TCP Maintenance and Minor F. Gont
Extensions (tcpm) UK CPNI
Internet-Draft January 21, 2011
Intended status: BCP
Expires: July 25, 2011
Security Assessment of the Transmission Control Protocol (TCP)
draft-ietf-tcpm-tcp-security-02.txt
Abstract
This document contains a security assessment of the specifications of
the Transmission Control Protocol (TCP), and of a number of
mechanisms and policies in use by popular TCP implementations.
Additionally, it contains best current practices for hardening a TCP
implementation.
Status of this Memo
This Internet-Draft is submitted to IETF in full conformance with the
provisions of BCP 78 and BCP 79.
Internet-Drafts are working documents of the Internet Engineering
Task Force (IETF). Note that other groups may also distribute
working documents as Internet-Drafts. The list of current Internet-
Drafts is at http://datatracker.ietf.org/drafts/current/.
Internet-Drafts are draft documents valid for a maximum of six months
and may be updated, replaced, or obsoleted by other documents at any
time. It is inappropriate to use Internet-Drafts as reference
material or to cite them other than as "work in progress."
This Internet-Draft will expire on July 25, 2011.
Copyright Notice
Copyright (c) 2011 IETF Trust and the persons identified as the
document authors. All rights reserved.
This document is subject to BCP 78 and the IETF Trust's Legal
Provisions Relating to IETF Documents
(http://trustee.ietf.org/license-info) in effect on the date of
publication of this document. Please review these documents
carefully, as they describe your rights and restrictions with respect
to this document. Code Components extracted from this document must
include Simplified BSD License text as described in Section 4.e of
the Trust Legal Provisions and are provided without warranty as
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described in the Simplified BSD License.
Table of Contents
1. Preface . . . . . . . . . . . . . . . . . . . . . . . . . . . 5
1.1. Introduction . . . . . . . . . . . . . . . . . . . . . . 5
1.2. Scope of this document . . . . . . . . . . . . . . . . . 6
1.3. Organization of this document . . . . . . . . . . . . . . 8
2. The Transmission Control Protocol . . . . . . . . . . . . . . 8
3. TCP header fields . . . . . . . . . . . . . . . . . . . . . . 9
3.1. Source Port and Destination Port . . . . . . . . . . . . 10
3.2. Sequence number . . . . . . . . . . . . . . . . . . . . . 12
3.3. Acknowledgement Number . . . . . . . . . . . . . . . . . 14
3.4. Data Offset . . . . . . . . . . . . . . . . . . . . . . . 15
3.5. Control bits . . . . . . . . . . . . . . . . . . . . . . 15
3.5.1. Reserved (four bits) . . . . . . . . . . . . . . . . 15
3.5.2. CWR (Congestion Window Reduced) . . . . . . . . . . . 16
3.5.3. ECE (ECN-Echo) . . . . . . . . . . . . . . . . . . . 16
3.5.4. URG . . . . . . . . . . . . . . . . . . . . . . . . . 17
3.5.5. ACK . . . . . . . . . . . . . . . . . . . . . . . . . 17
3.5.6. PSH . . . . . . . . . . . . . . . . . . . . . . . . . 17
3.5.7. RST . . . . . . . . . . . . . . . . . . . . . . . . . 19
3.5.8. SYN . . . . . . . . . . . . . . . . . . . . . . . . . 19
3.5.9. FIN . . . . . . . . . . . . . . . . . . . . . . . . . 20
3.6. Window . . . . . . . . . . . . . . . . . . . . . . . . . 20
3.7. Checksum . . . . . . . . . . . . . . . . . . . . . . . . 22
3.8. Urgent pointer . . . . . . . . . . . . . . . . . . . . . 23
3.9. Options . . . . . . . . . . . . . . . . . . . . . . . . . 24
3.10. Padding . . . . . . . . . . . . . . . . . . . . . . . . . 28
3.11. Data . . . . . . . . . . . . . . . . . . . . . . . . . . 28
4. Common TCP Options . . . . . . . . . . . . . . . . . . . . . 29
4.1. End of Option List (Kind = 0) . . . . . . . . . . . . . . 29
4.2. No Operation (Kind = 1) . . . . . . . . . . . . . . . . . 29
4.3. Maximum Segment Size (Kind = 2) . . . . . . . . . . . . . 29
4.4. Selective Acknowledgement Option . . . . . . . . . . . . 32
4.4.1. SACK-permitted Option (Kind = 4) . . . . . . . . . . 32
4.4.2. SACK Option (Kind = 5) . . . . . . . . . . . . . . . 33
4.5. MD5 Option (Kind=19) . . . . . . . . . . . . . . . . . . 35
4.6. Window scale option (Kind = 3) . . . . . . . . . . . . . 36
4.7. Timestamps option (Kind = 8) . . . . . . . . . . . . . . 37
4.7.1. Generation of timestamps . . . . . . . . . . . . . . 37
4.7.2. Vulnerabilities . . . . . . . . . . . . . . . . . . . 38
5. Connection-establishment mechanism . . . . . . . . . . . . . 39
5.1. SYN flood . . . . . . . . . . . . . . . . . . . . . . . . 40
5.2. Connection forgery . . . . . . . . . . . . . . . . . . . 44
5.3. Connection-flooding attack . . . . . . . . . . . . . . . 45
5.3.1. Vulnerability . . . . . . . . . . . . . . . . . . . . 45
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5.3.2. Countermeasures . . . . . . . . . . . . . . . . . . . 46
5.4. Firewall-bypassing techniques . . . . . . . . . . . . . . 48
6. Connection-termination mechanism . . . . . . . . . . . . . . 49
6.1. FIN-WAIT-2 flooding attack . . . . . . . . . . . . . . . 49
6.1.1. Vulnerability . . . . . . . . . . . . . . . . . . . . 49
6.1.2. Countermeasures . . . . . . . . . . . . . . . . . . . 50
7. Buffer management . . . . . . . . . . . . . . . . . . . . . . 52
7.1. TCP retransmission buffer . . . . . . . . . . . . . . . . 52
7.1.1. Vulnerability . . . . . . . . . . . . . . . . . . . . 52
7.1.2. Countermeasures . . . . . . . . . . . . . . . . . . . 53
7.2. TCP segment reassembly buffer . . . . . . . . . . . . . . 56
7.3. Automatic buffer tuning mechanisms . . . . . . . . . . . 59
7.3.1. Automatic send-buffer tuning mechanisms . . . . . . . 59
7.3.2. Automatic receive-buffer tuning mechanism . . . . . . 61
8. TCP segment reassembly algorithm . . . . . . . . . . . . . . 63
8.1. Problems that arise from ambiguity in the reassembly
process . . . . . . . . . . . . . . . . . . . . . . . . . 63
9. TCP Congestion Control . . . . . . . . . . . . . . . . . . . 64
9.1. Congestion control with misbehaving receivers . . . . . . 66
9.1.1. ACK division . . . . . . . . . . . . . . . . . . . . 66
9.1.2. DupACK forgery . . . . . . . . . . . . . . . . . . . 66
9.1.3. Optimistic ACKing . . . . . . . . . . . . . . . . . . 67
9.2. Blind DupACK triggering attacks against TCP . . . . . . . 68
9.2.1. Blind throughput-reduction attack . . . . . . . . . . 70
9.2.2. Blind flooding attack . . . . . . . . . . . . . . . . 70
9.2.3. Difficulty in performing the attacks . . . . . . . . 71
9.2.4. Modifications to TCP's loss recovery algorithms . . . 72
9.2.5. Countermeasures . . . . . . . . . . . . . . . . . . . 74
9.3. TCP Explicit Congestion Notification (ECN) . . . . . . . 79
9.3.1. Possible attacks by a compromised router . . . . . . 79
9.3.2. Possible attacks by a malicious TCP endpoint . . . . 80
10. TCP API . . . . . . . . . . . . . . . . . . . . . . . . . . . 81
10.1. Passive opens and binding sockets . . . . . . . . . . . . 81
10.2. Active opens and binding sockets . . . . . . . . . . . . 82
11. Blind in-window attacks . . . . . . . . . . . . . . . . . . . 84
11.1. Blind TCP-based connection-reset attacks . . . . . . . . 84
11.1.1. RST flag . . . . . . . . . . . . . . . . . . . . . . 85
11.1.2. SYN flag . . . . . . . . . . . . . . . . . . . . . . 86
11.1.3. Security/Compartment . . . . . . . . . . . . . . . . 88
11.1.4. Precedence . . . . . . . . . . . . . . . . . . . . . 89
11.1.5. Illegal options . . . . . . . . . . . . . . . . . . . 90
11.2. Blind data-injection attacks . . . . . . . . . . . . . . 90
12. Information leaking . . . . . . . . . . . . . . . . . . . . . 91
12.1. Remote Operating System detection via TCP/IP stack
fingerprinting . . . . . . . . . . . . . . . . . . . . . 91
12.1.1. FIN probe . . . . . . . . . . . . . . . . . . . . . . 91
12.1.2. Bogus flag test . . . . . . . . . . . . . . . . . . . 92
12.1.3. TCP ISN sampling . . . . . . . . . . . . . . . . . . 92
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12.1.4. TCP initial window . . . . . . . . . . . . . . . . . 92
12.1.5. RST sampling . . . . . . . . . . . . . . . . . . . . 93
12.1.6. TCP options . . . . . . . . . . . . . . . . . . . . . 94
12.1.7. Retransmission Timeout (RTO) sampling . . . . . . . . 94
12.2. System uptime detection . . . . . . . . . . . . . . . . . 94
13. Covert channels . . . . . . . . . . . . . . . . . . . . . . . 95
14. TCP Port scanning . . . . . . . . . . . . . . . . . . . . . . 95
14.1. Traditional connect() scan . . . . . . . . . . . . . . . 96
14.2. SYN scan . . . . . . . . . . . . . . . . . . . . . . . . 96
14.3. FIN, NULL, and XMAS scans . . . . . . . . . . . . . . . . 96
14.4. Maimon scan . . . . . . . . . . . . . . . . . . . . . . . 98
14.5. Window scan . . . . . . . . . . . . . . . . . . . . . . . 98
14.6. ACK scan . . . . . . . . . . . . . . . . . . . . . . . . 99
15. Processing of ICMP error messages by TCP . . . . . . . . . . 99
16. TCP interaction with the Internet Protocol (IP) . . . . . . . 99
16.1. TCP-based traceroute . . . . . . . . . . . . . . . . . . 99
16.2. Blind TCP data injection through fragmented IP traffic . 100
16.3. Broadcast and multicast IP addresses . . . . . . . . . . 102
17. Security Considerations . . . . . . . . . . . . . . . . . . . 102
18. Acknowledgements . . . . . . . . . . . . . . . . . . . . . . 102
19. References . . . . . . . . . . . . . . . . . . . . . . . . . 103
20. References . . . . . . . . . . . . . . . . . . . . . . . . . 113
20.1. Normative References . . . . . . . . . . . . . . . . . . 113
20.2. Informative References . . . . . . . . . . . . . . . . . 113
Appendix A. TODO list . . . . . . . . . . . . . . . . . . . . . 113
Appendix B. Change log (to be removed by the RFC Editor
before publication of this document as an RFC) . . . 113
B.1. Changes from draft-ietf-tcpm-tcp-security-01 . . . . . . 113
Author's Address . . . . . . . . . . . . . . . . . . . . . . . . 114
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1. Preface
1.1. Introduction
The TCP/IP protocol suite was conceived in an environment that was
quite different from the hostile environment they currently operate
in. However, the effectiveness of the protocols led to their early
adoption in production environments, to the point that, to some
extent, the current world's economy depends on them.
While many textbooks and articles have created the myth that the
Internet protocols were designed for warfare environments, the top
level goal for the DARPA Internet Program was the sharing of large
service machines on the ARPANET [Clark, 1988]. As a result, many
protocol specifications focus only on the operational aspects of the
protocols they specify, and overlook their security implications.
While the Internet technology evolved since it early inception, the
Internet's building blocks are basically the same core protocols
adopted by the ARPANET more than two decades ago. During the last
twenty years, many vulnerabilities have been identified in the TCP/IP
stacks of a number of systems. Some of them were based on flaws in
some protocol implementations, affecting only a reduced number of
systems, while others were based in flaws in the protocols
themselves, affecting virtually every existing implementation
[Bellovin, 1989]. Even in the last couple of years, researchers were
still working on security problems in the core protocols [NISCC,
2004] [NISCC, 2005].
The discovery of vulnerabilities in the TCP/IP protocol suite usually
led to reports being published by a number of CSIRTs (Computer
Security Incident Response Teams) and vendors, which helped to raise
awareness about the threats and the best mitigations known at the
time the reports were published. Unfortunately, this also led to the
documentation of the discovered protocol vulnerabilities being spread
among a large number of documents, which are sometimes difficult to
identify.
For some reason, much of the effort of the security community on the
Internet protocols did not result in official documents (RFCs) being
issued by the IETF (Internet Engineering Task Force). This basically
led to a situation in which "known" security problems have not always
been addressed by all vendors. In addition, in many cases vendors
have implemented quick "fixes" to the identified vulnerabilities
without a careful analysis of their effectiveness and their impact on
interoperability [Silbersack, 2005].
Producing a secure TCP/IP implementation nowadays is a very difficult
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task, in part because of the lack of a single document that serves as
a security roadmap for the protocols. Implementers are faced with
the hard task of identifying relevant documentation and
differentiating between that which provides correct advice, and that
which provides misleading advice based on inaccurate or wrong
assumptions.
There is a clear need for a companion document to the IETF
specifications that discusses the security aspects and implications
of the protocols, identifies the existing vulnerabilities, discusses
the possible countermeasures, and analyzes their respective
effectiveness.
This document is the result of a security assessment of the IETF
specifications of the Transmission Control Protocol (TCP), from a
security point of view. Possible threats are identified and, where
possible, countermeasures are proposed. Additionally, many
implementation flaws that have led to security vulnerabilities have
been referenced in the hope that future implementations will not
incur the same problems.
This document does not aim to be the final word on the security
aspects of TCP. On the contrary, it aims to raise awareness about a
number of TCP vulnerabilities that have been faced in the past, those
that are currently being faced, and some of those that we may still
have to deal with in the future.
Feedback from the community is more than encouraged to help this
document be as accurate as possible and to keep it updated as new
vulnerabilities are discovered.
This document is heavily based on the "Security Assessment of the
Transmission Control Protocol (TCP)" released by the UK Centre for
the Protection of National Infrastructure (CPNI), available at: http:
//www.cpni.gov.uk/Products/technicalnotes/
Feb-09-security-assessment-TCP.aspx .
1.2. Scope of this document
While there are a number of protocols that may affect the way TCP
operates, this document focuses only on the specifications of the
Transmission Control Protocol (TCP) itself.
The following IETF RFCs were selected for assessment as part of this
work:
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o RFC 793, "Transmission Control Protocol. DARPA Internet Program.
Protocol Specification" (91 pages)
o RFC 1122, "Requirements for Internet Hosts -- Communication
Layers" (116 pages)
o RFC 1191, "Path MTU Discovery" (19 pages)
o RFC 1323, "TCP Extensions for High Performance" (37 pages)
o RFC 1948, "Defending Against Sequence Number Attacks" (6 pages)
o RFC 1981, "Path MTU Discovery for IP version 6" (15 pages)
o RFC 2018, "TCP Selective Acknowledgment Options" (12 pages)
o RFC 2385, "Protection of BGP Sessions via the TCP MD5 Signature
Option" (6 pages)
o RFC 2581, "TCP Congestion Control" (14 pages)
o RFC 2675, "IPv6 Jumbograms" (9 pages)
o RFC 2883, "An Extension to the Selective Acknowledgement (SACK)
Option for TCP" (17 pages)
o RFC 2884, "Performance Evaluation of Explicit Congestion
Notification (ECN) in IP Networks" (18 pages)
o RFC 2988, "Computing TCP's Retransmission Timer" (8 pages)
o RFC 3168, "The Addition of Explicit Congestion Notification (ECN)
to IP" (63 pages)
o RFC 3465, "TCP Congestion Control with Appropriate Byte Counting
(ABC)" (10 pages)
o RFC 3517, "A Conservative Selective Acknowledgment (SACK)-based
Loss Recovery Algorithm for TCP" (13 pages)
o RFC 3540, "Robust Explicit Congestion Notification (ECN) Signaling
with Nonces" (13 pages)
o RFC 3782, "The NewReno Modification to TCP's Fast Recovery
Algorithm" (19 pages)
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1.3. Organization of this document
This document is basically organized in two parts. The first part
contains a discussion of each of the TCP header fields, identifies
their security implications, and discusses the possible
countermeasures. The second part contains an analysis of the
security implications of the mechanisms and policies implemented by
TCP, and of a number of implementation strategies in use by a number
of popular TCP implementations.
2. The Transmission Control Protocol
The Transmission Control Protocol (TCP) is a connection-oriented
transport protocol that provides a reliable byte-stream data transfer
service.
Very few assumptions are made about the reliability of underlying
data transfer services below the TCP layer. Basically, TCP assumes
it can obtain a simple, potentially unreliable datagram service from
the lower level protocols. Figure 1 illustrates where TCP fits in
the DARPA reference model.
+---------------+
| Application |
+---------------+
| TCP |
+---------------+
| IP |
+---------------+
| Network |
+---------------+
Figure 1: TCP in the DARPA reference model
TCP provides facilities in the following areas:
o Basic Data Transfer
o Reliability
o Flow Control
o Multiplexing
o Connections
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o Precedence and Security
o Congestion Control
The core TCP specification, RFC 793 [Postel, 1981c], dates back to
1981 and standardizes the basic mechanisms and policies of TCP. RFC
1122 [Braden, 1989] provides clarifications and errata for the
original specification. RFC 2581 [Allman et al, 1999] specifies TCP
congestion control and avoidance mechanisms, not present in the
original specification. Other documents specify extensions and
improvements for TCP.
The large amount of documents that specify extensions, improvements,
or modifications to existing TCP mechanisms has led the IETF to
publish a roadmap for TCP, RFC 4614 [Duke et al, 2006], that
clarifies the relevance of each of those documents.
3. TCP header fields
RFC 793 [Postel, 1981c] defines the syntax of a TCP segment, along
with the semantics of each of the header fields. Figure 2
illustrates the syntax of a TCP segment.
0 1 2 3
0 1 2 3 4 5 6 7 8 9 0 1 2 3 4 5 6 7 8 9 0 1 2 3 4 5 6 7 8 9 0 1
+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+
| Source Port | Destination Port |
+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+
| Sequence Number |
+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+
| Acknowledgment Number |
+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+
| Data | |C|E|U|A|P|R|S|F| |
| Offset|Resrved|W|C|R|C|S|S|Y|I| Window |
| | |R|E|G|K|H|T|N|N| |
+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+
| Checksum | Urgent Pointer |
+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+
| Options | Padding |
+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+
| data |
+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+-+
Note that one tick mark represents one bit position
Figure 2: Transmission Control Protocol header format
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The minimum TCP header size is 20 bytes, and corresponds to a TCP
segment with no options and no data. However, a TCP module might be
handed an (illegitimate) "TCP segment" of less than 20 bytes.
Therefore, before doing any processing of the TCP header fields, the
following check should be performed by TCP on the segments handed by
the internet layer:
Segment.Size >= 20
If a segment does not pass this check, it should be dropped.
The following subsections contain further sanity checks that should
be performed on TCP segments.
