NFSv4 S. Shepler
Internet-Draft Editor
Expires: June 15, 2006 December 12, 2005
NFSv4 Minor Version 1
draft-ietf-nfsv4-minorversion1-01.txt
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Copyright Notice
Copyright (C) The Internet Society (2005).
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
document are to be interpreted as described in RFC 2119 [1].
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Table of Contents
1. Security Negotiation . . . . . . . . . . . . . . . . . . . . 6
2. Clarification of Security Negotiation in NFSv4.1 . . . . . . 6
2.1 PUTFH + LOOKUP . . . . . . . . . . . . . . . . . . . . . . 6
2.2 PUTFH + LOOKUPP . . . . . . . . . . . . . . . . . . . . . 7
2.3 PUTFH + SECINFO . . . . . . . . . . . . . . . . . . . . . 7
2.4 PUTFH + Anything Else . . . . . . . . . . . . . . . . . . 7
3. NFSv4.1 Sessions . . . . . . . . . . . . . . . . . . . . . . 8
3.1 Sessions Background . . . . . . . . . . . . . . . . . . . 8
3.1.1 Introduction to Sessions . . . . . . . . . . . . . . . 8
3.1.2 Motivation . . . . . . . . . . . . . . . . . . . . . . 9
3.1.3 Problem Statement . . . . . . . . . . . . . . . . . . 10
3.1.4 NFSv4 Session Extension Characteristics . . . . . . . 11
3.2 Transport Issues . . . . . . . . . . . . . . . . . . . . . 12
3.2.1 Session Model . . . . . . . . . . . . . . . . . . . . 12
3.2.2 Connection State . . . . . . . . . . . . . . . . . . . 13
3.2.3 NFSv4 Channels, Sessions and Connections . . . . . . . 14
3.2.4 Reconnection, Trunking and Failover . . . . . . . . . 16
3.2.5 Server Duplicate Request Cache . . . . . . . . . . . . 17
3.3 Session Initialization and Transfer Models . . . . . . . . 18
3.3.1 Session Negotiation . . . . . . . . . . . . . . . . . 18
3.3.2 RDMA Requirements . . . . . . . . . . . . . . . . . . 19
3.3.3 RDMA Connection Resources . . . . . . . . . . . . . . 20
3.3.4 TCP and RDMA Inline Transfer Model . . . . . . . . . . 21
3.3.5 RDMA Direct Transfer Model . . . . . . . . . . . . . . 23
3.4 Connection Models . . . . . . . . . . . . . . . . . . . . 26
3.4.1 TCP Connection Model . . . . . . . . . . . . . . . . . 27
3.4.2 Negotiated RDMA Connection Model . . . . . . . . . . . 28
3.4.3 Automatic RDMA Connection Model . . . . . . . . . . . 29
3.5 Buffer Management, Transfer, Flow Control . . . . . . . . 29
3.6 Retry and Replay . . . . . . . . . . . . . . . . . . . . . 32
3.7 The Back Channel . . . . . . . . . . . . . . . . . . . . . 33
3.8 COMPOUND Sizing Issues . . . . . . . . . . . . . . . . . . 34
3.9 Data Alignment . . . . . . . . . . . . . . . . . . . . . . 34
3.10 NFSv4 Integration . . . . . . . . . . . . . . . . . . . 36
3.10.1 Minor Versioning . . . . . . . . . . . . . . . . . . 36
3.10.2 Slot Identifiers and Server Duplicate Request
Cache . . . . . . . . . . . . . . . . . . . . . . . 36
3.10.3 COMPOUND and CB_COMPOUND . . . . . . . . . . . . . . 40
3.10.4 eXternal Data Representation Efficiency . . . . . . 41
3.10.5 Effect of Sessions on Existing Operations . . . . . 41
3.10.6 Authentication Efficiencies . . . . . . . . . . . . 42
3.11 Sessions Security Considerations . . . . . . . . . . . . 43
3.11.1 Authentication . . . . . . . . . . . . . . . . . . . 44
4. Directory Delegations . . . . . . . . . . . . . . . . . . . 45
4.1 Introduction to Directory Delegations . . . . . . . . . . 45
4.2 Directory Delegation Design (in brief) . . . . . . . . . . 47
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4.3 Recommended Attributes in support of Directory
Delegations . . . . . . . . . . . . . . . . . . . . . . . 48
4.4 Delegation Recall . . . . . . . . . . . . . . . . . . . . 48
4.5 Delegation Recovery . . . . . . . . . . . . . . . . . . . 49
5. Introduction . . . . . . . . . . . . . . . . . . . . . . . . 49
6. General Definitions . . . . . . . . . . . . . . . . . . . . 51
6.1 Metadata Server . . . . . . . . . . . . . . . . . . . . . 52
6.2 Client . . . . . . . . . . . . . . . . . . . . . . . . . . 52
6.3 Storage Device . . . . . . . . . . . . . . . . . . . . . . 52
6.4 Storage Protocol . . . . . . . . . . . . . . . . . . . . . 52
6.5 Control Protocol . . . . . . . . . . . . . . . . . . . . . 53
6.6 Metadata . . . . . . . . . . . . . . . . . . . . . . . . . 53
6.7 Layout . . . . . . . . . . . . . . . . . . . . . . . . . . 53
7. pNFS protocol semantics . . . . . . . . . . . . . . . . . . 53
7.1 Definitions . . . . . . . . . . . . . . . . . . . . . . . 54
7.1.1 Layout Types . . . . . . . . . . . . . . . . . . . . . 54
7.1.2 Layout Iomode . . . . . . . . . . . . . . . . . . . . 54
7.1.3 Layout Segments . . . . . . . . . . . . . . . . . . . 55
7.1.4 Device IDs . . . . . . . . . . . . . . . . . . . . . . 56
7.1.5 Aggregation Schemes . . . . . . . . . . . . . . . . . 56
7.2 Guarantees Provided by Layouts . . . . . . . . . . . . . . 56
7.3 Getting a Layout . . . . . . . . . . . . . . . . . . . . . 58
7.4 Committing a Layout . . . . . . . . . . . . . . . . . . . 58
7.4.1 LAYOUTCOMMIT and mtime/atime/change . . . . . . . . . 59
7.4.2 LAYOUTCOMMIT and size . . . . . . . . . . . . . . . . 60
7.4.3 LAYOUTCOMMIT and layoutupdate . . . . . . . . . . . . 61
7.5 Recalling a Layout . . . . . . . . . . . . . . . . . . . . 61
7.5.1 Basic Operation . . . . . . . . . . . . . . . . . . . 61
7.5.2 Recall Callback Robustness . . . . . . . . . . . . . . 62
7.5.3 Recall/Return Sequencing . . . . . . . . . . . . . . . 63
7.6 Metadata Server Write Propagation . . . . . . . . . . . . 65
