treats the
request as a success. The request is presumed to be a duplicate of
an earlier, successful request for which the reply was lost and that
the server duplicate request cache mechanism did not detect. If the
verifiers do not match, the request is rejected with the status,
NFS4ERR_EXIST.
Once the client has performed a successful exclusive create, it must
issue a SETATTR to set the correct object attributes. Until it does
so, it should not rely upon any of the object attributes, since the
server implementation may need to overload object meta-data to store
the verifier. The subsequent SETATTR must not occur in the same
COMPOUND request as the OPEN. This separation will guarantee that
the exclusive create mechanism will continue to function properly in
the face of retransmission of the request.
Use of the GUARDED attribute does not provide exactly-once semantics.
In particular, if a reply is lost and the server does not detect the
retransmission of the request, the operation can fail with
NFS4ERR_EXIST, even though the create was performed successfully.
The client would use this behavior in the case that the application
has not requested an exclusive create but has asked to have the file
truncated when the file is opened. In the case of the client timing
out and retransmitting the create request, the client can use GUARDED
to prevent against a sequence like: create, write, create
(retransmitted) from occurring.
For SHARE reservations, the client must specify a value for
share_access that is one of READ, WRITE, or BOTH. For share_deny,
the client must specify one of NONE, READ, WRITE, or BOTH. If the
client fails to do this, the server must return NFS4ERR_INVAL.
Based on the share_access value (READ, WRITE, or BOTH) the client
should check that the requester has the proper access rights to
perform the specified operation. This would generally be the results
of applying the ACL access rules to the file for the current
requester. However, just as with the ACCESS operation, the client
should not attempt to second-guess the server's decisions, as access
rights may change and may be subject to server administrative
controls outside the ACL framework. If the requester is not
authorized to READ or WRITE (depending on the share_access value),
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the server must return NFS4ERR_ACCESS. Note that since the NFS
version 4 protocol does not impose any requirement that READs and
WRITEs issued for an open file have the same credentials as the OPEN
itself, the server still must do appropriate access checking on the
READs and WRITEs themselves.
If the component provided to OPEN is a symbolic link, the error
NFS4ERR_SYMLINK will be returned to the client. If the current
filehandle is not a directory, the error NFS4ERR_NOTDIR will be
returned.
"WARNING TO CLIENT IMPLEMENTORS"
OPEN resembles LOOKUP in that it generates a filehandle for the
client to use. Unlike LOOKUP though, OPEN creates server state on
the filehandle. In normal circumstances, the client can only release
this state with a CLOSE operation. CLOSE uses the current filehandle
to determine which file to close. Therefore the client MUST follow
every OPEN operation with a GETFH operation in the same COMPOUND
procedure. This will supply the client with the filehandle such that
CLOSE can be used appropriately.
Simply waiting for the lease on the file to expire is insufficient
because the server may maintain the state indefinitely as long as
another client does not attempt to make a conflicting access to the
same file.
22.16.6. ERRORS
NFS4ERR_ACCESS NFS4ERR_ADMIN_REVOKED NFS4ERR_ATTRNOTSUPP
NFS4ERR_BADCHAR NFS4ERR_BADHANDLE NFS4ERR_BADNAME NFS4ERR_BADOWNER
NFS4ERR_BAD_SEQID NFS4ERR_BADXDR NFS4ERR_DELAY NFS4ERR_DQUOT
NFS4ERR_EXIST NFS4ERR_EXPIRED NFS4ERR_FHEXPIRED NFS4ERR_GRACE
NFS4ERR_IO NFS4ERR_INVAL NFS4ERR_ISDIR NFS4ERR_LEASE_MOVED
NFS4ERR_MOVED NFS4ERR_NAMETOOLONG NFS4ERR_NOENT NFS4ERR_NOFILEHANDLE
NFS4ERR_NOSPC NFS4ERR_NOTDIR NFS4ERR_NO_GRACE NFS4ERR_PERM
NFS4ERR_RECLAIM_BAD NFS4ERR_RECLAIM_CONFLICT NFS4ERR_RESOURCE
NFS4ERR_ROFS NFS4ERR_SERVERFAULT NFS4ERR_SHARE_DENIED NFS4ERR_STALE
NFS4ERR_STALE_CLIENTID NFS4ERR_SYMLINK NFS4ERR_WRONGSEC
22.17. Operation 19: OPENATTR - Open Named Attribute Directory
22.17.1. SYNOPSIS
(cfh) createdir -> (cfh)
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22.17.2. ARGUMENTS
struct OPENATTR4args {
/* CURRENT_FH: object */
bool createdir;
};
22.17.3. RESULTS
struct OPENATTR4res {
/* CURRENT_FH: named attr directory*/
nfsstat4 status;
};
22.17.4. DESCRIPTION
The OPENATTR operation is used to obtain the filehandle of the named
attribute directory associated with the current filehandle. The
result of the OPENATTR will be a filehandle to an object of type
NF4ATTRDIR. From this filehandle, READDIR and LOOKUP operations can
be used to obtain filehandles for the various named attributes
associated with the original filesystem object. Filehandles returned
within the named attribute directory will have a type of
NF4NAMEDATTR.
The createdir argument allows the client to signify if a named
attribute directory should be created as a result of the OPENATTR
operation. Some clients may use the OPENATTR operation with a value
of FALSE for createdir to determine if any named attributes exist for
the object. If none exist, then NFS4ERR_NOENT will be returned. If
createdir has a value of TRUE and no named attribute directory
exists, one is created. The creation of a named attribute directory
assumes that the server has implemented named attribute support in
this fashion and is not required to do so by this definition.
22.17.5. IMPLEMENTATION
If the server does not support named attributes for the current
filehandle, an error of NFS4ERR_NOTSUPP will be returned to the
client.
22.17.6. ERRORS
NFS4ERR_ACCESS NFS4ERR_BADHANDLE NFS4ERR_BADXDR NFS4ERR_DELAY
NFS4ERR_DQUOT NFS4ERR_FHEXPIRED NFS4ERR_IO NFS4ERR_MOVED
NFS4ERR_NOENT NFS4ERR_NOFILEHANDLE NFS4ERR_NOSPC NFS4ERR_NOTSUPP
NFS4ERR_RESOURCE NFS4ERR_ROFS NFS4ERR_SERVERFAULT NFS4ERR_STALE
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22.18. Operation 20: OPEN_CONFIRM - Confirm Open
22.18.1. SYNOPSIS
(cfh), seqid, stateid-> stateid
22.18.2. ARGUMENTS
struct OPEN_CONFIRM4args {
/* CURRENT_FH: opened file */
stateid4 open_stateid;
seqid4 seqid;
};
22.18.3. RESULTS
struct OPEN_CONFIRM4resok {
stateid4 open_stateid;
};
union OPEN_CONFIRM4res switch (nfsstat4 status) {
case NFS4_OK:
OPEN_CONFIRM4resok resok4;
default:
void;
};
22.18.4. DESCRIPTION
This operation is used to confirm the sequence id usage for the first
time that a open_owner is used by a client. The stateid returned
from the OPEN operation is used as the argument for this operation
along with the next sequence id for the open_owner. The sequence id
passed to the OPEN_CONFIRM must be 1 (one) greater than the seqid
passed to the OPEN operation from which the open_confirm value was
obtained. If the server receives an unexpected sequence id with
respect to the original open, then the server assumes that the client
will not confirm the original OPEN and all state associated with the
original OPEN is released by the server.
On success, the current filehandle retains its value.
22.18.5. IMPLEMENTATION
A given client might generate many open_owner4 data structures for a
given clientid. The client will periodically either dispose of its
open_owner4s or stop using them for indefinite periods of time. The
latter situation is why the NFS version 4 protocol does not have an
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explicit operation to exit an open_owner4: such an operation is of no
use in that situation. Instead, to avoid unbounded memory use, the
server needs to implement a strategy for disposing of open_owner4s
that have no current lock, open, or delegation state for any files
and have not been used recently. The time period used to determine
when to dispose of open_owner4s is an implementation choice. The
time period should certainly be no less than the lease time plus any
grace period the server wishes to implement beyond a lease time. The
OPEN_CONFIRM operation allows the server to safely dispose of unused
open_owner4 data structures.
In the case that a client issues an OPEN operation and the server no
longer has a record of the open_owner4, the server needs ensure that
this is a new OPEN and not a replay or retransmission.
Servers must not require confirmation on OPENs that grant delegations
or are doing reclaim operations. See section "Use of Open
Confirmation" for details. The server can easily avoid this by
noting whether it has disposed of one open_owner4 for the given
clientid. If the server does not support delegation, it might simply
maintain a single bit that notes whether any open_owner4 (for any
client) has been disposed of.
The server must hold unconfirmed OPEN state until one of three events
occur. First, the client sends an OPEN_CONFIRM request with the
appropriate sequence id and stateid within the lease period. In this
case, the OPEN state on the server goes to confirmed, and the
open_owner4 on the server is fully established.
Second, the client sends another OPEN request with a sequence id that
is incorrect for the open_owner4 (out of sequence). In this case,
the server assumes the second OPEN request is valid and the first one
is a replay. The server cancels the OPEN state of the first OPEN
request, establishes an unconfirmed OPEN state for the second OPEN
request, and responds to the second OPEN request with an indication
that an OPEN_CONFIRM is needed. The process then repeats itself.
While there is a potential for a denial of service attack on the
client, it is mitigated if the client and server require the use of a
security flavor based on Kerberos V5, LIPKEY, or some other flavor
that uses cryptography.
What if the server is in the unconfirmed OPEN state for a given
open_owner4, and it receives an operation on the open_owner4 that has
a stateid but the operation is not OPEN, or it is OPEN_CONFIRM but
with the wrong stateid? Then, even if the seqid is correct, the
server returns NFS4ERR_BAD_STATEID, because the server assumes the
operation is a replay: if the server has no established OPEN state,
then there is no way, for example, a LOCK operation could be valid.
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Third, neither of the two aforementioned events occur for the
open_owner4 within the lease period. In this case, the OPEN state is
canceled and disposal of the open_owner4 can occur.
22.18.6. ERRORS
NFS4ERR_ADMIN_REVOKED NFS4ERR_BADHANDLE NFS4ERR_BAD_SEQID
NFS4ERR_BAD_STATEID NFS4ERR_BADXDR NFS4ERR_EXPIRED NFS4ERR_FHEXPIRED
NFS4ERR_INVAL NFS4ERR_ISDIR NFS4ERR_MOVED NFS4ERR_NOFILEHANDLE
NFS4ERR_OLD_STATEID NFS4ERR_RESOURCE NFS4ERR_SERVERFAULT
NFS4ERR_STALE NFS4ERR_STALE_STATEID
22.19. Operation 21: OPEN_DOWNGRADE - Reduce Open File Access
22.19.1. SYNOPSIS
(cfh), stateid, seqid, access, deny -> stateid
22.19.2. ARGUMENTS
struct OPEN_DOWNGRADE4args {
/* CURRENT_FH: opened file */
stateid4 stateid;
seqid4 seqid;
uint32_t share_access;
uint32_t share_deny;
};
22.19.3. RESULTS
struct OPEN_DOWNGRADE4resok {
stateid4 stateid;
};
union OPEN_DOWNGRADE4res switch(nfsstat4 status) {
case NFS4_OK:
OPEN_DOWNGRADE4resok resok4;
default:
void;
};
22.19.4. DESCRIPTION
This operation is used to adjust the share_access and share_deny bits
for a given open. This is necessary when a given lockowner opens the
same file multiple times with different share_access and share_deny
flags. In this situation, a close of one of the opens may change the
appropriate share_access and share_deny flags to remove bits
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associated with opens no longer in effect.
The share_access and share_deny bits specified in this operation
replace the current ones for the specified open file. The
share_access and share_deny bits specified must be exactly equal to
the union of the share_access and share_deny bits specified for some
subset of the OPENs in effect for current openowner on the current
file. If that constraint is not respected, the error NFS4ERR_INVAL
should be returned. Since share_access and share_deny bits are
subsets of those already granted, it is not possible for this request
to be denied because of conflicting share reservations.
On success, the current filehandle retains its value.
22.19.5. ERRORS
NFS4ERR_ADMIN_REVOKED NFS4ERR_BADHANDLE NFS4ERR_BAD_SEQID
NFS4ERR_BAD_STATEID NFS4ERR_BADXDR NFS4ERR_EXPIRED NFS4ERR_FHEXPIRED
NFS4ERR_INVAL NFS4ERR_MOVED NFS4ERR_NOFILEHANDLE NFS4ERR_OLD_STATEID
NFS4ERR_RESOURCE NFS4ERR_SERVERFAULT NFS4ERR_STALE
NFS4ERR_STALE_STATEID
22.20. Operation 22: PUTFH - Set Current Filehandle
22.20.1. SYNOPSIS
filehandle -> (cfh)
22.20.2. ARGUMENTS
struct PUTFH4args {
nfs_fh4 object;
};
22.20.3. RESULTS
struct PUTFH4res {
/* CURRENT_FH: */
nfsstat4 status;
};
22.20.4. DESCRIPTION
Replaces the current filehandle with the filehandle provided as an
argument.
If the security mechanism used by the requester does not meet the
requirements of the filehandle provided to this operation, the server
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MUST return NFS4ERR_WRONGSEC.
22.20.5. IMPLEMENTATION
Commonly used as the first operator in an NFS request to set the
context for following operations.
22.20.6. ERRORS
NFS4ERR_BADHANDLE NFS4ERR_BADXDR NFS4ERR_FHEXPIRED NFS4ERR_MOVED
NFS4ERR_RESOURCE NFS4ERR_SERVERFAULT NFS4ERR_STALE NFS4ERR_WRONGSEC
22.21. Operation 24: PUTROOTFH - Set Root Filehandle
22.21.1. SYNOPSIS
- -> (cfh)
22.21.2. ARGUMENTS
void;
22.21.3. RESULTS
struct PUTROOTFH4res {
/* CURRENT_FH: root fh */
nfsstat4 status;
};
22.21.4. DESCRIPTION
Replaces the current filehandle with the filehandle that represents
the root of the server's name space. From this filehandle a LOOKUP
operation can locate any other filehandle on the server. This
filehandle may be different from the "public" filehandle which may be
associated with some other directory on the server.
22.21.5. IMPLEMENTATION
Commonly used as the first operator in an NFS request to set the
context for following operations.
22.21.6. ERRORS
NFS4ERR_RESOURCE NFS4ERR_SERVERFAULT NFS4ERR_WRONGSEC
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22.22. Operation 25: READ - Read from File
22.22.1. SYNOPSIS
(cfh), stateid, offset, count -> eof, data
22.22.2. ARGUMENTS
struct READ4args {
/* CURRENT_FH: file */
stateid4 stateid;
offset4 offset;
count4 count;
};
22.22.3. RESULTS
struct READ4resok {
bool eof;
opaque data<>;
};
union READ4res switch (nfsstat4 status) {
case NFS4_OK:
READ4resok resok4;
default:
void;
};
22.22.4. DESCRIPTION
The READ operation reads data from the regular file identified by the
current filehandle.
The client provides an offset of where the READ is to start and a
count of how many bytes are to be read. An offset of 0 (zero) means
to read data starting at the beginning of the file. If offset is
greater than or equal to the size of the file, the status, NFS4_OK,
is returned with a data length set to 0 (zero) and eof is set to
TRUE. The READ is subject to access permissions checking.
If the client specifies a count value of 0 (zero), the READ succeeds
and returns 0 (zero) bytes of data again subject to access
permissions checking. The server may choose to return fewer bytes
than specified by the client. The client needs to check for this
condition and handle the condition appropriately.
The stateid value for a READ request represents a value returned from
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a previous record lock or share reservation request. The stateid is
used by the server to verify that the associated share reservation
and any record locks are still valid and to update lease timeouts for
the client.
If the read ended at the end-of-file (formally, in a correctly formed
READ request, if offset + count is equal to the size of the file), or
the read request extends beyond the size of the file (if offset +
count is greater than the size of the file), eof is returned as TRUE;
otherwise it is FALSE. A successful READ of an empty file will
always return eof as TRUE.
If the current filehandle is not a regular file, an error will be
returned to the client. In the case the current filehandle
represents a directory, NFS4ERR_ISDIR is return; otherwise,
NFS4ERR_INVAL is returned.
For a READ with a stateid value of all bits 0, the server MAY allow
the READ to be serviced subject to mandatory file locks or the
current share deny modes for the file. For a READ with a stateid
value of all bits 1, the server MAY allow READ operations to bypass
locking checks at the server.
On success, the current filehandle retains its value.
22.22.5. IMPLEMENTATION
It is possible for the server to return fewer than count bytes of
data. If the server returns less than the count requested and eof is
set to FALSE, the client should issue another READ to get the
remaining data. A server may return less data than requested under
several circumstances. The file may have been truncated by another
client or perhaps on the server itself, changing the file size from
what the requesting client believes to be the case. This would
reduce the actual amount of data available to the client. It is
possible that the server may back off the transfer size and reduce
the read request return. Server resource exhaustion may also occur
necessitating a smaller read return.
If mandatory file locking is on for the file, and if the region
corresponding to the data to be read from file is write locked by an
owner not associated the stateid, the server will return the
NFS4ERR_LOCKED error. The client should try to get the appropriate
read record lock via the LOCK operation before re-attempting the
READ. When the READ completes, the client should release the record
lock via LOCKU.