3.1. Source Port and Destination Port
The Source Port field contains a 16-bit number that identifies the
TCP end-point that originated this TCP segment. The TCP Destination
Port contains a 16-bit number that identifies the destination TCP
end-point of this segment. In most of the discussion we refer to
client-side (or "ephemeral") port-numbers and server-side port
numbers, since that distinction is what usually affects the
interpretation of a port number.
TCP SHOULD randomize its ephemeral (client-side) ports, to improve
its resistance to off-path attacks. For the purpose of ephemeral
port selection, the largest posible port range SHOULD be used
(ideally 1024-65535) I-D.ietf-tsvwg-port-randomization.
DISCUSSION:
[I-D.ietf-tsvwg-port-randomization] provides advice on port
randomization.
TCP MUST NOT allocate port number 0, as its use could lead to
interoperability problems. If a segment is received with port 0 as
the Source Port or the Destination Port, a RST segment SHOULD be sent
in response (provided that the incomming segment does not have the
RST flag set).
DISCUSSION:
While port 0 is a legitimate port number, it has a special meaning
in the UNIX Sockets API. For example, when a TCP port number of 0
is passed as an argument to the bind() function, rather than
binding port 0, an ephemeral port is selected for the
corresponding TCP end-point. As a result, the TCP port number 0
is never actually used in TCP segments.
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Different implementations have been found to respond differently
to TCP segments that have a port number of 0 as the Source Port
and/or the Destination Port. As a result, TCP segments with a
port number of 0 are usually employed for remote OS detection via
TCP/IP stack fingerprinting [Jones, 2003].
Since in practice TCP port 0 is not used by any legitimate
application and is only used for fingerprinting purposes, a number
of host implementations already reject TCP segments that use 0 as
the Source Port and/or the Destination Port. Also, a number
firewalls filter (by default) any TCP segments that contain a port
number of zero for the Source Port and/or the Destination Port.
We therefore recommend that TCP implementations respond to
incoming TCP segments that have a Source Port or a Destination
Port of 0 with an RST (provided these incoming segments do not
have the RST bit set).
Responding with an RST segment to incoming segments that have the
RST bit would open the door to RST-war attacks.
TCP MUST be able to grecefully handle the case where the source end-
point (IP Source Address, TCP Source Port) is the same as the
destination end-point (IP Destination Address, TCP Destination Port).
DISCUSSION:
Some systems have been found to be unable to process TCP segments
in which the source endpoint {Source Address, Source Port} is the
same than the destination end-point {Destination Address,
Destination Port}. Such TCP segments have been reported to cause
malfunction of a number of implementations [CERT, 1996], and have
been exploited in the past to perform Denial of Service (DoS)
attacks [Meltman, 1997]. While these packets are very very
unlikely to exist in real and legitimate scenarios, TCP should
nevertheless be able to process them without the need of any
"extra" code.
A SYN segment in which the source end-point {Source Address,
Source Port} is the same as the destination end-point {Destination
Address, Destination Port} will result in a "simultaneous open"
scenario, such as the one described in page 32 of RFC 793 [Postel,
1981c]. Therefore, those TCP implementations that correctly
handle simultaneous opens should already be prepared to handle
these unusual TCP segments.
TCP SHOULD NOT allocate of port numbers that are in use by a TCP that
is in the LISTEN or CLOSED states for use as ephemeral ports, as this
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could allow attackers on the local system to "steal" incomming TCP
connections.
DISCUSSION:
While the only requirement for a selected ephemeral port is that
the resulting four-tuple (connection-id) is unique (i.e., not
currently in use by any other TCP connection), in practice it may
be necessary to not allow the allocation of port numbers that are
in use by a TCP that is in the LISTEN or CLOSED states for use as
ephemeral ports, as this might allow an attacker to "steal"
incoming connections from a local server application. Therefore,
TCP SHOULD NOT allocate port numbers that are in use by a TCP in
the LISTEN or CLOSED states for use as ephemeral ports. Section
10.2 of this document provides a detailed discussion of this
issue.
While some systems restrict use of the port numbers in the range
0-1024 to privileged users, applications SHOULD NOT grant any trust
based on the port numbers used for a TCP connection.
DISCUSSION:
Not all systems require superuser privileges to bind port numbers
in that range. Besides, with desktop computers such "distinction"
has generally become irrelevant.
Middle-boxes such as packet filters MUST NOT assume that clients use
port numbers from only the Dynamic or Registered port ranges.
DISCUSSION:
It should also be noted that some clients, such as DNS resolvers,
are known to use port numbers from the "Well Known Ports" range.
Therefore, middle-boxes such as packet filters MUST NOT assume
that clients use port number from only the Dynamic or Registered
port ranges.
3.2. Sequence number
TCP SHOULD select its Initial Sequence Numbers (ISNs) with the
following expression:
ISN = M + F(localhost, localport, remotehost, remoteport, secret_key)
where M is a monotonically increasing counter maintained within TCP,
and F() is a Pseudo-Random Function (PRF). As it is vital that F()
not be computable from the outside, F() could be a PRF of the
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connection-id and some secret data. HMAC-SHA-256 would be a good
choice for F()
DISCUSSION:
The choice of the Initial Sequence Number of a connection is not
arbitrary, but aims to minimize the chances of a stale segment
from being accepted by a new incarnation of a previous connection.
RFC 793 [Postel, 1981c] suggests the use of a global 32-bit ISN
generator, whose lower bit is incremented roughly every 4
microseconds.
However, use of such an ISN generator makes it trivial to predict
the ISN that a TCP will use for new connections, thus allowing a
variety of attacks against TCP, such as those described in Section
5.2 and Section 11 of this document. This vulnerability was first
described in [Morris, 1985], and its exploitation was widely
publicized about 10 years later [Shimomura, 1995].
As a matter of fact, protection against old stale segments from a
previous incarnation of the connection comes from allowing the
creation of a new incarnation of a previous connection only after
2*MSL have passed since a segment corresponding to the old
incarnation was last seen. This is accomplished by the TIME-WAIT
state, and TCP's "quiet time" concept. However, as discussed in
Section 3.1 and Section 11.1.2 of this document, the ISN can be
used to perform some heuristics meant to avoid an interoperability
problem that may arise when two systems establish connections at a
high rate. In order for such heuristics to work, the ISNs
generated by a TCP should be monotonically increasing.
The ISN generation scheme recommended in this section was
originally proposed in RFC 1948 [Bellovin, 1996], such that the
chances of an attacker from guessing the ISN of a TCP are reduced,
while still producing a monotonically-increasing sequence that
allows implementation of the optimization described in Section 3.1
and Section 11.1.2 of this document.
[CERT, 2001] and [US-CERT, 2001] are advisories about the security
implications of weak ISN generators. [Zalewski, 2001a] and
[Zalewski, 2002] contain a detailed analysis of ISN generators,
and a survey of the algorithms in use by popular TCP
implementations.
Another security consideration that should be made about TCP
sequence numbers is that they might allow an attacker to count the
number of systems behind a Network Address Translator (NAT)
[Srisuresh and Egevang, 2001]. Depending on the ISN generators
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implemented by each of the systems behind the NAT, an attacker
might be able to count the number of systems behind the NAT by
establishing a number of TCP connections (using the public address
of the NAT) and indentifying the number of different sequence
number "spaces". This information leakage could be eliminated by
rewriting the contents of all those header fields and options that
make use of sequence numbers (such as the Sequence Number and the
Acknowledgement Number fields, and the SACK Option) at the NAT.
[Gont and Srisuresh, 2008] provides a detailed discussion of the
security implications of NATs and of the possible mitigations for
this and other issues.
3.3. Acknowledgement Number
TCP SHOULD set the Acknowledgement Number to zero when sending a TCP
segment that does not have the ACK bit set (i.e., a SYN segment).
TCP MUST check that, on segments that have the ACK bit set, the
Acknowledgment Number satisfies the expression:
SND.UNA - SND.MAX.WND <= SEG.ACK <= SND.NXT
If a TCP segment does not pass this check, the segment MUST be
dropped, and an ACK segment SHOULD be sent in response.
DISCUSSION:
If the ACK bit is on, the Acknowledgement Number contains the
value of the next sequence number the sender of this segment is
expecting to receive. According to RFC 793, the Acknowledgement
Number is considered valid as long as it does not acknowledge the
receipt of data that has not yet been sent.
However, as a result of recent concerns on forgery attacks against
TCP (see Section 11 of this document), ongoing work at the IETF
[Ramaiah et al, 2008] has proposed to enforce a more strict check
on the Acknowledgement Number of segments that have the ACK bit
set:
SND.UNA - SND.MAX.WND <= SEG.ACK <= SND.NXT
If the ACK bit is off, the Acknowledgement Number field is not
valid. We recommend TCP implementations to set the
Acknowledgement Number to zero when sending a TCP segment that
does not have the ACK bit set (i.e., a SYN segment). Some TCP
implementations have been known to fail to set the Acknowledgement
Number to zero, thus leaking information.
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TCP Acknowledgements are also used to perform heuristics for loss
recovery and congestion control. Section 9 of this document
describes a number of ways in which these mechanisms can be
exploited.
3.4. Data Offset
TCP MUST enforce the following checks on the Data Offset field:
Data Offset >= 5
Data Offset * 4 <= TCP segment length
If a TCP segment does not pass these checks, it should be silently
dropped.
The TCP segment length should be obtained from the IP layer, as
TCP does not include a TCP segment length field.
DISCUSSION:
The Data Offset field indicates the length of the TCP header in
32-bit words. As the minimum TCP header size is 20 bytes, the
minimum legal value for this field is 5.
For obvious reasons, the TCP header cannot be larger than the
whole TCP segment it is part of.
3.5. Control bits
The following subsections provide a discussion of the different
control bits in the TCP header. TCP segments with unusual
combinations of flags set have been known in the past to cause
malfunction of some implementations, sometimes to the extent of
causing them to crash [Postel, 1987] [Braden, 1992]. These packets
are still usually employed for the purpose of TCP/IP stack
fingerprinting. Section 12.1 contains a discussion of TCP/IP stack
fingerprinting.
3.5.1. Reserved (four bits)
TCP MUST ignore the Reserved field of incoming TCP segments.
DISCUSSION:
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These four bits are reserved for future use, and must be zero. As
with virtually every field, the Reserved field could be used as a
covert channel. While there exist intermediate devices such as
protocol scrubbers that clear these bits, and firewalls that drop/
reject segments with any of these bits set, these devices should
consider the impact of these policies on TCP interoperability.
For example, as TCP continues to evolve, all or part of the bits
in the Reserved field could be used to implement some new
functionality. If some middle-box or end-system implementation
were to drop a TCP segment merely because some of these bits are
not set to zero, interoperability problems would arise.
3.5.2. CWR (Congestion Window Reduced)
DISCUSSION:
The CWR flag, defined in RFC 3168 [Ramakrishnan et al, 2001], is
used as part of the Explicit Congestion Notification (ECN)
mechanism. For connections in any of the synchronized states,
this flag indicates, when set, that the TCP sending this segment
has reduced its congestion window.
An analysis of the security implications of ECN can be found in
Section 9.3 of this document.
3.5.3. ECE (ECN-Echo)
DISCUSSION:
The ECE flag, defined in RFC 3168 [Ramakrishnan et al, 2001], is
used as part of the Explicit Congestion Notification (ECN)
mechanism.
Once a TCP connection has been established, an ACK segment with
the ECE bit set indicates that congestion was encountered in the
network on the path from the sender to the receiver. This
indication of congestion should be treated just as a congestion
loss in non-ECN-capable TCP [Ramakrishnan et al, 2001].
Additionally, TCP should not increase the congestion window (cwnd)
in response to such an ACK segment that indicates congestion, and
should also not react to congestion indications more than once
every window of data (or once per round-trip time).
An analysis of the security implications of ECN can be found in
Section 9.3 of this document.
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3.5.4. URG
DISCUSSION:
When the URG flag is set, the Urgent Pointer field contains the
current value of the urgent pointer.
Receipt of an "urgent" indication generates, in a number of
implementations (such as those in UNIX-like systems), a software
interrupt (signal) that is delivered to the corresponding process.
In UNIX-like systems, receipt of an urgent indication causes a
SIGURG signal to be delivered to the corresponding process.
A number of applications handle TCP urgent indications by
installing a signal handler for the corresponding signal (e.g.,
SIGURG). As discussed in [Zalewski, 2001b], some signal handlers
can be maliciously exploited by an attacker, for example to gain
remote access to a system. While secure programming of signal
handlers is out of the scope of this document, we nevertheless
raise awareness that TCP urgent indications might be exploited to
abuse poorly-written signal handlers.
Section 3.9 discusses the security implications of the TCP urgent
mechanism.
3.5.5. ACK
DISCUSSION:
When the ACK bit is one, the Acknowledgment Number field contains
the next sequence number expected, cumulatively acknowledging the
receipt of all data up to the sequence number in the
Acknowledgement Number, minus one. Section 3.4 of this document
describes sanity checks that should be performed on the
Acknowledgement Number field.
TCP Acknowledgements are also used to perform heuristics for loss
recovery and congestion control. Section 9 of this document
describes a number of ways in which these mechanisms can be
exploited.
3.5.6. PSH
As a result of a SEND call, TCP SHOULD send all queued data (provided
that TCP's flow control and congestion control algorithms allow it).
Received data SHOULD be immediately delivered to an application
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calling the RECEIVE function, even if the data already available are
less than those requested by the application.
DISCUSSION:
RFC 793 [Postel, 1981c] contains (in pages 54-64) a functional
description of a TCP Application Programming Interface (API). One
of the parameters of the SEND function is the PUSH flag which,
when set, signals the local TCP that it must send all unsent data.
The TCP PSH (PUSH) flag will be set in the last outgoing segment,
to signal the push function to the receiving TCP. Upon receipt of
a segment with the PSH flag set, the receiving user's buffer is
returned to the user, without waiting for additional data to
arrive.
There are two security considerations arising from the PUSH
function. On the sending side, an attacker could cause a large
amount of data to be queued for transmission without setting the
PUSH flag in the SEND call. This would prevent the local TCP from
sending the queued data, causing system memory to be tied to those
data for an unnecessarily long period of time.
An analogous consideration should be made for the receiving TCP.
TCP is allowed to buffer incoming data until the receiving user's
buffer fills or a segment with the PSH bit set is received. If
the receiving TCP implements this policy, an attacker could send a
large amount of data, slightly less than the receiving user's
buffer size, to cause system memory to be tied to these data for
an unnecessarily long period of time. Both of these issues are
discussed in Section 4.2.2.2 of RFC 1122 [Braden, 1989].
In order to mitigate these potential vulnerabilities, we suggest
assuming an implicit "PUSH" in every SEND call. On the sending
side, this means that as a result of a SEND call TCP should try to
send all queued data (provided that TCP's flow control and
congestion control algorithms allow it). On the receiving side,
this means that the received data will be immediately delivered to
an application calling the RECEIVE function, even if the data
already available are less than those requested by the
application.
It is interesting to note that popular TCP APIs (such as
"sockets") do not provide a PUSH flag in any of the interfaces
they define, but rather perform some kind of "heuristics" to set
the PSH bit in outgoing segments. As a result, the value of the
PSH bit in the received TCP segments is usually a policy of the
sending TCP, rather than a policy of the sending application. All
robust applications that make use of those APIs (such as the
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sockets API) properly handle the case of a RECEIVE call returning
less data (e.g., zero) than requested, usually by performing
subsequent RECEIVE calls.
Another potential malicious use of the PSH bit would be for an
attacker to send small TCP segments (probably with zero bytes of
data payload) to cause the receiving application to be
unnecessarily woken up (increasing the CPU load), or to cause
malfunction of poorly-written applications that may not handle
well the case of RECEIVE calls returning less data than requested.
3.5.7. RST
TCP MUST process RST segments (i.e., segments with the RST bit set)
as follows:
o If the Sequence Number of the RST segment is not valid (i.e.,
falls outside of the receive window), silently drop the segment.
o If the Sequence Number of the RST segment matches the next
expected sequence number (RCV.NXT), abort the corresponding
connection.
o If the Sequence Number is valid (i.e., falls within the receive
window) but is not exactly RCV.NXT, send an ACK segment (a
"challenge ACK") of the form: <SEQ=SND.NXT><ACK=RCV.NXT><CTL=ACK>.
TCP SHOULD rate-limit these challenge ACK segments.
DISCUSSION:
The RST bit is used to request the abortion (abnormal close) of a
TCP connection. RFC 793 [Postel, 1981c] suggests that an RST
segment should be considered valid if its Sequence Number is valid
(i.e., falls within the receive window). However, in response to
the security concerns raised by [Watson, 2004] and [NISCC, 2004],
[Ramaiah et al, 2008] proposec the aforementioned stricter
validity checks.
Section 11.1 of this document describes TCP-based connection-reset
attacks, along with a number of countermeasures to mitigate their
impact.
3.5.8. SYN
DISCUSSION:
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The SYN bit is used during the connection-establishment phase, to
request the synchronization of sequence numbers.
There are basically four different vulnerabilities that make use
of the SYN bit: SYN-flooding attacks, connection forgery attacks,
connection flooding attacks, and connection-reset attacks. They
are described in Section 5.1, Section 5.2, Section 5.3, and
Section 11.1.2, respectively, along with the possible
countermeasures.
3.5.9. FIN
DISCUSSION:
The FIN flag is used to signal the remote end-point the end of the
data transfer in this direction. Receipt of a valid FIN segment
(i.e., a TCP segment with the FIN flag set) causes the transition
in the connection state, as part of what is usually referred to as
the "connection termination phase".
The connection-termination phase can be exploited to perform a
number of resource-exhaustion attacks. Section 6 of this document
describes a number of attacks that exploit the connection-
termination phase along with the possible countermeasures.
3.6. Window
DISCUSSION:
The TCP Window field advertises how many bytes of data the remote
peer is allowed to send before a new advertisement is made.
Theoretically, the maximum transfer rate that can be achieved by
TCP is limited to:
Maximum Transfer Rate = Window / RTT
This means that, under ideal network conditions (e.g., no packet
loss), the TCP Window in use should be at least:
Window = 2 * Bandwidth * Delay
Using a larger Window than that resulting from the previous
equation will not provide any improvements in terms of
performance.
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In practice, selection of the most convenient Window size may also
depend on a number of other parameters, such as: packet loss rate,
loss recovery mechanisms in use, etc.
Security implications of the maximum TCP window size
An aspect of the TCP Window that is usually overlooked is the
security implications of its size. Increasing the TCP window
increases the sequence number space that will be considered
"valid" for incoming segments. Thus, use of unnecessarily large
TCP Window sizes increases TCP's vulnerability to forgery attacks
unnecessarily.
In those scenarios in which the network conditions are known
and/or can be easily predicted, it is recommended that the TCP
Window is never set to a value larger than that resulting from the
equations above. Additionally, the nature of the application
running on top of TCP should be considered when tuning the TCP
window. As an example, an H.245 signaling application certainly
does not have high requirements on throughput, and thus a window
size of around 4 KBytes will usually fulfill its needs, while
keeping TCP's resistance to off-path forgery attacks at a decent
level. Some rough measurements seem to indicate that a TCP window
of 4Kbytes is common practice for TCP connections servicing
applications such as BGP.
In principle, a possible approach to avoid requiring
administrators to manually set the TCP window would be to
implement an automatic buffer tuning mechanism, such as that
described in [Heffner, 2002]. However, as discussed in Section
7.3.2 of this document these mechanisms can be exploited to
perform other types of attacks.