7.7 Crash Recovery . . . . . . . . . . . . . . . . . . . . . . 66
7.7.1 Leases . . . . . . . . . . . . . . . . . . . . . . . . 66
7.7.2 Client Recovery . . . . . . . . . . . . . . . . . . . 67
7.7.3 Metadata Server Recovery . . . . . . . . . . . . . . . 68
7.7.4 Storage Device Recovery . . . . . . . . . . . . . . . 70
8. Security Considerations . . . . . . . . . . . . . . . . . . 71
8.1 File Layout Security . . . . . . . . . . . . . . . . . . . 72
8.2 Object Layout Security . . . . . . . . . . . . . . . . . . 72
8.3 Block/Volume Layout Security . . . . . . . . . . . . . . . 73
9. The NFSv4 File Layout Type . . . . . . . . . . . . . . . . . 74
9.1 File Striping and Data Access . . . . . . . . . . . . . . 74
9.1.1 Sparse and Dense Storage Device Data Layouts . . . . . 75
9.1.2 Metadata and Storage Device Roles . . . . . . . . . . 77
9.1.3 Device Multipathing . . . . . . . . . . . . . . . . . 78
9.1.4 Operations Issued to Storage Devices . . . . . . . . . 79
9.2 Global Stateid Requirements . . . . . . . . . . . . . . . 79
9.3 The Layout Iomode . . . . . . . . . . . . . . . . . . . . 80
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9.4 Storage Device State Propagation . . . . . . . . . . . . . 80
9.4.1 Lock State Propagation . . . . . . . . . . . . . . . . 80
9.4.2 Open-mode Validation . . . . . . . . . . . . . . . . . 81
9.4.3 File Attributes . . . . . . . . . . . . . . . . . . . 81
9.5 Storage Device Component File Size . . . . . . . . . . . . 82
9.6 Crash Recovery Considerations . . . . . . . . . . . . . . 83
9.7 Security Considerations . . . . . . . . . . . . . . . . . 83
9.8 Alternate Approaches . . . . . . . . . . . . . . . . . . . 84
10. pNFS Typed Data Structures . . . . . . . . . . . . . . . . . 85
10.1 pnfs_layouttype4 . . . . . . . . . . . . . . . . . . . . 85
10.2 pnfs_deviceid4 . . . . . . . . . . . . . . . . . . . . . 85
10.3 pnfs_deviceaddr4 . . . . . . . . . . . . . . . . . . . . 86
10.4 pnfs_devlist_item4 . . . . . . . . . . . . . . . . . . . 86
10.5 pnfs_layout4 . . . . . . . . . . . . . . . . . . . . . . 87
10.6 pnfs_layoutupdate4 . . . . . . . . . . . . . . . . . . . 87
10.7 pnfs_layouthint4 . . . . . . . . . . . . . . . . . . . . 88
10.8 pnfs_layoutiomode4 . . . . . . . . . . . . . . . . . . . 88
11. pNFS File Attributes . . . . . . . . . . . . . . . . . . . . 88
11.1 pnfs_layouttype4<> FS_LAYOUT_TYPES . . . . . . . . . . . 88
11.2 pnfs_layouttype4<> FILE_LAYOUT_TYPES . . . . . . . . . . 88
11.3 pnfs_layouthint4 FILE_LAYOUT_HINT . . . . . . . . . . . 89
11.4 uint32_t FS_LAYOUT_PREFERRED_BLOCKSIZE . . . . . . . . . 89
11.5 uint32_t FS_LAYOUT_PREFERRED_ALIGNMENT . . . . . . . . . 89
12. pNFS Error Definitions . . . . . . . . . . . . . . . . . . . 89
13. Layouts and Aggregation . . . . . . . . . . . . . . . . . . 90
13.1 Simple Map . . . . . . . . . . . . . . . . . . . . . . . 90
13.2 Block Extent Map . . . . . . . . . . . . . . . . . . . . 90
13.3 Striped Map (RAID 0) . . . . . . . . . . . . . . . . . . 90
13.4 Replicated Map . . . . . . . . . . . . . . . . . . . . . 91
13.5 Concatenated Map . . . . . . . . . . . . . . . . . . . . 91
13.6 Nested Map . . . . . . . . . . . . . . . . . . . . . . . 91
14. NFSv4.1 Operations . . . . . . . . . . . . . . . . . . . . . 91
14.1 LOOKUPP - Lookup Parent Directory . . . . . . . . . . . 91
14.2 SECINFO -- Obtain Available Security . . . . . . . . . . 93
14.3 SECINFO_NO_NAME - Get Security on Unnamed Object . . . . 96
14.4 CREATECLIENTID - Instantiate Clientid . . . . . . . . . 98
14.5 CREATESESSION - Create New Session and Confirm
Clientid . . . . . . . . . . . . . . . . . . . . . . . . 104
14.6 BIND_BACKCHANNEL - Create a callback channel binding . . 109
14.7 DESTROYSESSION - Destroy existing session . . . . . . . 112
14.8 SEQUENCE - Supply per-procedure sequencing and control . 113
14.9 CB_RECALLCREDIT - change flow control limits . . . . . . 114
14.10 CB_SEQUENCE - Supply callback channel sequencing and
control . . . . . . . . . . . . . . . . . . . . . . . . 115
14.11 GET_DIR_DELEGATION - Get a directory delegation . . . . 117
14.12 CB_NOTIFY - Notify directory changes . . . . . . . . . . 120
14.13 CB_RECALL_ANY - Keep any N delegations . . . . . . . . . 124
14.14 LAYOUTGET - Get Layout Information . . . . . . . . . . . 126
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14.15 LAYOUTCOMMIT - Commit writes made using a layout . . . . 128
14.16 LAYOUTRETURN - Release Layout Information . . . . . . . 131
14.17 GETDEVICEINFO - Get Device Information . . . . . . . . . 133
14.18 GETDEVICELIST - Get List of Devices . . . . . . . . . . 134
14.19 CB_LAYOUTRECALL . . . . . . . . . . . . . . . . . . . . 136
14.20 CB_SIZECHANGED . . . . . . . . . . . . . . . . . . . . . 138
15. References . . . . . . . . . . . . . . . . . . . . . . . . . 139
15.1 Normative References . . . . . . . . . . . . . . . . . . 139
15.2 Informative References . . . . . . . . . . . . . . . . . 139
Author's Address . . . . . . . . . . . . . . . . . . . . . . 139
A. Acknowledgments . . . . . . . . . . . . . . . . . . . . . . 139
Intellectual Property and Copyright Statements . . . . . . . 141
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1. 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.
2. 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.
2.1 PUTFH + LOOKUP
The server implementation may decide whether to impose any
restrictions on export security administration. There are at least