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22.22.6. ERRORS
NFS4ERR_ACCESS NFS4ERR_ADMIN_REVOKED NFS4ERR_BADHANDLE
NFS4ERR_BAD_STATEID NFS4ERR_BADXDR NFS4ERR_DELAY NFS4ERR_EXPIRED
NFS4ERR_FHEXPIRED NFS4ERR_GRACE NFS4ERR_IO NFS4ERR_INVAL
NFS4ERR_ISDIR NFS4ERR_LEASE_MOVED NFS4ERR_LOCKED NFS4ERR_MOVED
NFS4ERR_NOFILEHANDLE NFS4ERR_NXIO NFS4ERR_OLD_STATEID
NFS4ERR_OPENMODE NFS4ERR_RESOURCE NFS4ERR_SERVERFAULT NFS4ERR_STALE
NFS4ERR_STALE_STATEID
22.23. Operation 26: READDIR - Read Directory
22.23.1. SYNOPSIS
(cfh), cookie, cookieverf, dircount, maxcount, attr_request ->
cookieverf { cookie, name, attrs }
22.23.2. ARGUMENTS
struct READDIR4args {
/* CURRENT_FH: directory */
nfs_cookie4 cookie;
verifier4 cookieverf;
count4 dircount;
count4 maxcount;
bitmap4 attr_request;
};
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22.23.3. RESULTS
struct entry4 {
nfs_cookie4 cookie;
component4 name;
fattr4 attrs;
entry4 *nextentry;
};
struct dirlist4 {
entry4 *entries;
bool eof;
};
struct READDIR4resok {
verifier4 cookieverf;
dirlist4 reply;
};
union READDIR4res switch (nfsstat4 status) {
case NFS4_OK:
READDIR4resok resok4;
default:
void;
};
22.23.4. DESCRIPTION
The READDIR operation retrieves a variable number of entries from a
filesystem directory and returns client requested attributes for each
entry along with information to allow the client to request
additional directory entries in a subsequent READDIR.
The arguments contain a cookie value that represents where the
READDIR should start within the directory. A value of 0 (zero) for
the cookie is used to start reading at the beginning of the
directory. For subsequent READDIR requests, the client specifies a
cookie value that is provided by the server on a previous READDIR
request.
The cookieverf value should be set to 0 (zero) when the cookie value
is 0 (zero) (first directory read). On subsequent requests, it
should be a cookieverf as returned by the server. The cookieverf
must match that returned by the READDIR in which the cookie was
acquired. If the server determines that the cookieverf is no longer
valid for the directory, the error NFS4ERR_NOT_SAME must be returned.
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The dircount portion of the argument is a hint of the maximum number
of bytes of directory information that should be returned. This
value represents the length of the names of the directory entries and
the cookie value for these entries. This length represents the XDR
encoding of the data (names and cookies) and not the length in the
native format of the server.
The maxcount value of the argument is the maximum number of bytes for
the result. This maximum size represents all of the data being
returned within the READDIR4resok structure and includes the XDR
overhead. The server may return less data. If the server is unable
to return a single directory entry within the maxcount limit, the
error NFS4ERR_TOOSMALL will be returned to the client.
Finally, attr_request represents the list of attributes to be
returned for each directory entry supplied by the server.
On successful return, the server's response will provide a list of
directory entries. Each of these entries contains the name of the
directory entry, a cookie value for that entry, and the associated
attributes as requested. The "eof" flag has a value of TRUE if there
are no more entries in the directory.
The cookie value is only meaningful to the server and is used as a
"bookmark" for the directory entry. As mentioned, this cookie is
used by the client for subsequent READDIR operations so that it may
continue reading a directory. The cookie is similar in concept to a
READ offset but should not be interpreted as such by the client.
Ideally, the cookie value should not change if the directory is
modified since the client may be caching these values.
In some cases, the server may encounter an error while obtaining the
attributes for a directory entry. Instead of returning an error for
the entire READDIR operation, the server can instead return the
attribute 'fattr4_rdattr_error'. With this, the server is able to
communicate the failure to the client and not fail the entire
operation in the instance of what might be a transient failure.
Obviously, the client must request the fattr4_rdattr_error attribute
for this method to work properly. If the client does not request the
attribute, the server has no choice but to return failure for the
entire READDIR operation.
For some filesystem environments, the directory entries "." and ".."
have special meaning and in other environments, they may not. If the
server supports these special entries within a directory, they should
not be returned to the client as part of the READDIR response. To
enable some client environments, the cookie values of 0, 1, and 2 are
to be considered reserved. Note that the UNIX client will use these
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values when combining the server's response and local representations
to enable a fully formed UNIX directory presentation to the
application.
For READDIR arguments, cookie values of 1 and 2 should not be used
and for READDIR results cookie values of 0, 1, and 2 should not be
returned.
On success, the current filehandle retains its value.
22.23.5. IMPLEMENTATION
The server's filesystem directory representations can differ greatly.
A client's programming interfaces may also be bound to the local
operating environment in a way that does not translate well into the
NFS protocol. Therefore the use of the dircount and maxcount fields
are provided to allow the client the ability to provide guidelines to
the server. If the client is aggressive about attribute collection
during a READDIR, the server has an idea of how to limit the encoded
response. The dircount field provides a hint on the number of
entries based solely on the names of the directory entries. Since it
is a hint, it may be possible that a dircount value is zero. In this
case, the server is free to ignore the dircount value and return
directory information based on the specified maxcount value.
The cookieverf may be used by the server to help manage cookie values
that may become stale. It should be a rare occurrence that a server
is unable to continue properly reading a directory with the provided
cookie/cookieverf pair. The server should make every effort to avoid
this condition since the application at the client may not be able to
properly handle this type of failure.
The use of the cookieverf will also protect the client from using
READDIR cookie values that may be stale. For example, if the file
system has been migrated, the server may or may not be able to use
the same cookie values to service READDIR as the previous server
used. With the client providing the cookieverf, the server is able
to provide the appropriate response to the client. This prevents the
case where the server may accept a cookie value but the underlying
directory has changed and the response is invalid from the client's
context of its previous READDIR.
Since some servers will not be returning "." and ".." entries as has
been done with previous versions of the NFS protocol, the client that
requires these entries be present in READDIR responses must fabricate
them.
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22.23.6. ERRORS
NFS4ERR_ACCESS NFS4ERR_BADHANDLE NFS4ERR_BAD_COOKIE NFS4ERR_BADXDR
NFS4ERR_DELAY NFS4ERR_FHEXPIRED NFS4ERR_INVAL NFS4ERR_IO
NFS4ERR_MOVED NFS4ERR_NOFILEHANDLE NFS4ERR_NOTDIR NFS4ERR_RESOURCE
NFS4ERR_SERVERFAULT NFS4ERR_STALE NFS4ERR_TOOSMALL
22.24. Operation 27: READLINK - Read Symbolic Link
22.24.1. SYNOPSIS
(cfh) -> linktext
22.24.2. ARGUMENTS
/* CURRENT_FH: symlink */
void;
22.24.3. RESULTS
struct READLINK4resok {
linktext4 link;
};
union READLINK4res switch (nfsstat4 status) {
case NFS4_OK:
READLINK4resok resok4;
default:
void;
};
22.24.4. DESCRIPTION
READLINK reads the data associated with a symbolic link. The data is
a UTF-8 string that is opaque to the server. That is, whether
created by an NFS client or created locally on the server, the data
in a symbolic link is not interpreted when created, but is simply
stored.
On success, the current filehandle retains its value.
22.24.5. IMPLEMENTATION
A symbolic link is nominally a pointer to another file. The data is
not necessarily interpreted by the server, just stored in the file.
It is possible for a client implementation to store a path name that
is not meaningful to the server operating system in a symbolic link.
A READLINK operation returns the data to the client for
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interpretation. If different implementations want to share access to
symbolic links, then they must agree on the interpretation of the
data in the symbolic link.
The READLINK operation is only allowed on objects of type NF4LNK.
The server should return the error, NFS4ERR_INVAL, if the object is
not of type, NF4LNK.
22.24.6. ERRORS
NFS4ERR_ACCESS NFS4ERR_BADHANDLE NFS4ERR_DELAY NFS4ERR_FHEXPIRED
NFS4ERR_INVAL NFS4ERR_IO NFS4ERR_ISDIR NFS4ERR_MOVED
NFS4ERR_NOFILEHANDLE NFS4ERR_NOTSUPP NFS4ERR_RESOURCE
NFS4ERR_SERVERFAULT NFS4ERR_STALE
22.25. Operation 28: REMOVE - Remove Filesystem Object
22.25.1. SYNOPSIS
(cfh), filename -> change_info
22.25.2. ARGUMENTS
struct REMOVE4args {
/* CURRENT_FH: directory */
component4 target;
};
22.25.3. RESULTS
struct REMOVE4resok {
change_info4 cinfo;
}
union REMOVE4res switch (nfsstat4 status) {
case NFS4_OK:
REMOVE4resok resok4;
default:
void;
}
22.25.4. DESCRIPTION
The REMOVE operation removes (deletes) a directory entry named by
filename from the directory corresponding to the current filehandle.
If the entry in the directory was the last reference to the
corresponding filesystem object, the object may be destroyed.
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For the directory where the filename was removed, the server returns
change_info4 information in cinfo. With the atomic field of the
change_info4 struct, the server will indicate if the before and after
change attributes were obtained atomically with respect to the
removal.
If the target has a length of 0 (zero), or if target does not obey
the UTF-8 definition, the error NFS4ERR_INVAL will be returned.
On success, the current filehandle retains its value.
22.25.5. IMPLEMENTATION
NFS versions 2 and 3 required a different operator RMDIR for
directory removal and REMOVE for non-directory removal. This allowed
clients to skip checking the file type when being passed a non-
directory delete system call (e.g. unlink() in POSIX) to remove a
directory, as well as the converse (e.g. a rmdir() on a non-
directory) because they knew the server would check the file type.
NFS version 4 REMOVE can be used to delete any directory entry
independent of its file type. The implementor of an NFS version 4
client's entry points from the unlink() and rmdir() system calls
should first check the file type against the types the system call is
allowed to remove before issuing a REMOVE. Alternatively, the
implementor can produce a COMPOUND call that includes a LOOKUP/VERIFY
sequence to verify the file type before a REMOVE operation in the
same COMPOUND call.
The concept of last reference is server specific. However, if the
numlinks field in the previous attributes of the object had the value
1, the client should not rely on referring to the object via a
filehandle. Likewise, the client should not rely on the resources
(disk space, directory entry, and so on) formerly associated with the
object becoming immediately available. Thus, if a client needs to be
able to continue to access a file after using REMOVE to remove it,
the client should take steps to make sure that the file will still be
accessible. The usual mechanism used is to RENAME the file from its
old name to a new hidden name.
If the server finds that the file is still open when the REMOVE
arrives: .in 7 .IP o The server SHOULD NOT delete the file's
directory entry if the file was opened with OPEN4_SHARE_DENY_WRITE or
OPEN4_SHARE_DENY_BOTH. .IP o If the file was not opened with
OPEN4_SHARE_DENY_WRITE or OPEN4_SHARE_DENY_BOTH, the server SHOULD
delete the file's directory entry. However, until last CLOSE of the
file, the server MAY continue to allow access to the file via its
filehandle. .in 5
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22.25.6. ERRORS
NFS4ERR_ACCESS NFS4ERR_BADCHAR NFS4ERR_BADHANDLE NFS4ERR_BADNAME
NFS4ERR_BADXDR NFS4ERR_DELAY NFS4ERR_FHEXPIRED NFS4ERR_FILE_OPEN
NFS4ERR_INVAL NFS4ERR_IO NFS4ERR_MOVED NFS4ERR_NAMETOOLONG
NFS4ERR_NOENT NFS4ERR_NOFILEHANDLE NFS4ERR_NOTDIR NFS4ERR_NOTEMPTY
NFS4ERR_RESOURCE NFS4ERR_ROFS NFS4ERR_SERVERFAULT NFS4ERR_STALE
22.26. Operation 29: RENAME - Rename Directory Entry
22.26.1. SYNOPSIS
(sfh), oldname, (cfh), newname -> source_change_info,
target_change_info
22.26.2. ARGUMENTS
struct RENAME4args {
/* SAVED_FH: source directory */
component4 oldname;
/* CURRENT_FH: target directory */
component4 newname;
};
22.26.3. RESULTS
struct RENAME4resok {
change_info4 source_cinfo;
change_info4 target_cinfo;
};
union RENAME4res switch (nfsstat4 status) {
case NFS4_OK:
RENAME4resok resok4;
default:
void;
};
22.26.4. DESCRIPTION
The RENAME operation renames the object identified by oldname in the
source directory corresponding to the saved filehandle, as set by the
SAVEFH operation, to newname in the target directory corresponding to
the current filehandle. The operation is required to be atomic to
the client. Source and target directories must reside on the same
filesystem on the server. On success, the current filehandle will
continue to be the target directory.
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If the target directory already contains an entry with the name,
newname, the source object must be compatible with the target: either
both are non-directories or both are directories and the target must
be empty. If compatible, the existing target is removed before the
rename occurs (See the IMPLEMENTATION subsection of the section
"Operation 28: REMOVE - Remove Filesystem Object" for client and
server actions whenever a target is removed). If they are not
compatible or if the target is a directory but not empty, the server
will return the error, NFS4ERR_EXIST.
If oldname and newname both refer to the same file (they might be
hard links of each other), then RENAME should perform no action and
return success.
For both directories involved in the RENAME, the server returns
change_info4 information. With the atomic field of the change_info4
struct, the server will indicate if the before and after change
attributes were obtained atomically with respect to the rename.
If the oldname refers to a named attribute and the saved and current
filehandles refer to different filesystem objects, the server will
return NFS4ERR_XDEV just as if the saved and current filehandles
represented directories on different filesystems.
If the oldname or newname has a length of 0 (zero), or if oldname or
newname does not obey the UTF-8 definition, the error NFS4ERR_INVAL
will be returned.
22.26.5. IMPLEMENTATION
The RENAME operation must be atomic to the client. The statement
"source and target directories must reside on the same filesystem on
the server" means that the fsid fields in the attributes for the
directories are the same. If they reside on different filesystems,
the error, NFS4ERR_XDEV, is returned.
Based on the value of the fh_expire_type attribute for the object,
the filehandle may or may not expire on a RENAME. However, server
implementors are strongly encouraged to attempt to keep filehandles
from expiring in this fashion.
On some servers, the file names "." and ".." are illegal as either
oldname or newname, and will result in the error NFS4ERR_BADNAME. In
addition, on many servers the case of oldname or newname being an
alias for the source directory will be checked for. Such servers
will return the error NFS4ERR_INVAL in these cases.
If either of the source or target filehandles are not directories,
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the server will return NFS4ERR_NOTDIR.
22.26.6. ERRORS
NFS4ERR_ACCESS NFS4ERR_BADCHAR NFS4ERR_BADHANDLE NFS4ERR_BADNAME
NFS4ERR_BADXDR NFS4ERR_DELAY NFS4ERR_DQUOT NFS4ERR_EXIST
NFS4ERR_FHEXPIRED NFS4ERR_FILE_OPEN NFS4ERR_INVAL NFS4ERR_IO
NFS4ERR_MOVED NFS4ERR_NAMETOOLONG NFS4ERR_NOENT NFS4ERR_NOFILEHANDLE
NFS4ERR_NOSPC NFS4ERR_NOTDIR NFS4ERR_NOTEMPTY NFS4ERR_RESOURCE
NFS4ERR_ROFS NFS4ERR_SERVERFAULT NFS4ERR_STALE NFS4ERR_WRONGSEC
NFS4ERR_XDEV
22.27. Operation 30: RENEW - Renew a Lease
22.27.1. SYNOPSIS
clientid -> ()
22.27.2. ARGUMENTS
struct RENEW4args {
clientid4 clientid;
};
22.27.3. RESULTS
struct RENEW4res {
nfsstat4 status;
};
22.27.4. DESCRIPTION
The RENEW operation is used by the client to renew leases which it
currently holds at a server. In processing the RENEW request, the
server renews all leases associated with the client. The associated
leases are determined by the clientid provided via the SETCLIENTID
operation.
22.27.5. IMPLEMENTATION
When the client holds delegations, it needs to use RENEW to detect
when the server has determined that the callback path is down. When
the server has made such a determination, only the RENEW operation
will renew the lease on delegations. If the server determines the
callback path is down, it returns NFS4ERR_CB_PATH_DOWN. Even though
it returns NFS4ERR_CB_PATH_DOWN, the server MUST renew the lease on
the record locks and share reservations that the client has
established on the server. If for some reason the lock and share
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reservation lease cannot be renewed, then the server MUST return an
error other than NFS4ERR_CB_PATH_DOWN, even if the callback path is
also down.
The client that issues RENEW MUST choose the principal, RPC security
flavor, and if applicable, GSS-API mechanism and service via one of
the following algorithms:
The client uses the same principal, RPC security flavor -- and if
the flavor was RPCSEC_GSS -- the same mechanism and service that
was used when the client id was established via
SETCLIENTID_CONFIRM.
The client uses any principal, RPC security flavor mechanism and
service combination that currently has an OPEN file on the server.
I.e. the same principal had a successful OPEN operation, the file
is still open by that principal, and the flavor, mechanism, and
service of RENEW match that of the previous OPEN.
The server MUST reject a RENEW that does not use one the
aforementioned algorithms, with the error NFS4ERR_ACCESS.
22.27.6. ERRORS
NFS4ERR_ACCESS NFS4ERR_ADMIN_REVOKED NFS4ERR_BADXDR
NFS4ERR_CB_PATH_DOWN NFS4ERR_EXPIRED NFS4ERR_LEASE_MOVED
NFS4ERR_RESOURCE NFS4ERR_SERVERFAULT NFS4ERR_STALE_CLIENTID
22.28. Operation 31: RESTOREFH - Restore Saved Filehandle
22.28.1. SYNOPSIS
(sfh) -> (cfh)
22.28.2. ARGUMENTS
/* SAVED_FH: */
void;
22.28.3. RESULTS
struct RESTOREFH4res {
/* CURRENT_FH: value of saved fh */
nfsstat4 status;
};
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22.28.4. DESCRIPTION
Set the current filehandle to the value in the saved filehandle. If
there is no saved filehandle then return the error NFS4ERR_RESTOREFH.
22.28.5. IMPLEMENTATION
Operations like OPEN and LOOKUP use the current filehandle to
represent a directory and replace it with a new filehandle. Assuming
the previous filehandle was saved with a SAVEFH operator, the
previous filehandle can be restored as the current filehandle. This
is commonly used to obtain post-operation attributes for the
directory, e.g.