Security implications arising from closed windows
The TCP window is a flow-control mechanism that prevents a fast
data sender application from overwhelming a "slow" receiver. When
a TCP end-point is not willing to receive any more data (before
some of the data that have already been received are consumed), it
will advertise a TCP window of zero bytes. This will effectively
stop the sender from sending any new data to the TCP receiver.
Transmission of new data will resume when the TCP receiver
advertises a nonzero TCP window, usually with a TCP segment that
contains no data ("an ACK").
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This segment is usually referred to as a "window update", as the
only purpose of this segment is to update the server regarding the
new window.
To accommodate those scenarios in which the ACK segment that
"opens" the window is lost, TCP implements a "persist timer" that
causes the TCP sender to query the TCP receiver periodically if
the last segment received advertised a window of zero bytes. This
probe simply consists of sending one byte of new data that will
force the TCP receiver to send an ACK segment back to the TCP
sender, containing the current TCP window. Similarly to the
retransmission timeout timer, an exponential back-off is used when
calculating the retransmission timer, so that the spacing between
probes increases exponentially.
A fundamental difference between the "persist timer" and the
retransmission timer is that there is no limit on the amount of
time during which a TCP can advertise a zero window. This means
that a TCP end-point could potentially advertise a zero window
forever, thus keeping kernel memory at the TCP sender tied to the
TCP retransmission buffer. This could clearly be exploited as a
vector for performing a Denial of Service (DoS) attack against
TCP, such as that described in Section 7.1 of this document.
Section 7.1 of this document describes a Denial of Service attack
that aims at exhausting the kernel memory used for the TCP
retransmission buffer, along with possible countermeasures.
3.7. Checksum
Middleboxes that process TCP segments MUST validate the Checksum
field, and silently discard the TCP segment if such validation fails.
DISCUSSION:
The Checksum field is an error detection mechanism meant for the
contents of the TCP segment and a number of important fields of
the IP header. It is computed over the full TCP header pre-pended
with a pseudo header that includes the IP Source Address, the IP
Destination Address, the Protocol number, and the TCP segment
length. While in principle there should not be security
implications arising from this field, due to non-RFC-compliant
implementations, the Checksum can be exploited to detect
firewalls, evade network intrusion detection systems (NIDS),
and/or perform Denial of Service attacks.
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If a stateful firewall does not check the TCP Checksum in the
segments it processes, an attacker can exploit this situation to
perform a variety of attacks. For example, he could send a flood
of TCP segments with invalid checksums, which would nevertheless
create state information at the firewall. When each of these
segments is received at its intended destination, the TCP checksum
will be found to be incorrect, and the corresponding will be
silently discarded. As these segments will not elicit a response
(e.g., an RST segment) from the intended recipients, the
corresponding connection state entries at the firewall will not be
removed. Therefore, an attacker may end up tying all the state
resources of the firewall to TCP connections that will never
complete or be terminated, probably leading to a Denial of Service
to legitimate users, or forcing the firewall to randomly drop
connection state entries.
If a NIDS does not check the Checksum of TCP segments, an attacker
may send TCP segments with an invalid checksum to cause the NIDS
to obtain a TCP data stream different from that obtained by the
system being monitored. In order to "confuse" the NIDS, the
attacker would send TCP segments with an invalid Checksum and a
Sequence Number that would overlap the sequence number space being
used for his malicious activity. FTester [Barisani, 2006] is a
tool that can be used to assess NIDS on this issue.
Finally, an attacker performing port-scanning could potentially
exploit intermediate systems that do not check the TCP Checksum to
detect whether a given TCP port is being filtered by an
intermediate firewall, or the port is actually closed by the host
being port-scanned. If a given TCP port appeared to be closed,
the attacker would then send a SYN segment with an invalid
Checksum. If this segment elicited a response (either an ICMP
error message or a TCP RST segment) to this packet, then that
response should come from a system that does not check the TCP
checksum. Since normal host implementations of the TCP protocol
do check the TCP checksum, such a response would most likely come
from a firewall or some other middle-box.
[Ed3f, 2002] describes the exploitation of the TCP checksum for
performing the above activities. [US-CERT, 2005d] provides an
example of a TCP implementation that failed to check the TCP
checksum.
3.8. Urgent pointer
Segment.Size - Data Offset * 4 > 0
If a TCP segment with the URG bit set does not pass this check, it
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MUST be silently dropped.
For TCP segments that have the URG bit set to zero, sending TCP TCP
SHOULD set the Urgent Pointer to zero.
A receiving TCP MUST ignore the Urgent Pointer field of TCP segments
for which the URG bit is zero.
DISCUSSION:
Section 3.7 of RFC 793 [Postel, 1981c] states (in page 42) that to
send an urgent indication the user must also send at least one
byte of data.
If the URG bit is zero, the Urgent Pointer is not valid, and thus
should not be processed by the receiving TCP. Nevertheless, we
recommend TCP implementations to set the Urgent Pointer to zero
when sending a TCP segment that does not have the URG bit set, and
to ignore the Urgent Pointer (as required by RFC 793) when the URG
bit is zero.
Some stacks have been known to fail to set the Urgent Pointer to
zero when the URG bit is zero, thus leaking out the corresponding
system memory contents. [Zalewski, 2008] provides further details
about this issue.
Some implementations have been found to be unable to process TCP
urgent indications correctly. [Myst, 1997] originally described
how TCP urgent indications could be exploited to perform a Denial
of Service (DoS) attack against some TCP/IP implementations,
usually leading to a system crash.
3.9. Options
[IANA, 2007] contains the official list of the assigned option
numbers. TCP Options have been specified in the past both within the
IETF and by other groups. [Hnes, 2007] contains an un-official
updated version of the IANA list of assigned option numbers. The
following table contains a summary of the assigned TCP option
numbers, which is based on [Hnes, 2007].
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+--------+----------------------+-----------------------------------+
| Kind | Meaning | Summary |
+--------+----------------------+-----------------------------------+
| 0 | End of Option List | Discussed in Section 4.1 |
+--------+----------------------+-----------------------------------+
| 1 | No-Operation | Discussed in Section 4.2 |
+--------+----------------------+-----------------------------------+
| 2 | Maximum Segment Size | Discussed in Section 4.3 |
+--------+----------------------+-----------------------------------+
| 3 | WSOPT - Window Scale | Discussed in Section 4.6 |
+--------+----------------------+-----------------------------------+
| 4 | SACK Permitted | Discussed in Section 4.4.1 |
+--------+----------------------+-----------------------------------+
| 5 | SACK | Discussed in Section 4.4.2 |
+--------+----------------------+-----------------------------------+
| 6 | Echo (obsoleted by | Obsolete. Specified in RFC 1072 |
| | option 8) | [Jacobson and Braden, 1988] |
+--------+----------------------+-----------------------------------+
| 7 | Echo Reply | Obsolete. Specified in RFC 1072 |
| | (obsoleted by option | [Jacobson and Braden, 1988] |
| | 8) | |
+--------+----------------------+-----------------------------------+
| 8 | TSOPT - Time Stamp | Discussed in Section 4.7 |
| | Option | |
+--------+----------------------+-----------------------------------+
| 9 | Partial Order | Historic. Specified in RFC 1693 |
| | Connection Permitted | [Connolly et al, 1994] |
+--------+----------------------+-----------------------------------+
| 10 | Partial Order | Historic. Specified in RFC 1693 |
| | Service Profile | [Connolly et al, 1994] |
+--------+----------------------+-----------------------------------+
| 11 | CC | Historic. Specified in RFC 1644 |
| | | [Braden, 1994] |
+--------+----------------------+-----------------------------------+
| 12 | CC.NEW | Historic. Specified in RFC 1644 |
| | | [Braden, 1994] |
+--------+----------------------+-----------------------------------+
| 13 | CC.ECHO | Historic. Specified in RFC 1644 |
| | | [Braden, 1994] |
+--------+----------------------+-----------------------------------+
| 14 | TCP Alternate | Historic. Specified in RFC 1146 |
| | Checksum Request | [Zweig and Partridge, 1990] |
+--------+----------------------+-----------------------------------+
| 15 | TCP Alternate | Historic. Specified in RFC 1145 |
| | Checksum Data | [Zweig and Partridge, 1990] |
+--------+----------------------+-----------------------------------+
| 16 | Skeeter | Historic |
+--------+----------------------+-----------------------------------+
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+--------+----------------------+-----------------------------------+
| 17 | Bubba | Historic |
+--------+----------------------+-----------------------------------+
| 18 | Trailer Checksum | Historic |
| | Option | |
+--------+----------------------+-----------------------------------+
| 19 | MD5 Signature Option | Discussed in Section 4.5 |
+--------+----------------------+-----------------------------------+
| 20 | SCPS Capabilities | Specified in [CCSDS, 2006] |
+--------+----------------------+-----------------------------------+
| 21 | Selective Negative | Specified in [CCSDS, 2006] |
| | Acknowledgements | |
+--------+----------------------+-----------------------------------+
| 22 | Record Boundaries | Specified in [CCSDS, 2006] |
+--------+----------------------+-----------------------------------+
| 23 | Corruption | Specified in [CCSDS, 2006] |
| | experienced | |
+--------+----------------------+-----------------------------------+
| 24 | SNAP | Historic |
+--------+----------------------+-----------------------------------+
| 25 | Unassigned (released | Unassigned |
| | 2000-12-18) | |
+--------+----------------------+-----------------------------------+
| 26 | TCP Compression | Historic |
| | Filter | |
+--------+----------------------+-----------------------------------+
| 27 | Quick-Start Response | Specified in RFC 4782 [Floyd et |
| | | al, 2007] |
+--------+----------------------+-----------------------------------+
| 28-252 | Unassigned | Unassigned |
+--------+----------------------+-----------------------------------+
| 253 | RFC3692-style | Described by RFC 4727 [Fenner, |
| | Experiment 1 | 2006] |
+--------+----------------------+-----------------------------------+
| 254 | RFC3692-style | Described by RFC 4727 [Fenner, |
| | Experiment 2 | 2006] |
+--------+----------------------+-----------------------------------+
Table 1: TCP Options
There are two cases for the format of a TCP option:
o Case 1: A single byte of option-kind.
o Case 2: An option-kind byte, followed by an option-length byte,
and the actual option-data bytes.
In options of the Case 2 above, the option-length byte counts the
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option-kind byte and the option-length byte, as well as the actual
option-data bytes.
All options except "End of Option List" (Kind = 0) and "No Operation"
(Kind = 1), are of "Case 2".
For options that belong to the "Case 2" described above, the
following checks MUST be performed:
option-length >= 2
option-offset + option-length <= Data Offset * 4
Where option-offset is the offset of the first byte of the option
within the TCP header, with the first byte of the TCP header being
assigned an offset of 0.
If a TCP segment fails to pass any of these checks, it SHOULD be
silently dropped.
TCP MUST ignore unknown TCP options, provided they pass the
validation checks specified above. In the same way, middle-boxes
such as packet filters SHOULD NOT reject TCP segments containing
"unknown" TCP options that pass the validation checks described
earlier in this Section.
DISCUSSION:
The value "2" in the first equation accounts for the option-kind
byte and the option-length byte, and assumes zero bytes of option-
data. This check prevents, among other things, loops in option
processing that may arise from incorrect option lengths.
The second equation takes into account the limit on the legitimate
option length imposed by the syntax of the TCP header, and is
meant to detect forged option-length values that might make an
option overlap with the TCP payload, or even go past the actual
end of the TCP segment carrying the option.
Middle-boxes such as packet filters should not reject TCP segments
containing unknown options solely because these options have not been
present in the SYN/SYN-ACK handshake.
DISCUSSION:
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There is renewed interest in defining new TCP options for purposes
like improved connection management and maintenance, advanced
congestion control schemes, and security features. The evolution
of the TCP/IP protocol suite would be severely impacted by
obstacles to deploying such new protocol mechanisms.
Middle-boxes such as packet filters SHOULD NOT reject TCP segments
containing unknown options solely because these options have not been
present in the SYN/SYN-ACK handshake.
DISCUSSION:
In the past, TCP enhancements based on TCP options regularly have
specified the exchange of a specific "enabling" option during the
initial SYN/SYN-ACK handshake. Due to the severely limited TCP
option space which has already become a concern, it should be
expected that future specifications might introduce new options
not negotiated or enabled in this way. Therefore, middle-boxes
such as packet filters should not reject TCP segments containing
unknown options solely because these options have not been present
in the SYN/SYN-ACK handshake.
TCP MUST NOT "echo" in any way unknown TCP options received in
inbound TCP segments.
DISCUSSION:
Some TCP implementations have been known to "echo" unknown TCP
options received in incoming segments. Here we stress that TCP
must not "echo" in any way unknown TCP options received in inbound
TCP segments. This is at the foundation for the introduction of
new TCP options, ensuring unambiguous behavior of systems not
supporting a new specification.
Section 4 discusses the security implications of common TCP options.
3.10. Padding
The TCP header padding is used to ensure that the TCP header ends and
data begins on a 32-bit boundary. The padding is composed of zeros.
3.11. Data
The data field contains the upper-layer packet being transmitted by
means of TCP. This payload is processed by the application process
making use of the transport services of TCP. Therefore, the security
implications of this field are out of the scope of this document.
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4. Common TCP Options
4.1. End of Option List (Kind = 0)
TCP implementations MUST be able to gracefully handle those TCP
segments in which the End of Option List should have been present,
but is missing.
DISCUSSION:
This option is used to indicate the "end of options" in those
cases in which the end of options would not coincide with the end
of the TCP header.
TCP implementations are required to ignore those options they do
not implement, and to be able to handle options with illegal
lengths. Therefore, TCP implementations should be able to
gracefully handle those TCP segments in which the End of Option
List should have been present, but is missing.
It is interesting to note that some TCP implementations do not use
the "End of Option List" option for indicating the "end of
options", but simply pad the TCP header with several "No
Operation" (Kind = 1) options to meet the header length specified
by the Data Offset header field.
4.2. No Operation (Kind = 1)
The no-operation option is basically used to allow the sending system
to align subsequent options in, for example, 32-bit boundaries.
This option does not have any known security implications.
4.3. Maximum Segment Size (Kind = 2)
The Maximum Segment Size (MSS) option is used to indicate to the
remote TCP endpoint the maximum segment size this TCP is willing to
receive.
The following check MUST be performed on a TCP segment that carries a
MSS option:
SYN == 1
If the segment does not pass this check, it MUST be silently dropped.
DISCUSSION:
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As stated in Section 3.1 of RFC 793 [Postel, 1981c], this option
can only be sent in the initial connection request (i.e., in
segments with the SYN control bit set).
TCP MUST check that the option length is 4. If the option does not
pass this check, it MUST be dropped.
The received MSS SHOULD be sanitized as follows:
Sanitized_MSS = max(MSS, 536)
This "sanitized" MSS value SHOULD be used to compute the "effective
send MSS" by the expression included in Section 4.2.2.6 of RFC 1122
[Braden, 1989], as follows:
Eff.snd.MSS = min(Sanitized_MSS+20, MMS_S) - TCPhdrsize - IPoptionsize
where:
Sanitized_MSS:
sanitized MSS value (the value received in the MSS option, with an
enforced minimum value)
MMS_S:
maximum size for a transport-layer message that TCP may send
TCPhdrsize:
size of the TCP header, which typically was 20, but may be larger
if TCP options are to be sent.
IPoptionsize
size of any IP options that TCP will pass to the IP layer with the
current message.
DISCUSSION:
The advertised maximum segment size may be the result of the
consideration of a number of factors. Firstly, if fragmentation
is employed, the size of the IP reassembly buffer may impose a
limit on the maximum TCP segment size that can be received.
Considering that the minimum IP reassembly buffer size is 576
bytes, if an MSS option is not present included in the connection-
establishment phase, an MSS of 536 bytes should be assumed.
Secondly, if Path-MTU Discovery (specified in RFC 1191 [Mogul and
Deering, 1990] and RFC 1981 [McCann et al, 1996]) is expected to
be used for the connection, an artificial maximum segment size may
be enforced by a TCP to prevent the remote peer from sending TCP
segments which would be too large to be transmitted without
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fragmentation. Finally, a system connected by a low-speed link
may choose to introduce an artificial maximum segment size to
enforce an upper limit on the network latency that would otherwise
negatively affect its interactive applications [Stevens, 1994].
The TCP specifications do not impose any requirements on the
maximum segment size value that is included in the MSS option.
However, there are a number of values that may cause undesirable
results. Firstly, an MSS of 0 could possible "freeze" the TCP
connection, as it would not allow data to be included in the
payload of the TCP segments. Secondly, low values other than 0
would degrade the performance of the TCP connection (wasting more
bandwidth in protocol headers than in actual data), and could
potentially exhaust processing cycles at the sending TCP and/or
the receiving TCP by producing an increase in the interrupt rate
caused by the transmitted (or received) packets.
The problems that might arise from low MSS values were first
described by [Reed, 2001]. However, the community did not reach
consensus on how to deal with these issues at that point.
RFC 791 [Postel, 1981a] requires IP implementations to be able to
receive IP datagrams of at least 576 bytes. Assuming an IPv4
header of 20 bytes, and a TCP header of 20 bytes, there should be
room in each IP packet for 536 application data bytes.
There are two cases to analyze when considering the possible
interoperability impact of sanitizing the received MSS value: TCP
connections relying on IP fragmentation and TCP connections
implementing Path-MTU Discovery. In case the corresponding TCP
connection relies on IP fragmentation, given that the minimum
reassembly buffer size is required to be 576 bytes by RFC 791
[Postel, 1981a], the adoption of 536 bytes as a lower limit is
safe.
In case the TCP connection relies on Path-MTU Discovery, imposing
a lower limit on the adopted MSS may ignore the advice of the
remote TCP on the maximum segment size that can possibly be
transmitted without fragmentation. As a result, this could lead
to the first TCP data segment to be larger than the Path-MTU.
However, in such a scenario, the TCP segment should elicit an ICMP
Unreachable "fragmentation needed and DF bit set" error message
that would cause the "effective send MSS" (E_MSS) to be decreased
appropriately. Thus, imposing a lower limit on the accepted MSS
will not cause any interoperability problems.
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A possible scenario exists in which the proposed enforcement of a
lower limit in the received MSS might lead to an interoperability
problem. If a system was attached to the network by means of a
link with an MTU of less than 576 bytes, and there was some
intermediate system which either silently dropped (i.e., without
sending an ICMP error message) those packets equal to or larger
than that 576 bytes, or some intermediate system simply filtered
ICMP "fragmentation needed and DF bit set" error messages, the
proposed behavior would not lead to an interoperability problem,
when communication could have otherwise succeeded. However, the
interoperability problem would really be introduced by the network
setup (e.g., the middle-box silently dropping packets), rather
than by the mechanism proposed in this section. In any case, TCP
should nevertheless implement a mechanism such as that specified
by RFC 4821 [Mathis and Heffner, 2007] to deal with this type of
"network black-holes".
4.4. Selective Acknowledgement Option
The Selective Acknowledgement option provides an extension to allow
the acknowledgement of individual segments, to enhance TCP's loss
recovery.
Two options are involved in the SACK mechanism. The "Sack-permitted
option" is sent during the connections-establishment phase, to
advertise that SACK is supported. If both TCP peers agree to use
selective acknowledgements, the actual selective acknowledgements are
sent, if needed, by means of "SACK options".
4.4.1. SACK-permitted Option (Kind = 4)
The SACK-permitted option is meant to advertise that the TCP sending
this segment supports Selective Acknowledgements.
The following check MUST be performed on a TCP segment that carries a
MSS option:
SYN == 1
If a segment does not pass this check, it MUST be silently dropped.