three approaches (Sc is the flavor set of the child export, Sp that
of the parent),
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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.
2.2 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).
2.3 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.
2.4 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
to recover from NFS4ERR_WRONGSEC.
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3. NFSv4.1 Sessions
3.1 Sessions Background
3.1.1 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:
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.
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.
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
referenced for details of RPC RDMA support.
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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 |
+-------------------------------+-----------------------+
3.1.2 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.
NFSv4 advances the state of high-performance local sharing, by virtue
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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.
3.1.3 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
requirement.
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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.
3.1.4 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
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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).
3.2 Transport Issues
The Transport Issues section of the document explores the details of
utilizing the various supported transports.
3.2.1 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
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.
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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.
3.2.2 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
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
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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.
3.2.3 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
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
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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:
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
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|
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.
3.2.4 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.
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.
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3.2.5 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.
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.
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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.
3.3 Session Initialization and Transfer Models
Session initialization issues, and data transfer models relevant to
both TCP and RDMA are discussed in this section.
3.3.1 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
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
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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:
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.
3.3.2 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
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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.
3.3.3 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
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.
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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.
3.3.4 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.
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
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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
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
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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
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.
3.3.5 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.
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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
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
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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
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 : :
3.4 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.
3.4.1 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.
3.4.2 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.
3.4.3 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> :
: ------------------------------> :
: <------------------------------ :
: : :
3.5 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.
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RDMA operations are performed solely by the server in this proposal,
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.
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Because the client requests an arbitrary credit count in each
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
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granted, it must send zero-length Send RDMA operations, or NULL NFSv4
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.