PUTFH (directory filehandle)
SAVEFH
GETATTR attrbits (pre-op dir attrs)
CREATE optbits "foo" attrs
GETATTR attrbits (file attributes)
RESTOREFH
GETATTR attrbits (post-op dir attrs)
22.28.6. ERRORS
NFS4ERR_BADHANDLE NFS4ERR_FHEXPIRED NFS4ERR_MOVED NFS4ERR_RESOURCE
NFS4ERR_RESTOREFH NFS4ERR_SERVERFAULT NFS4ERR_STALE NFS4ERR_WRONGSEC
22.29. Operation 32: SAVEFH - Save Current Filehandle
22.29.1. SYNOPSIS
(cfh) -> (sfh)
22.29.2. ARGUMENTS
/* CURRENT_FH: */
void;
22.29.3. RESULTS
struct SAVEFH4res {
/* SAVED_FH: value of current fh */
nfsstat4 status;
};
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22.29.4. DESCRIPTION
Save the current filehandle. If a previous filehandle was saved then
it is no longer accessible. The saved filehandle can be restored as
the current filehandle with the RESTOREFH operator.
On success, the current filehandle retains its value.
22.29.5. IMPLEMENTATION
22.29.6. ERRORS
NFS4ERR_BADHANDLE NFS4ERR_FHEXPIRED NFS4ERR_MOVED
NFS4ERR_NOFILEHANDLE NFS4ERR_RESOURCE NFS4ERR_SERVERFAULT
NFS4ERR_STALE
22.30. Operation 33: SECINFO - Obtain Available Security
22.30.1. SYNOPSIS
(cfh), name -> { secinfo }
22.30.2. ARGUMENTS
struct SECINFO4args {
/* CURRENT_FH: directory */
component4 name;
};
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22.30.3. RESULTS
enum rpc_gss_svc_t { /* From RFC 2203 */
RPC_GSS_SVC_NONE = 1,
RPC_GSS_SVC_INTEGRITY = 2,
RPC_GSS_SVC_PRIVACY = 3
};
struct rpcsec_gss_info {
sec_oid4 oid;
qop4 qop;
rpc_gss_svc_t service;
};
union secinfo4 switch (uint32_t flavor) {
case RPCSEC_GSS:
rpcsec_gss_info flavor_info;
default:
void;
};
typedef secinfo4 SECINFO4resok<>;
union SECINFO4res switch (nfsstat4 status) {
case NFS4_OK:
SECINFO4resok resok4;
default:
void;
};
22.30.4. DESCRIPTION
The SECINFO operation is used by the client to obtain a list of valid
RPC authentication flavors for a specific directory filehandle, file
name pair. SECINFO should apply the same access methodology used for
LOOKUP when evaluating the name. Therefore, if the requester does
not have the appropriate access to LOOKUP the name then SECINFO must
behave the same way and return NFS4ERR_ACCESS.
The result will contain an array which represents the security
mechanisms available, with an order corresponding to the server's
preferences, the most preferred being first in the array. The client
is free to pick whatever security mechanism it both desires and
supports, or to pick in the server's preference order the first one
it supports. The array entries are represented by the secinfo4
structure. The field 'flavor' will contain a value of AUTH_NONE,
AUTH_SYS (as defined in [RFC1831]), or RPCSEC_GSS (as defined in
[RFC2203]). The field flavor can also any other security flavor
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registered with IANA.
For the flavors AUTH_NONE and AUTH_SYS, no additional security
information is returned. The same is true of many (if not most)
other security flavors, including AUTH_DH. For a return value of
RPCSEC_GSS, a security triple is returned that contains the mechanism
object id (as defined in [RFC2743]), the quality of protection (as
defined in [RFC2743]) and the service type (as defined in [RFC2203]).
It is possible for SECINFO to return multiple entries with flavor
equal to RPCSEC_GSS with different security triple values.
On success, the current filehandle retains its value.
If the name has a length of 0 (zero), or if name does not obey the
UTF-8 definition, the error NFS4ERR_INVAL will be returned.
22.30.5. IMPLEMENTATION
The SECINFO operation is expected to be used by the NFS client when
the error value of NFS4ERR_WRONGSEC is returned from another NFS
operation. This signifies to the client that the server's security
policy is different from what the client is currently using. At this
point, the client is expected to obtain a list of possible security
flavors and choose what best suits its policies.
As mentioned, the server's security policies will determine when a
client request receives NFS4ERR_WRONGSEC. The operations which may
receive this error are: LINK, LOOKUP, LOOKUPP, OPEN, PUTFH, PUTPUBFH,
PUTROOTFH, RESTOREFH, RENAME, and indirectly READDIR. LINK and
RENAME will only receive this error if the security used for the
operation is inappropriate for saved filehandle. With the exception
of READDIR, these operations represent the point at which the client
can instantiate a filehandle into the "current filehandle" at the
server. The filehandle is either provided by the client (PUTFH,
PUTPUBFH, PUTROOTFH) or generated as a result of a name to filehandle
translation (LOOKUP and OPEN). RESTOREFH is different because the
filehandle is a result of a previous SAVEFH. Even though the
filehandle, for RESTOREFH, might have previously passed the server's
inspection for a security match, the server will check it again on
RESTOREFH to ensure that the security policy has not changed.
If the client wants to resolve an error return of NFS4ERR_WRONGSEC,
the following will occur:
o For LOOKUP and OPEN, the client will use SECINFO with the same
current filehandle and name as provided in the original LOOKUP or
OPEN to enumerate the available security triples.
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o For LINK, PUTFH, PUTROOTFH, PUTPUBFH, RENAME, and RESTOREFH, the
client will use SECINFO_NO_NAME { style = current_fh }. The
client will prefix the SECINFO_NO_NAME operation with the
appropriate PUTFH, PUTPUBFH, or PUTROOTFH operation that provides
the file handled originally provided by the PUTFH, PUTPUBFH,
PUTROOTFH, or RESTOREFH, or for the failed LINK or RENAME, the
SAVEFH.
o NOTE: In NFSv4.0, the client was required to use SECINFO, and had
to reconstruct the parent of the original file handle, and the
component name of the original filehandle.
o For LOOKUPP, the client will use SECINFO_NO_NAME { style = parent
} and provide the filehandle with equals the filehandle originally
provided to LOOKUPP.
The READDIR operation will not directly return the NFS4ERR_WRONGSEC
error. However, if the READDIR request included a request for
attributes, it is possible that the READDIR request's security triple
did not match that of a directory entry. If this is the case and the
client has requested the rdattr_error attribute, the server will
return the NFS4ERR_WRONGSEC error in rdattr_error for the entry.
See the section "Security Considerations" for a discussion on the
recommendations for security flavor used by SECINFO and
SECINFO_NO_NAME.
22.30.6. ERRORS
NFS4ERR_ACCESS NFS4ERR_BADCHAR NFS4ERR_BADHANDLE NFS4ERR_BADNAME
NFS4ERR_BADXDR NFS4ERR_FHEXPIRED NFS4ERR_INVAL NFS4ERR_MOVED
NFS4ERR_NAMETOOLONG NFS4ERR_NOENT NFS4ERR_NOFILEHANDLE NFS4ERR_NOTDIR
NFS4ERR_RESOURCE NFS4ERR_SERVERFAULT NFS4ERR_STALE
22.31. Operation 34: SETATTR - Set Attributes
22.31.1. SYNOPSIS
(cfh), stateid, attrmask, attr_vals -> attrsset
22.31.2. ARGUMENTS
struct SETATTR4args {
/* CURRENT_FH: target object */
stateid4 stateid;
fattr4 obj_attributes;
};
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22.31.3. RESULTS
struct SETATTR4res {
nfsstat4 status;
bitmap4 attrsset;
};
22.31.4. DESCRIPTION
The SETATTR operation changes one or more of the attributes of a
filesystem object. The new attributes are specified with a bitmap
and the attributes that follow the bitmap in bit order.
The stateid argument for SETATTR is used to provide file locking
context that is necessary for SETATTR requests that set the size
attribute. Since setting the size attribute modifies the file's
data, it has the same locking requirements as a corresponding WRITE.
Any SETATTR that sets the size attribute is incompatible with a share
reservation that specifies DENY_WRITE. The area between the old end-
of-file and the new end-of-file is considered to be modified just as
would have been the case had the area in question been specified as
the target of WRITE, for the purpose of checking conflicts with
record locks, for those cases in which a server is implementing
mandatory record locking behavior. A valid stateid should always be
specified. When the file size attribute is not set, the special
stateid consisting of all bits zero should be passed.
On either success or failure of the operation, the server will return
the attrsset bitmask to represent what (if any) attributes were
successfully set. The attrsset in the response is a subset of the
bitmap4 that is part of the obj_attributes in the argument.
On success, the current filehandle retains its value.
22.31.5. IMPLEMENTATION
If the request specifies the owner attribute to be set, the server
should allow the operation to succeed if the current owner of the
object matches the value specified in the request. Some servers may
be implemented in a way as to prohibit the setting of the owner
attribute unless the requester has privilege to do so. If the server
is lenient in this one case of matching owner values, the client
implementation may be simplified in cases of creation of an object
followed by a SETATTR.
The file size attribute is used to request changes to the size of a
file. A value of 0 (zero) causes the file to be truncated, a value
less than the current size of the file causes data from new size to
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the end of the file to be discarded, and a size greater than the
current size of the file causes logically zeroed data bytes to be
added to the end of the file. Servers are free to implement this
using holes or actual zero data bytes. Clients should not make any
assumptions regarding a server's implementation of this feature,
beyond that the bytes returned will be zeroed. Servers must support
extending the file size via SETATTR.
SETATTR is not guaranteed atomic. A failed SETATTR may partially
change a file's attributes.
Changing the size of a file with SETATTR indirectly changes the
time_modify. A client must account for this as size changes can
result in data deletion.
The attributes time_access_set and time_modify_set are write-only
attributes constructed as a switched union so the client can direct
the server in setting the time values. If the switched union
specifies SET_TO_CLIENT_TIME4, the client has provided an nfstime4 to
be used for the operation. If the switch union does not specify
SET_TO_CLIENT_TIME4, the server is to use its current time for the
SETATTR operation.
If server and client times differ, programs that compare client time
to file times can break. A time maintenance protocol should be used
to limit client/server time skew.
Use of a COMPOUND containing a VERIFY operation specifying only the
change attribute, immediately followed by a SETATTR, provides a means
whereby a client may specify a request that emulates the
functionality of the SETATTR guard mechanism of NFS version 3. Since
the function of the guard mechanism is to avoid changes to the file
attributes based on stale information, delays between checking of the
guard condition and the setting of the attributes have the potential
to compromise this function, as would the corresponding delay in the
NFS version 4 emulation. Therefore, NFS version 4 servers should
take care to avoid such delays, to the degree possible, when
executing such a request.
If the server does not support an attribute as requested by the
client, the server should return NFS4ERR_ATTRNOTSUPP.
A mask of the attributes actually set is returned by SETATTR in all
cases. That mask must not include attributes bits not requested to
be set by the client, and must be equal to the mask of attributes
requested to be set only if the SETATTR completes without error.
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22.31.6. ERRORS
NFS4ERR_ACCESS NFS4ERR_ADMIN_REVOKED NFS4ERR_ATTRNOTSUPP
NFS4ERR_BADCHAR NFS4ERR_BADHANDLE NFS4ERR_BADOWNER
NFS4ERR_BAD_STATEID NFS4ERR_BADXDR NFS4ERR_DELAY NFS4ERR_DQUOT
NFS4ERR_EXPIRED NFS4ERR_FBIG NFS4ERR_FHEXPIRED NFS4ERR_GRACE
NFS4ERR_INVAL NFS4ERR_IO NFS4ERR_ISDIR NFS4ERR_LOCKED NFS4ERR_MOVED
NFS4ERR_NOFILEHANDLE NFS4ERR_NOSPC NFS4ERR_OLD_STATEID
NFS4ERR_OPENMODE NFS4ERR_PERM NFS4ERR_RESOURCE NFS4ERR_ROFS
NFS4ERR_SERVERFAULT NFS4ERR_STALE NFS4ERR_STALE_STATEID
22.32. Operation 35: SETCLIENTID - Negotiate Clientid
22.32.1. SYNOPSIS
client, callback, callback_ident -> clientid, setclientid_confirm
22.32.2. ARGUMENTS
struct SETCLIENTID4args {
nfs_client_id4 client;
cb_client4 callback;
uint32_t callback_ident;
};
22.32.3. RESULTS
struct SETCLIENTID4resok {
clientid4 clientid;
verifier4 setclientid_confirm;
};
union SETCLIENTID4res switch (nfsstat4 status) {
case NFS4_OK:
SETCLIENTID4resok resok4;
case NFS4ERR_CLID_INUSE:
clientaddr4 client_using;
default:
void;
};
22.32.4. DESCRIPTION
The client uses the SETCLIENTID operation to notify the server of its
intention to use a particular client identifier, callback, and
callback_ident for subsequent requests that entail creating lock,
share reservation, and delegation state on the server. Upon
successful completion the server will return a short hand clientid
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which, if confirmed via a separate step, will be used in subsequent
file locking and file open requests. Confirmation of the clientid
must be done via the SETCLIENTID_CONFIRM operation to return the
clientid and setclientid_confirm values, as verifiers, to the server.
The reason why two verifiers are necessary is that it is possible to
use SETCLIENTID and SETCLIENTID_CONFIRM to modify the callback and
callback_ident information but not the short hand clientid. In that
event, the setclientid_confirm value is effectively the only
verifier.
The callback information provided in this operation will be used if
the client is provided an open delegation at a future point.
Therefore, the client must correctly reflect the program and port
numbers for the callback program at the time SETCLIENTID is used.
The callback_ident value is used by the server on the callback. The
client can use leverage the callback_ident eliminate the need for
more than one callback RPC program number while still being able to
determine which server is initiating the callback.
22.32.5. IMPLEMENTATION
To understand how to implement SETCLIENTID, make the following
notations. Let:
x be the value of the client.id subfield of the SETCLIENTID4args
structure.
v be the value of the client.verifier subfield of the
SETCLIENTID4args structure.
c be the value of the clientid field returned in the
SETCLIENTID4resok structure.
k represent the value combination of the fields callback and
callback_ident fields of the SETCLIENTID4args structure.
s be the setclientid_confirm value returned in the SETCLIENTID4resok
structure.
{ v, x, c, k, s } be a quintuple for a client record. A client
record is confirmed if there has been a SETCLIENTID_CONFIRM
operation to confirm it. Otherwise it is unconfirmed. An
unconfirmed record is established by a SETCLIENTID call.
Since SETCLIENTID is a non-idempotent operation, let us assume that
the server is implementing the duplicate request cache (DRC).
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When the server gets a SETCLIENTID { v, x, k } request, it processes
it in the following manner.
o It first looks up the request in the DRC. If there is a hit, it
returns the result cached in the DRC. The server does NOT remove
client state (locks, shares, delegations) nor does it modify any
recorded callback and callback_ident information for client { x }.
For any DRC miss, the server takes the client id string x, and
searches for client records for x that the server may have
recorded from previous SETCLIENTID calls. For any confirmed
record with the same id string x, if the recorded principal does
not match that of SETCLIENTID call, then the server returns a
NFS4ERR_CLID_INUSE error.
For brevity of discussion, the remaining description of the
processing assumes that there was a DRC miss, and that where the
server has previously recorded a confirmed record for client x,
the aforementioned principal check has successfully passed.
o The server checks if it has recorded a confirmed record for { v,
x, c, l, s }, where l may or may not equal k. If so, and since
the id verifier v of the request matches that which is confirmed
and recorded, the server treats this as a probable callback
information update and records an unconfirmed { v, x, c, k, t }
and leaves the confirmed { v, x, c, l, s } in place, such that t
!= s. It does not matter if k equals l or not. Any pre-existing
unconfirmed { v, x, c, *, * } is removed.
The server returns { c, t }. It is indeed returning the old
clientid4 value c, because the client apparently only wants to
update callback value k to value l. It's possible this request is
one from the Byzantine router that has stale callback information,
but this is not a problem. The callback information update is
only confirmed if followed up by a SETCLIENTID_CONFIRM { c, t }.
The server awaits confirmation of k via SETCLIENTID_CONFIRM { c, t
}.
The server does NOT remove client (lock/share/delegation) state
for x.
o The server has previously recorded a confirmed { u, x, c, l, s }
record such that v != u, l may or may not equal k, and has not
recorded any unconfirmed { *, x, *, *, * } record for x. The
server records an unconfirmed { v, x, d, k, t } (d != c, t != s).
The server returns { d, t }.
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The server awaits confirmation of { d, k } via SETCLIENTID_CONFIRM
{ d, t }.
The server does NOT remove client (lock/share/delegation) state
for x.
o The server has previously recorded a confirmed { u, x, c, l, s }
record such that v != u, l may or may not equal k, and recorded an
unconfirmed { w, x, d, m, t } record such that c != d, t != s, m
may or may not equal k, m may or may not equal l, and k may or may
not equal l. Whether w == v or w != v makes no difference. The
server simply removes the unconfirmed { w, x, d, m, t } record and
replaces it with an unconfirmed { v, x, e, k, r } record, such
that e != d, e != c, r != t, r != s.
The server returns { e, r }.
The server awaits confirmation of { e, k } via SETCLIENTID_CONFIRM
{ e, r }.
The server does NOT remove client (lock/share/delegation) state
for x.
o The server has no confirmed { *, x, *, *, * } for x. It may or
may not have recorded an unconfirmed { u, x, c, l, s }, where l
may or may not equal k, and u may or may not equal v. Any
unconfirmed record { u, x, c, l, * }, regardless whether u == v or
l == k, is replaced with an unconfirmed record { v, x, d, k, t }
where d != c, t != s.
The server returns { d, t }.
The server awaits confirmation of { d, k } via SETCLIENTID_CONFIRM
{ d, t }. The server does NOT remove client (lock/share/
delegation) state for x. .in -2
The server generates the clientid and setclientid_confirm values and
must take care to ensure that these values are extremely unlikely to
ever be regenerated.
22.32.6. ERRORS
NFS4ERR_BADXDR NFS4ERR_CLID_INUSE NFS4ERR_INVAL NFS4ERR_RESOURCE
NFS4ERR_SERVERFAULT
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22.33. Operation 36: SETCLIENTID_CONFIRM - Confirm Clientid
22.33.1. SYNOPSIS
clientid, verifier -> -
22.33.2. ARGUMENTS
struct SETCLIENTID_CONFIRM4args {
clientid4 clientid;
verifier4 setclientid_confirm;
};
22.33.3. RESULTS
struct SETCLIENTID_CONFIRM4res {
nfsstat4 status;
};
22.33.4. DESCRIPTION
This operation is used by the client to confirm the results from a
previous call to SETCLIENTID. The client provides the server
supplied (from a SETCLIENTID response) clientid. The server responds
with a simple status of success or failure.