DISCUSSION:
The SACK-permitted option can be sent only in SYN segments.
TCP MUST check that the option length is 2. If the option does not
pass this check it MUST be silently dropped.
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4.4.2. SACK Option (Kind = 5)
The SACK option is used to convey extended acknowledgment information
from the receiver to the sender over an established TCP connection.
The option consists of an option-kind byte (which must be 5), an
option-length byte, and a variable number of SACK blocks.
TCP MUST silently discard those TCP segments carrying a SACK option
that does not pass the following check:
option-offset + option-length <= Data Offset * 4
TCP MUST silently discard those TCP segments carrying a SACK option
that does not pass the following check:
option-length >= 10
DISCUSSION:
A SACK Option with zero SACK blocks is nonsensical. The value
"10" accounts for the option-kind byte, the option-length byte, a
4-byte left-edge field, and a 4-byte right-edge field.
TCP MUST silently discard those TCP segments carrying a SACK option
that does not pass the following check:
(option-length - 2) % 8 == 0
DISCUSSION:
As stated in Section 3 of RFC 2018 [Mathis et al, 1996], a SACK
option that specifies n blocks will have a length of 8*n+2.
TCP MUST silently discard those TCP segments carrying a SACK option
that contains a SACK block that does not pass the following check:
Left Edge of Block < Right Edge of Block
As in all the other occurrences in this document, all comparisons
between sequence numbers should be performed using sequence number
arithmetic.
DISCUSSION:
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Each block included in a SACK option represents a number of
received data bytes that are contiguous and isolated; that is, the
bytes just below the block, (Left Edge of Block - 1), and just
above the block, (Right Edge of Block), have not yet been
received.
TCP MUST enforce a limit on the number of SACK blocks that a TCP will
store in memory for each connection at any time.
DISCUSSION:
The TCP receiving a SACK option is expected to keep track of the
selectively-acknowledged blocks. Even when space in the TCP
header is limited (and thus each TCP segment can selectively-
acknowledge at most four blocks of data), an attacker could try to
perform a buffer overflow or a resource-exhaustion attack by
sending a large number of SACK options.
For example, an attacker could send a large number of SACK
options, each of them acknowledging one byte of data.
Additionally, for the purpose of wasting resources on the attacked
system, each of these blocks would be separated from each other by
one byte, to prevent the attacked system from coalescing two (or
more) contiguous SACK blocks into a single SACK block. If the
attacked system kept track of each SACKed block by storing both
the Left Edge and the Right Edge of the block, then for each
window of data, the attacker could waste up to 4 * Window bytes of
memory at the attacked TCP.
The value "4 * Window" results from the expression "(Window / 2) *
8", in which the value "2" accounts for the 1-byte block
selectively-acknowledged by each SACK block and 1 byte that would
be used to separate each SACK blocks from each other, and the
value "8" accounts for the 8 bytes needed to store the Left Edge
and the Right Edge of each SACKed block.
Therefore, it is clear that a limit should be imposed on the
number of SACK blocks that a TCP will store in memory for each
connection at any time. Measurements in [Dharmapurikar and
Paxson, 2005] indicate that in the vast majority of cases
connections have a single hole in the data stream at any given
time. Thus, a limit of 16 SACK blocks for each connection would
handle even most of the more unusual cases in which there is more
than one simultaneous hole at a time.
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4.5. MD5 Option (Kind=19)
The TCP MD5 option provides a mechanism for authenticating TCP
segments with a 18-byte digest produced by the MD5 algorithm. The
option consists of an option-kind byte (which must be 19), an option-
length byte (which must be 18), and a 16-byte MD5 digest.
TCP MUST silently drop a TCP segment that carries a TCP MD5 option
that does not pass the following checks:
option-offset + option-length <= Data Offset * 4
option-length == 18
DISCUSSION:
The TCP MD5 option is of "Case 2", and has a fixed length.
DISCUSSION:
A basic weakness on the TCP MD5 option is that the MD5 algorithm
itself has been known (for a long time) to be vulnerable to
collision search attacks.
[Bellovin, 2006] argues that it has two other weaknesses, namely
that it does not provide a key identifier, and that it has no
provision for automated key management. However, it is generally
accepted that while a Key-ID field can be a good approach for
providing smooth key rollover, it is not actually a requirement.
For instance, most systems implementing the TCP MD5 option include
a "keychain" mechanism that fully supports smooth key rollover.
Additionally, with some further work, ISAKMP/IKE could be used to
configure the MD5 keys.
It is interesting to note that while the TCP MD5 option, as
specified by RFC 2385 [Heffernan, 1998], addresses the TCP-based
forgery attacks against TCP discussed in Section 11, it does not
address the ICMP-based connection-reset attacks discussed in
Section 15. As a result, while a TCP connection may be protected
from TCP-based forgery attacks by means of the MD5 option, an
attacker might still be able to successfully perform the ICMP-
based counter-part.
The TCP MD5 option has been obsoleted by the TCP-AO.
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4.6. Window scale option (Kind = 3)
The window scale option provides a mechanism to expand the definition
of the TCP window to 32 bits, such that the performance of TCP can be
improved in some network scenarios. The Window scale option consists
of an option-kind byte (which must be 3), followed by an option-
length byte (which must be 3), and a shift count (shift.cnt) byte
(the actual option-data).
The option may be sent only in the initial SYN segment, but may also
be sent in a SYN/ACK segment if the option was received in the
initial SYN segment. If the option is received in any other segment,
it MUST be silently dropped.
TCP MUST silently discard TCP segments that contain a Window scale
option whose option-length is not 3.
DISCUSSION:
This option has a fixed length.
TCP MUST silently discard TCP segments that contain a Window scale
option that does not pass the following check:
shift.cnt <= 14
DISCUSSION:
As discussed in Section 2.3 of RFC 1323 [Jacobson et al, 1992], in
order to prevent new data from being mistakenly considered as old
and vice versa, the resulting window should be equal to or smaller
than 2^32.
DISCUSSION:
[Welzl, 2008] describes major problems with the use of the Window
scale option in the Internet due to faulty equipment.
While there are not known security implications arising from the
window scale mechanism itself, the size of the TCP window has a
number of security implications. In general, larger window sizes
increase the chances of an attacker from successfully performing
forgery attacks against TCP, such as those described in Section 11
of this document. Additionally, large windows can exacerbate the
impact of resource exhaustion attacks such as those described in
Section 7 of this document.
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Section 3.7 provides a general discussion of the security
implications of the TCP window size. Section 7.3.2 discusses the
security implications of Automatic receive-buffer tuning
mechanisms.
4.7. Timestamps option (Kind = 8)
The Timestamps option, specified in RFC 1323 [Jacobson et al, 1992],
is used to perform two functions: Round-Trip Time Measurement (RTTM),
and Protection Against Wrapped Sequence Numbers (PAWS).
TCP MUST silently discard TCP segments that contain a Timestamps
option that does not pass the following check:
option-length == 10
DISCUSSION:
As specified by RFC 1323, the option-length must be 10.
4.7.1. Generation of timestamps
TCP SHOULD generate timestamps with the following expression:
timestamp = T() + F(localhost, localport, remotehost, remoteport, secret_key)
where the result of T() is a global system clock that complies with
the requirements of Section 4.2.2 of RFC 1323 [Jacobson et al, 1992],
and F() is a function that should not be computable from the outside.
Therefore, we suggest F() to be a cryptographic hash function of the
connection-id and some secret data.
DISCUSSION:
For the purpose of PAWS, the timestamps sent on a connection are
required to be monotonically increasing. While there is no
requirement that timestamps are monotonically increasing across
TCP connections, the generation of timestamps such that they are
monotonically increasing across connections between the same two
endpoints allows the use of timestamps for improving the handling
of SYN segments that are received while the corresponding four-
tuple is in the TIME-WAIT state. This is discussed in Section
11.1.2 of this document.
F() provides an offset that will be the same for all incarnations
of a connection between the same two endpoints, while T() provides
the monotonically increasing values that are needed for PAWS.
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Further discussion about this algorithm is available in
[I-D.gont-timestamps-generation].
TCP SHOULD NOT initialize a global timestamp counter to a fixed value
when the system is bootstrapped.
DISCUSSION:
Some implementations are known to initialize their global
timestamp clock to zero when the system is bootstrapped. This is
undesirable, as the timestamp clock would disclose the system
uptime.
TCP SHOULD set the Timestamp Echo Reply (TSecr) field to zero when
sending a TCP segment that does not have the ACK bit set (i.e., a SYN
segment).
DISCUSSION:
Some TCP implementations have been found to fail to set the
Timestamp Echo Reply field (TSecr) to zero in TCP segments that do
not have the ACK bit set, thus potentially leaking information.
4.7.2. Vulnerabilities
Blind In-Window Attacks
Segments that contain a timestamp option smaller than the last
timestamp option recorded by TCP are silently dropped. This allows
for a subtle attack against TCP that would allow an attacker to cause
one direction of data transfer of the attacked connection to freeze
[US-CERT, 2005c]. An attacker could forge a TCP segment that
contains a timestamp that is much larger than the last timestamp
recorded for that direction of the data transfer of the connection.
The offending segment would cause the recorded timestamp (TS.Recent)
to be updated and, as a result, subsequent segments sent by the
impersonated TCP peer would be simply dropped by the receiving TCP.
This vulnerability has been documented in [US-CERT, 2005d]. However,
it is worth noting that exploitation of this vulnerability requires
an attacker to guess (or know) the four-tuple {IP Source Address, IP
Destination Address, TCP Source Port, TCP Destination Port}, as well
a valid Sequence Number and a valid Acknowledgement Number. If an
attacker has such detailed knowledge about a TCP connection, unless
TCP segments are protected by proper authentication mechanisms (such
as IPsec [Kent and Seo, 2005]), he can perform a variety of attacks
against the TCP connection, even more devastating than the one just
described.
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Information leaking
Some implementations are known to maintain a global timestamp clock,
which is used for all connections. This is undesirable, as an
attacker that can establish a connection with a host would learn the
timestamp used for all the other connections maintained by that host,
which could be useful for performing any attacks that require the
attacker to forge TCP segments. A timestamps generator such as the
one recommended in Section 4.7.1 of this document would prevent this
information leakage, as it separates the "timestamps space" among the
different TCP connections.
Some implementations are known to initialize their global timestamp
clock to zero when the system is bootstrapped. This is undesirable,
as the timestamp clock would disclose the system uptime. A
timestamps generator such as the one recommended in Section 4.7.1 of
this document would prevent this information leakage, as the function
F() introduces an "offset" that does not disclose the system uptime.
As discussed in Section 3.2 of RFC 1323 [Jacobson et al, 1992], the
Timestamp Echo Reply field (TSecr) is only valid if the ACK bit of
the TCP header is set, and its value must be zero when it is not
valid. However, some TCP implementations have been found to fail to
set the Timestamp Echo Reply field (TSecr) to zero in TCP segments
that do not have the ACK bit set, thus potentially leaking
information. We stress that TCP implementations should comply with
RFC 1323 by setting the Timestamp Echo Reply field (TSecr) to zero in
those TCP segments that do not have the ACK bit set, thus eliminating
this potential information leakage.
Finally, it should be noted that the Timestamps option can be
exploited to count the number of systems behind NATs (Network Address
Translators) [Srisuresh and Egevang, 2001]. An attacker could count
the number of systems behind a NAT by establishing a number of TCP
connections (using the public address of the NAT) and indentifying
the number of different timestamp sequences. This information
leakage could be eliminated by rewriting the contents of the
Timestamps option at the NAT. [Gont and Srisuresh, 2008] provides a
detailed discussion of the security implications of NATs, and
proposes mitigations for this and other issues.
5. Connection-establishment mechanism
The following subsections describe a number of attacks that can be
performed against TCP by exploiting its connection-establishment
mechanism.
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5.1. SYN flood
TCP SHOULD implement (and enable by default) a syn-cache [Lemon,
2002].
TCP SHOULD implement syn-cookies, and SHOULD enable them only after a
specified number of TCBs has been allocated for connections in the
SYN-RECEIVED state.
DISCUSSION:
TCP uses a mechanism known as the "three-way handshake" for the
establishment of a connection between two TCP peers. RFC 793
[Postel, 1981c] states that when a TCP that is in the LISTEN state
receives a SYN segment (i.e., a TCP segment with the SYN flag
set), it must transition to the SYN-RECEIVED state, record the
control information (e.g., the ISN) contained in the SYN segment
in a Transmission Control Block (TCB), and respond with a SYN/ACK
segment.
A Transmission Control Block is the data structure used to store
(usually within the kernel) all the information relevant to a TCP
connection. The concept of "TCB" is introduced in the core TCP
specification RFC 793 [Postel, 1981c].
In practice, virtually all existing implementations do not modify
the state of the TCP that was in the LISTEN state, but rather
create a new TCP (i.e., a new "protocol machine"), and perform all
the state transitions on this newly-created TCP. This allows the
application running on top of TCP to service to more than one
client at the same time. As a result, each connection request
results in the allocation of system memory to store the TCB
associated with the newly created TCB.
If TCP was implemented strictly as described in RFC 793, the
application running on top of TCP would have to finish servicing
the current client before being able to service the next one in
line, or should instead be able to perform some kind of connection
hand-off.
An attacker could exploit TCP's connection-establishment mechanism
to perform a Denial of Service (DoS) attack, by sending a large
number of connection requests to the target system, with the
intent of exhausting the system memory destined for storing TCBs
(or related kernel data structures), thus preventing the attacked
system from establishing new connections with legitimate users.
This attack is widely known as "SYN flood", and has received a lot
of attention during the late 90's [CERT, 1996].
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Given that the attacker does not need to complete the three-way
handshake for the attacked system to tie system resources to the
newly created TCBs, he will typically forge the source IP address
of the malicious SYN segments he sends, thus concealing his own IP
address.
If the forged IP addresses corresponded to some reachable system,
the impersonated system would receive the SYN/ACK segment sent by
the attacked host (in response to the forged SYN segment), which
would elicit an RST segment. This RST segment would be delivered
to the attacked system, causing the corresponding connection to be
aborted, and the corresponding TCB to be removed.
As the impersonated host would not have any state information for
the TCP connection being referred to by the SYN/ACK segment, it
would respond with a RST segment, as specified by the TCP segment
processing rules of RFC 793 [Postel, 1981c].
However, if the forged IP source addresses were unreachable, the
attacked TCP would continue retransmitting the SYN/ACK segment
corresponding to each connection request, until timing out and
aborting the connection. For this reason, a number of widely
available attack tools first check whether each of the (forged) IP
addresses are reachable by sending an ICMP echo request to them.
The receipt of an ICMP echo response is considered an indication
of the IP address being reachable (and thus results in the
corresponding IP address not being used for performing the
attack), while the receipt of an ICMP unreachable error message is
considered an indication of the IP address being unreachable (and
thus results in the corresponding IP address being used for
performing the attack).
[Gont, 2008b] describes how the so-called ICMP soft errors could
be used by TCP to abort connections in any of the non-synchronized
states. While implementation of the mechanism described in that
document would certainly not eliminate the vulnerability of TCP to
SYN flood attacks (as the attacker could use addresses that are
simply "black-holed"), it provides an example of how signaling
information such as that provided by means of ICMP error messages
can provide valuable information that a transport protocol could
use to perform heuristics.
In order to mitigate the impact of this attack, the amount of
information stored for non-established connections should be
reduced (ideally, non-synchronized connections should not require
any state information to be maintained at the TCP performing the
passive OPEN). There are basically two mitigation techniques for
this vulnerability: a syn-cache and syn-cookies.
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[Borman, 1997] and RFC 4987 [Eddy, 2007] contain a general
discussion of SYN-flooding attacks and common mitigation
approaches.
The syn-cache [Lemon, 2002] approach aims at reducing the amount
of state information that is maintained for connections in the
SYN-RECEIVED state, and allocates a full TCB only after the
connection has transited to the ESTABLISHED state.
The syn-cookie [Bernstein, 1996] approach aims at completely
eliminating the need to maintain state information at the TCP
performing the passive OPEN, by encoding the most elementary
information required to complete the three-way handshake in the
Sequence Number of the SYN/ACK segment that is sent in response to
the received SYN segment. Thus, TCP is relieved from keeping
state for connections in the SYN-RECEIVED state.
The syn-cookie approach has a number of drawbacks:
* Firstly, given the limited space in the Sequence Number field,
it is not possible to encode all the information included in
the initial segment, such as, for example, support of Selective
Acknowledgements (SACK).
* Secondly, in the event that the Acknowledgement segment sent in
response to the SYN/ACK sent by the TCP that performed the
passive OPEN (i.e., the TCP server) were lost, the connection
would end up in the ESTABLISHED state on the client-side, but
in the CLOSED state on the server side. This scenario is
normally handled in TCP by having the TCP server retransmit its
SYN/ACK. However, if syn-cookies are enabled, there would be
no connection state information on the server side, and thus
the SYN/ACK would never be retransmitted. This could lead to a
scenario in which the connection could remain in the
ESTABLISHED state on the client side, but in the CLOSED state
at the server side, indefinitely. If the application protocol
was such that it required the client to wait for some data from
the server (e.g., a greeting message) before sending any data
to the server, a deadlock would take place, with the client
application waiting for such server data, and the server
waiting for the TCP three-way handshake to complete.
* Thirdly, unless the function used to encode information in the
SYN/ACK packet is cryptographically strong, an attacker could
forge TCP connections in the ESTABLISHED state by forging ACK
segments that would be considered as "legitimate" by the
receiving TCP.
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* Fourthly, in those scenarios in which establishment of new
connections is blocked by simply dropping segments with the SYN
bit set, use of SYN cookies could allow an attacker to bypass
the firewall rules, as a connection could be established by
forging an ACK segment with the correct values, without the
need of setting the SYN bit.
As a result, syn-cookies are usually not employed as a first line
of defense against SYN-flood attacks, but are only as the last
resort to cope with them. For example, some TCP implementations
enable syn-cookies only after a certain number of TCBs has been
allocated for connections in the SYN-RECEIVED state. We recommend
this implementation technique, with a syn-cache enabled by
default, and use of syn-cookies triggered, for example, when the
limit of TCBs for non-synchronized connections with a given port
number has been reached.
It is interesting to note that a SYN-flood attack should only
affect the establishment of new connections. A number of books
and online documents seem to assume that TCP will not be able to
respond to any TCP segment that is meant for a TCP port that is
being SYN-flooded (e.g., respond with an RST segment upon receipt
of a TCP segment that refers to a non-existent TCP connection).
While SYN-flooding attacks have been successfully exploited in the
past for achieving such a goal [Shimomura, 1995], as clarified by
RFC 1948 [Bellovin, 1996] the effectiveness of SYN flood attacks
to silence a TCP implementation arose as a result of a bug in the
4.4BSD TCP implementation [Wright and Stevens, 1994], rather than
from a theoretical property of SYN-flood attacks themselves.
Therefore, those TCP implementations that do not suffer from such
a bug should not be silenced as a result of a SYN-flood attack.
[Zquete, 2002] describes a mechanism that could theoretically
improve the functionality of SYN cookies. It exploits the TCP
"simultaneous open" mechanism, as illustrated in Figure 5.
See Figure 5, in page 46 of the UK CPNI document.