3.6 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
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outside that of a single connection. Therefore, handles used by the
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.
3.7 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.
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This can be convenient for the NFSv4.1 client; if the client expects
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.
3.8 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.
3.9 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
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RDMA receives preserve alignment of data, even when they are placed
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.
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Padding is negotiated by the session creation operation, and
subsequently used by the RPC RDMA layer, as described in [RPCRDMA].
3.10 NFSv4 Integration
The following section discusses the integration of the proposed RDMA
extensions with NFSv4.0.
3.10.1 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.
3.10.2 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
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to be unique, is not especially convenient for tracking requests.
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:
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.
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A retransmitted request, in which the sequenceid is equal to that
last seen in the slot. Note that this request may be either
complete, or in progress. The server performs replay processing
in these cases.
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
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the port number in the server DRC implementation, and the session
replaces the IP address. This approach is considerably more portable
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
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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
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.
3.10.3 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:
+-----+--------------+-----------+------------+-----------+----
| 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 | ...
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//-----------------------+----
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.
3.10.4 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.
3.10.5 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
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
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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.
3.10.6 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
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.
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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.
3.11 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
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
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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.
3.11.1 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
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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 clientid and sessionid must be 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
NFSv4.0 with clients who may perform SETCLIENTID, then never perform
SETCLIENTID_CONFIRM.) Careful authentication coupled with resource
checks is highly recommended.
4. Directory Delegations
4.1 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
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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
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
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benefits for the client, and possibly the server, under some common
operating conditions such as slowly-changing and/or very large
directories.
4.2 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
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.
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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.
4.3 Recommended Attributes in support of Directory Delegations
supp_dir_attr_notice - notification delays on directory attributes
supp_child_attr_notice - 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 supp_dir_attr_notice value covers all attribute changes to the
directory and the supp_child_attr_notice 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 will request a notification delay
that is within the server's supported range. If the client violates
what supp_attr_file_notice or supp_attr_dir_notice values are, the
server should not commit to sending notifications for that change
event.
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. A value of N means that the server will make a best effort
guarentee that attribute notification are not delayed by more than
that. nfstime4 values that compute to negative values are illegal.
4.4 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
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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.
4.5 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.
5. Introduction
The NFSv4 protocol [2] 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.
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:
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+-----------+
|+-----------+ +-----------+
||+-----------+ | |
||| | NFSv4 + pNFS | |
+|| Clients |<------------------------------>| Server |
+| | | |
+-----------+ | |
||| +-----------+
||| |
||| |
||| Storage +-----------+ |
||| Protocol |+-----------+ |
||+----------------||+-----------+ Control|
|+-----------------||| | Protocol|
+------------------+|| Storage |------------+
+| Devices |
+-----------+
Figure 9
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
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:
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o Block/volume protocols such as iSCSI ([3]), and FCP ([4]). 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 [5].
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
file access in such situations. See Section 8 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.
6. General Definitions
This protocol extension partitions the NFSv4 file system protocol
into two parts, the control path and the data path. The control path
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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.
6.1 Metadata Server
A pNFS "server" or "metadata server" is a server as defined by
RFC3530 [2], 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.
6.2 Client
A pNFS "client" is a client as defined by RFC3530 [2], 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.
6.3 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.
6.4 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 9.
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6.5 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.
6.6 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.
6.7 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.
7. 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.
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7.1 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.
7.1.1 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 10.1 for its RPC definition). The layout
type allows for variants to handle different storage protocols (e.g.,
block/volume [6], object [7], and file [Section 9] 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 9). Adjunct specifications are being
drafted that precisely define other layout formats (e.g., block/
volume [6], and object [7] layouts) to allow interoperability among
clients and metadata servers.
7.1.2 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.
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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.
7.1.3 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
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.
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7.1.4 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.
7.1.5 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.
7.2 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
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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.
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.
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7.3 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
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.
7.4 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
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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 7.6 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.
7.4.1 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
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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.
7.4.2 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.
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
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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.
7.4.3 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.
7.5 Recalling a Layout
7.5.1 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 14.19). 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 7.5.3 discusses sequencing
issues surrounding the getting, returning, and recalling of layouts.
The iomode is also specified when recalling a layout or layout
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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
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.
7.5.2 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:
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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
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.
7.5.3 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
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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.
7.5.3.1 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 the client cannot distinguish when the callback
arrives:
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. 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
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errors for them.
Now the client can wait for the LAYOUTGET response, as it will be
received in both cases.
7.5.3.2 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:
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.
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).]
7.6 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
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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
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.
7.7 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.
7.7.1 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.
7.7.1.1 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.
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7.7.1.2 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
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.
7.7.2 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.
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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
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.
7.7.3 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 are not reclaimable in the same sense as data delegations;
there is no reclaim bit, thus no guarantee of continuity between the
previous and new layout. This is not necessarily required since a
layout is not required to perform I/O; I/O can always be performed
through the metadata server.