22.33.5. IMPLEMENTATION
The client must use the SETCLIENTID_CONFIRM operation to confirm the
following two distinct cases:
o The client's use of a new shorthand client identifier (as returned
from the server in the response to SETCLIENTID), a new callback
value (as specified in the arguments to SETCLIENTID) and a new
callback_ident (as specified in the arguments to SETCLIENTID)
value. The client's use of SETCLIENTID_CONFIRM in this case also
confirms the removal of any of the client's previous relevant
leased state. Relevant leased client state includes record locks,
share reservations, and where the server does not support the
CLAIM_DELEGATE_PREV claim type, delegations. If the server
supports CLAIM_DELEGATE_PREV, then SETCLIENTID_CONFIRM MUST NOT
remove delegations for this client; relevant leased client state
would then just include record locks and share reservations.
o The client's re-use of an old, previously confirmed, shorthand
client identifier, a new callback value, and a new callback_ident
value. The client's use of SETCLIENTID_CONFIRM in this case MUST
NOT result in the removal of any previous leased state (locks,
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share reservations, and delegations)
We use the same notation and definitions for v, x, c, k, s, and
unconfirmed and confirmed client records as introduced in the
description of the SETCLIENTID operation. The arguments to
SETCLIENTID_CONFIRM are indicated by the notation { c, s }, where c
is a value of type clientid4, and s is a value of type verifier4
corresponding to the setclientid_confirm field.
As with SETCLIENTID, SETCLIENTID_CONFIRM is a non-idempotent
operation, and we assume that the server is implementing the
duplicate request cache (DRC).
When the server gets a SETCLIENTID_CONFIRM { c, s } request, it
processes it in the following manner.
o It first looks up the request in the DRC. If there is a hit, it
returns the result cached in the DRC. The server does not remove
any relevant leased client state nor does it modify any recorded
callback and callback_ident information for client { x } as
represented by the short hand value c.
For a DRC miss, the server checks for client records that match the
short hand value c. The processing cases are as follows:
o The server has recorded an unconfirmed { v, x, c, k, s } record
and a confirmed { v, x, c, l, t } record, such that s != t. If
the principals of the records do not match that of the
SETCLIENTID_CONFIRM, the server returns NFS4ERR_CLID_INUSE, and no
relevant leased client state is removed and no recorded callback
and callback_ident information for client { x } is changed.
Otherwise, the confirmed { v, x, c, l, t } record is removed and
the unconfirmed { v, x, c, k, s } is marked as confirmed, thereby
modifying recorded and confirmed callback and callback_ident
information for client { x }.
The server does not remove any relevant leased client state.
The server returns NFS4_OK.
o The server has not recorded an unconfirmed { v, x, c, *, * } and
has recorded a confirmed { v, x, c, *, s }. If the principals of
the record and of SETCLIENTID_CONFIRM do not match, the server
returns NFS4ERR_CLID_INUSE without removing any relevant leased
client state and without changing recorded callback and
callback_ident values for client { x }.
If the principals match, then what has likely happened is that the
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client never got the response from the SETCLIENTID_CONFIRM, and
the DRC entry has been purged. Whatever the scenario, since the
principals match, as well as { c, s } matching a confirmed record,
the server leaves client x's relevant leased client state intact,
leaves its callback and callback_ident values unmodified, and
returns NFS4_OK.
o The server has not recorded a confirmed { *, *, c, *, * }, and has
recorded an unconfirmed { *, x, c, k, s }. Even if this is a
retry from client, nonetheless the client's first
SETCLIENTID_CONFIRM attempt was not received by the server. Retry
or not, the server doesn't know, but it processes it as if were a
first try. If the principal of the unconfirmed { *, x, c, k, s }
record mismatches that of the SETCLIENTID_CONFIRM request the
server returns NFS4ERR_CLID_INUSE without removing any relevant
leased client state.
Otherwise, the server records a confirmed { *, x, c, k, s }. If
there is also a confirmed { *, x, d, *, t }, the server MUST
remove the client x's relevant leased client state, and overwrite
the callback state with k. The confirmed record { *, x, d, *, t }
is removed.
Server returns NFS4_OK.
o The server has no record of a confirmed or unconfirmed { *, *, c,
*, s }. The server returns NFS4ERR_STALE_CLIENTID. The server
does not remove any relevant leased client state, nor does it
modify any recorded callback and callback_ident information for
any client.
The server needs to cache unconfirmed { v, x, c, k, s } client
records and await for some time their confirmation. As should be
clear from the record processing discussions for SETCLIENTID and
SETCLIENTID_CONFIRM, there are cases where the server does not
deterministically remove unconfirmed client records. To avoid
running out of resources, the server is not required to hold
unconfirmed records indefinitely. One strategy the server might use
is to set a limit on how many unconfirmed client records it will
maintain, and then when the limit would be exceeded, remove the
oldest record. Another strategy might be to remove an unconfirmed
record when some amount of time has elapsed. The choice of the
amount of time is fairly arbitrary but it is surely no higher than
the server's lease time period. Consider that leases need to be
renewed before the lease time expires via an operation from the
client. If the client cannot issue a SETCLIENTID_CONFIRM after a
SETCLIENTID before a period of time equal to that of a lease expires,
then the client is unlikely to be able maintain state on the server
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during steady state operation.
If the client does send a SETCLIENTID_CONFIRM for an unconfirmed
record that the server has already deleted, the client will get
NFS4ERR_STALE_CLIENTID back. If so, the client should then start
over, and send SETCLIENTID to reestablish an unconfirmed client
record and get back an unconfirmed clientid and setclientid_confirm
verifier. The client should then send the SETCLIENTID_CONFIRM to
confirm the clientid.
SETCLIENTID_CONFIRM does not establish or renew a lease. However, if
SETCLIENTID_CONFIRM removes relevant leased client state, and that
state does not include existing delegations, the server MUST allow
the client a period of time no less than the value of lease_time
attribute, to reclaim, (via the CLAIM_DELEGATE_PREV claim type of the
OPEN operation) its delegations before removing unreclaimed
delegations.
22.33.6. ERRORS
NFS4ERR_BADXDR NFS4ERR_CLID_INUSE NFS4ERR_RESOURCE
NFS4ERR_SERVERFAULT NFS4ERR_STALE_CLIENTID
22.34. Operation 37: VERIFY - Verify Same Attributes
22.34.1. SYNOPSIS
(cfh), fattr -> -
22.34.2. ARGUMENTS
struct VERIFY4args {
/* CURRENT_FH: object */
fattr4 obj_attributes;
};
22.34.3. RESULTS
struct VERIFY4res {
nfsstat4 status;
};
22.34.4. DESCRIPTION
The VERIFY operation is used to verify that attributes have a value
assumed by the client before proceeding with following operations in
the compound request. If any of the attributes do not match then the
error NFS4ERR_NOT_SAME must be returned. The current filehandle
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retains its value after successful completion of the operation.
22.34.5. IMPLEMENTATION
One possible use of the VERIFY operation is the following compound
sequence. With this the client is attempting to verify that the file
being removed will match what the client expects to be removed. This
sequence can help prevent the unintended deletion of a file.
PUTFH (directory filehandle)
LOOKUP (file name)
VERIFY (filehandle == fh)
PUTFH (directory filehandle)
REMOVE (file name)
This sequence does not prevent a second client from removing and
creating a new file in the middle of this sequence but it does help
avoid the unintended result.
In the case that a recommended attribute is specified in the VERIFY
operation and the server does not support that attribute for the
filesystem object, the error NFS4ERR_ATTRNOTSUPP is returned to the
client.
When the attribute rdattr_error or any write-only attribute (e.g.
time_modify_set) is specified, the error NFS4ERR_INVAL is returned to
the client.
22.34.6. ERRORS
NFS4ERR_ACCESS NFS4ERR_ATTRNOTSUPP NFS4ERR_BADCHAR NFS4ERR_BADHANDLE
NFS4ERR_BADXDR NFS4ERR_DELAY NFS4ERR_FHEXPIRED NFS4ERR_INVAL
NFS4ERR_MOVED NFS4ERR_NOFILEHANDLE NFS4ERR_NOT_SAME NFS4ERR_RESOURCE
NFS4ERR_SERVERFAULT NFS4ERR_STALE
22.35. Operation 38: WRITE - Write to File
22.35.1. SYNOPSIS
(cfh), stateid, offset, stable, data -> count, committed, writeverf
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22.35.2. ARGUMENTS
enum stable_how4 {
UNSTABLE4 = 0,
DATA_SYNC4 = 1,
FILE_SYNC4 = 2
};
struct WRITE4args {
/* CURRENT_FH: file */
stateid4 stateid;
offset4 offset;
stable_how4 stable;
opaque data<>;
};
22.35.3. RESULTS
struct WRITE4resok {
count4 count;
stable_how4 committed;
verifier4 writeverf;
};
union WRITE4res switch (nfsstat4 status) {
case NFS4_OK:
WRITE4resok resok4;
default:
void;
};
22.35.4. DESCRIPTION
The WRITE operation is used to write data to a regular file. The
target file is specified by the current filehandle. The offset
specifies the offset where the data should be written. An offset of
0 (zero) specifies that the write should start at the beginning of
the file. The count, as encoded as part of the opaque data
parameter, represents the number of bytes of data that are to be
written. If the count is 0 (zero), the WRITE will succeed and return
a count of 0 (zero) subject to permissions checking. The server may
choose to write fewer bytes than requested by the client.
Part of the write request is a specification of how the write is to
be performed. The client specifies with the stable parameter the
method of how the data is to be processed by the server. If stable
is FILE_SYNC4, the server must commit the data written plus all
filesystem metadata to stable storage before returning results. This
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corresponds to the NFS version 2 protocol semantics. Any other
behavior constitutes a protocol violation. If stable is DATA_SYNC4,
then the server must commit all of the data to stable storage and
enough of the metadata to retrieve the data before returning. The
server implementor is free to implement DATA_SYNC4 in the same
fashion as FILE_SYNC4, but with a possible performance drop. If
stable is UNSTABLE4, the server is free to commit any part of the
data and the metadata to stable storage, including all or none,
before returning a reply to the client. There is no guarantee
whether or when any uncommitted data will subsequently be committed
to stable storage. The only guarantees made by the server are that
it will not destroy any data without changing the value of verf and
that it will not commit the data and metadata at a level less than
that requested by the client.
The stateid value for a WRITE request represents a value returned
from a previous record lock or share reservation request. The
stateid is used by the server to verify that the associated share
reservation and any record locks are still valid and to update lease
timeouts for the client.
Upon successful completion, the following results are returned. The
count result is the number of bytes of data written to the file. The
server may write fewer bytes than requested. If so, the actual
number of bytes written starting at location, offset, is returned.
The server also returns an indication of the level of commitment of
the data and metadata via committed. If the server committed all
data and metadata to stable storage, committed should be set to
FILE_SYNC4. If the level of commitment was at least as strong as
DATA_SYNC4, then committed should be set to DATA_SYNC4. Otherwise,
committed must be returned as UNSTABLE4. If stable was FILE4_SYNC,
then committed must also be FILE_SYNC4: anything else constitutes a
protocol violation. If stable was DATA_SYNC4, then committed may be
FILE_SYNC4 or DATA_SYNC4: anything else constitutes a protocol
violation. If stable was UNSTABLE4, then committed may be either
FILE_SYNC4, DATA_SYNC4, or UNSTABLE4.
The final portion of the result is the write verifier. The write
verifier is a cookie that the client can use to determine whether the
server has changed instance (boot) state between a call to WRITE and
a subsequent call to either WRITE or COMMIT. This cookie must be
consistent during a single instance of the NFS version 4 protocol
service and must be unique between instances of the NFS version 4
protocol server, where uncommitted data may be lost.
If a client writes data to the server with the stable argument set to
UNSTABLE4 and the reply yields a committed response of DATA_SYNC4 or
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UNSTABLE4, the client will follow up some time in the future with a
COMMIT operation to synchronize outstanding asynchronous data and
metadata with the server's stable storage, barring client error. It
is possible that due to client crash or other error that a subsequent
COMMIT will not be received by the server.
For a WRITE with a stateid value of all bits 0, the server MAY allow
the WRITE to be serviced subject to mandatory file locks or the
current share deny modes for the file. For a WRITE with a stateid
value of all bits 1, the server MUST NOT allow the WRITE operation to
bypass locking checks at the server and are treated exactly the same
as if a stateid of all bits 0 were used.
On success, the current filehandle retains its value.
22.35.5. IMPLEMENTATION
It is possible for the server to write fewer bytes of data than
requested by the client. In this case, the server should not return
an error unless no data was written at all. If the server writes
less than the number of bytes specified, the client should issue
another WRITE to write the remaining data.
It is assumed that the act of writing data to a file will cause the
time_modified of the file to be updated. However, the time_modified
of the file should not be changed unless the contents of the file are
changed. Thus, a WRITE request with count set to 0 should not cause
the time_modified of the file to be updated.
The definition of stable storage has been historically a point of
contention. The following expected properties of stable storage may
help in resolving design issues in the implementation. Stable
storage is persistent storage that survives:
1. Repeated power failures.
2. Hardware failures (of any board, power supply, etc.).
3. Repeated software crashes, including reboot cycle.
This definition does not address failure of the stable storage module
itself.
The verifier is defined to allow a client to detect different
instances of an NFS version 4 protocol server over which cached,
uncommitted data may be lost. In the most likely case, the verifier
allows the client to detect server reboots. This information is
required so that the client can safely determine whether the server
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could have lost cached data. If the server fails unexpectedly and
the client has uncommitted data from previous WRITE requests (done
with the stable argument set to UNSTABLE4 and in which the result
committed was returned as UNSTABLE4 as well) it may not have flushed
cached data to stable storage. The burden of recovery is on the
client and the client will need to retransmit the data to the server.
A suggested verifier would be to use the time that the server was
booted or the time the server was last started (if restarting the
server without a reboot results in lost buffers).
The committed field in the results allows the client to do more
effective caching. If the server is committing all WRITE requests to
stable storage, then it should return with committed set to
FILE_SYNC4, regardless of the value of the stable field in the
arguments. A server that uses an NVRAM accelerator may choose to
implement this policy. The client can use this to increase the
effectiveness of the cache by discarding cached data that has already
been committed on the server.
Some implementations may return NFS4ERR_NOSPC instead of
NFS4ERR_DQUOT when a user's quota is exceeded. In the case that the
current filehandle is a directory, the server will return
NFS4ERR_ISDIR. If the current filehandle is not a regular file or a
directory, the server will return NFS4ERR_INVAL.
If mandatory file locking is on for the file, and corresponding
record of the data to be written file is read or write locked by an
owner that is not associated with the stateid, the server will return
NFS4ERR_LOCKED. If so, the client must check if the owner
corresponding to the stateid used with the WRITE operation has a
conflicting read lock that overlaps with the region that was to be
written. If the stateid's owner has no conflicting read lock, then
the client should try to get the appropriate write record lock via
the LOCK operation before re-attempting the WRITE. When the WRITE
completes, the client should release the record lock via LOCKU.
If the stateid's owner had a conflicting read lock, then the client
has no choice but to return an error to the application that
attempted the WRITE. The reason is that since the stateid's owner
had a read lock, the server either attempted to temporarily
effectively upgrade this read lock to a write lock, or the server has
no upgrade capability. If the server attempted to upgrade the read
lock and failed, it is pointless for the client to re-attempt the
upgrade via the LOCK operation, because there might be another client
also trying to upgrade. If two clients are blocked trying upgrade
the same lock, the clients deadlock. If the server has no upgrade
capability, then it is pointless to try a LOCK operation to upgrade.
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22.35.6. ERRORS
NFS4ERR_ACCESS NFS4ERR_ADMIN_REVOKED NFS4ERR_BADHANDLE
NFS4ERR_BAD_STATEID NFS4ERR_BADXDR NFS4ERR_DELAY NFS4ERR_DQUOT
NFS4ERR_EXPIRED NFS4ERR_FBIG NFS4ERR_FHEXPIRED NFS4ERR_GRACE
NFS4ERR_INVAL NFS4ERR_IO NFS4ERR_ISDIR NFS4ERR_LEASE_MOVED
NFS4ERR_LOCKED NFS4ERR_MOVED NFS4ERR_NOFILEHANDLE NFS4ERR_NOSPC
NFS4ERR_NXIO NFS4ERR_OLD_STATEID NFS4ERR_OPENMODE NFS4ERR_RESOURCE
NFS4ERR_ROFS NFS4ERR_SERVERFAULT NFS4ERR_STALE NFS4ERR_STALE_STATEID
22.36. Operation 39: RELEASE_LOCKOWNER - Release Lockowner State
22.36.1. SYNOPSIS
lockowner -> ()
22.36.2. ARGUMENTS
struct RELEASE_LOCKOWNER4args {
lock_owner4 lock_owner;
};
22.36.3. RESULTS
struct RELEASE_LOCKOWNER4res {
nfsstat4 status;
};
22.36.4. DESCRIPTION
This operation is used to notify the server that the lock_owner is no
longer in use by the client. This allows the server to release
cached state related to the specified lock_owner. If file locks,
associated with the lock_owner, are held at the server, the error
NFS4ERR_LOCKS_HELD will be returned and no further action will be
taken.
22.36.5. IMPLEMENTATION
The client may choose to use this operation to ease the amount of
server state that is held. Depending on behavior of applications at
the client, it may be important for the client to use this operation
since the server has certain obligations with respect to holding a
reference to a lock_owner as long as the associated file is open.
Therefore, if the client knows for certain that the lock_owner will
no longer be used under the context of the associated open_owner4, it
should use RELEASE_LOCKOWNER.