Use of TCP simultaneous open for handling SYN floods
In line 1, TCP A initiates the connection-establishment phase by
sending a SYN segment to TCP B. In line 2, TCP B creates a SYN
cookie as described by [Bernstein, 1996], but does not set the ACK
bit of the segment it sends (thus really sending a SYN segment,
rather than a SYN/ACK). This "fools" TCP A into thinking that
both SYN segments "have crossed each other in the network" as if a
"simultaneous open" scenario had taken place. As a result, in
line 3 TCP A sends a SYN/ACK segment containing the same options
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that were contained in the original SYN segment. In line 4, upon
receipt of this segment, TCP processes the cookie encoded in the
ACK field as if it had been the result of a traditional SYN cookie
scenario, and moves the connection into the ESTABLISHED state. In
line 5, TCP B sends a SYN/ACK segment, which causes the connection
at TCP A to move into the ESTABLISHED state. In line 6, TCP A
sends a data segment on the connection.
While this mechanism would work in theory, unfortunately there are
a number of factors that prevent it from being usable in real
network environments:
* Some systems are not able to perform the "simultaneous open"
operation specified in RFC 793, and thus the connection
establishment will fail.
* Some firewalls might prevent the establishment of TCP
connections that rely on the "simultaneous open" mechanism
(e.g., a given firewall might be allowing incoming SYN/ACK
segments, but not outgoing SYN/ACK segments).
Therefore, we do not recommend implementation of this mechanism
for mitigating SYN-flood attacks.
5.2. Connection forgery
The process of causing a TCP connection to be illegitimately
established between two arbitrary remote peers is usually referred to
as "connection spoofing" or "connection forgery". This can have a
great negative impact when systems establish some sort of trust
relationships based on the IP addresses used to establish a TCP
connection [daemon9 et al, 1996].
It should be stressed that hosts should not establish trust
relationships based on the IP addresses [CPNI, 2008] or on the TCP
ports in use for the TCP connection (see Section 3.1 and Section 3.2
of this document).
One of the underlying weaknesses that allow this vulnerability to be
more easily exploited is the use of an inadequate Initial Sequence
Number (ISN) generator, as explained back in the 80's in [Morris,
1985]. As discussed in Section 3.3.1 of this document, any TCP
implementation that makes use of an inadequate ISN generator will be
more vulnerable to this type of attack. A discussion of approaches
for a more careful generation of Initial Sequence Numbers (ISNs) can
be found in Section 3.3.1 of this document.
Another attack vector for performing connection-forgery attacks is
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the use of IP source routing. By forging the Source Address of the
IP packets that encapsulate the TCP segments of a connection, and
carefully crafting an IP source route option (i.e., either LSSR or
SSRR) that includes a system whose traffic he can monitor, an
attacker could cause the packets sent by the attacked system (e.g.,
the SYN/ACK segment sent in response to the attacker's SYN segment)
to be illegitimately directed to him [CPNI, 2008]. Thus, the
attacker would not even need to guess valid sequence numbers for
forging a TCP connection, as he would simply have direct access to
all this information. As discussed in [CPNI, 2008], it is strongly
recommended that systems disable IP Source Routing by default, or at
the very least, they disable source routing for IP packets that
encapsulate TCP segments.
The IPv6 Routing Header Type 0, which provides a similar
functionality to that provided by IPv4 source routing, has been
officially deprecated by RFC 5095 [Abley et al, 2007].
5.3. Connection-flooding attack
5.3.1. Vulnerability
The creation and maintenance of a TCP connection requires system
memory to maintain shared state between the local and the remote TCP.
As system memory is a finite resource, there is a limit on the number
of TCP connections that a system can maintain at any time. When the
TCP API is employed to create a TCP connection with a remote peer, it
allocates system memory for maintaining shared state with the remote
TCP peer, and thus the resulting connection would tie a similar
amount of resources at the remote host as at the local host.
However, if special packet-crafting tools are employed to forge TCP
segments to establish TCP connections with a remote peer, the local
kernel implementation of TCP can be bypassed, and the allocation of
resources on the attacker's system for maintaining shared state can
be avoided. Thus, a malicious user could create a large number of
TCP connections, and subsequently abandon them, thus tying system
resources only at the remote peer. This allows an attacker to create
a large number of TCP connections at the attacked system with the
intent of exhausting its kernel memory, without exhausting the
attacker's own resources. [CERT, 2000] discusses this vulnerability,
which is usually referred to as the "Naptha attack".
This attack is similar in nature to the "Netkill" attack discussed in
Section 7.1.1. However, while Netkill ties both TCBs and TCP send
buffers to the abandoned connections, Naptha only ties TCBs (and
related kernel structures), as it doesn't issue any application
requests.
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The symptom of this attack is an extremely large number of TCP
connections in the ESTABLISHED state, which would tend to exhaust
system resources and deny service to new clients (or possibly cause
the system to crash).
It should be noted that it is possible for an attacker to perform the
same type of attack causing the abandoned connections to remain in
states other than ESTABLISHED. This might be interesting for an
attacker, as it is usually the case that connections in states other
than ESTABLISHED usually have no controlling user-space process (that
is, the former controlling process for the connection has already
closed the corresponding file descriptor).
A particularly interesting case of a connection-flooding attack that
aims at abandoning connections in a state other than ESTABLISHED is
discussed in Section 6.1 of this document.
5.3.2. Countermeasures
As with many other resource exhaustion attacks, the problem in
generating countermeasures for this attack is that it may be
difficult to differentiate between an actual attack and a legitimate
high-load scenario. However, there are a number of countermeasures
which, when tuned for each particular network environment, could
allow a system to resist this attack and continue servicing
legitimate clients.
Hosts SHOULD enforce limits on the number of TCP connections with no
user-space controlling process.
DISCUSSION:
Connections in states other than ESTABLISHED usually have no user-
space controlling process. This prevents the application making
use of those connections from enforcing limits on the maximum
number of ongoing connections (either on a global basis or a
per-IP address basis). When resource exhaustion is imminent or
some threshold of ongoing connections is reached, the operating
system should consider freeing system resources by aborting
connections that have no user-space controlling process. A number
of such connections could be aborted on a random basis, or based
on some heuristics performed by the operating system (e.g., first
abort connections with peers that have the largest number of
ongoing connections with no user-space controlling process).
Hosts SHOULD enforce per-process and per-user limits on maximum
kernel memory that can be used at any time.
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Hosts SHOULD enforce per-process and per-user limits on the number of
existent TCP connections at any time.
DISCUSSION:
While the Naphta attack is usually targeted at a service such as
HTTP, its impact is usually system-wide. This is particularly
undesirable, as an attack against a single service might affect
the system as a whole (for example, possibly precluding remote
system administration).
In order to avoid an attack to a single service from affecting
other services, we advise TCP implementations to enforce per-
process and per-user limits on maximum kernel memory that can be
used at any time. Additionally, we recommend implementations to
enforce per-process and per-user limits on the number of existent
TCP connections at any time.
Applications SHOULD enforce limits on the number of simultaneous
connections that can be established from a single IP address or
network prefix at any given time.
DISCUSSION:
An application could limit the number of simultaneous connections
that can be established from a single IP address or network prefix
at any given time. Once that limit has been reached, some other
connection from the same IP address or network prefix would be
aborted, thus allowing the application to service this new
incoming connection.
There are a number of factors that should be taken into account
when defining the specific limit to enforce. For example, in the
case of protocols that have an authentication phase (e.g., SSH,
POP3, etc.), this limit could be applied to sessions that have not
yet been authenticated. Additionally, depending on the nature and
use of the application, it might or might not be normal for a
single system to have multiple connections to the same server at
the same time.
For many network services, the limit of maximum simultaneous
connections could be kept very low. For example, an SMTP server
could limit the number of simultaneous connections from a single
IP address to 10 or 20 connections.
While this limit could work in many network scenarios, we
recommend network operators to measure the maximum number of
concurrent connections from a single IP address during normal
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operation, and set the limit accordingly.
In the case of web servers, this limit will usually need to be set
much higher, as it is common practice for web clients to establish
multiple simultaneous connections with a single web server to
speed up the process of loading a web page (e.g., multiple graphic
files can be downloaded simultaneously using separate TCP
connections).
NATs (Network Address Translators) [Srisuresh and Egevang, 2001]
are widely deployed in the Internet, and may exacerbate this
situation, as a large number of clients behind a NAT might each
establish multiple TCP connections with a given web server, which
would all appear to be originate from the same IP address (that of
the NAT box).
Firewalls MAY enforce limits on the number of simultaneous
connections that can be established from a single IP address or
network prefix at any given time.
DISCUSSION:
Some firewalls can be configured to limit the number of
simultaneous connections that any system can maintain with a
specific system and/or service at any given time. Limiting the
number of simultaneous connections that each system can establish
with a specific system and service would effectively limit the
possibility of an attacker that controls a single IP address to
exhaust system resources at the attacker system/service.
5.4. Firewall-bypassing techniques
TCP MUST silently drop those TCP segments that have both the SYN and
the RST flags set.
DISCUSSION:
Some firewalls block incoming TCP connections by blocking only
incoming SYN segments. However, there are inconsistencies in how
different TCP implementations handle SYN segments that have
additional flags set, which may allow an attacker to bypass
firewall rules [US-CERT, 2003b].
For example, some firewalls have been known to mistakenly allow
incoming SYN segments if they also have the RST bit set. As some
TCP implementations will create a new connection in response to a
TCP segment with both the SYN and RST bits set, an attacker could
bypass the firewall rules and establish a connection with a
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"protected" system by setting the RST bit in his SYN segments.
Here we advise TCP implementations to silently drop those TCP
segments that have both the SYN and the RST flags set.
6. Connection-termination mechanism
6.1. FIN-WAIT-2 flooding attack
6.1.1. Vulnerability
TCP implements a connection-termination mechanism that is employed
for the graceful termination of a TCP connection. This mechanism
usually consists of the exchange of four-segments. Figure 6
illustrates the usual segment exchange for this mechanism.
Figure 6: TCP connection-termination mechanism
See Figure 6, in page 50 of the UK CPNI document.
TCP connection-termination mechanism
A potential problem may arise as a result of the FIN-WAIT-2 state:
there is no limit on the amount of time that a TCP can remain in the
FIN-WAIT-2 state. Furthermore, no segment exchange is required to
maintain the connection in that state.
As a result, an attacker could establish a large number of
connections with the target system, and cause it close each of them.
For each connection, once the target system has sent its FIN segment,
the attacker would acknowledge the receipt of this segment, but would
send no further segments on that connection. As a result, an
attacker could cause the corresponding system resources (e.g., the
system memory used for storing the TCB) without the need to send any
further packets.
While the CLOSE command described in RFC 793 [Postel, 1981c] simply
signals the remote TCP end-point that this TCP has finished sending
data (i.e., it closes only one direction of the data transfer), the
close() system-call available in most operating systems has different
semantics: it marks the corresponding file descriptor as closed (and
thus it is no longer usable), and assigns the operating system the
responsibility to deliver any queued data to the remote TCP peer and
to terminate the TCP connection. This makes the FIN-WAIT-2 state
particularly attractive for performing memory exhaustion attacks, as
even if the application running on top of TCP were imposing limits on
the maximum number of ongoing connections, and/or time limits on the
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function calls performed on TCP connections, that application would
be unable to enforce these limits on the FIN-WAIT-2 state.
6.1.2. Countermeasures
A number of countermeasures can be implemented to mitigate FIN-WAIT-2
flooding attacks. Some of these countermeasures require changes in
the TCP implementations, while others require changes in the
applications running on top of TCP.
TCP SHOULD enforce limits on the duration of the FIN-WAIT-2 state.
DISCUSSION:
In order to avoid the risk of having connections stuck in the FIN-
WAIT-2 state indefinitely, a number of systems incorporate a
timeout for the FIN-WAIT-2 state. For example, the Linux kernel
version 2.4 enforces a timeout of 60 seconds [Linux, 2008]. If
the connection-termination mechanism does not complete before that
timeout value, it is aborted.
Enabling applications to enforce limits on ongoing connections
As discussed in Section 6.1.1, the fact that the close() system call
marks the corresponding file descriptor as closed prevents the
application running on top of TCP from enforcing limits on the
corresponding connection.
While it is common practice for applications to terminate their
connections by means of the close() system call, it is possible for
an application to initiate the connection-termination phase without
closing the corresponding file descriptor (hence keeping control of
the connection).
In order to achieve this, an application performing an active close
(i.e., initiating the connection-termination phase) should replace
the system-call close(sockfd) with the following code sequence:
o A call to shutdown(sockfd, SHUT_WR), to close the sending
direction of this connection
o Successive calls to read(), until it returns 0, thus indicating
that the remote TCP peer has finished sending data.
o A call to setsockopt(sockfd, SOL_SOCKET, SO_LINGER, &l,
sizeof(l)), where l is of type struct linger (with its members
l.l_onoff=1 and l.l_linger=90).
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o A call to close(sockfd), to close the corresponding file
descriptor.
The call to shutdown() (instead of close()) allows the application to
retain control of the underlying TCP connection while the connection
transitions through the FIN-WAIT-1 and FIN-WAIT-2 states. However,
the application will not retain control of the connection while it
transitions through the CLOSING and TIME-WAIT states.
It should be noted that, strictly speaking, close(sockfd) decrements
the reference count for the descriptor sockfd, and initiates the
connection termination phase only when the reference count reaches 0.
On the other hand, shutdown(sockfd, SHUT_WR) initiates the
connection-termination phase, regardless of the reference count for
the sockfd descriptor. This should be taken into account when
performing the code replacement described above. For example, it
would be a bug for two processes (e.g., parent and child) that share
a descriptor to both call shutdown(sockfd, SHUT_WR).
An application performing a passive close should replace the call to
close(sockfd) with the following code sequence:
o A call to setsockopt(sockfd, SOL_SOCKET, SO_LINGER, &l,
sizeof(l)), where l is of type struct linger (with its members
l.l_onoff=1 and l.l_linger=90).
o A call to close(sockfd), to close the corresponding file
descriptor.
It is assumed that if the application is performing a passive close,
the application already detected that the remote TCP peer finished
sending data by means as a result of a call to read() returning 0.
In this scenario, the application will not retain control of the
underlying connection when it transitions through the LAST_ACK state.
Enforcing limits on the number of connections with no user-space
controlling process
The considerations and recommendations in Section 5.3.2 for enforcing
limits on the number of connections with no user-space controlling
process are applicable to mitigate this vulnerability.
Limiting the number of simultaneous connections at the application
The considerations and recommendations in Section 5.3.2 for limiting
the number of simultaneous connections at the application are to
mitigate this vulnerability. We note, however, that unless
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applications are implemented to retain control of the underlying TCP
connection while the connection transitions through the FIN-WAIT-1
and FIN-WAIT-2 states, enforcing such limits may prove to be a
difficult task.
Limiting the number of simultaneous connections at firewalls
The considerations and recommendations in Section 5.3.2 for enforcing
limiting the number of simultaneous connections at firewalls are
applicable to mitigate this vulnerability.
7. Buffer management
7.1. TCP retransmission buffer
7.1.1. Vulnerability
[Shalunov, 2000] describes a resource exhaustion attack (Netkill)
that can be performed against TCP. The attack aims at exhausting
system memory by creating a large number of TCP connections which are
then abandoned. The attack is usually performed as follows:
o The attacker creates a TCP connection to a service in which a
small client request can result in a large server response (e.g.,
HTTP). Rather than relying on his kernel implementation of TCP,
the attacker creates his TCP connections by means of a specialized
packet-crafting tool. This allows the attacker to create the TCP
connections and later abandon them, exhausting the resources at
the attacked system, while not tying his own system resources to
the abandoned connections.
o When the connection is established (i.e., the three-way handshake
has completed), an application request is sent, and the TCP
connection is subsequently abandoned. At this point, any state
information kept by the attack tool is removed.
o The attacked server allocates TCP send buffers for transmitting
the response to the client's request. This causes the victim TCP
to tie resources not only for the Transmission Control Block
(TCB), but also for the application data that needs to be
transferred.
o Once the application response is queued for transmission, the
application closes the TCP connection, and thus TCP takes the
responsibility to deliver the queued data. Having the application
close the connection has the benefit for the attacker that the
application is not able to keep track of the number of TCP
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connections in use, and thus it is not able to enforce limits on
the number of connections.
o The attacker repeats the above steps a large number of times, thus
causing a large amount of system memory at the victim host to be
tied to the abandoned connections. When the system memory is
exhausted, the victim host denies service to new connections, or
possibly crashes.
There are a number of factors that affect the effectiveness of this
attack that are worth considering. Firstly, while the attack is
typically targeted at a service such as HTTP, the consequences of the
attack are usually system-wide. Secondly, depending on the size of
the server's response, the underlying TCP connection may or may not
be closed: if the response is larger than the TCP send buffer size at
the server, the application will usually block in a call to write()
or send(), and would therefore not close the TCP connection, thus
allowing the application to enforce limits on the number of ongoing
connections. Consequently, the attacker will usually try to elicit a
response that is equal to or slightly smaller than the send buffer of
the attacked TCP. Thirdly, while [Shalunov, 2000] notes that one
visible effect of this attack is a large number of connections in the
FIN-WAIT-1 state, this will not usually be the case. Given that the
attacker never acknowledges any segment other than the SYN/ACK
segment that is part of the three-way handshake, at the point in
which the attacked TCP tries to send the application's response the
congestion window (cwnd) will usually be 4*SMSS (four maximum-sized
segments). If the application's response were larger than 4*SMSS,
even if the application had closed the connection, the FIN segment
would never be sent, and thus the connection would still remain in
the ESTABLISHED state (rather than transit to the FIN-WAIT-1 state).
7.1.2. Countermeasures
The resource exhaustion attack described in Section 7.1.1 does not
necessarily differ from a legitimate high-load scenario, and
therefore is hard to mitigate without negatively affecting the
robustness of TCP. However, complementary mitigations can still be
implemented to limit the impact of these attacks.
Enforcing limits on the number of connections with no user-space
controlling process
The considerations and recommendations in Section 5.3.2 for enforcing
limits on the number of connections with no user-space controlling
process are applicable to mitigate this vulnerability.
Enforcing per-user and per-process limits
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While the Netkill attack is usually targeted at a service such as
HTTP, its impact is usually system-wide. This is particularly
undesirable, as an attack against a single service might affect the
system as a whole (for example possibly precluding remote system
administration).
In order to avoid an attack against a single service from affecting
other services, we advise TCP implementations to enforce per-process
and per-user limits on maximum kernel memory that can be used at any
time. Additionally, we recommend implementations to enforce per-
process and per-user limits on the number of existent TCP connections
at any time.
Limiting the number of ongoing connections at the application
The considerations and recommendations in Section 5.3.2 for enforcing
limits on the number of ongoing connections at the application are
applicable to mitigate this vulnerability.
Enabling applications to enforce limits on ongoing connections
As discussed in Section 6.1.1, the fact that the close() system call
marks the corresponding file descriptor as closed prevents the
application running on top of TCP from enforcing limits on the
corresponding connection.
While it is common practice for applications to terminate their
connections by means of the close() system call, it is possible for
an application to initiate the connection-termination phase without
closing the corresponding file descriptor (hence keeping control of
the connection).
In order to achieve this, an application performing an active close
(i.e., initiating the connection-termination phase) should replace
the call to close(sockfd) with the following code sequence:
o A call to shutdown(sockfd, SHUT_WR), to close the sending
direction of this connection
o Successive calls to read(), until it returns 0, thus indicating
that the remote TCP peer has finished sending data.
o A call to setsockopt(sockfd, SOL_SOCKET, SO_LINGER, &l,
sizeof(l)), where l is of type struct linger (with its members
l.l_onoff=1 and l.l_linger=90).
o A call to close(sockfd), to close the corresponding file
descriptor.