[NOTE: there is no reclaim bit for getting a layout. Thus, in the
case of reclaiming an old layout obtained through LAYOUTGET, there is
no guarantee of continuity. If a reclaim bit existed a block/volume
layout type might be happier knowing it got the layout back with the
assurance of continuity. However, this would require the metadata
server trusting the client in telling it the exact layout it had
(i.e., the full block-list); however, divergence is avoided by having
the server tell the client what is contained within the layout.]
If the client has dirty data that it needs to write out, or an
outstanding LAYOUTCOMMIT, the client should try to obtain a new
layout segment covering the byte range covered by the previous layout
segment. However, the client might not not get the same layout
segment it had. 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; in
the interest of simplicity, the client might decide to always write
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them again. Alternatively, the client might be unable to obtain a
new layout and thus, must write the data using normal NFSv4 through
the metadata server.
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
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 7.7.1.
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 9.6); 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.
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7.7.4 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
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.
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8. 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 9) 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
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.
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[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.]
8.1 File Layout Security
A NFSv4 file layout type is defined in Section 9; see Section 9.7 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.
8.2 Object Layout Security
The object storage protocol MUST implement the security aspects
described in version 1 of the T10 OSD protocol definition [5]. The
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
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{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
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.
8.3 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
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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.
9. The NFSv4 File Layout Type
This section describes the semantics and format of NFSv4 file-based
layouts.
9.1 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
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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 9.1.3. 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).
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 10.7. 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.
9.1.1 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:
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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
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
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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 9.5 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
the end of one of these component files.
9.1.2 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.
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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
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.
9.1.3 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
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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.]
9.1.4 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 7.2, 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.
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 7.4.2). Section 9.5, describes the
mechanism by which the client is to handle storage device file's that
do not reflect the metadata server's size.
9.2 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 9.4.
Depending on when stateids are propagated, the existence of a valid
stateid on the storage device may act as proof of a valid layout.
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[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.]
9.3 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.
9.4 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.
9.4.1 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
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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.
9.4.2 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.
9.4.3 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 7.4
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.
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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 [2] 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
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.]
9.5 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;
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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.
9.6 Crash Recovery Considerations
As described in Section 7.7, 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.
9.7 Security Considerations
The NFSv4 file layout type MUST adhere to the security considerations
outlined in Section 8. More specifically, storage devices must make
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all of the required access checks on each READ or WRITE I/O as
determined by the NFSv4 protocol [2]. This impacts the control
protocol and the propagation of state from the metadata server to the
storage devices; see Section 9.4 for more details.
9.8 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
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
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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 9.4).
10. pNFS Typed Data Structures
10.1 pnfs_layouttype4
enum pnfs_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.
[OPEN ISSUE: 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 [7], is to be used.
Similarly, the LAYOUT_BLOCK_VOLUME enumeration that the block/volume
layout, as defined in [6], is to be used.
10.2 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 7.1.4 for more
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details.
10.3 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
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.]
10.4 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.
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10.5 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 9.1 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.
[OPEN ISSUE: 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
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.]
10.6 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.
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10.7 pnfs_layouthint4
struct pnfs_layouthint4 {
pnfs_layouttype4 type;
opaque layouthint_data<>;
};
The pnfs_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 9.1.
10.8 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
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.
11. pNFS File Attributes
11.1 pnfs_layouttype4<> FS_LAYOUT_TYPES
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.
11.2 pnfs_layouttype4<> FILE_LAYOUT_TYPES
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.,
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specific device information) in order to perform I/O.
11.3 pnfs_layouthint4 FILE_LAYOUT_HINT
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.
[OPEN ISSUE: it has been suggested that the HINT is a well defined
type other than pnfs_layoutdata4, similar to pnfs_layoutupdate4.]
11.4 uint32_t FS_LAYOUT_PREFERRED_BLOCKSIZE
This attribute is a file system wide attribute and indicates the
preferred block size for direct storage device access.
11.5 uint32_t FS_LAYOUT_PREFERRED_ALIGNMENT
This attribute is a file system wide attribute and indicates the
preferred alignment for direct storage device access.
12. pNFS Error Definitions
NFS4ERR_BADLAYOUT Layout specified is invalid.
NFS4ERR_BADIOMODE Layout iomode is invalid.
NFS4ERR_LAYOUTUNAVAILABLE Layouts are not available for the file or
its containing file system.
NFS4ERR_LAYOUTTRYLATER Layouts are temporarily unavailable for the
file, client should retry later.
NFS4ERR_NOMATCHING_LAYOUT Client has no matching layout (segment) to
return.
NFS4ERR_RECALLCONFLICT Layout is unavailable due to a conflicting
LAYOUTRECALL that is in progress.
NFS4ERR_UNKNOWN_LAYOUTTYPE Layout type is unknown.
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13. 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.
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.
13.1 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).
13.2 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.
13.3 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
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map is an ordered list of (deviceID, blockID, blocksize), where the
deviceID is rotated among a set of devices to achieve striping.
13.4 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
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.
13.5 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.
13.6 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.
14. NFSv4.1 Operations
14.1 LOOKUPP - Lookup Parent Directory
If the NFSv4 minor version is 1, then following replaces section
14.2.14 of the NFSv4.0 specification. The LOOKUPP operation's "over
the wire" format is not altered, but the semantics are slightly
modified to account for the addition of SECINFO_NO_NAME.