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22.36.6. ERRORS
NFS4ERR_ADMIN_REVOKED NFS4ERR_BADXDR NFS4ERR_EXPIRED
NFS4ERR_LEASE_MOVED NFS4ERR_LOCKS_HELD NFS4ERR_RESOURCE
NFS4ERR_SERVERFAULT NFS4ERR_STALE_CLIENTID
22.37. Operation 10044: ILLEGAL - Illegal operation
22.37.1. SYNOPSIS
-> ()
22.37.2. ARGUMENTS
void;
22.37.3. RESULTS
struct ILLEGAL4res {
nfsstat4 status;
};
22.37.4. DESCRIPTION
This operation is a placeholder for encoding a result to handle the
case of the client sending an operation code within COMPOUND that is
not supported. See the COMPOUND procedure description for more
details.
The status field of ILLEGAL4res MUST be set to NFS4ERR_OP_ILLEGAL.
22.37.5. IMPLEMENTATION
A client will probably not send an operation with code OP_ILLEGAL but
if it does, the response will be ILLEGAL4res just as it would be with
any other invalid operation code. Note that if the server gets an
illegal operation code that is not OP_ILLEGAL, and if the server
checks for legal operation codes during the XDR decode phase, then
the ILLEGAL4res would not be returned.
22.37.6. ERRORS
NFS4ERR_OP_ILLEGAL
22.38. SECINFO_NO_NAME - Get Security on Unnamed Object
Obtain available security mechanisms with the use of the parent of an
object or the current filehandle.
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22.38.1. SYNOPSIS
(cfh), secinfo_style -> { secinfo }
22.38.2. ARGUMENT
enum secinfo_style_4 {
current_fh = 0,
parent = 1
};
typedef secinfo_style_4 SECINFO_NO_NAME4args;
22.38.3. RESULT
typedef SECINFO4res SECINFO_NO_NAME4res;
22.38.4. DESCRIPTION
Like the SECINFO operation, SECINFO_NO_NAME is used by the client to
obtain a list of valid RPC authentication flavors for a specific file
object. Unlike SECINFO, SECINFO_NO_NAME only works with objects are
accessed by file handle.
There are two styles of SECINFO_NO_NAME, as determined by the value
of the secinfo_style_4 enumeration. If "current_fh" is passed, then
SECINFO_NO_NAME is querying for the required security for the current
filehandle. If "parent" is passed, then SECINFO_NO_NAME is querying
for the required security of the current filehandles's parent. If
the style selected is "parent", then SECINFO should apply the same
access methodology used for LOOKUPP when evaluating the traversal to
the parent directory. Therefore, if the requester does not have the
appropriate access to LOOKUPP the parent then SECINFO_NO_NAME must
behave the same way and return NFS4ERR_ACCESS.
Note that if PUTFH, PUTPUBFH, or PUTROOTFH return NFS4ERR_WRONGSEC,
this is tantamount to the server asserting that the client will have
to guess what the required security is, because there is no way to
query. Therefore, the client must iterate through the security
triples available at the client and reattempt the PUTFH, PUTROOTFH or
PUTPUBFH operation. In the unfortunate event none of the MANDATORY
security triples are supported by the client and server, the client
SHOULD try using others that support integrity. Failing that, the
client can try using other forms (e.g. AUTH_SYS and AUTH_NONE), but
because such forms lack integrity checks, this puts the client at
risk.
The server implementor should pay particular attention to the section
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"Clarification of Security Negotiation in NFSv4.1" for implementation
suggestions for avoiding NFS4ERR_WRONGSEC error returns from PUTFH,
PUTROOTFH or PUTPUBFH.
Everything else about SECINFO_NO_NAME is the same as SECINFO. See
the previous discussion on SECINFO.
22.38.5. IMPLEMENTATION
See the previous dicussion on SECINFO.
22.38.6. ERRORS
NFS4ERR_ACCESS NFS4ERR_BADCHAR NFS4ERR_BADHANDLE NFS4ERR_BADNAME
NFS4ERR_BADXDR NFS4ERR_FHEXPIRED NFS4ERR_INVAL NFS4ERR_MOVED
NFS4ERR_NAMETOOLONG NFS4ERR_NOENT NFS4ERR_NOFILEHANDLE NFS4ERR_NOTDIR
NFS4ERR_RESOURCE NFS4ERR_SERVERFAULT NFS4ERR_STALE
22.39. CREATECLIENTID - Instantiate Clientid
Create a clientid
22.39.1. SYNOPSIS
client -> clientid
22.39.2. ARGUMENT
struct CREATECLIENTID4args {
nfs_client_id4 clientdesc;
};
22.39.3. RESULT
struct CREATECLIENTID4resok {
clientid4 clientid;
verifier4 clientid_confirm;
};
union SETCLIENTID4res switch (nfsstat4 status) {
case NFS4_OK:
CREATECLIENTID4resok resok4;
case NFS4ERR_CLID_INUSE:
void;
default:
void;
};
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22.39.4. DESCRIPTION
The client uses the CREATECLIENTID operation to register a particular
client identifier with the server. The clientid returned from this
operation will be necessary for requests that create state on the
server and will serve as a parent object to sessions created by the
client. In order to verify the clientid it must first be used as an
argument to CREATESESSION.
22.39.5. IMPLEMENTATION
A server's client record is a 5-tuple:
1. clientdesc.id:
The long form client identifier, sent via the client.id
subfield of the CREATECLIENTID4args structure
2. clientdesc.verifier:
A client-specific value used to indicate reboots, sent via the
clientdesc.verifier subfield of the CREATECLIENTID4args
structure
3. principal:
The RPCSEC_GSS principal sent via the RPC headers
4. clientid:
The shorthand client identifier, generated by the server and
returned via the clientid field in the CREATECLIENTID4resok
structure
5. confirmed:
A private field on the server indicating whether or not a
client record has been confirmed. A client record is
confirmed if there has been a successful CREATESESSION
operation to confirm it. Otherwise it is unconfirmed. An
unconfirmed record is established by a CREATECLIENTID call.
Any unconfirmed record that is not confirmed within a lease
period may be removed.
The following identifiers represent special values for the fields in
the records.
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id_arg:
The value of the clientdesc.id subfield of the CREATECLIENTID4args
structure of the current request.
verifier_arg:
The value of the clientdesc.verifier subfield of the
CREATECLIENTID4args structure of the current request.
old_verifier_arg:
A value of the clientdesc.verifier field of a client record
received in a previous request; this is distinct from
verifier_arg.
principal_arg:
The value of the RPCSEC_GSS principal for the current request.
old_principal_arg:
A value of the RPCSEC_GSS principal received for a previous
request. This is distinct from principal_arg.
clientid_ret:
The value of the clientid field the server will return in the
CREATECLIENTID4resok structure for the current request.
old_clientid_ret:
The value of the clientid field the server returned in the
CREATECLIENTID4resok structure for a previous request. This is
distinct from clientid_ret.
Since CREATECLIENTID is a non-idempotent operation, we must consider
the possibility that replays may occur as a result of a client
reboot, network partition, malfunctioning router, etc. Replays are
identified by the value of the client field of CREATECLIENTID4args
and the method for dealing with them is outlined in the scenarios
below.
The scenarios are described in terms of what client records whose
clientdesc.id subfield have value equal to id_arg exist in the
server's set of client records. Any cases in which there is more
than one record with identical values for id_arg represent a server
implementation error. Operation in the potential valid cases is
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summarized as follows.
1. Common case
If no client records with clientdesc.id matching id_arg exist,
a new shorthand client identifier clientid_ret is generated,
and the following unconfirmed record is added to the server's
state.
{ id_arg, verifier_arg, principal_arg, clientid_ret, FALSE }
Subsequently, the server returns clientid_ret.
2. Router Replay
If the server has the following confirmed record, then this
request is likely the result of a replayed request due to a
faulty router or lost connection.
{ id_arg, verifier_arg, principal_arg, clientid_ret, TRUE }
Since the record has been confirmed, the client must have
received the server's reply from the initial CREATECLIENTID
request. Since this is simply a spurious request, there is no
modification to the server's state, and the server makes no
reply to the client.
3. Client Collision
If the server has the following confirmed record, then this
request is likely the result of a chance collision between the
values of the clientdesc.id subfield of CREATECLIENTID4args
for two different clients.
{ id_arg, *, old_principal_arg, clientid_ret, TRUE }
Since the value of the clientdesc.id subfield of each client
record must be unique, there is no modification of the
server's state, and NFS4ERR_CLID_INUSE is returned to indicate
the client should retry with a different value for the
clientdesc.id subfield of CREATECLIENTID4args.
This scenario may also represent a malicious attempt to
destroy a client's state on the server. For security reasons,
the server MUST NOT remove the client's state when there is a
principal mismatch.
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4. Replay
If the server has the following unconfirmed record then this
request is likely the result of a client replay due to a
network partition or some other connection failure.
{ id_arg, verifier_arg, principal_arg, clientid_ret, FALSE }
Since the response to the CREATECLIENTID request that created
this record may have been lost, it is not acceptable to drop
this duplicate request. However, rather than processing it
normally, the existing record is left unchanged and
clientid_ret, which was generated for the previous request, is
returned.
5. Change of Principal
If the server has the following unconfirmed record then this
request is likely the result of a client which has for
whatever reasons changed principals (possibly to change
security flavor) after calling CREATECLIENTID, but before
calling CREATESESSION.
{ id_arg, verifier_arg, old_principal_arg, clientid_ret,
FALSE}
Since the client has not changed, the principal field of the
unconfirmed record is updated to principal_arg and
clientid_ret is again returned. There is a small possibility
that this is merely a collision on the client field of
CREATECLIENTID4args between unrelated clients, but since that
is unlikely, and an unconfirmed record does not generally have
any filesystem pertinent state, we can assume it is the same
client without risking loss of any important state.
After processing, the following record will exist on the
server.
{ id_arg, verifier_arg, principal_arg, clientid_ret, FALSE}
6. Client Reboot
If the server has the following confirmed client record, then
this request is likely from a previously confirmed client
which has rebooted.
{ id_arg, old_verifier_arg, principal_arg, clientid_ret, TRUE
}
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Since the previous incarnation of the same client will no
longer be making requests, lock and share reservations should
be released immediately rather than forcing the new
incarnation to wait for the lease time on the previous
incarnation to expire. Furthermore, session state should be
removed since if the client had maintained that information
across reboot, this request would not have been issued. If
the server does not support the CLAIM_DELEGATE_PREV claim
type, associated delegations should be purged as well;
otherwise, delegations are retained and recovery proceeds
according to RFC3530. The client record is updated with the
new verifier and its status is changed to unconfirmed.
After processing, clientid_ret is returned to the client and
the following record will exist on the server.
{ id_arg, verifier_arg, principal_arg, clientid_ret, FALSE }
7. Reboot before confirmation
If the server has the following unconfirmed record, then this
request is likely from a client which rebooted before sending
a CREATESESSION request.
{ id_arg, old_verifier_arg, *, clientid_ret, FALSE }
Since this is believed to be a request from a new incarnation
of the original client, the server updates the value of
clientdesc.verifier and returns the original clientid_ret.
After processing, the following state exists on the server.
{ id_arg, verifier_arg, *, clientid_ret, FALSE }
22.39.6. ERROR
NFS4ERR_BADXDR NFS4ERR_CLID_INUSE NFS4ERR_INVAL NFS4ERR_RESOURCE
NFS4ERR_SERVERFAULT
22.40. CREATESESSION - Create New Session and Confirm Clientid
Start up session and confirm clientid.
22.40.1. SYNOPSIS
clientid, session_args -> sessionid, session_args
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22.40.2. ARGUMENT
struct CREATESESSION4args {
clientid4 clientid;
bool persist;
count4 maxrequestsize;
count4 maxresponsesize;
count4 maxrequests;
count4 headerpadsize;
switch (bool clientid_confirm) {
case TRUE:
verifier4 setclientid_confirm;
case FALSE:
void;
}
switch (channelmode4 mode) {
case DEFAULT:
void;
case STREAM:
streamchannelattrs4 streamchanattrs;
case RDMA:
rdmachannelattrs4 rdmachanattrs;
};
};
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22.40.3. RESULT
typedef opaque sessionid4[16];
struct CREATESESSION4resok {
sessionid4 sessionid;
bool persist;
count4 maxrequestsize;
count4 maxresponsesize;
count4 maxrequests;
count4 headerpadsize;
switch (channelmode4 mode) {
case DEFAULT:
void;
case STREAM:
streamchannelattrs4 streamchanattrs;
case RDMA:
rdmachannelattrs4 rdmachanattrs;
};
};
union CREATESESSION4res switch (nfsstat4 status) {
case NFS4_OK:
CREATESESSION4resok resok4;
default:
void;
};
22.40.4. DESCRIPTION
This operation is used by the client to create new session objects on
the server. Additionally the first session created with a new
shorthand client identifier serves to confirm the creation of that
client's state on the server. The server returns the parameter
values for the new session.
22.40.5. IMPLEMENTATION
To describe the implementation, the same notation for client records
introduced in the description of CREATECLIENTID is used with the
following addition.
clientid_arg: The value of the clientid field of the
CREATESESSION4args structure of the current request.
Since CREATESESSION is a non-idempotent operation, we must consider
the possibility that replays may occur as a result of a client
reboot, network partition, malfunctioning router, etc. Replays are
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identified by the value of the clientid and sessionid fields of
CREATESESSION4args and the method for dealing with them is outlined
in the scenarios below.
The processing of this operation is divided into two phases: clientid
confirmation and session creation. In case the state for the
provided clientid has not been verified, it is confirmed before the
session is created. Otherwise the clientid confirmation phase is
skipped and only the session creation phase occurs. Note that since
only confirmed clients may create sessions, the clientid confirmation
stage does not depend upon sessionid_arg.
CLIENTID CONFIRMATION
The operational cases are described in terms of what client records
whose clientid field have value equal to clientid_arg exist in the
server's set of client records. Any cases in which there is more
than one record with identical values for clientid represent a server
implementation error. Operation in the potential valid cases is
summarized as follows.
1. Common Case
If the server has the following unconfirmed record, then this
is the expected confirmation of an unconfirmed record.
{ *, *, principal_arg, clientid_arg, FALSE }
The confirmed field of the record is set to TRUE and
processing of the operation continues normally.
2. Stale Clientid
If the server contains no records with clientid equal to
clientid_arg, then most likely the client's state has been
purged during a period of inactivity, possibly due to a loss
of connectivity. NFS4ERR_STALE_CLIENTID is returned, and no
changes are made to any client records on the server.
3. Principal Change or Collision
If the server has the following record, then the client has
changed principals after the previous CREATECLIENTID request,
or there has been a chance collision between shortand client
identifiers.
{ *, *, old_principal_arg, clientid_arg, * }
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Neither of these cases are permissible. Processing stops and
NFS4ERR_CLID_INUSE is returned to the client. No changes are
made to any client records on the server.
SESSION CREATION
To determine whether this request is a replay, the server examines
the sessionid argument provided by the client. If the sessionid
matches the identifier of a previously created session, then this
request must be interpreted as a replay. No new state is created and
a reply with the parameters of the existing session is returned to
the client. If a session corresponding to the sessionid does not
already exist, then the request is not a replay and is processed as
follows.
NOTE: It is the responsibility of the client to generate appropriate
values for sessionid. Since the ordering of messages sent on
different transport connections is not guaranteed, immediately
reusing the sessionid of a previously destroyed session may yield
unpredictable results. Client implementations should avoid recently
used sessionids to ensure correct behavior.
The server examines the persist, maxrequestsize, maxresponsesize,
maxrequests and headerpadsize arguments. For each argument, if the
value is acceptable to the server, it is recommended that the server
use the provided value to create the new session. If it is not
acceptable, the server may use a different value, but must return the
value used to the client. These parameters have the following
interpretation.
persist:
True if the client desires server support for "reliable"
semantics. For sessions in which only idempotent operations will
be used (e.g. a read-only session), clients should set this value
to false. If the server does not or cannot provide "reliable"
semantics this value must be set to false on return.
maxrequestsize:
The maximum size of a COMPOUND request that will be sent by the
client including RPC headers.
maxresponsesize:
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The maximum size of a COMPOUND reply that the client will accept
from the server including RPC headers. The server must not
increase the value of this parameter. If a client sends a
COMPOUND request for which the size of the reply would exceed this
value, the server will return NFS4ERR_RESOURCE.
maxrequests:
The maximum number of concurrent COMPOUND requests that the client
will issue on the session. Subsequent COMPOUND requests will each
be assigned a slot identifier by the client on the range 0 to
maxrequests - 1 inclusive. A slot id cannot be reused until the
previous request on that slot has completed.
headerpadsize:
The maximum amount of padding the client is willing to apply to
ensure that write payloads are aligned on some boundary at the
server. The server should reply with its preferred value, or zero
if padding is not in use. The server may decrease this value but
must not increase it.
The server creates the session by recording the parameter values used
and if the persist parameter is true and has been accepted by the
server, allocating space for the duplicate request cache (DRC).
If the session state is created successfully, the server associates
it with the session identifier provided by the client. This
identifier must be unique among the client's active sessions but
there is no need for it to be globally unique. Finally, the server
returns the negotiated values used to create the session to the
client.
22.40.6. ERRORS
NFS4ERR_BADXDR NFS4ERR_CLID_INUSE NFS4ERR_RESOURCE
NFS4ERR_SERVERFAULT NFS4ERR_STALE_CLIENTID
22.41. BIND_BACKCHANNEL - Create a callback channel binding
Establish a callback channel on the connection.
22.41.1. SYNOPSIS
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22.41.2. ARGUMENT
struct BIND_BACKCHANNEL4args {
clientid4 clientid;
uint32_t callback_program;
uint32_t callback_ident;
count4 maxrequestsize;
count4 maxresponsesize;
count4 maxrequests;
switch (channelmode4 mode) {
case DEFAULT:
void;
case STREAM:
streamchannelattrs4 streamchanattrs;
case RDMA:
rdmachannelattrs4 rdmachanattrs;
};
};
22.41.3. RESULT
struct BIND_BACKCHANNEL4resok {
count4 maxrequestsize;
count4 maxresponsesize;
count4 maxrequests;
switch (channelmode4 mode) {
case DEFAULT:
void;
case STREAM:
streamchannelattrs4 streamchanattrs;
case RDMA:
rdmachannelattrs4 rdmachanattrs;
};
};
union BIND_BACKCHANNEL4res switch (nfsstat4 status) {
case NFS4_OK:
BIND_BACKCHANNEL4resok resok4;
default:
void;
};
22.41.4. DESCRIPTION
The BIND_BACKCHANNEL operation serves to establish the current
connection as a designated callback channel for the specified
session. Normally, only one callback channel is bound, however if
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more than one are established, they are used at the server's
prerogative, no affinity or preference is specified by the client.