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The call to shutdown() (instead of close()) allows the application to
retain control of the underlying TCP connection while the connection
transitions through the FIN-WAIT-1 and FIN-WAIT-2 states. However,
the application will not retain control of the connection while it
transitions through the CLOSING and TIME-WAIT states. Nevertheless,
in these states TCP should not have any pending data to send to the
remote TCP peer or to be received by the application running on top
of it, and thus these states are less of a concern for this
particular vulnerability (Netkill).
It should be noted that, strictly speaking, close(sockfd) decrements
the reference count for the descriptor sockfd, and initiates the
connection termination phase only when the reference count reaches 0.
On the other hand, shutdown(sockfd, SHUT_WR) initiates the
connection-termination phase, regardless of the reference count for
the sockfd descriptor. This should be taken into account when
performing the code replacement described above. For example, it
would be a bug for two processes (e.g., parent and child) that share
a descriptor to both call shutdown(sockfd, SHUT_WR).
An application performing a passive close should replace the call to
close(sockfd) with the following code sequence:
o A call to setsockopt(sockfd, SOL_SOCKET, SO_LINGER, &l,
sizeof(l)), where l is of type struct linger (with its members
l.l_onoff=1 and l.l_linger=90).
o A call to close(sockfd), to close the corresponding file
descriptor.
It is assumed that if the application is performing a passive close,
the application already detected that the remote TCP peer finished
sending data by means as a result of a call to read() returning 0.
In this scenario, the application will not retain control of the
underlying connection when it transitions through the LAST_ACK state.
However, in this state TCP should not have any pending data to send
to the remote TCP peer or to be received by the application running
on top of TCP, and thus this state is less of a concern for this
particular vulnerability (Netkill).
Limiting the number of simultaneous connections at firewalls
The considerations and recommendations in Section 5.3.2 for enforcing
limiting the number of simultaneous connections at firewalls are
applicable to mitigate this vulnerability.
Performing heuristics on ongoing TCP connections
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Some heuristics could be performed on TCP connections that may
possibly help if scarce system requirements such as memory become
exhausted. A number of parameters may be useful to perform such
heuristics.
In the case of the Netkill attack described in [Shalunov, 2000],
there are two parameters that are characteristic of a TCP being
attacked:
o A large amount of data queued in the TCP retransmission buffer
(e.g., the socket send buffer).
o Only small amount of data has been successfully transferred to the
remote peer.
Clearly, these two parameters do not necessarily indicate an ongoing
attack. However, if exhaustion of the corresponding system resources
was imminent, these two parameters (among others) could be used to
perform heuristics when considering aborting ongoing connections.
It should be noted that while an attacker could advertise a zero
window to cause the target system to tie system memory to the TCP
retransmission buffer, it is hard to perform any useful statistics
from the advertised window. While it is tempting to enforce a limit
on the length of the persist state (see Section 3.7.2 of this
document), an attacker could simply open the window (i.e., advertise
a TCP window larger than zero) from time to time to prevent this
enforced limit from causing his malicious connections to be aborted.
7.2. TCP segment reassembly buffer
TCP MAY discard out-of-order data when system-memory exhaustion is
imminent.
DISCUSSION:
TCP buffers out-of-order segments to more efficiently handle the
occurrence of packet reordering and segment loss. When out-of-
order data are received, a "hole" momentarily exists in the data
stream which must be filled before the received data can be
delivered to the application making use of TCP's services. This
situation can be exploited by an attacker, which could
intentionally create a hole in the data stream by sending a number
of segments with a sequence number larger than the next sequence
number expected (RCV.NXT) by the attacked TCP. Thus, the attacked
TCP would tie system memory to buffer the out-of-order segments,
without being able to hand the received data to the corresponding
application.
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If a large number of such connections were created, system memory
could be exhausted, precluding the attacked TCP from servicing new
connections and/or continue servicing TCP connections previously
established.
Fortunately, these attacks can be easily mitigated, at the expense
of degrading the performance of possibly legitimate connections.
When out-of-order data is received, an Acknowledgement segment is
sent with the next sequence number expected (RCV.NXT). This means
that receipt of the out-of-order data will not be actually
acknowledged by the TCP's cumulative Acknowledgement Number. As a
result, a TCP is free to discard any data that have been received
out-of-order, without affecting the reliability of the data
transfer. Given the performance implications of discarding out-
of-order segments for legitimate connections, this pruning policy
should be applied only if memory exhaustion is imminent.
As a result of discarding the out-of-order data, these data will
need to be unnecessarily retransmitted. Additionally, a loss
event will be detected by the sending TCP, and thus the slow start
phase of TCP's congestion control will be entered, thus reducing
the data transfer rate of the connection.
It is interesting to note that this pruning policy could be
applied even if Selective Acknowledgements (SACK) (specified in
RFC 2018 [Mathis et al, 1996]) are in use, as SACK provides only
advisory information, and does not preclude the receiving TCP from
discarding data that have been previously selectively-acknowledged
by means of TCP's SACK option, but not acknowledged by TCP's
cumulative Acknowledgement Number.
There are a number of ways in which the pruning policy could be
triggered. For example, when out of order data are received, a
timer could be set, and the sequence number of the out-of-order
data could be recorded. If the hole were filled before the timer
expires, the timer would be turned off. However, if the timer
expired before the hole were filled, all the out-of-order segments
of the corresponding connection would be discarded. This would be
a proactive counter-measure for attacks that aim at exhausting the
receive buffers.
In addition, an implementation could incorporate reactive
mechanisms for more carefully controlling buffer allocation when
some predefined buffer allocation threshold was reached. At such
point, pruning policies would be applied.
A number of mechanisms can aid in the process of freeing system
resources. For example, a table of network prefixes corresponding
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to the IP addresses of TCP peers that have ongoing TCP connections
could record the aggregate amount of out-of-order data currently
buffered for those connections. When the pruning policy was
triggered, TCP connections with hosts that have network prefixes
with large aggregate out-of-order buffered data could be selected
first for pruning the out-of-order segments.
Alternatively, if TCP segments were de-multiplexed by means of a
hash table (as it is currently the case in many TCP
implementations), a counter could be held at each entry of the
hash table that would record the aggregate out-of-order data
currently buffered for those connections belonging to that hash
table entry. When the pruning policy is triggered, the out-of-
order data corresponding to those connections linked by the hash
table entry with largest amount of aggregate out-of-order data
could be pruned first. It is important that this hash is not
computable by an attacker, as this would allow him to maliciously
cause the performance of specific connections to be degraded.
That is, given a four-tuple that identifies a connection, an
attacker should not be able to compute the corresponding hash
value used by the target system to de-multiplex incoming TCP
segments to that connection.
Another variant of a resource exhaustion attack against TCP's
segment reassembly mechanism would target the data structures used
to link the different holes in a data stream. For example, an
attacker could send a burst of 1 byte segments, leaving a one-byte
hole between each of the data bytes sent. Depending on the data
structures used for holding and linking together each of the data
segments, such an attack might waste a large amount of system
memory by exploiting the overhead needed store and link together
each of these one-byte segments.
For example, if a linked-list is used for holding and linking each
of the data segments, each of the involved data structures could
involve one byte of kernel memory for storing the received data
byte (the TCP payload), plus 4 bytes (32 bits) for storing a
pointer to the next node in the linked-list. Additionally, while
such a data structure would require only a few bytes of kernel
memory, it could result in the allocation of a whole memory page,
thus consuming much more memory than expected.
Therefore, implementations should enforce a limit on the number of
holes that are allowed in the received data stream at any given
time. When such a limit is reached, incoming TCP segments which
would create new holes would be silently dropped. Measurements in
[Dharmapurikar and Paxson, 2005] indicate that in the vast
majority of TCP connections have at most a single hole at any
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given time. A limit of 16 holes for each connection would
accommodate even most of the very unusual cases in which there can
be more than hole in the data stream at a given time.
[US-CERT, 2004a] is a security advisory about a Denial of Service
vulnerability resulting from a TCP implementation that did not
enforce limits on the number of segments stored in the TCP
reassembly buffer.
Section 8 of this document describes the security implications of
the TCP segment reassembly algorithm.
7.3. Automatic buffer tuning mechanisms
7.3.1. Automatic send-buffer tuning mechanisms
A TCP implementing an automatic send-buffer tuning mechanism SHOULD
enforce the following limit on the size of the send buffer of each
TCP connection:
send_buffer_size <= send_buffer_pool / (min_buffer_size * max_connections)
where
send_buffer_size:
Maximum send buffer size to be used for this connection
send_buffer_pool:
Total amount of system memory meant for TCP send buffers
min_buffer_size:
Minimum send buffer size for each TCP connection
max_connections:
Maximum number of TCP connections this system is expected to
handle at a time
max_connections may be an artificial limit enforced by the system
administrator specifically on the number of TCP connections, or may
be derived from some other system limit (e.g., the maximum number of
file descriptors)
DISCUSSION:
A number of TCP implementations incorporate automatic tuning
mechanisms for the TCP send buffer size. In most of them, the
underlying idea is to set the send buffer to some multiple of the
congestion window (cwnd). This type of mechanism usually improves
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TCP's performance, by preventing the socket send buffer from
becoming a bottleneck, while avoiding the need to simply
overestimate the TCP send buffer size (i.e., make it arbitrarily
large). [Semke et al, 1998] discusses such an automatic buffer
tuning mechanism.
Unfortunately, automatic tuning mechanisms can be exploited by
attackers to amplify the impact of other resource exhaustion
attacks. For example, an attacker could establish a TCP
connection with a victim host, and cause the congestion window to
be increased (either legitimately or illegitimately). Once the
congestion window (and hence the TCP send buffer) is increased, he
could cause the corresponding system memory to be tied up by
advertising a zero-byte TCP window (see Section 3.7) or simply not
acknowledging any data, thus amplifying the effect of resource
exhaustion attacks such as that discussed in Section 7.1.1.
When an automatic buffer tuning mechanism is implemented, a number
of countermeasures should be incorporated to prevent the mechanism
from being exploited to amplify other resource exhaustion attacks.
Firstly, appropriate policies should be applied to guarantee fair
use of the available system memory by each of the established TCP
connections. Secondly, appropriate policies should be applied to
avoid existing TCP connections from consuming all system
resources, thus preventing service to new TCP connections.
Appendix A of [Semke et al, 1998] proposes an algorithm for the
fair share of the available system memory among the established
connections. However, there are a number of limits that should be
enforced on the system memory assigned for the send buffer of each
connection. Firstly, each connection should always be assigned
some minimum send buffer space that would enable TCP to perform at
an acceptable performance. Secondly, some system memory should be
reserved for future connections, according to the maximum number
of concurrent TCP connections that are expected to be successfully
handled at any given time.
These limits preclude the automatic tuning algorithm from
assigning all the available memory buffers to ongoing connections,
thus preventing the establishment of new connections.
Even if these limits are enforced, an attacker could still create
a large number of TCP connections, each of them tying valuable
system resources. Therefore, in scenarios in which most of the
system memory reserved for TCP send buffers is allocated to
ongoing connections, it may be necessary for TCP to enforce some
policy to free resources to either service more TCP connections,
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or to be able to improve the performance of other existing
connections, by allocating more resources to them.
When needing to free memory in use for send buffers, particular
attention should be paid to TCP's that have a large amount of data
in the socket send buffer, and that at the same time fall into any
of these categories:
* The remote TCP peer that has been advertising a small (possibly
zero) window for a considerable period of time.
* There have been a large number of retransmissions of segments
corresponding to the first few windows of data.
* Connections that fall into one of the previous categories, for
which only a reduced amount of data have been successfully
transferred to the peer TCP since the connection was
established.
Unfortunately, all these cases are valid scenarios for the TCP
protocol, and thus aborting connections that fall in any of these
categories has the potential of causing interoperability problems.
However, in scenarios in which all system resources are allocated,
it may make sense to free resources allocated to TCP connections
which are tying a considerable amount of system resources and that
have not made progress in a considerable period of time.
7.3.2. Automatic receive-buffer tuning mechanism
A number of TCP implementations include automatic tuning mechanisms
for the receive buffer size. These mechanisms aim at setting the
socket buffer to a size that is large enough to avoid the TCP window
from becoming a bottleneck that would limit TCP's throughput, without
wasting system memory by over-sizing it.
[Heffner, 2002] describes a mechanism for the automatic tuning of the
socket receive buffer. Basically, the mechanism aims at measuring
the amount of data received during a RTT (Round-Trip Time), and
setting the socket receive buffer to some multiple of that value.
A TCP implementing an automatic receive-buffer tuning mechanism
SHOULD enforce the following limit on the size of the receive buffer
of each TCP connection:
recv_buffer_size <= recv_buffer_pool / (min_buffer_size * max_connections)
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where:
recv_buffer_size:
Maximum receive buffer size to be used for this connection
recv_buffer_pool:
Total amount of system memory meant for TCP receive buffers
min_buffer_size:
Minimum receive buffer size for each TCP connection
max_connections:
Maximum number of TCP connections this system is expected to
handle at a time
max_connections may be an artificial limit enforced by the system
administrator specifically on the number of TCP connections, or may
be derived from some other system limit (e.g., the maximum number of
file descriptors).
DISCUSSION:
Unfortunately, automatic tuning mechanisms for the socket receive
buffer can be exploited to perform a resource exhaustion attack.
An attacker willing to exploit the automatic buffer tuning
mechanism would first establish a TCP connection with the victim
host. Subsequently, he would start a bulk data transfer to the
victim host. By carefully responding to the peer's TCP segments,
the attacker could cause the peer TCP to measure a large data/RTT
value, which would lead to the adoption of an unnecessarily large
socket receive buffer. For example, the attacker could
optimistically send more data than those allowed by the TCP window
advertised by the remote TCP. Those extra data would cross in the
network with the window updates sent by the remote TCP, and could
lead the TCP receiver to measure a data/RTT twice as big as the
real one. Alternatively, if the TCP timestamp option (specified
in RFC 1323 [Jacobson et al, 1992]) is used for RTT measurement,
the attacker could lead the TCP receiver to measure a small RTT
(and hence a large Data/RTT rate) by "optimistically" echoing
timestamps that have not yet been received.
Finally, once the TCP receiver is led to increase the size of its
receive buffer, the attacker would transmit a large amount of
data, filling the whole peer's receive buffer except for a few
bytes at the beginning of the window (RCV.NXT). This gap would
prevent the peer application from reading the data queued by TCP,
thus tying system memory to the received data segments until (if
ever) the peer application times out.
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A number of limits should be enforced on the amount of system
memory assigned to any given connection. Firstly, each connection
should always be assigned some minimum receive buffer space that
would enable TCP to perform at a minimum acceptable performance.
Additionally, some system memory should be reserved for future
connections, according to the maximum number of concurrent TCP
connections that are expected to be successfully handled at any
given time.
These limits preclude the automatic tuning algorithm from
assigning all the available memory buffers to existing
connections, thus preventing the establishment of new connections.
It is interesting to note that a TCP sender will always try to
retransmit any data that have not been acknowledged by TCP's
cumulative acknowledgement. Therefore, if memory exhaustion is
imminent, a system should consider freeing those memory buffers
used for TCP segments that were received out of order,
particularly when a given connection has been keeping a large
number of out-of-order segments in the receive buffer for a
considerable period of time.
It is worth noting that TCP Selective Acknowledgements (SACK) are
advisory, in the sense that a TCP that has SACKed (but not ACKed)
a block of data is free to discard that block, and expect the TCP
sender to retransmit them when the retransmission timer of the
peer TCP expires.
8. TCP segment reassembly algorithm
8.1. Problems that arise from ambiguity in the reassembly process
If a TCP segment is received containing some data bytes that had
already been received, the first copy of those data SHOULD be used
for reassembling the application data stream.
DISCUSSION:
A security consideration that should be made for the TCP segment
reassembly algorithm is that of data stream consistency between
the host performing the TCP segment reassembly, and a Network
Intrusion Detection System (NIDS) being employed to monitor the
host in question.
In the event a TCP segment was unnecessarily retransmitted, or
there was packet duplication in any of the intervening networks, a
TCP might get more than one copy of the same data. Also, as TCP
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segments can be re-packetized when they are retransmitted, a given
TCP segment might partially overlap data already received in
earlier segments. In all these cases, the question arises about
which of the copies of the received data should be used when
reassembling the data stream. In legitimate and normal
circumstances, all copies would be identical, and the same data
stream would be obtained regardless of which copy of the data was
used. However, an attacker could maliciously send overlapping
segments containing different data, with the intent of evading a
Network Intrusion Detection Systems (NIDS), which might reassemble
the received TCP segments differently than the monitored system.
[Ptacek and Newsham, 1998] provides a detailed discussion of these
issues.
As suggested in Section 3.9 of RFC 793 [Postel, 1981c], if a TCP
segment arrives containing some data bytes that have already been
received, the first copy of those data should be used for
reassembling the application data stream. It should be noted that
while convergence to this policy might prevent some cases of
ambiguity in the reassembly process, there are a number of other
techniques that an attacker could still exploit to evade a NIDS
[CPNI, 2008]. These techniques can generally be defeated if the
NIDS is placed in-line with the monitored system, thus allowing
the NIDS to normalize the network traffic or apply some other
policy that could ensure consistency between the result of the
segment reassembly process obtained by the monitored host and that
obtained by the NIDS.
[CERT, 2003] and [CORE, 2003] are advisories about a heap buffer
overflow in a popular Network Intrusion Detection System resulting
from incorrect sequence number calculations in its TCP stream-
reassembly module.
9. TCP Congestion Control
TCP implements two algorithms, "slow start" and "congestion
avoidance", for controlling the rate at which data is transmitted on
a TCP connection [Allman et al, 1999]. These algorithms require the
addition of two variables as part of TCP per-connection state: cwnd
and ssthresh.
The congestion window (cwnd) is a sender-side limit on the amount of
outstanding data that the sender can have at any time, while the
receiver's advertised window (rwnd) is a receiver-side limit on the
amount of outstanding data. The minimum of cwnd and rwnd governs
data transmission.
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Another state variable, the slow-start threshold (ssthresh), is used
to determine whether it is the slow start or the congestion avoidance
algorithm that should control data transmission. When cwnd <
ssthresh, "slow start" governs data transmission, and the congestion
window (cwnd) is exponentially increased. When cwnd > ssthresh,
"congestion avoidance" governs data transmission, and the congestion
window (cwnd) is only linearly increased.
As specified in RFC 2581 [Allman et al, 1999], when cwnd and ssthresh
are equal the sender may use either slow start or congestion
avoidance.
During slow start, TCP increments cwnd by at most SMSS bytes for each
ACK received that acknowledges new data. During congestion
avoidance, cwnd is incremented by 1 full-sized segment per round-trip
time (RTT), until congestion is detected.
Additionally, TCP uses two algorithms, Fast Retransmit and Fast
Recovery, to mitigate the effects of packet loss. The "Fast
Retransmit" algorithm infers packet loss when three Duplicate
Acknowledgements (DupACKs) are received.
The value "three" is meant to allow for fast-retransmission of
"missing" data, while avoiding network packet reordering from
triggering loss recovery.
Once packet loss is detected by the receipt of three duplicate-ACKs,
the "Fast Recovery" algorithm governs the transfer of new data until
a non-duplicate ACK is received that acknowledges the receipt of new
data. The Fast Retransmit and Fast Recovery algorithms are usually
implemented together, as follows (from RFC 2581):
o When the third duplicate ACK is received, set ssthresh to no more
than the value given in the equation: ssthresh = max (FlightSize /
2, 2*SMSS)
o Retransmit the lost segment and set cwnd to ssthresh plus 3*SMSS.