SYNOPSIS
(cfh) -> (cfh)
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ARGUMENT
/* CURRENT_FH: object */
void;
RESULT
struct LOOKUPP4res {
/* CURRENT_FH: directory */
nfsstat4 status;
};
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.
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
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14.2 SECINFO -- Obtain Available Security
If the NFSv4 minor version is 1, then following replaces section
14.2.31 of the NFSv4.0 specification. The SECINFO operation's "over
the wire" format is not altered, but the semantics are slightly
modified to account for the addition of SECINFO_NO_NAME.
SYNOPSIS
(cfh), name -> { secinfo }
ARGUMENT
struct SECINFO4args {
/* CURRENT_FH: directory */
component4 name;
};
RESULT
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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;
};
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 registered with IANA.
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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.
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:
* For LOOKUP and OPEN, the client will use SECINFO with the same
current filehandle and name as provided in the original LOOKUP
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or OPEN to enumerate the available security triples.
* 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.
* ========================================================= 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.
========================================================
* 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.
ERRORS
14.3 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.
SYNOPSIS
(cfh), secinfo_style -> { secinfo }
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ARGUMENT
enum secinfo_style_4 {
current_fh = 0,
parent = 1
};
typedef secinfo_style_4 SECINFO_NO_NAME4args;
RESULT
typedef SECINFO4res SECINFO_NO_NAME4res;
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.
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The server implementor should pay particular attention to the
section "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.
IMPLEMENTATION
See the previous dicussion on SECINFO.
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
14.4 CREATECLIENTID - Instantiate Clientid
Create a clientid
SYNOPSIS
client -> clientid
ARGUMENT
struct CREATECLIENTID4args {
nfs_client_id4 clientdesc;
};
RESULT
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struct CREATECLIENTID4resok {
clientid4 clientid;
verifier4 clientid_confirm;
};
union SETCLIENTID4res switch (nfsstat4 status) {
case NFS4_OK:
CREATECLIENTID4resok resok4;
case NFS4ERR_CLID_INUSE:
void;
default:
void;
};
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.
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
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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.
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.
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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 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.
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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.
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}
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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 }
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
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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 }
ERRORS
NFS4ERR_BADXDR NFS4ERR_CLID_INUSE NFS4ERR_INVAL NFS4ERR_RESOURCE
NFS4ERR_SERVERFAULT
14.5 CREATESESSION - Create New Session and Confirm Clientid
Start up session and confirm clientid.
SYNOPSIS
clientid, session_args -> sessionid, session_args
ARGUMENT
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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;
};
};
RESULT
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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;
};
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.
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 identified by the value of the clientid and sessionid
fields of CREATESESSION4args and the method for dealing with them
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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.
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{ *, *, old_principal_arg, clientid_arg, * }
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.
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maxresponsesize:
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.
ERRORS
NFS4ERR_BADXDR NFS4ERR_CLID_INUSE NFS4ERR_RESOURCE
NFS4ERR_SERVERFAULT NFS4ERR_STALE_CLIENTID
14.6 BIND_BACKCHANNEL - Create a callback channel binding
Establish a callback channel on the connection.
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SYNOPSIS
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;
};
};
RESULT
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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;
};
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
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).
IMPLEMENTATION
No discussion at this time.
ERRORS
TBD
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14.7 DESTROYSESSION - Destroy existing session
Destroy existing session.
SYNOPSIS
void -> status
ARGUMENT
struct DESTROYSESSION4args {
sessionid4 sessionid;
};
RESULT
struct SESSION_DESTROYres {
nfsstat status;
};
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 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.
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IMPLEMENTATION
No discussion at this time.
ERRORS
TBD
14.8 SEQUENCE - Supply per-procedure sequencing and control
Supply per-procedure sequencing and control
SYNOPSIS
control -> control
ARGUMENT
typedef uint32_t sequenceid4;
typedef uint32_t slotid4;
struct SEQUENCE4args {
clientid4 clientid;
sessionid4 sessionid;
sequenceid4 sequenceid;
slotid4 slotid;
slotid4 maxslot;
};
RESULT
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struct SEQUENCE4resok {
clientid4 clientid;
sessionid4 sessionid;
sequenceid4 sequenceid;
slotid4 slotid;
slotid4 maxslot;
slotid4 target_maxslot;
};
union SEQUENCE4res switch (nfsstat4 status) {
case NFS4_OK:
SEQUENCE4resok resok4;
default:
void;
};
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.
IMPLEMENTATION
No discussion at this time.
ERRORS
NFS4ERR_BADSESSION NFS4ERR_BADSLOT
14.9 CB_RECALLCREDIT - change flow control limits
Change flow control limits
SYNOPSIS
targetcount -> status
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ARGUMENT
struct CB_RECALLCREDIT4args {
sessionid4 sessionid;
uint32_t target;
};
RESULT
struct CB_RECALLCREDIT4res {
nfsstat4 status;
};
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.