The arguments and results of the BIND_BACKCHANNEL call are a subset
of the session parameters, and used identically to those values on
the callback channel only. However, not all session operation
channel parameters are relevant to the callback channel, for example
header padding (since writes of bulk data are not performed in
callbacks).
22.41.5. IMPLEMENTATION
No discussion at this time.
22.41.6. ERRORS
TBD
22.42. DESTROYSESSION - Destroy existing session
Destroy existing session.
22.42.1. SYNOPSIS
void -> status
22.42.2. ARGUMENT
struct DESTROYSESSION4args {
sessionid4 sessionid;
};
22.42.3. RESULT
struct SESSION_DESTROYres {
nfsstat status;
};
22.42.4. DESCRIPTION
The SESSION_DESTROY operation closes the session and discards any
active state such as locks, leases, and server duplicate request
cache entries. Any remaining connections bound to the session are
immediately unbound and may additionally be closed by the server.
This operation must be the final, or only operation in any request.
Because the operation results in destruction of the session, any
duplicate request caching for this request, as well as previously
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completed requests, will be lost. For this reason, it is advisable
to not place this operation in a request with other state-modifying
operations. In addition, a SEQUENCE operation is not required in the
request.
Note that because the operation will never be replayed by the server,
a client that retransmits the request may receive an error in
response, even though the session may have been successfully
destroyed.
22.42.5. IMPLEMENTATION
No discussion at this time.
22.42.6. ERRORS
TBD
22.43. SEQUENCE - Supply per-procedure sequencing and control
Supply per-procedure sequencing and control
22.43.1. SYNOPSIS
control -> control
22.43.2. ARGUMENT
typedef uint32_t sequenceid4;
typedef uint32_t slotid4;
struct SEQUENCE4args {
sessionid4 sessionid;
sequenceid4 sequenceid;
slotid4 slotid;
slotid4 maxslot;
};
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22.43.3. RESULT
struct SEQUENCE4resok {
sessionid4 sessionid;
sequenceid4 sequenceid;
slotid4 slotid;
slotid4 maxslot;
slotid4 target_maxslot;
};
union SEQUENCE4res switch (nfsstat4 status) {
case NFS4_OK:
SEQUENCE4resok resok4;
default:
void;
};
22.43.4. DESCRIPTION
The SEQUENCE operation is used to manage operational accounting for
the session on which the operation is sent. The contents include the
client and session to which this request belongs, slotid and
sequenceid, used by the server to implement session request control
and the duplicate reply cache semantics, and exchanged slot counts
which are used to adjust these values. This operation must appear
once as the first operation in each COMPOUND sent after the channel
is successfully bound, or a protocol error must result.
22.43.5. IMPLEMENTATION
No discussion at this time.
22.43.6. ERRORS
NFS4ERR_BADSESSION NFS4ERR_BADSLOT
22.44. GET_DIR_DELEGATION - Get a directory delegation
Obtain a directory delegation.
22.44.1. SYNOPSIS
(cfh), requested notification ->
(cfh), cookieverf, stateid, supported notification
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22.44.2. ARGUMENT
/*
* Notification types.
*/
const DIR_NOTIFICATION_NONE = 0x00000000;
const DIR_NOTIFICATION_CHANGE_CHILD_ATTRIBUTES = 0x00000001;
const DIR_NOTIFICATION_CHANGE_DIR_ATTRIBUTES = 0x00000002;
const DIR_NOTIFICATION_REMOVE_ENTRY = 0x00000004;
const DIR_NOTIFICATION_ADD_ENTRY = 0x00000008;
const DIR_NOTIFICATION_RENAME_ENTRY = 0x00000010;
const DIR_NOTIFICATION_CHANGE_COOKIE_VERIFIER = 0x00000020;
typedef uint32_t dir_notification_type4;
typedef nfstime4 attr_notice4;
struct GET_DIR_DELEGATION4args {
dir_notification_type4 notification_type;
attr_notice4 child_attr_delay;
attr_notice4 dir_attr_delay;
};
22.44.3. RESULT
struct GET_DIR_DELEGATION4resok {
verifier4 cookieverf;
/* Stateid for get_dir_delegation */
stateid4 stateid;
/* Which notifications can the server support */
dir_notification_type4 notification;
bitmap4 child_attributes;
bitmap4 dir_attributes;
};
union GET_DIR_DELEGATION4res switch (nfsstat4 status) {
case NFS4_OK:
/* CURRENT_FH: delegated dir */
GET_DIR_DELEGATION4resok resok4;
default:
void;
};
22.44.4. DESCRIPTION
The GET_DIR_DELEGATION operation is used by a client to request a
directory delegation. The directory is represented by the current
filehandle. The client also specifies whether it wants the server to
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notify it when the directory changes in certain ways by setting one
or more bits in a bitmap. The server may also choose not to grant
the delegation. In that case the server will return
NFS4ERR_DIRDELEG_UNAVAIL. If the server decides to hand out the
delegation, it will return a cookie verifier for that directory. If
the cookie verifier changes when the client is holding the
delegation, the delegation will be recalled unless the client has
asked for notification for this event. In that case a notification
will be sent to the client.
The server will also return a directory delegation stateid in
addition to the cookie verifier as a result of the GET_DIR_DELEGATION
operation. This stateid will appear in callback messages related to
the delegation, such as notifications and delegation recalls. The
client will use this stateid to return the delegation voluntarily or
upon recall. A delegation is returned by calling the DELEGRETURN
operation.
The server may not be able to support notifications of certain
events. If the client asks for such notifications, the server must
inform the client of its inability to do so as part of the
GET_DIR_DELEGATION reply by not setting the appropriate bits in the
supported notifications bitmask contained in the reply.
The GET_DIR_DELEGATION operation can be used for both normal and
named attribute directories. It covers all the entries in the
directory except the ".." entry. That means if a directory and its
parent both hold directory delegations, any changes to the parent
will not cause a notification to be sent for the child even though
the child's ".." entry points to the parent.
22.44.5. IMPLEMENTATION
Directory delegation provides the benefit of improving cache
consistency of namespace information. This is done through
synchronous callbacks. A server must support synchronous callbacks
in order to support directory delegations. In addition to that,
asynchronous notifications provide a way to reduce network traffic as
well as improve client performance in certain conditions.
Notifications would not be requested when the goal is just cache
consitency.
Notifications are specified in terms of potential changes to the
directory. A client can ask to be notified whenever an entry is
added to a directory by setting notification_type to
DIR_NOTIFICATION_ADD_ENTRY. It can also ask for notifications on
entry removal, renames, directory attribute changes and cookie
verifier changes by setting notification_type flag appropriately. In
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addition to that, the client can also ask for notifications upon
attribute changes to children in the directory to keep its attribute
cache up to date. However any changes made to child attributes do
not cause the delegation to be recalled. If a client is interested
in directory entry caching, or negative name caching, it can set the
notification_type appropriately and the server will notify it of all
changes that would otherwise invalidate its name cache. The kind of
notification a client asks for may depend on the directory size, its
rate of change and the applications being used to access that
directory. However, the conditions under which a client might ask
for a notification, is out of the scope of this specification.
The client will set one or more bits in a bitmap (notification_type)
to let the server know what kind of notification(s) it is interested
in. For attribute notifications it will set bits in another bitmap
to indicate which attributes it wants to be notified of. If the
server does not support notifications for changes to a certain
attribute, it should not set that attribute in the supported
attribute bitmap (notification) specified in the reply.
In addition to that, the client will also let the server know if it
wants to get the notification as soon as the attribute change occurs
or after a certain delay by setting a delay factor, child_attr_delay
for attribute changes to children and dir_attr_delay for attribute
changes to the directory. If this delay factor is set to zero, that
indicates to the server that the client wants to be notified of any
attribute changes as soon as they occur. If the delay factor is set
to N, the server will make a best effort guarantee that attribute
updates are not out of sync by more than that. One value covers all
attribute changes for the directory and another value covers all
attribute changes for all children in the directory. If the client
asks for a delay factor that the server does not support or that may
cause significant resource consumption on the server by causing the
server to send a lot of notifications, the server should not commit
to sending out notifications for that attribute and therefore must
not set the approprite bit in the child_attributes and dir_attributes
bitmaps in the response.
The server will let the client know about which notifications it can
support by setting appropriate bits in a bitmap. If it agrees to
send attribute notifications, it will also set two attribute masks
indicating which attributes it will send change notifications for.
One of the masks covers changes in directory attributes and the other
covers atttribute changes to any files in the directory.
The client should use a security flavor that the filesystem is
exported with. If it uses a different flavor, the server should
return NFS4ERR_WRONGSEC.
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22.44.6. ERRORS
NFS4ERR_ACCESS NFS4ERR_BADHANDLE NFS4ERR_BADXDR NFS4ERR_FHEXPIRED
NFS4ERR_INVAL NFS4ERR_MOVED NFS4ERR_NOFILEHANDLE NFS4ERR_NOTDIR
NFS4ERR_RESOURCE NFS4ERR_SERVERFAULT NFS4ERR_STALE
NFS4ERR_DIRDELEG_UNAVAIL NFS4ERR_WRONGSEC NFS4ERR_EIO NFS4ERR_NOTSUPP
22.45. LAYOUTGET - Get Layout Information
22.45.1. SYNOPSIS
(cfh), clientid, layout_type, iomode, offset,
length, minlength, maxcount -> layout example synopsis
22.45.2. ARGUMENT
struct LAYOUTGET4args {
/* CURRENT_FH: file */
clientid4 clientid;
pnfs_layouttype4 layout_type;
pnfs_layoutiomode4 iomode;
offset4 offset;
length4 length;
length4 minlength;
count4 maxcount;
};
22.45.3. RESULT
struct LAYOUTGET4resok {
pnfs_layout4 layout;
};
union LAYOUTGET4res switch (nfsstat4 status) {
case NFS4_OK:
LAYOUTGET4resok resok4;
default:
void;
};
22.45.4. DESCRIPTION
Requests a layout for reading or writing (and reading) the file given
by the filehandle at the byte range specified by offset and length.
Layouts are identified by the clientid, filehandle, and layout type.
The use of the iomode depends upon the layout type, but should
reflect the client's data access intent.
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The LAYOUTGET operation returns layout information for the specified
byte range, a layout segment. To get a layout segment from a
specific offset through the end-of-file, regardless of the file's
length, a length field with all bits set to 1 (one) should be used.
If the length is zero, or if a length which is not all bits set to
one is specified, and length when added to the offset exceeds the
maximum 64-bit unsigned integer value, the error NFS4ERR_INVAL will
result.
The "minlength" field specifies the minimum size overlap with the
requested offset and length that is to be returned. If this
requirement cannot be met, no layout must be returned; the error
NFS4ERR_LAYOUTTRYLATER can be returned.
The "maxcount" field specifies the maximum layout size (in bytes)
that the client can handle. If the size of the layout structure
exceeds the size specified by maxcount, the metadata server will
return the NFS4ERR_TOOSMALL error.
As well, the metadata server may adjust the range of the returned
layout segment based on striping patterns and usage implied by the
iomode. The client must be prepared to get a layout that does not
line up exactly with their request; there MUST be at least an overlap
of "minlength" between the layout returned by the server and the
client's request, or the server SHOULD reject the request. See
Section 14.3 for more details.
The metadata server may also return a layout segment with an iomode
other than that requested by the client. If it does so, it must
ensure that the iomode is more permissive than the iomode requested.
E.g., this allows an implementation to upgrade read-only requests to
read/write requests at its discretion, within the limits of the
layout type specific protocol. An iomode of either LAYOUTIOMODE_READ
or LAYOUTIOMODE_RW must be returned.
The format of the returned layout is specific to the underlying file
system. Layout types other than the NFSv4 file layout type should be
specified outside of this document.
If layouts are not supported for the requested file or its containing
file system the server SHOULD return NFS4ERR_LAYOUTUNAVAILABLE. If
the layout type is not supported, the metadata server should return
NFS4ERR_UNKNOWN_LAYOUTTYPE. If layouts are supported but no layout
matches the client provided layout identification, the server should
return NFS4ERR_BADLAYOUT. If an invalid iomode is specified, or an
iomode of LAYOUTIOMODE_ANY is specified, the server should return
NFS4ERR_BADIOMODE.
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If the layout for the file is unavailable due to transient
conditions, e.g. file sharing prohibits layouts, the server must
return NFS4ERR_LAYOUTTRYLATER.
If the layout request is rejected due to an overlapping layout
recall, the server must return NFS4ERR_RECALLCONFLICT. See
Section 14.5.3 for details.
If the layout conflicts with a mandatory byte range lock held on the
file, and if the storage devices have no method of enforcing
mandatory locks, other than through the restriction of layouts, the
metadata server should return NFS4ERR_LOCKED.
On success, the current filehandle retains its value.
22.45.5. IMPLEMENTATION
Typically, LAYOUTGET will be called as part of a compound RPC after
an OPEN operation and results in the client having location
information for the file; a client may also hold a layout across
multiple OPENs. The client specifies a layout type that limits what
kind of layout the server will return. This prevents servers from
issuing layouts that are unusable by the client.
22.45.6. ERRORS
NFS4ERR_BADLAYOUT NFS4ERR_BADIOMODE NFS4ERR_FHEXPIRED NFS4ERR_INVAL
NFS4ERR_LAYOUTUNAVAILABLE NFS4ERR_LAYOUTTRYLATER NFS4ERR_LOCKED
NFS4ERR_NOFILEHANDLE NFS4ERR_NOTSUPP NFS4ERR_RECALLCONFLICT
NFS4ERR_STALE NFS4ERR_STALE_CLIENTID NFS4ERR_TOOSMALL
NFS4ERR_UNKNOWN_LAYOUTTYPE
22.46. LAYOUTCOMMIT - Commit writes made using a layout
22.46.1. SYNOPSIS
(cfh), clientid, offset, length, reclaim, last_write_offset,
time_modify, time_access, layoutupdate -> newsize
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22.46.2. ARGUMENT
union newtime4 switch (bool timechanged) {
case TRUE:
nfstime4 time;
case FALSE:
void;
};
union newsize4 switch (bool sizechanged) {
case TRUE:
length4 size;
case FALSE:
void;
};
struct LAYOUTCOMMIT4args {
/* CURRENT_FH: file */
clientid4 clientid;
offset4 offset;
length4 length;
bool reclaim;
length4 last_write_offset;
newtime4 time_modify;
newtime4 time_access;
pnfs_layoutupdate4 layoutupdate;
};
22.46.3. RESULT
union LAYOUTCOMMIT4res switch (nfsstat4 status) {
case NFS4_OK:
LAYOUTCOMMIT4resok resok4;
default:
void;
};
struct LAYOUTCOMMIT4resok {
newsize4 newsize;
};
22.46.4. DESCRIPTION
Commits changes in the layout segment represented by the current
filehandle, clientid, and byte range. Since layouts are sub-
dividable, a smaller portion of a layout, retrieved via LAYOUTGET,
may be committed. The region being committed is specified through
the byte range (length and offset). Note: the "layoutupdate"
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structure does not include the length and offset, as they are already
specified in the arguments.
The LAYOUTCOMMIT operation indicates that the client has completed
writes using a layout obtained by a previous LAYOUTGET. The client
may have only written a subset of the data range it previously
requested. LAYOUTCOMMIT allows it to commit or discard provisionally
allocated space and to update the server with a new end of file. The
layout referenced by LAYOUTCOMMIT is still valid after the operation
completes and can be continued to be referenced by the clientid,
filehandle, byte range, and layout type.
The "reclaim" field set to "true" in a LAYOUTCOMMIT request specifies
that the client is attempting to commit changes to a layout after the
reboot of the metadata server during the metadata server's recovery
grace period. This type of request may be necessary when the client
has uncommitted writes to provisionally allocated regions of a file
which were sent to the storage devices before the reboot of the
metadata server. In this case the layout provided by the client MUST
be a subset of a writable layout that the client held immediately
before the reboot of the metadata server. The metadata server is
free to accept or reject this request based on its own internal
metadata consistency checks. If the metadata server finds that the
layout provided by the client does not pass its consistency checks,
it MUST reject the request with the status NFS4ERR_RECLAIM_BAD. The
successful completion of the LAYOUTCOMMIT request with "reclaim" set
to true does NOT provide the client with a layout for the file. It
simply commits the changes to the file layout specified in the
"layoutupdate" field. To obtain a layout for the file the client
must issue a LAYOUTGET request to the server after the server's grace
period has expired. If the metadata server receives a LAYOUTCOMMIT
request with "reclaim" set to true when the metadata server is not in
its recovery grace period, it MUST reject the request with the status
NFS4ERR_NO_GRACE.
Setting the "reclaim" field to "true" is required if and only if the
committed layout was acquired before the metadata server reboot.
Committing layouts that were acquired during the metadata server's
grace period MUST set the "reclaim" field to "false".
The "last_write_offset" field specifies the offset of the last byte
written by the client previous to the LAYOUTCOMMIT. Note: this value
is never equal to the file's size (at most it is one byte less than
the file's size). The metadata server may use this information to
determine whether the file's size needs to be updated. If the
metadata server updates the file's size as the result of the
LAYOUTCOMMIT operation, it must return the new size as part of the
results.
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The "time_modify" and "time_access" fields allow the client to
suggest times it would like the metadata server to set. The metadata
server may use these time values or it may use the time of the
LAYOUTCOMMIT operation to set these time values. If the metadata
server uses the client provided times, it should sanity check the
values (e.g., to ensure time does not flow backwards). If the client
wants to force the metadata server to set an exact time, the client
should use a SETATTR operation in a compound right after
LAYOUTCOMMIT. See Section 14.4 for more details. If the new client
desires the resultant mtime or atime, it should issue a GETATTR
following the LAYOUTCOMMIT; e.g., later in the same compound.
The "layoutupdate" argument to LAYOUTCOMMIT provides a mechanism for
a client to provide layout specific updates to the metadata server.
For example, the layout update can describe what regions of the
original layout have been used and what regions can be deallocated.