This artificially "inflates" the congestion window by the number
of segments (three) that have left the network and which the
receiver has buffered.
o For each additional duplicate ACK received, increment cwnd by
SMSS. This artificially inflates the congestion window in order
to reflect the additional segment that has left the network.
o Transmit a segment, if allowed by the new value of cwnd and the
receiver's advertised window.
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o When the next ACK arrives that acknowledges new data, set cwnd to
ssthresh (the value set in step 1). This is termed "deflating"
the window.
9.1. Congestion control with misbehaving receivers
[Savage et al, 1999] describes a number of ways in which TCP's
congestion control mechanisms can be exploited by a misbehaving TCP
receiver to obtain more than its fair share of bandwidth. The
following subsections provide a brief discussion of these
vulnerabilities, along with the possible countermeasures.
9.1.1. ACK division
TCP SHOULD increase cwnd by one SMSS only when a valid ACK covers the
entire data segment sent
(note: or should we recommend the other counter-measure (i.e.,
implementation of ABC?)
DISCUSSION:
Given that TCP updates cwnd based on the number of duplicate ACKs
it receives, rather than on the amount of data that each ACK is
actually acknowledging, a malicious TCP receiver could cause the
TCP sender to illegitimately increase its congestion window by
acknowledging a data segment with a number of separate
Acknowledgements, each covering a distinct piece of the received
data segment.
See Figure 7, in page 64 of the UK CPNI document.
ACK division attack
[Savage et al, 1999] describes two possible countermeasures for
this vulnerability. One of them is to increment cwnd not by a
full SMSS, but proportionally to the amount of data being
acknowledged by the received ACK, similarly to the policy
described in RFC 3465 [Allman, 2003]. Another alternative is to
increase cwnd by one SMSS only when a valid ACK covers the entire
data segment sent.
9.1.2. DupACK forgery
TCP SHOULD keep track of the number of outstanding segments (o_seg),
and accept only up to (o_seg -1) duplicate Acknowledgements.
DISCUSSION:
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The second vulnerability discussed in [Savage et al, 1999] allows
an attacker to cause the TCP sender to illegitimately increase its
congestion window by forging a number of duplicate
Acknowledgements (DupACKs). Figure 8 shows a sample scenario.
The first three DupACKs trigger the Fast Recovery mechanism, while
the rest of them cause the congestion window at the TCP sender to
be illegitimately inflated. Thus, the attacker is able to
illegitimately cause the TCP sender to increase its data
transmission rate.
See Figure 8, in page 65 of the UK CPNI document.
DupACK forgery attack
Fortunately, a number of sender-side heuristics can be implemented
to mitigate this vulnerability. First, the TCP sender could keep
track of the number of outstanding segment (o_seg), and accept
only up to (o_seg -1) DupACKs. Secondly, a TCP sender might, for
example, refuse to enter Fast Recovery multiple times in some
period of time (e.g., one RTT).
[Savage et al, 1999] also describes a modification to TCP to
implement a nonce protocol that would eliminate this
vulnerability. However, this would require modification of all
implementations, which makes this counter-measure hard to deploy.
9.1.3. Optimistic ACKing
Another alternative for an attacker to exploit TCP's congestion
control mechanisms is to acknowledge data that has not yet been
received, thus causing the congestion window at the TCP sender to be
incremented faster than it should.
See Figure 9, in page 66 of the UK CPNI document.
Optimistic ACKing attack
[Savage et al, 1999] describes a number of mitigations for this
vulnerability. Firstly, it describes a countermeasure based on the
concept of "cumulative nonce", which would allow a receiver to prove
that it has received all the segments it is acknowledging. However,
this countermeasure requires the introduction of two new fields to
the TCP header, thus requiring a modification to all the
communicating TCPs, makes this counter-measure hard to deploy.
Secondly, it describes a possible way to encode the nonce in a TCP
segment by carefully modifying its size. While this countermeasure
could be easily deployed (as it is just sender side policy), we
believe that middle-boxes such as protocol-scrubbers might prevent
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this counter-measure from working as expected. Finally, it suggests
that a TCP sender might penalize a TCP receiver that acknowledges
data not yet sent by resetting the corresponding connection. Here we
discourage the implementation of this policy, as it would provide an
attack vector for a TCP-based connection-reset attack, similar to
those described in Section 11.
[US-CERT, 2005a] is a vulnerability advisory about this issue.
9.2. Blind DupACK triggering attacks against TCP
While all of the attacks discussed in [Savage et al, 1999] have the
goal of increasing the performance of the attacker's TCP connections,
TCP congestion control mechanisms can be exploited with a variety of
goals.
Firstly, if bursts of many duplicate-ACKs are sent to the "sending
TCP", the third duplicate-ACK will cause the "lost" segment to be
retransmitted, and each subsequent duplicate-ACK will cause cwnd to
be artificially inflated. Thus, the "sending TCP" might end up
injecting more packets into the network than it really should, with
the potential of causing network congestion. This is a potential
consequence of the "Duplicate-ACK spoofing attack" described in
[Savage et al, 1999].
Secondly, if bursts of three duplicate ACKs are sent to the TCP
sender, the attacked system would infer packet loss, and ssthresh and
cwnd would be reduced. As noted in RFC 2581 [Allman et al, 1999],
causing two congestion control events back-to-back will often cut
ssthresh and cwnd to their minimum value of 2*SMSS, with the
connection immediately entering the slower-performing congestion
avoidance phase. While it would not be attractive for an attacker to
perform this attack against one of his TCP connections, the attack
might be attractive when the TCP connection to be attacked is
established between two other parties.
It is usually assumed that in order for an off-path attacker to
perform attacks against a third-party TCP connection, he should be
able to guess a number of values, including a valid TCP Sequence
Number and a valid TCP Acknowledgement Number. While this is true if
the attacker tries to "inject" valid packets into the connection by
himself, a feature of TCP can be exploited to fool one of the TCP
endpoints to transmit valid duplicate Acknowledgements on behalf of
the attacker, hence relieving the attacker of the hard task of
forging valid values for the Sequence Number and Acknowledgement
Number TCP header fields.
Section 3.9 of RFC 793 [Postel, 1981c] describes the processing of
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incoming TCP segments as a function of the connection state and the
contents of the various header fields of the received segment. For
connections in the ESTABLISHED state, the first check that is
performed on incoming segments is that they contain "in window" data.
That is,
RCV.NXT <= SEG.SEQ <= RCV.NXT+RCV.WND, or
RCV.NXT <= SEG.SEQ+SEG.LEN-1 < RCV.NXT+RCV.WND
If a segment does not pass this check, it is dropped, and an
Acknowledgement is sent in response:
<SEQ=SND.NXT><ACK=RCV.NXT><CTL=ACK>
The goal of this behavior is that, in the event data segments are
received by the TCP receiver, but all the corresponding
Acknowledgements are lost, when the TCP sender retransmits the
supposedly lost data, the TCP receiver will send an Acknowledgement
reflecting all the data received so far. If "old" TCP segments were
silently dropped, the scenario just described would lead to a
"frozen" TCP connection, with the TCP sender retransmitting the data
for which it has not yet received an Acknowledgement, and the TCP
receiver silently ignoring these segments. Additionally, it helps
TCP to detect half-open connections.
This feature implies that, provided the four-tuple that identifies a
given TCP connection is known or can be easily guessed, an attacker
could send a TCP segment with an "out of window" Sequence Number to
one of the endpoints of the TCP connection to cause it to send a
valid ACK to the other endpoint of the connection. Figure 10
illustrates such a scenario.
See Figure 10, in page 68 of the UK CPNI document.
Blind Dup-ACK forgery attack
As discussed in [Watson, 2004] and RFC 4953 [Touch, 2007], there are
a number of scenarios in which the four-tuple that identifies a TCP
connection is known or can be easily guessed. In those scenarios, an
attacker could perform any of the "blind" attacks described in the
following subsections by exploiting the technique described above.
The following subsections describe blind DupACK-triggering attacks
that aim at either degrading the performance of an arbitrary
connection, or causing a TCP sender to illegitimately increase the
rate at which it transmits data, potentially leading to network
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congestion.
9.2.1. Blind throughput-reduction attack
As discussed in Section 9, when three duplicate Acknowledgements are
received, the congestion window is reduced to half the current amount
of outstanding data (FlightSize). Additionally, the slow-start
threshold (ssthresh) is reduced to the same value, causing the
connection to enter the slower-performing congestion avoidance phase.
If two congestion-control events occur back to back, ssthresh and
cwnd will often be reduced to their minimum value of 2*SMSS.
An attacker could exploit the technique described in Section 9.2 to
cause the throughput of the attacked TCP connection to be reduced, by
eliciting three duplicate acknowledgements from the TCP receiver,
which would cause the TCP sender to reduce its congestion window. In
principle, the attacker would need to send a burst of only three out-
of-window segments. However, in case the TCP receiver implements an
acknowledgement policy such as "ACK every other segment", four out-
of-window segments might be needed. The first segment would cause
the pending (delayed) Acknowledgement to be sent, and the next three
segments would elicit the actual duplicate Acknowledgements.
Figure 11 shows a time-line graph of a sample scenario. The burst of
DupACKs (in green) elicited by the burst of out-of-window segments
(in red) sent by the attacker causes the TCP sender to retransmit the
missing segment (in blue) and enter the loss recovery phase. Once a
segment that acknowledges new data is received by the TCP sender, the
loss recovery phase ends, and cwnd and ssthresh are set to half the
number of segments that were outstanding when the loss recovery phase
was entered.
See Figure 11, in page 69 of the UK CPNI document.
Blind throughput-reduction attack (time-line graph)
The graphic assumes that the TCP receiver sends an Acknowledgement
for every other data segment it receives, and that the TCP sender
implements Appropriate Byte Counting (specified in RFC 3465 [Allman,
2003]) on the received Acknowledgement segments. However,
implementation of these policies is not required for the attack to
succeed.
9.2.2. Blind flooding attack
As discussed in Section 9, when three duplicate Acknowledgements are
received, the "lost" segment is retransmitted, and the congestion
window is artificially inflated for each DupACK received, until the
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loss recovery phase ends. By sending a long burst of out-of-window
segments to the TCP receiver of the attacked connection, an attacker
could elicit a long burst of valid duplicate acknowledgements that
would illegitimately cause the TCP sender of the attacked TCP
connection to increase its data transmission rate.
Figure 12 shows a time-line graph for this attack. The long burst of
DupACKs (in green) elicited by the long burst of out-of-window
segments (in red) sent by the attacker causes the TCP sender to enter
the loss recovery phase and illegitimately inflate the congestion
window, leading to an increase in the data transmission rate. Once a
segment that acknowledges new data is received by the TCP sender, the
loss recovery phase ends, and the data transmission rate is reduced.
See Figure 12, in page 70 of the UK CPNI document.
Blind flooding attack (time-line graph)
Figure 13 is a time-sequence graph produced from packet logs obtained
from tests of the described attack in a real network. A burst of
segments is sent upon receipt of the burst of Duplicate
Acknowledgements illegitimately elicited by the attacker. Figure 14
is an averaged-throughput graphic for the same time frame, which
clearly shows the effect of the attack in terms of throughput.
See Figure 13, in page 71 of the UK CPNI document.
Blind flooding attack (time sequence graph)
See Figure 14, in page 71 of the UK CPNI document.
Blind flooding attack (averaged throughput graph)
These graphics were produced with Shawn Ostermann's tcptrace tool
[Ostermann, 2008]. An explanation of the format of the graphics can
be found in tcptrace's manual (available at the project's web site:
http://www.tcptrace.org).
9.2.3. Difficulty in performing the attacks
In order to exploit the technique described in Section 9.2 of this
document, an attacker would need to know the four-tuple {IP Source
Address, TCP Source Port, IP Destination Address, TCP Destination
Port} that identifies the connection to be attacked. As discussed by
[Watson, 2004] and RFC 4953 [Touch, 2007], there are a number of
scenarios in which these values may be known or easily guessed.
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It is interesting to note that the attacks described in Section 9.2
of this document will typically require a much smaller number of
packets than other "blind" attacks against TCP, such as those
described in [Watson, 2004] and RFC 4953 [Touch, 2007], as the
technique discussed in Section 9.2 relieves the attacker from having
to guess valid TCP Sequence Numbers and a TCP Acknowledgement
numbers.
The attacks described in Section 9.2.1 and Section 9.2.2 of this
document require the attacker to forge the source address of the
packets it sends. Therefore, if ingress/egress filtering is
performed by intermediate systems, the attacker's packets would not
get to the intended recipient, and thus the attack would not succeed.
However, we consider that ingress/egress filtering cannot be relied
upon as the first line of defense against these attacks.
Finally, it is worth noting that in order to successfully perform the
blind attacks discussed in Section 9.2.1 and Section 9.2.2 of this
document, the burst of out-of-sequence segments sent by the attacker
should not be intermixed with valid data segments sent by the TCP
sender, or else the Acknowledgement number of the illegitimately-
elicited ACK segments would change, and the Acknowledgements would
not be considered "Duplicate Acknowledgements" by the TCP sender.
Tests performed in real networks seem to suggest that this
requirement is not hard to fulfill, though.
9.2.4. Modifications to TCP's loss recovery algorithms
There are a number of algorithms that augment TCP's loss recovery
mechanism that have been suggested by TCP researchers and have been
specified by the IETF in the RFC series. This section describes a
number of these algorithms, and discusses how their implementation
affects (or not) the vulnerability of TCP to the attacks discussed in
Section 9.2.1 and Section 9.2.2 of this document.
NewReno
RFC 3782 [Floyd et al, 2004] specifies the NewReno algorithm, which
is meant to improve TCP's performance in the presence of multiple
losses in a single window of data. The implication of this algorithm
with respect to the attacks discussed in the previous sections is
that whenever either of the attacks is performed against a connection
with a NewReno TCP sender, a full-window (or half a window) of data
will be unnecessarily retransmitted. This is particularly
interesting in the case of the blind-flooding attack, as the attack
would elicit even more packets from the TCP sender.
Whether a full-window or just half a window of data is retransmitted
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depends on the Acknowledgement policy at the TCP receiver. If the
TCP receiver sends an Acknowledgement (ACK) for every segment, a
full-window of data will be retransmitted. If the TCP receiver sends
an Acknowledgement (ACK) for every other segment, then only half a
window of data will be retransmitted.
Figure 15 is a time-sequence graph produced from packet logs obtained
from tests performed in a real network. Once loss recovery is
illegitimately triggered by the duplicate-ACKs elicited by the
attacker, an entire flight of data is unnecessarily retransmitted.
Figure 16 is an averaged-throughput graphic for the same time-frame,
which shows an increase in the throughput of the connection resulting
from the retransmission of segments governed by NewReno's loss
recovery.
See Figure 15, in page 73 of the UK CPNI document.
NewReno loss recovery (time-sequence graph)
See Figure 16, in page 74 of the UK CPNI document.
NewReno loss recovery (averaged throughput graph)
Limited Transmit
RFC 3042 [Allman et al, 2001] proposes an enhancement to TCP to more
effectively recover lost segments when a connection's congestion
window is small, or when a large number of segments are lost in a
single transmission window. The "Limited Transmit" algorithm calls
for sending a new data segment in response to each of the first two
Duplicate Acknowledgements that arrive at the TCP sender. This would
provide two additional transmitted packets that may be useful for the
attacker in the case of the blind flooding attack described in
Section 9.2.2 is performed.
SACK-based loss recovery
RFC 3517 [Blanton et al, 2003] specifies a conservative loss-recovery
algorithm that is based on the use of the selective acknowledgement
(SACK) TCP option. The algorithm uses DupACKs as an indication of
congestion, as specified in RFC 2581 [Allman et al, 1999]. However,
a difference between this algorithm and the basic algorithm described
in RFC 2581 is that it clocks out segments only with the SACK
information included in the DupACKs. That is, during the loss
recovery phase, segments will be injected in the network only if the
SACK information included in the received DupACKs indicates that one
or more segments have left the network. As a result, those systems
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that implement SACK-based loss recovery will not be vulnerable to the
blind flooding attack described in Section 9.2.2. However, as RFC
3517 does not actually require DupACKs to include new SACK
information (corresponding to data that has not yet been acknowledged
by TCP's cumulative Acknowledgement), systems that implement SACK-
based loss-recovery may still remain vulnerable to the blind
throughput-reduction attack described in Section 9.2.1. SACK-based
loss recovery implementations should be updated to implement the
countermeasure ("Use of SACK information to validate DupACKs")
described in Section 9.2.5.
9.2.5. Countermeasures
TCP SHOULD validate the Sequence Number of an incomming TCP segment
as follows:
RCV.NXT - MAX.RCV.WND <= SEG.SEQ <= RCV.NXT + RCV.WND
where MAX.RCV.WND is the largest TCP window that has so far been
advertised to the remote endpoint.
If a segment passes this check, the processing rules specified in RFC
793 [Postel, 1981c] MUST applied. Otherwise, TCP SHOULD send an ACK
(as specified by the processing rules in RFC 793 [Postel, 1981c]),
applying rate-limiting to the Acknowledgement segments sent in
response to out-of-window segments.
DISCUSSION:
As discussed in Section 9.2, TCP responds with an ACK when an out-
of-window segment is received, to accommodate those scenarios in
which the Acknowledgement segments that correspond to some
received data are lost in the network, and to help discover half-
open TCP connections.
However, it is possible to restrict the sequence numbers that are
considered acceptable, and have TCP respond with ACKs only when it
is strictly necessary.
A feature of TCP is that, in some scenarios, it can detect half-
open connections. If an implementation chose to silently drop
those TCP segments that do not pass the check enforced by the
equation above, it could prevent TCP from detecting half-open
connections. Figure 17 shows a scenario in which, provided that
"TCP B" behaves as specified in RFC 793, a half-open connection
would be discovered and aborted.
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An established connection is said to be "half open" if one of the
TCPs has closed or aborted the connection at its end without the
knowledge of the other, or if the two ends of the connection have
become desynchronized owing to a crash that resulted in loss of
memory.
See Figure 17, in page 76 of the UK CPNI document.
Half-Open Connection Discovery
In the scenario illustrated by Figure 17, TCP A crashes losing the
connection-state information of the TCP connection with TCP B. In
line 3, TCP A tries to establish a new connection with TCP B,
using the same four-tuple {IP Source Address, TCP source port, IP
Destination Address, TCP destination port}. In line 4, as the SYN
segment is out of window, TCP B responds with an ACK. This ACK
elicits an RST segment from TCP A, which causes the half-open
connection at TCP B to be aborted.
If the SYN segment had been "in window", TCP B would have sent an
RST segment instead, which would have closed the half-open
connection. Ongoing work at the TCPM WG of the IETF proposes to
change this behavior, and make TCP respond to a SYN segment
received for any of the synchronized states with an ACK segment,
to avoid in-window SYN segments from being used to perform
connection-reset attacks [Ramaiah et al, 2008].
However, in case the out-of-window segment was silently dropped,
the scenario in Figure 17 would change into that in Figure 18.
See Figure 18, in page 76 of the UK CPNI document.
Half-Open Connection Discovery with the proposed counter-measure
In line 3, the SYN segment sent by TCP A is silently dropped by
TCP B because it does not pass the check enforced by the equation
above (i.e., it contains an out-of-window sequence number). As a
result, some time later (an RTO) TCP A retransmits its SYN
segment. Even after TCP A times out, the half-open connection at
TCP B will remain in the same state.