IMPLEMENTATION
No discussion at this time.
ERRORS
NONE
14.10 CB_SEQUENCE - Supply callback channel sequencing and control
Sequence and control
SYNOPSIS
control -> control
ARGUMENT
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typedef uint32_t sequenceid4;
typedef uint32_t slotid4;
struct CB_SEQUENCE4args {
clientid4 clientid;
sessionid4 sessionid;
sequenceid4 sequenceid;
slotid4 slotid;
slotid4 maxslot;
};
RESULT
struct CB_SEQUENCE4resok {
clientid4 clientid;
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;
};
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.
IMPLEMENTATION
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No discussion at this time.
ERRORS
NFS4ERR_BADSESSION NFS4ERR_BADSLOT
14.11 GET_DIR_DELEGATION - Get a directory delegation
Obtain a directory delegation.
SYNOPSIS
(cfh), requested notification -> (cfh), cookieverf, stateid,
supported notification
ARGUMENT
struct GET_DIR_DELEGATION4args {
dir_notification_type4 notification_type;
attr_notice4 child_attr_delay;
attr_notice4 dir_attr_delay;
};
/*
* 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;
RESULT
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struct GET_DIR_DELEGATION4resok {
verifier4 cookieverf;
/* Stateid for get_dir_delegation */
stateid4 stateid;
/* Which notifications can the server support */
dir_notification_type4 supp_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;
};
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 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.
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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.
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 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
(supp_notification) specified in the reply.
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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.
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
14.12 CB_NOTIFY - Notify directory changes
Tell the client of directory changes.
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SYNOPSIS
stateid, notification -> {}
ARGUMENT
struct CB_NOTIFY4args {
stateid4 stateid;
dir_notification4 changes<>;
};
/*
* 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;
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nfs_cookie4 old_entry_cookie;
};
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;
};
RESULT
struct CB_NOTIFY4res {
nfsstat4 status;
};
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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
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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 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 supp_dir_attr_notice and
supp_child_attr_notice 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.
IMPLEMENTATION
ERRORS
NFS4ERR_BAD_STATEID NFS4ERR_INVAL NFS4ERR_BADXDR
NFS4ERR_SERVERFAULT
14.13 CB_RECALL_ANY - Keep any N delegations
Notify client to return delegation and keep N of them.
SYNOPSIS
N -> {}
ARGUMENT
struct CB_RECALLANYY4args {
uint4 dlgs_to_keep;
}
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RESULT
struct CB_RECALLANY4res {
nfsstat4 status;
};
DESCRIPTION
The server may decide that it can not hold all the delegation
state without running out of resources. Since the server has no
knowledge of which delegations are being used more than others, it
can not implement an effective reclaim scheme that avoids
reclaiming frequently used delegations. In that case the server
may issue a CB_RECALL_ANY callback to the client asking it to keep
N delegations and return the rest. The reason why CB_RECALL_ANY
specifies a count of delegations the client may keep as opposed to
a count of delegations the client must yield is as follows. Were
it otherwise, there is a potential for a race between a
CB_RECALL_ANY that had a count of delegations to free with a set
of client originated operations to return delegations. As a
result of the race the client and server would have differing
ideas as to how many delegations to return. Hence the client
could mistakenly free too many delegations. This operation
applies to delegations for a regular file (read or write) as well
as for a directory.
The client can choose to return any type of delegation as a result
of this callback i.e. read, write or directory delegation. The
client can also choose to keep more delegations than what the
server asked for and it is up to the server to handle this
situation. The server must give the client enough time to return
the delegations. This time should not be less than the lease
period.
IMPLEMENTATION
ERRORS
NFS4ERR_RESOURCE
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14.14 LAYOUTGET - Get Layout Information
SYNOPSIS
(cfh), clientid, layout_type, iomode, offset, length,
minlength, maxcount -> layout
ARGUMENT
struct LAYOUTGET4args {
/* CURRENT_FH: file */
clientid4 clientid;
pnfs_layouttype4 layout_type;
pnfs_layoutiomode4 iomode;
offset4 offset;
length4 length;
length4 minlength;
count4 maxcount;
};
RESULT
struct LAYOUTGET4resok {
pnfs_layout4 layout;
};
union LAYOUTGET4res switch (nfsstat4 status) {
case NFS4_OK:
LAYOUTGET4resok resok4;
default:
void;
};
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.
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
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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 7.3 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.
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 7.5.3 for details.
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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.
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.
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
14.15 LAYOUTCOMMIT - Commit writes made using a layout
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SYNOPSIS
(cfh), clientid, offset, length, last_write_offset,
time_modify, time_access, layoutupdate -> newsize
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;
length4 last_write_offset;
newtime4 time_modify;
newtime4 time_access;
pnfs_layoutupdate4 layoutupdate;
};
RESULT
struct LAYOUTCOMMIT4resok {
newsize4 newsize;
};
union LAYOUTCOMMIT4res switch (nfsstat4 status) {
case NFS4_OK:
LAYOUTCOMMIT4resok resok4;
default:
void;
};
DESCRIPTION
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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"
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 "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.