There is no NFSv4 file layout specific layoutupdate structure.
The layout information is more verbose for block devices than for
objects and files because the latter hide the details of block
allocation behind their storage protocols. At the minimum, the
client needs to communicate changes to the end of file location back
to the server, and, if desired, its view of the file modify and
access time. For block/volume layouts, it needs to specify precisely
which blocks have been used.
If the layout identified in the arguments does not exist, the error
NFS4ERR_BADLAYOUT is returned. The layout being committed may also
be rejected if it does not correspond to an existing layout with an
iomode of RW.
If the LAYOUTCOMMIT request sets the "reclaim" field to "true" after
the metadata server's grace period, NFS4ERR_NO_GRACE is returned.
On success, the current filehandle retains its value.
22.46.5. IMPLEMENTATION
Optionally, the client can also use LAYOUTCOMMIT with the "reclaim"
field set to "true" to convey hints to modified file attributes or to
report layout-type specific information such as I/O errors for
object-based storage layouts, as normally done during normal
operation. Doing so may help the metadata server to recover files
more efficiently after reboot. For example, some file system
implementations may require expansive recovery of filesystem objects
if the metadata server does not get a positive indication from all
clients holding a write layout that they have successfully completed
all their writes. Sending a LAYOUTCOMMIT (if required) and then
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following with LAYOUTRETURN can provide such an indication and allow
for graceful and efficient recovery.
22.46.6. ERRORS
NFS4ERR_BADLAYOUT NFS4ERR_BADIOMODE NFS4ERR_FHEXPIRED NFS4ERR_INVAL
NFS4ERR_NOFILEHANDLE NFS4ERR_NO_GRACE NFS4ERR_RECLAIM_BAD
NFS4ERR_STALE NFS4ERR_STALE_CLIENTID NFS4ERR_UNKNOWN_LAYOUTTYPE
22.47. LAYOUTRETURN - Release Layout Information
22.47.1. SYNOPSIS
(cfh), clientid, offset, length, reclaim, iomode, layout_type -> -
22.47.2. ARGUMENT
struct LAYOUTRETURN4args {
/* CURRENT_FH: file */
clientid4 clientid;
offset4 offset;
length4 length;
bool reclaim;
pnfs_layoutiomode4 iomode;
pnfs_layouttype4 layout_type;
};
22.47.3. RESULT
struct LAYOUTRETURN4res {
nfsstat4 status;
};
22.47.4. DESCRIPTION
Returns the layout segment represented by the current filehandle,
clientid, byte range, iomode, and layout type. After this call, the
client MUST NOT use the layout and the associated storage protocol to
access the file data. The layout being returned may be a subdivision
of a layout previously fetched through LAYOUTGET. As well, it may be
a subset or superset of a layout specified by CB_LAYOUTRECALL.
However, if it is a subset, the recall is not complete until the full
byte range has been returned. It is also permissible, and no error
should result, for a client to return a byte range covering a layout
it does not hold. If the length is all 1s, the layout covers the
range from offset to EOF. An iomode of ANY specifies that all
layouts that match the other arguments to LAYOUTRETURN (i.e.,
clientid, byte range, and type) are being returned.
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The "reclaim" field set to "true" in a LAYOUTRETURN request specifies
that the client is attempting to return a layout that was acquired
before the reboot of the metadata server during the metadata server's
grace period. Returning layouts that were acquired during the
metadata server's grace period MUST set the "reclaim" field to
"false". See LAYOUTCOMMIT (Section 22.46) for more details.
Layouts may be returned when recalled or voluntarily (i.e., before
the server has recalled them). In either case the client must
properly propagate state changed under the context of the layout to
storage or to the server before returning the layout.
If a client fails to return a layout in a timely manner, then the
metadata server should use its control protocol with the storage
devices to fence the client from accessing the data referenced by the
layout. See Section 14.5 for more details.
If the layout identified in the arguments does not exist, the error
NFS4ERR_BADLAYOUT is returned. If a layout exists, but the iomode
does not match, NFS4ERR_BADIOMODE is returned.
If the LAYOUTRETURN request sets the "reclaim" field to "true" after
the metadata server's grace period, NFS4ERR_NO_GRACE is returned.
On success, the current filehandle retains its value.
[[Comment.6: Should LAYOUTRETURN be modified to handle FSID
callbacks?]]
22.47.5. IMPLEMENTATION
22.47.6. ERRORS
NFS4ERR_BADLAYOUT NFS4ERR_BADIOMODE NFS4ERR_FHEXPIRED NFS4ERR_INVAL
NFS4ERR_NOFILEHANDLE NFS4ERR_NO_GRACE NFS4ERR_STALE
NFS4ERR_STALE_CLIENTID NFS4ERR_UNKNOWN_LAYOUTTYPE
22.48. GETDEVICEINFO - Get Device Information
22.48.1. SYNOPSIS
(cfh), device_id, layout_type, maxcount -> device_addr
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22.48.2. ARGUMENT
struct GETDEVICEINFO4args {
/* CURRENT_FH: file */
pnfs_deviceid4 device_id;
pnfs_layouttype4 layout_type;
count4 maxcount;
};
22.48.3. RESULT
struct GETDEVICEINFO4resok {
pnfs_deviceaddr4 device_addr;
};
union GETDEVICEINFO4res switch (nfsstat4 status) {
case NFS4_OK:
GETDEVICEINFO4resok resok4;
default:
void;
};
22.48.4. DESCRIPTION
Returns device type and device address information for a specified
device. The returned device_addr includes a type that indicates how
to interpret the addressing information for that device. The current
filehandle (cfh) is used to identify the file system; device IDs are
unique per file system (FSID) and are qualified by the layout type.
See Section 14.1.4 for more details on device ID assignment.
If the size of the device address exceeds maxcount bytes, the
metadata server will return the error NFS4ERR_TOOSMALL. If an
invalid device ID is given, the metadata server will respond with
NFS4ERR_INVAL.
22.48.5. IMPLEMENTATION
22.48.6. ERRORS
NFS4ERR_FHEXPIRED NFS4ERR_INVAL NFS4ERR_TOOSMALL
NFS4ERR_UNKNOWN_LAYOUTTYPE
22.49. GETDEVICELIST
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22.49.1. SYNOPSIS
(cfh), layout_type, maxcount, cookie, cookieverf ->
cookie, cookieverf, device_addrs<>
22.49.2. ARGUMENT
struct GETDEVICELIST4args {
/* CURRENT_FH: file */
pnfs_layouttype4 layout_type;
count4 maxcount;
nfs_cookie4 cookie;
verifier4 cookieverf;
};
22.49.3. RESULT
struct GETDEVICELIST4resok {
nfs_cookie4 cookie;
verifier4 cookieverf;
pnfs_devlist_item4 device_addrs<>;
};
union GETDEVICELIST4res switch (nfsstat4 status) {
case NFS4_OK:
GETDEVICELIST4resok resok4;
default:
void;
};
22.49.4. DESCRIPTION
In some applications, especially SAN environments, it is convenient
to find out about all the devices associated with a file system.
This lets a client determine if it has access to these devices, e.g.,
at mount time.
This operation returns an array of items (pnfs_devlist_item4) that
establish the association between the short pnfs_deviceid4 and the
addressing information for that device, for a particular layout type.
This operation may not be able to fetch all device information at
once, thus it uses a cookie based approach, similar to READDIR, to
fetch additional device information (see [6], section 14.2.24). As
in GETDEVICEINFO, the current filehandle (cfh) is used to identify
the file system.
As in GETDEVICEINFO, maxcount specifies the maximum number of bytes
to return. If the metadata server is unable to return a single
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device address, it will return the error NFS4ERR_TOOSMALL. If an
invalid device ID is given, the metadata server will respond with
NFS4ERR_INVAL.
22.49.5. IMPLEMENTATION
22.49.6. ERRORS
NFS4ERR_BAD_COOKIE NFS4ERR_FHEXPIRED NFS4ERR_INVAL NFS4ERR_TOOSMALL
NFS4ERR_UNKNOWN_LAYOUTTYPE
23. NFS version 4.1 Callback Procedures
The procedures used for callbacks are defined in the following
sections. In the interest of clarity, the terms "client" and
"server" refer to NFS clients and servers, despite the fact that for
an individual callback RPC, the sense of these terms would be
precisely the opposite.
23.1. Procedure 0: CB_NULL - No Operation
23.1.1. SYNOPSIS
23.1.2. ARGUMENTS
void;
23.1.3. RESULTS
void;
23.1.4. DESCRIPTION
Standard NULL procedure. Void argument, void response. Even though
there is no direct functionality associated with this procedure, the
server will use CB_NULL to confirm the existence of a path for RPCs
from server to client.
23.1.5. ERRORS
None.
23.2. Procedure 1: CB_COMPOUND - Compound Operations
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23.2.1. SYNOPSIS
compoundargs -> compoundres
23.2.2. ARGUMENTS
enum nfs_cb_opnum4 {
OP_CB_GETATTR = 3,
OP_CB_RECALL = 4,
OP_CB_ILLEGAL = 10044
};
union nfs_cb_argop4 switch (unsigned argop) {
case OP_CB_GETATTR: CB_GETATTR4args opcbgetattr;
case OP_CB_RECALL: CB_RECALL4args opcbrecall;
case OP_CB_ILLEGAL: void opcbillegal;
};
struct CB_COMPOUND4args {
utf8str_cs tag;
uint32_t minorversion;
uint32_t callback_ident;
nfs_cb_argop4 argarray<>;
};
23.2.3. RESULTS
union nfs_cb_resop4 switch (unsigned resop){
case OP_CB_GETATTR: CB_GETATTR4res opcbgetattr;
case OP_CB_RECALL: CB_RECALL4res opcbrecall;
};
struct CB_COMPOUND4res {
nfsstat4 status;
utf8str_cs tag;
nfs_cb_resop4 resarray<>;
};
23.2.4. DESCRIPTION
The CB_COMPOUND procedure is used to combine one or more of the
callback procedures into a single RPC request. The main callback RPC
program has two main procedures: CB_NULL and CB_COMPOUND. All other
operations use the CB_COMPOUND procedure as a wrapper.
In the processing of the CB_COMPOUND procedure, the client may find
that it does not have the available resources to execute any or all
of the operations within the CB_COMPOUND sequence. In this case, the
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error NFS4ERR_RESOURCE will be returned for the particular operation
within the CB_COMPOUND procedure where the resource exhaustion
occurred. This assumes that all previous operations within the
CB_COMPOUND sequence have been evaluated successfully.
Contained within the CB_COMPOUND results is a 'status' field. This
status must be equivalent to the status of the last operation that
was executed within the CB_COMPOUND procedure. Therefore, if an
operation incurred an error then the 'status' value will be the same
error value as is being returned for the operation that failed.
For the definition of the "tag" field, see the section "Procedure 1:
COMPOUND - Compound Operations".
The value of callback_ident is supplied by the client during
SETCLIENTID. The server must use the client supplied callback_ident
during the CB_COMPOUND to allow the client to properly identify the
server.
Illegal operation codes are handled in the same way as they are
handled for the COMPOUND procedure.
23.2.5. IMPLEMENTATION
The CB_COMPOUND procedure is used to combine individual operations
into a single RPC request. The client interprets each of the
operations in turn. If an operation is executed by the client and
the status of that operation is NFS4_OK, then the next operation in
the CB_COMPOUND procedure is executed. The client continues this
process until there are no more operations to be executed or one of
the operations has a status value other than NFS4_OK.
23.2.6. ERRORS
NFS4ERR_BADHANDLE NFS4ERR_BAD_STATEID NFS4ERR_BADXDR
NFS4ERR_OP_ILLEGAL NFS4ERR_RESOURCE NFS4ERR_SERVERFAULT
24. NFS version 4.1 Callback Operations
24.1. Operation 3: CB_GETATTR - Get Attributes
24.1.1. SYNOPSIS
fh, attr_request -> attrmask, attr_vals
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24.1.2. ARGUMENTS
struct CB_GETATTR4args {
nfs_fh4 fh;
bitmap4 attr_request;
};
24.1.3. RESULTS
struct CB_GETATTR4resok {
fattr4 obj_attributes;
};
union CB_GETATTR4res switch (nfsstat4 status) {
case NFS4_OK:
CB_GETATTR4resok resok4;
default:
void;
};
24.1.4. DESCRIPTION
The CB_GETATTR operation is used by the server to obtain the current
modified state of a file that has been write delegated. The
attributes size and change are the only ones guaranteed to be
serviced by the client. See the section "Handling of CB_GETATTR" for
a full description of how the client and server are to interact with
the use of CB_GETATTR.
If the filehandle specified is not one for which the client holds a
write open delegation, an NFS4ERR_BADHANDLE error is returned.
24.1.5. IMPLEMENTATION
The client returns attrmask bits and the associated attribute values
only for the change attribute, and attributes that it may change
(time_modify, and size).
24.1.6. ERRORS
NFS4ERR_BADHANDLE NFS4ERR_BADXDR NFS4ERR_RESOURCE NFS4ERR_SERVERFAULT
24.2. Operation 4: CB_RECALL - Recall an Open Delegation
24.2.1. SYNOPSIS
stateid, truncate, fh -> ()
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24.2.2. ARGUMENTS
struct CB_RECALL4args {
stateid4 stateid;
bool truncate;
nfs_fh4 fh;
};
24.2.3. RESULTS
struct CB_RECALL4res {
nfsstat4 status;
};
24.2.4. DESCRIPTION
The CB_RECALL operation is used to begin the process of recalling an
open delegation and returning it to the server.
The truncate flag is used to optimize recall for a file which is
about to be truncated to zero. When it is set, the client is freed
of obligation to propagate modified data for the file to the server,
since this data is irrelevant.
If the handle specified is not one for which the client holds an open
delegation, an NFS4ERR_BADHANDLE error is returned.
If the stateid specified is not one corresponding to an open
delegation for the file specified by the filehandle, an
NFS4ERR_BAD_STATEID is returned.
24.2.5. IMPLEMENTATION
The client should reply to the callback immediately. Replying does
not complete the recall except when an error was returned. The
recall is not complete until the delegation is returned using a
DELEGRETURN.
24.2.6. ERRORS
NFS4ERR_BADHANDLE NFS4ERR_BAD_STATEID NFS4ERR_BADXDR NFS4ERR_RESOURCE
NFS4ERR_SERVERFAULT
24.3. Operation 10044: CB_ILLEGAL - Illegal Callback Operation
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24.3.1. SYNOPSIS
-> ()
24.3.2. ARGUMENTS
void;
24.3.3. RESULTS
struct CB_ILLEGAL4res {
nfsstat4 status;
};
24.3.4. DESCRIPTION
This operation is a placeholder for encoding a result to handle the
case of the client sending an operation code within COMPOUND that is
not supported. See the COMPOUND procedure description for more
details.
The status field of CB_ILLEGAL4res MUST be set to NFS4ERR_OP_ILLEGAL.
24.3.5. IMPLEMENTATION
A server will probably not send an operation with code OP_CB_ILLEGAL
but if it does, the response will be CB_ILLEGAL4res just as it would
be with any other invalid operation code. Note that if the client
gets an illegal operation code that is not OP_ILLEGAL, and if the
client checks for legal operation codes during the XDR decode phase,
then the CB_ILLEGAL4res would not be returned.
24.3.6. ERRORS
NFS4ERR_OP_ILLEGAL
24.4. CB_RECALLCREDIT - change flow control limits
Change flow control limits
24.4.1. SYNOPSIS
targetcount -> status
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24.4.2. ARGUMENT
struct CB_RECALLCREDIT4args {
sessionid4 sessionid;
uint32_t target;
};
24.4.3. RESULT
struct CB_RECALLCREDIT4res {
nfsstat4 status;
};
24.4.4. DESCRIPTION
The CB_RECALLCREDIT operation requests the client to return session
and transport credits to the server, by zero-length RDMA Sends or
NULL NFSv4 operations.
24.4.5. IMPLEMENTATION
No discussion at this time.
24.4.6. ERRORS
NONE
24.5. CB_SEQUENCE - Supply callback channel sequencing and control
Sequence and control
24.5.1. SYNOPSIS
control -> control
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24.5.2. ARGUMENT
typedef uint32_t sequenceid4;
typedef uint32_t slotid4;
struct CB_SEQUENCE4args {
sessionid4 sessionid;
sequenceid4 sequenceid;
slotid4 slotid;
slotid4 maxslot;
sequenceid4 referring_sequenceid;
slotid4 referring_slotid;
};
24.5.3. RESULT
struct CB_SEQUENCE4resok {
sessionid4 sessionid;
sequenceid4 sequenceid;
slotid4 slotid;
slotid4 maxslot;
slotid4 target_maxslot;
};
union CB_SEQUENCE4res switch (nfsstat4 status) {
case NFS4_OK:
CB_SEQUENCE4resok resok4;
default:
void;
};
24.5.4. DESCRIPTION
The CB_SEQUENCE operation is used to manage operational accounting
for the callback channel of the session on which the operation is
sent. The contents include the client and session to which this
request belongs, slotid and sequenceid, used by the server to
implement session request control and the duplicate reply cache
semantics, and exchanged slot counts which are used to adjust these
values. This operation must appear once as the first operation in
each CB_COMPOUND sent after the callback channel is successfully
bound, or a protocol error must result.
24.5.5. IMPLEMENTATION
No discussion at this time.
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24.5.6. ERRORS
NFS4ERR_BADSESSION NFS4ERR_BADSLOT
24.6. CB_NOTIFY - Notify directory changes
Tell the client of directory changes.
24.6.1. SYNOPSIS
stateid, notification -> {}
24.6.2. ARGUMENT
/*
* Notification information sent to the client.
*/
union dir_notification4
switch (dir_notification_type4 notification_type) {
case DIR_NOTIFICATION_CHANGE_CHILD_ATTRIBUTES:
dir_notification_attribute4 change_child_attributes;
case DIR_NOTIFICATION_CHANGE_DIR_ATTRIBUTES:
fattr4 change_dir_attributes;
case DIR_NOTIFICATION_REMOVE_ENTRY:
dir_notification_remove4 remove_notification;
case DIR_NOTIFICATION_ADD_ENTRY:
dir_notification_add4 add_notification;
case DIR_NOTIFICATION_RENAME_ENTRY:
dir_notification_rename4 rename_notification;
case DIR_NOTIFICATION_CHANGE_COOKIE_VERIFIER:
dir_notification_verifier4 verf_notification;
};
/*
* Changed entry information.