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 7.4 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
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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.
On success, the current filehandle retains its value.
ERRORS
NFS4ERR_BADLAYOUT
NFS4ERR_BADIOMODE
NFS4ERR_FHEXPIRED
NFS4ERR_INVAL
NFS4ERR_NOFILEHANDLE
NFS4ERR_STALE
NFS4ERR_STALE_CLIENTID
NFS4ERR_UNKNOWN_LAYOUTTYPE
14.16 LAYOUTRETURN - Release Layout Information
SYNOPSIS
(cfh), clientid, offset, length, iomode, layout_type -> -
ARGUMENT
struct LAYOUTRETURN4args {
/* CURRENT_FH: file */
clientid4 clientid;
offset4 offset;
length4 length;
pnfs_layoutiomode4 iomode;
pnfs_layouttype4 layout_type;
};
RESULT
struct LAYOUTRETURN4res {
nfsstat4 status;
};
DESCRIPTION
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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.
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 7.5 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.
On success, the current filehandle retains its value.
[OPEN ISSUE: Should LAYOUTRETURN be modified to handle FSID
callbacks?]
ERRORS
NFS4ERR_BADLAYOUT
NFS4ERR_BADIOMODE
NFS4ERR_FHEXPIRED
NFS4ERR_INVAL
NFS4ERR_NOFILEHANDLE
NFS4ERR_STALE
NFS4ERR_STALE_CLIENTID
NFS4ERR_UNKNOWN_LAYOUTTYPE
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14.17 GETDEVICEINFO - Get Device Information
SYNOPSIS
(cfh), device_id, layout_type, maxcount -> device_addr
ARGUMENT
struct GETDEVICEINFO4args {
/* CURRENT_FH: file */
pnfs_deviceid4 device_id;
pnfs_layouttype4 layout_type;
count4 maxcount;
};
RESULT
struct GETDEVICEINFO4resok {
pnfs_deviceaddr4 device_addr;
};
union GETDEVICEINFO4res switch (nfsstat4 status) {
case NFS4_OK:
GETDEVICEINFO4resok resok4;
default:
void;
};
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 7.1.4 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.
ERRORS
NFS4ERR_FHEXPIRED
NFS4ERR_INVAL
NFS4ERR_TOOSMALL
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NFS4ERR_UNKNOWN_LAYOUTTYPE
14.18 GETDEVICELIST - Get List of Devices
SYNOPSIS
(cfh), layout_type, maxcount, cookie, cookieverf ->
cookie, cookieverf, device_addrs<>
ARGUMENT
struct GETDEVICELIST4args {
/* CURRENT_FH: file */
pnfs_layouttype4 layout_type;
count4 maxcount;
nfs_cookie4 cookie;
verifier4 cookieverf;
};
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;
};
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 [2], section 14.2.24). As
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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
device address, it will return the error NFS4ERR_TOOSMALL. If an
invalid device ID is given, the metadata server will respond with
NFS4ERR_INVAL.
ERRORS
NFS4ERR_BAD_COOKIE
NFS4ERR_FHEXPIRED
NFS4ERR_INVAL
NFS4ERR_TOOSMALL
NFS4ERR_UNKNOWN_LAYOUTTYPE
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14.19 CB_LAYOUTRECALL
SYNOPSIS
layout_type, iomode, layoutchanged, layoutrecall -> -
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;
};
RESULT
struct CB_LAYOUTRECALLres {
nfsstat4 status;
};
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
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
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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.
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
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 7.5.2 for
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details).
ERRORS
NFS4ERR_NOMATCHING_LAYOUT
14.20 CB_SIZECHANGED
SYNOPSIS
fh, size -> -
ARGUMENT
struct CB_SIZECHANGEDargs {
nfs_fh4 fh;
length4 size;
};
RESULT
struct CB_SIZECHANGEDres {
nfsstat4 status;
};
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 9.5). For other layout types see Section 7.4.2 for more
details.
If the handle specified is not one for which the client holds a
layout, an NFS4ERR_BADHANDLE error is returned.
ERRORS
NFS4ERR_BADHANDLE
15. References
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15.1 Normative References
[1] Bradner, S., "Key words for use in RFCs to Indicate Requirement
Levels", March 1997.
[2] 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.
15.2 Informative References
[3] Satran, J., Meth, K., Sapuntzakis, C., Chadalapaka, M., and E.
Zeidner, "Internet Small Computer Systems Interface (iSCSI)",
RFC 3720, April 2004.
[4] Snively, R., "Fibre Channel Protocol for SCSI, 2nd Version
(FCP-2)", ANSI/INCITS 350-2003, Oct 2003.
[5] 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>.
[6] Black, D., "pNFS Block/Volume Layout", July 2005, <ftp://
www.ietf.org/internet-drafts/draft-black-pnfs-block-01.txt>.
[7] Zelenka, J., Welch, B., and B. Halevy, "Object-based pNFS
Operations", July 2005, <ftp://www.ietf.org/internet-drafts/
draft-zelenka-pnfs-obj-01.txt>.
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
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
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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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