*/
struct dir_entry {
component4 file;
fattr4 attrs;
};
struct dir_notification_attribute4 {
dir_entry changed_entry;
};
struct dir_notification_remove4 {
dir_entry old_entry;
nfs_cookie4 old_entry_cookie;
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};
struct dir_notification_rename4 {
dir_entry old_entry;
dir_notification_add4 new_entry;
};
struct dir_notification_verifier4 {
verifier4 old_cookieverf;
verifier4 new_cookieverf;
};
struct dir_notification_add4 {
dir_entry new_entry;
/* what READDIR would have returned for this entry */
nfs_cookie4 new_entry_cookie;
bool last_entry;
prev_entry_info4 prev_info;
};
union prev_entry_info4 switch (bool isprev) {
case TRUE: /* A previous entry exists */
prev_entry4 prev_entry_info;
case FALSE: /* we are adding to an empty
directory */
void;
};
/*
* Previous entry information
*/
struct prev_entry4 {
dir_entry prev_entry;
/* what READDIR returned for this entry */
nfs_cookie4 prev_entry_cookie;
};
struct CB_NOTIFY4args {
stateid4 stateid;
dir_notification4 changes<>;
};
24.6.3. RESULT
struct CB_NOTIFY4res {
nfsstat4 status;
};
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24.6.4. DESCRIPTION
The CB_NOTIFY operation is used by the server to send notifications
to clients about changes in a delegated directory. These
notifications are sent over the callback path. The notification is
sent once the original request has been processed on the server. The
server will send an array of notifications for all changes that might
have occurred in the directory. The dir_notification_type4 can only
have one bit set for each notification in the array. If the client
holding the delegation makes any changes in the directory that cause
files or sub directories to be added or removed, the server will
notify that client of the resulting change(s). If the client holding
the delegation is making attribute or cookie verifier changes only,
the server does not need to send notifications to that client. The
server will send the following information for each operation:
ADDING A FILE The server will send information about the new entry
being created along with the cookie for that entry. The entry
information contains the nfs name of the entry and attributes. If
this entry is added to the end of the directory, the server will
set a last_entry flag to true. If the file is added such that
there is atleast one entry before it, the server will also return
the previous entry information along with its cookie. This is to
help clients find the right location in their DNLC or directory
caches where this entry should be cached.
REMOVING A FILE The server will send information about the directory
entry being deleted. The server will also send the cookie value
for the deleted entry so that clients can get to the cached
information for this entry.
RENAMING A FILE The server will send information about both the old
entry and the new entry. This includes name and attributes for
each entry. This notification is only sent if both entries are in
the same directory. If the rename is across directories, the
server will send a remove notification to one directory and an add
notification to the other directory, assuming both have a
directory delegation.
FILE/DIR ATTRIBUTE CHANGE The client will use the attribute mask to
inform the server of attributes for which it wants to receive
notifications. This change notification can be requested for both
changes to the attributes of the directory as well as changes to
any file attributes in the directory by using two separate
attribute masks. The client can not ask for change attribute
notification per file. One attribute mask covers all the files in
the directory. Upon any attribute change, the server will send
back the values of changed attributes. Notifications might not
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make sense for some filesystem wide attributes and it is up to the
server to decide which subset it wants to support. The client can
negotiate the frequency of attribute notifications by letting the
server know how often it wants to be notified of an attribute
change. The server will return supported notification frequencies
or an indication that no notification is permitted for directory
or child attributes by setting the dir_notif_delay and
dir_entry_notif_delay attributes respectively.
COOKIE VERIFIER CHANGE If the cookie verifier changes while a client
is holding a delegation, the server will notify the client so that
it can invalidate its cookies and reissue a READDIR to get the new
set of cookies.
24.6.5. IMPLEMENTATION
24.6.6. ERRORS
NFS4ERR_BAD_STATEID NFS4ERR_INVAL NFS4ERR_BADXDR NFS4ERR_SERVERFAULT
24.7. CB_RECALL_ANY - Keep any N delegations
Notify client to return delegation and keep N of them.
24.7.1. SYNOPSIS
N, type_mask -> {}
24.7.2. ARGUMENT
const TYPE_MASK_RDATA_DLG = 0;
const TYPE_MASK_WDATA_DLG = 1;
const TYPE_MASK_DIR_DLG = 2;
const TYPE_MASK_FILE_LAYOUT = 3;
const TYPE_MASK_BLK_LAYOUT_MIN = 4;
const TYPE_MASK_BLK_LAYOUT_MAX = 7;
const TYPE_MASK_OBJ_LAYOUT_MIN = 8;
const TYPE_MASK_OBJ_LAYOUT_MAX = 11;
const TYPE_MASK_OTHER_LAYOUT_MIN = 12;
const TYPE_MASK_OTHER_LAYOUT_MAX = 15;
struct CB_RECALLANYY4args {
uint4 objects_to_keep;
bitmap4 type_mask;
}
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24.7.3. RESULT
struct CB_RECALLANY4res {
nfsstat4 status;
};
24.7.4. DESCRIPTION
The server may decide that it cannot hold all of the state for
recallable objects, such as delegations and layouts, without running
out of resources. In such a case, it is free to recall individual
objects to reduce the load but this would be far from optimal.
Because the general purpose of such recallable objects as delegations
is to eliminate client interaction with the server, the server cannot
interpret lack of recent use as indicating that the object is no
longer useful. The absence of visible use may be the result of a
large number of potential operations eliminated. In the case of
layouts, the layout will be used explicitly but the meta-data server
does not have direct knowledge of such use.
In order to implement an effective reclaim scheme for such objects,
the server's knowledge of available resources must be used to
determine when objects must be recalled with the clients selecting
the actual objects to be returned.
Server implementations may differ in their resource allocation
requirements. For example, one server may share resources among all
classes of recallable objects whereas another may use separate
resource pools for layouts and for delegations, or further separate
resources by types of delegations.
When a given resource pool is over-utilized, the server can issue a
CB_RECALL_ANY to clients holding recallable objects of the types
involved, allowing it to keep a certain number of such objects and
return any excess. A mask specifies which types of objects are to be
limited. The client chooses, based on its own knowledge of current
usefulness, which of the objects in that class should be returned.
For NFSv4.1, sixteen bits are defined. For some of these, ranges are
defined and it is up to the definition of the storage protocol to
specify how these are to be used. There are ranges for blocks-based
storage protocols, for object-based storage protocols and a reserved
range for other experimental storage protocols. The RFC defining
such a storage protocol needs to specify how particular bits within
its range are to be used. For example, it may specify a mapping
between attributes of the layout (read vs. write, size of area) and
the bit to be used or it may define a field in the layout where the
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associated bit position is made available by the server to the
client.
When an undefined bit is set in the type mask, NFS4ERR_INVAL should
be returned. However even if a client does not support an object of
the specified type, if the bit is defined, NFS4ERR_INVAL should not
be returned. Future minor versions of NFSv4 may expand the set of
valid type mask bits.
CB_RECALL_ANY specifies a count of objects that the client may keep
as opposed to a count that the client must return. This is to avoid
potential race between a CB_RECALL_ANY that had a count of objects to
free with a set of client-originated operations to return layouts or
delegations. As a result of the race, the client and server would
have differing ideas as to how many objects to return. Hence the
client could mistakenly free too many.
If resource demands prompt it, the server may send another
CB_RECALL_ANY with a lower count, even it has not yet received an
acknowledgement from the client for a previous CB_RECALL_ANY with the
same type mask. Although the possibility exists that these will be
received by the client in a order different from the order in which
they were sent, any such permutation of the callback stream is
harmless. It is the job of the client to bring down the size of the
recallable object set in line with each CB_RECALL_ANY received and
until that obligation is met it cannot be canceled or modified by any
subsequent CB_RECALL_ANY for the same type mask. Thus if the server
sends two CB_RECALL_ANY's, the effect will be the same as if the
lower count was sent, whatever the order of recall receipt. Note
that this means that a server may not cancel the effect of a
CB_RECALL_ANY by sending another recall with a higher count. When a
CB_RECALL_ANY is received and the count is already within the limit
set or is above a limit that the client is working to get down to,
that callback has no effect.
The client can choose to return any type of object specified by the
mask. If a server wishes to limit use of objects of a specific type,
it should only specify that type in the mask sent. The client may
not return requested objects and it is up to the server to handle
this situation, typically by doing specific recalls to properly limit
resource usage. The server should give the client enough time to
return objects before proceeding to specific recalls. This time
should not be less than the lease period.
Servers are generally free not to give out recallable objects when
insufficient resources are available. Note that the effect of such a
policy is implicitly to give precedence to existing objects relative
to requested ones, with the result that resources might not be
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optimally used. To prevent this, servers are well advised to make
the point at which they start issuing CB_RECALL_ANY callbacks
somewhat below that at which they cease to give out new delegations
and layouts. This allows the client to purge its less-used objects
whenever appropriate and so continue to have its subsequent requests
given new resources freed up by object returns.
24.7.5. IMPLEMENTATION
24.7.6. ERRORS
NFS4ERR_RESOURCE NFS4ERR_INVAL
24.8. CB_SIZECHANGED
24.8.1. SYNOPSIS
fh, size -> -
24.8.2. ARGUMENT
struct CB_SIZECHANGEDargs {
nfs_fh4 fh;
length4 size;
};
24.8.3. RESULT
struct CB_SIZECHANGEDres {
nfsstat4 status;
};
24.8.4. DESCRIPTION
The CB_SIZECHANGED operation is used to notify the client that the
size pertaining to the filehandle associated with "fh", has changed.
The new size is specified. Upon reception of this notification
callback, the client should update its internal size for the file.
If the layout being held for the file is of the NFSv4 file layout
type, then the size field within that layout should be updated (see
Section 16.5). For other layout types see Section 14.4.2 for more
details.
If the handle specified is not one for which the client holds a
layout, an NFS4ERR_BADHANDLE error is returned.
24.8.5. IMPLEMENTATION
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24.8.6. ERRORS
NFS4ERR_BADHANDLE
24.9. CB_LAYOUTRECALL
24.9.1. SYNOPSIS
layout_type, iomode, layoutchanged, layoutrecall -> -
24.9.2. ARGUMENT
enum layoutrecall_type4 {
RECALL_FILE = 1,
RECALL_FSID = 2
};
struct layoutrecall_file4 {
nfs_fh4 fh;
offset4 offset;
length4 length;
};
union layoutrecall4 switch(layoutrecall_type4 recalltype) {
case RECALL_FILE:
layoutrecall_file4 layout;
case RECALL_FSID:
fsid4 fsid;
};
struct CB_LAYOUTRECALLargs {
pnfs_layouttype4 layout_type;
pnfs_layoutiomode4 iomode;
bool layoutchanged;
layoutrecall4 layoutrecall;
};
24.9.3. RESULT
struct CB_LAYOUTRECALLres {
nfsstat4 status;
};
24.9.4. DESCRIPTION
The CB_LAYOUTRECALL operation is used to begin the process of
recalling a layout, a portion thereof, or all layouts pertaining to a
particular file system (FSID). If RECALL_FILE is specified, the
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offset and length fields specify the portion of the layout to be
returned. The iomode specifies the set of layouts to be returned.
An iomode of ANY specifies that all matching layouts, regardless of
iomode, must be returned; otherwise, only layouts that exactly match
the iomode must be returned.
If the "layoutchanged" field is TRUE, then the client SHOULD not
flush its dirty data to the devices specified by the layout being
recalled. Instead, it is preferable for the client to flush the
dirty data through the metadata server. Alternatively, the client
may attempt to obtain a new layout. Note: in order to obtain a new
layout the client must first return the old layout. Since obtaining
a new layout is not guaranteed to succeed, the client must be ready
to flush its dirty data through the metadata server.
If RECALL_FSID is specified, the fsid specifies the file system for
which any outstanding layouts must be returned. Layouts are returned
through the LAYOUTRETURN operation.
If the client does not hold any layout segment either matching or
overlapping with the requested layout, it returns
NFS4ERR_NOMATCHING_LAYOUT. If a length of all 1s is specified then
the layout corresponding to the byte range from "offset" to the end-
of-file MUST be returned.
24.9.5. IMPLEMENTATION
The client should reply to the callback immediately. Replying does
not complete the recall except when an error is returned. The recall
is not complete until the layout(s) are returned using a
LAYOUTRETURN.
The client should complete any in-flight I/O operations using the
recalled layout(s) before returning it/them via LAYOUTRETURN. If the
client has buffered dirty data there are a number of options for
flushing that data. If "layoutchanged" is false, the client may
choose to write dirty data directly to storage before calling
LAYOUTRETURN. However, if "layoutchanged" is true, the client may
either choose to write it later using normal NFSv4 WRITE operations
to the metadata server or it may attempt to obtain a new layout,
after first returning the recalled layout, using the new layout to
flush the dirty data. Regardless of whether the client is holding a
layout, it may always write data through the metadata server.
If dirty data is flushed while the layout is held, the client must
still issue LAYOUTCOMMIT operations at the appropriate time,
especially before issuing the LAYOUTRETURN. If a large amount of
dirty data is outstanding, the client may issue LAYOUTRETURNs for
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portions of the layout being recalled; this allows the server to
monitor the client's progress and adherence to the callback.
However, the last LAYOUTRETURN in a sequence of returns, SHOULD
specify the full range being recalled (see Section 14.5.2 for
details).
24.9.6. ERRORS
NFS4ERR_NOMATCHING_LAYOUT
25. References
25.1. Normative References
[1] Bradner, S., "Key words for use in RFCs to Indicate Requirement
Levels", March 1997.
[2] Srinivasan, R., "XDR: External Data Representation Standard",
RFC 1832, August 1995.
[3] Srinivasan, R., "RPC: Remote Procedure Call Protocol
Specification Version 2", RFC 1831, August 1995.
[4] Linn, J., "Generic Security Service Application Program
Interface Version 2, Update 1", RFC 2743, January 2000.
[5] Hinden, R. and S. Deering, "IP Version 6 Addressing
Architecture", RFC 1884, December 1995.
[6] Shepler, S., Callaghan, B., Robinson, D., Thurlow, R., Beame,
C., Eisler, M., and D. Noveck, "Network File System (NFS)
version 4 Protocol", RFC 3530, April 2003.
[7] International Organization for Standardization, "Information
Technology - Universal Multiple-octet coded Character Set (UCS)
- Part 1: Architecture and Basic Multilingual Plane",
ISO Standard 10646-1, May 1993.
[8] Alvestrand, H., "IETF Policy on Character Sets and Languages",
BCP 18, RFC 2277, January 1998.
25.2. Informative References
[9] Srinivasan, R., "Binding Protocols for ONC RPC Version 2",
RFC 1833, August 1995.
[10] Zelenka, J., Welch, B., and B. Halevy, "Object-based pNFS
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Operations", July 2005, <ftp://www.ietf.org/internet-drafts/
draft-zelenka-pnfs-obj-01.txt>.
[11] Black, D., "pNFS Block/Volume Layout", July 2005, <ftp://
www.ietf.org/internet-drafts/draft-black-pnfs-block-01.txt>.
[12] Satran, J., Meth, K., Sapuntzakis, C., Chadalapaka, M., and E.
Zeidner, "Internet Small Computer Systems Interface (iSCSI)",
RFC 3720, April 2004.
[13] Snively, R., "Fibre Channel Protocol for SCSI, 2nd Version
(FCP-2)", ANSI/INCITS 350-2003, Oct 2003.
[14] Weber, R., "Object-Based Storage Device Commands (OSD)", ANSI/
INCITS 400-2004, July 2004,
<http://www.t10.org/ftp/t10/drafts/osd/osd-r10.pdf>.
Appendix A. Acknowledgments
The initial drafts for the SECINFO extensions were edited by Mike
Eisler with contributions from Tom Talpey, Saadia Khan, and Jon
Bauman.
The initial drafts for the SESSIONS extensions were edited by Tom
Talpey, Spencer Shepler, Jon Bauman with contributions from Charles
Antonelli, Brent Callaghan, Mike Eisler, John Howard, Chet Juszczak,
Trond Myklebust, Dave Noveck, John Scott, Mike stolarchuk and Mark
Wittle.
The initial drafts for the Directory Delegations support were
contributed by Saadia Khan with input from Dave Noveck, Mike Eisler,
Carl Burnett, Ted Anderson and Tom Talpey.
The initial drafts for the parellel NFS support were edited by Brent
Welch and Garth Goodson. Additional authors for those documents were
Benny Halevy, David Black, and Andy Adamson. Additional input came
from the informal group which contributed to the construction of the
initial pNFS drafts; specific acknowledgement goes to Gary Grider,
Peter Corbett, Dave Noveck, and Peter Honeyman. The pNFS work was
inspired by the NASD and OSD work done by Garth Gibson. Gary Grider
of the national labs (LANL) has also been a champion of high-
performance parallel I/O.
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Author's Address
Spencer Shepler
Sun Microsystems, Inc.
7808 Moonflower Drive
Austin, TX 78750
USA
Phone: +1-512-349-9376
Email: spencer.shepler@sun.com
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Full Copyright Statement
Copyright (C) The Internet Society (2006).
This document is subject to the rights, licenses and restrictions
contained in BCP 78, and except as set forth therein, the authors
retain all their rights.
This document and the information contained herein are provided on an
"AS IS" basis and THE CONTRIBUTOR, THE ORGANIZATION HE/SHE REPRESENTS
OR IS SPONSORED BY (IF ANY), THE INTERNET SOCIETY AND THE INTERNET
ENGINEERING TASK FORCE DISCLAIM ALL WARRANTIES, EXPRESS OR IMPLIED,
INCLUDING BUT NOT LIMITED TO ANY WARRANTY THAT THE USE OF THE
INFORMATION HEREIN WILL NOT INFRINGE ANY RIGHTS OR ANY IMPLIED
WARRANTIES OF MERCHANTABILITY OR FITNESS FOR A PARTICULAR PURPOSE.
Intellectual Property
The IETF takes no position regarding the validity or scope of any
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Acknowledgment
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Shepler Expires September 7, 2006 [Page 360]