Thus, a conservative reaction to those segments that do not pass
the check enforced by the equation above would be to respond with
an Acknowledgement segment (as specified by RFC 793), applying
rate-limiting to those Acknowledgement segments sent in response
to segments that do not pass the check enforced by that equation.
An implementation might choose to enforce a rate-limit of, e.g.,
one ACK per five seconds, as a single ACK segment is needed for
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the Half-Open Connection Discovery mechanism to work.
As the only reason to respond with an ACK to those segments that
do not pass the check enforced by the equation above is to allow
TCP to discover half-open connections, an aggressive rate-limit
can be enforced. As long as the rate-limit prevents out-of-window
segments from eliciting three Acknowledgment segments in a Round-
trip Time (RTT), an attacker would not be able to trigger TCP's
loss-recovery, and thus would not be able to perform the attacks
described in the previous sections.
It is interesting to note that RFC 793 [Postel, 1981c] itself
states that half-open connections are expected to be unusual.
Additionally, given that in many scenarios it may be unlikely for
a TCP connection request to be issued with the same four-tuple as
that of the half-open connection, a complete solution for the
discovery of half-open connections cannot rely on the mechanism
illustrated by Figure 17, either. Therefore, some implementations
might choose to sacrifice TCP's ability to detect half-open
connections, and have a more aggressive reaction to those segments
that do not pass the check enforced by the equation above by
silently dropping them.
This validation check can also help to avoid ACK wars in some
scenarios that may arise from the use of transparent proxies. In
those scenarios, when the transparent proxy fails to wire (i.e.,
is disabled), the sequence numbers of the two end-points of the
TCP connection become desynchronized, and both TCPs begin to send
duplicate Acknowledgements to each other, with the intention of
re-synchronizing them. As the sequence numbers never get re-
synchronized, the ACK war can only be stopped by an external
agent.
TCP SHOULD limit the number of duplicate acknowledgements it will
honour to:
Max_DupACKs = (FlightSize / SMSS) - 1
Where FlightSize and SMSS are the values defined in RFC 2581 [Allman
et al, 1999]. When more than Max_DupACKs duplicate acknowledgements
are received, the exceeding DupACKs should be silently dropped.
DISCUSSION:
Note that duplicate acknowledgements should be elicited by out-of-
order segments.
In the case of TCP connections that have agreed to employ SACK, TCP
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SHOULD validate duplicate ACKs with the following criteria: Valid
Duplicate ACKs MUST contain new SACK information. The SACK
information MUST refer to data that has already been sent, but that
has not yet been acknowledged by TCP's cumulative Acknowledgement. A
TCP segment that does not pass this check SHOULD NOT be considered as
"duplicate Acknowledgement".
DISCUSSION:
SACK, specified in 2018 [Mathis et al, 1996], provides a mechanism
for TCP to be able to acknowledge the receipt of out-of-order TCP
segments. For connections that have agreed to use SACK, each
legitimate DupACK will contain new SACK information that reflects
the data bytes contained in the out-of-order data segment that
elicited the DupACK.
RFC 3517 [Blanton et al, 2003] specifies a SACK-based loss
recovery algorithm for TCP. However, it does recommend TCP
implementations to validate DupACKs by requiring that they contain
new SACK information. Results obtained from auditing a number of
TCP implementations seem to indicate that most TCP implementations
do not enforce this validation check on incoming DupACKs, either.
In the case of TCP connections that have agreed to use SACK, a
validation check should be performed on incoming ACK segments to
completely eliminate the attacks described in Section 9.2.1 and
Section 9.2.2 of this document: "Duplicate ACKs should contain new
SACK information. The SACK information should refer to data that
has already been sent, but that has not yet been acknowledged by
TCP's cumulative Acknowledgement".
Those ACK segments that do not comply with this validation check
should not be considered "duplicate ACKs", and thus should not
trigger the loss-recovery phase.
In case at least one segment in a window of data has been lost,
the successive segments will elicit the generation of Duplicate
ACKs containing new SACK information. This SACK information will
indicate the receipt of these successive segments by the TCP
receiver.
In the case of pure ACKs illegitimately elicited by out-of-window
segments, however, the ACKs will not contain any SACK information.
If DSACK (specified in 2883 [Floyd et al, 2000]) were implemented
by the TCP receiver, then the illegitimately elicited DupACKs
might contain out-of-window SACK information if the sequence
number of the forged TCP segment (SEG.SEQ) is lower than the next
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expected sequence number (RECV.NXT) at the TCP receiver. Such
segments should be considered to indicate the receipt of duplicate
data, rather than an indication of lost data, and therefore should
not trigger loss recovery.
Other possible general mitigations are discussed in the following
paragraphs:
TCP port number randomization
As in order to perform the blind attacks described in Section 9.2.1
and Section 9.2.2 the attacker needs to know the TCP port numbers in
use by the connection to be attacked, obfuscating the TCP source port
used for outgoing TCP connections will increase the number of packets
required to successfully perform these attacks. Section 3.1 of this
document discusses the use of port randomization.
It must be noted that given that these blind DupACK triggering
attacks do not require the attacker to forge valid TCP Sequence
numbers and TCP Acknowledgement numbers, port randomization should
not be relied upon as a first line of defense.
Ingress and Egress filtering
Ingress and Egress filtering reduces the number of systems in the
global Internet that can perform attacks that rely on forged source
IP addresses. While protection from the blind attacks discussed in
Section 9.2 should not rely only on Ingress and Egress filtering, its
deployment is recommended to help prevent all attacks that rely on
forged IP addresses. RFC 3704 [Baker and Savola, 2004], RFC 2827
[Ferguson and Senie, 2000], and [NISCC, 2006] provide advice on
Ingress and Egress filtering.
Generalized TTL Security Mechanism (GTSM)
RFC 5082 [Gill et al, 2007] proposes a check on the TTL field of the
IP packets that correspond to a given TCP connection to reduce the
number of systems that could successfully attack the protected TCP
connection. It provides for the attacks discussed in this document
the same level of protection than for the attacks described in
[Watson, 2004] and RFC 4953 [Touch, 2007]. While implementation of
this mechanism may be useful in some scenarios, it should be clear
that countermeasures discussed in the previous sections provide a
more effective and simpler solution than that provided by the GTSM.
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9.3. TCP Explicit Congestion Notification (ECN)
ECN (Explicit Congestion Notification) provides a mechanism for
intermediate systems to signal congestion to the communicating
endpoints that in some scenarios can be used as an alternative to
dropping packets.
RFC 3168 [Ramakrishnan et al, 2001] contains a detailed discussion of
the possible ways and scenarios in which ECN could be exploited by an
attacker.
RFC 3540 [Spring et al, 2003] specifies an improvement to ECN based
on nonces, that protects against accidental or malicious concealment
of marked packets from the TCP sender. The specified mechanism
defines a "NS" ("Nonce Sum") field in the TCP header that makes use
of one bit from the Reserved field, and requires a modification in
both of the endpoints of a TCP connection to process this new field.
This mechanism is still in "Experimental" status, and since it might
suffer from the behavior of some middle-boxes such as firewalls or
packet-scrubbers, we defer a recommendation of this mechanism until
more experience is gained.
There also is ongoing work in the research community and the IETF to
define alternate semantics for the ECN field of the IP header (e.g.,
see [PCNWG, 2009]).
The following subsections try to summarize the security implications
of ECN.
9.3.1. Possible attacks by a compromised router
Firstly, a router controlled by a malicious user could erase the CE
codepoint (either by replacing it with the ECT(0), ECT(1), or non-ECT
codepoints), effectively eliminating the congestion indication. As a
result, the corresponding TCP sender would not reduce its data
transmission rate, possibly leading to network congestion. This
could also lead to unfairness, as this flow could experience better
performance than other flows for which the congestion indication is
not erased (and thus their transmission rate is reduced).
Secondly, a router controlled by a malicious user could
illegitimately set the CE codepoint, falsely indicating congestion,
to cause the TCP sender to reduce its data transmission rate.
However, this particular attack is no worse than the malicious router
simply dropping the packets rather setting their CE codepoint.
Thirdly, a malicious router could turn off the ECT codepoint of a
packet, thus disabling ECN support. As a result, if the packet later
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arrives at a router that is experiencing congestion, it may be
dropped rather than marked. As with the previous scenario, though,
this is no worse than the malicious router simply dropping the
corresponding packet.
It should be noted that a compromised on-path IP router could engage
in a much broader range of attacks, with broader impacts, and at much
lower attacker cost than the ones described here. Such a compromised
router is extremely unlikely to engage in the attack vectors
discussed in this section, given the existence of more effective
attack vectors that have lower attacker cost.
9.3.2. Possible attacks by a malicious TCP endpoint
If a packet with the ECT codepoint set arrives at an ECN-capable
router that is experiencing moderate congestion, the router may
decide to set its CE codepoint instead of dropping it. If either of
the TCP endpoints do not honour the congestion indication provided by
an ECN-capable router, this would result in unfairness, as other
(legitimate) ECN-capable flows would still reduce their sending rate
in response to the ECN marking of packets. Furthermore, under
moderate congestion, non-ECN-capable flows would be subject to packet
drops by the same router. As a result, the flow with a malicious TCP
end-point would obtain better service than the legitimate flows.
As noted in RFC 3168 [Ramakrishnan et al, 2001], a TCP endpoint
falsely indicating ECN capability could lead to unfairness, allowing
the mis-beheaving flow to get more than its fair share of the
bandwidth. This could be the result of the mis-behavior of either of
the TCP endpoints. For example, the sending TCP could indicate ECN
capability, but then send a CWR in response to an ECE without
actually reducing its congestion window. Alternatively (or in
addition), the receiving TCP could simply ignore those packets with
the CE codepoint set, thus avoiding the sending TCP from receiving
the congestion indication.
In the case of the sending TCP ignoring the ECN congestion
indication, this would be no worse than the sending TCP ignoring the
congestion indication provided by a lost segment. However, the case
of a TCP receiver ignoring the CE codepoint allows the TCP receiver
to get more than its fair share of bandwidth in a way that was
previously unavailable. If congestion was kept "moderate", then the
malicious TCP receiver could maintain the unfairness, as the router
experiencing congestion would mark the offending packets of the
misbehaving flow rather than dropping them. At the same time,
legitimate ECN-capable flows would respond to the congestion
indication provided by the CE codepoint, while legitimate non-ECN-
capable flows would be subject of packet dropping. However, if
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congestion turned to sufficiently heavy, the router experiencing
congestion would switch from marking packets to dropping packets, and
at that point the attack vector provided by ECN could no longer be
exploited (until congestion returns to moderate state).
RFC 3168 [Ramakrishnan et al, 2001] describes the use of "penalty
boxes" which would act on flows that do not respond appropriately to
congestion indications. Section 10 of RFC 3168 suggests that a first
action taken at a penalty box for an ECN-capable flow would be to
switch to dropping packets (instead of marking them), and, if the
flow does not respond appropriately to the congestion indication, the
penalty box could reset the misbehaving connection. Here we
discourage implementation of such a policy, as it would create a
vector for connection-reset attacks. For example, an attacker could
forge TCP segments with the same four-tuple as the targeted
connection and cause them to transit the penalty box. The penalty
box would first switch from marking to dropping packets. However,
the attacker would continue sending forged segments, at a steady
rate. As a result, if the penalty box implemented such a severe
policy of resetting connections for flows that still do not respond
to end-to-end congestion control after switching from marking to
dropping, the attacked connection would be reset.
10. TCP API
Section 3.8 of RFC 793 [Postel, 1981c] describes the minimum set of
TCP User Commands required of all TCP Implementations. Most
operating systems provide an Application Programming Interface (API)
that allows applications to make use of the services provided by TCP.
One of the most popular APIs is the Sockets API, originally
introduced in the BSD networking package [McKusick et al, 1996].
10.1. Passive opens and binding sockets
When there is already a pending passive OPEN for some local port
number, TCP SHOULD NOT allow processes that do not belong to the same
user to "reuse" the local port for another passive OPEN.
Additionally, reuse of a local port SHOULD default to "off", and be
enabled only by an explicit command (e.g., the setsockopt() function
of the Sockets API).
DISCUSSION:
RFC 793 specifies the syntax of the "OPEN" command, which can be
used to perform both passive and active opens. The syntax of this
command is as follows:
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OPEN (local port, foreign socket, active/passive [, timeout] [,
precedence] [, security/compartment] [, options]) -> local
connection name
When this command is used to perform a passive open (i.e., the
active/passive flag is set to passive), the foreign socket
parameter may be either fully-specified (to wait for a particular
connection) or unspecified (to wait for any call).
As discussed in Section 2.7 of RFC 793 [Postel, 1981c], if there
are several passive OPENs with the same local socket (recorded in
the corresponding TCB), an incoming connection will be matched to
the TCB with the more specific foreign socket. This means that
when the foreign socket of a passive OPEN matches that of the
incoming connection request, that passive OPEN takes precedence
over those passive OPENs with an unspecified foreign socket.
Popular implementations such as the Sockets API let the user
specify the local socket as fully-specified {local IP address,
local TCP port} pair, or as just the local TCP port (leaving the
local IP address unspecified). In the former case, only those
connection requests sent to {local port, local IP address} will be
accepted. In the latter case, connection requests sent to any of
the system's IP addresses will be accepted. In a similar fashion
to the generic API described in Section 2.7 of RFC 793, if there
is a pending passive OPEN with a fully-specified local socket that
matches that for which a connection establishment request has been
received, that local socket will take precedence over those which
have left the local IP address unspecified. The implication of
this is that an attacker could "steal" incoming connection
requests meant for a local application by performing a passive
OPEN that is more specific than that performed by the legitimate
application.
10.2. Active opens and binding sockets
TCP SHOULD NOT allow port numbers that have been allocated for a TCP
that is the LISTEN or CLOSED states to be specified as the "local
port" argument of the "OPEN" command.
An implementation MAY relax the aforementioned restriction when the
process or system user requesting allocation of such a port number is
the same that the process or system user controlling the TCP in the
CLOSED or LISTEN states with the same port number.
DISCUSSION:
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As discussed in Section 10.1, the "OPEN" command specified in
Section 3.8 of RFC 793 [Postel, 1981c] can be used to perform
active opens. In case of active opens, the parameter "local port"
will contain a so-called "ephemeral port". While the only
requirement for such an ephemeral port is that the resulting
connection-id is unique, port numbers that are currently in use by
a TCP in the LISTEN state should not be allowed for use as
ephemeral ports. If this rule is not complied, an attacker could
potentially "steal" an incoming connection to a local server
application by issuing a connection request to the victim client
at roughly the same time the client tries to connect to the victim
server application. If the SYN segment corresponding to the
attacker's connection request and the SYN segment corresponding to
the victim client "cross each other in the network", and provided
the attacker is able to know or guess the ephemeral port used by
the client, a TCP simultaneous open scenario would take place, and
the incoming connection request sent by the client would be
matched with the attacker's socket rather than with the victim
server application's socket.
As already noted, in order for this attack to succeed, the
attacker should be able to guess or know (in advance) the
ephemeral port selected by the victim client, and be able to know
the right moment to issue a connection request to the victim
client. While in many scenarios this may prove to be a difficult
task, some factors such as an inadequate ephemeral port selection
policy at the victim client could make this attack feasible.
It should be noted that most applications based on popular
implementations of TCP API (such as the Sockets API) perform
"passive opens" in three steps. Firstly, the application obtains
a file descriptor to be used for inter-process communication
(e.g., by issuing a socket() call). Secondly, the application
binds the file descriptor to a local TCP port number (e.g., by
issuing a bind() call), thus creating a TCP in the fictional
CLOSED state. Thirdly, the aforementioned TCP is put in the
LISTEN state (e.g., by issuing a listen() call). As a result,
with such an implementation of the TCP API, even if port numbers
in use for TCPs in the LISTEN state were not allowed for use as
ephemeral ports, there is a window of time between the second and
the third steps in which an attacker could be allowed to select a
port number that would be later used for listening to incoming
connections. Therefore, these implementations of the TCP API
should enforce a stricter requirement for the allocation of port
numbers: port numbers that are in use by a TCP in the LISTEN or
CLOSED states should not be allowed for allocation as ephemeral
ports.
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An implementation might choose to relax the aforementioned
restriction when the process or system user requesting allocation
of such a port number is the same that the process or system user
controlling the TCP in the CLOSED or LISTEN states with the same
port number.
11. Blind in-window attacks
In the last few years awareness has been raised about a number of
"blind" attacks that can be performed against TCP by forging TCP
segments that fall within the receive window [NISCC, 2004] [Watson,
2004].
The term "blind" refers to the fact that the attacker does not have
access to the packets that belong to the attacked connection.
The effects of these attacks range from connection resets to data
injection. While these attacks were known in the research community,
they were generally considered unfeasible. However, increases in
bandwidth availability and the use of larger TCP windows raised
concerns in the community. The following subsections discuss a
number of forgery attacks against TCP, along with the possible
countermeasures to mitigate their impact.
11.1. Blind TCP-based connection-reset attacks
Blind connection-reset attacks have the goal of causing a TCP
connection maintained between two TCP endpoints to be aborted. The
level of damage that the attack may cause usually depends on the
application running on top of TCP, with the more vulnerable
applications being those that rely on long-lived TCP connections.
An interesting case of such applications is BGP [Rekhter et al,
2006], in which a connection-reset usually results in the
corresponding entries of the routing table being flushed.
There are a variety of vectors for performing TCP-based connection-
reset attacks against TCP. [Watson, 2004] and [NISCC, 2004] raised
awareness about connection-reset attacks that exploit the RST flag of
TCP segments. [Ramaiah et al, 2008] noted that carefully crafted SYN
segments could also be used to perform connection-reset attacks.
This document describes yet two previously undocumented vectors for
performing connection-reset attacks: the Precedence field of IP
packets that encapsulate TCP segments, and illegal TCP options.
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11.1.1. RST flag
TCP SHOULD implement the mitigation for RST-based attacks specified
in [Ramaiah et al, 2008].
DISCUSSION:
The RST flag signals a TCP peer that the connection should be
aborted. In contrast with the FIN handshake (which gracefully
terminates a TCP connection), an RST segment causes the connection
to be abnormally closed.
As stated in Section 3.4 of RFC 793 [Postel, 1981c], all reset
segments are validated by checking their Sequence Numbers, with
the Sequence Number considered valid if it is within the receive
window. In the SYN-SENT state, however, an RST is valid if the
Acknowledgement Number acknowledges the SYN segment that
supposedly elicited the reset.
[Ramaiah et al, 2008] proposes a modification to TCP's transition
diagram to address this attack vector. The counter-measure is a
combination of enforcing a more strict validation check on the
sequence number of reset segments, and the addition of a
"challenge" mechanism. With the implementation of the proposed
mechanism, TCP would behave as follows:
If the Sequence Number of an RST segment is outside the receive
window, the segment is silently dropped (as stated by RFC 793).
That is, a reset segment is discarded unless it passes the
following check:
RCV.NXT <= Sequence Number < RCV.NXT+RCV.WND
If the sequence number falls exactly on the left-edge of the
receive window, the reset is honoured. That is, the connection is
reset if the following condition is true:
Sequence Number == RCV.NXT
If an RST segment passes the first check (i.e., it is within the
receive window) but does not pass the second check (i.e., it does
not fall exactly on the left edge of the receive window), an
Acknowledgement segment ("challenge ACK") is set in response:
<SEQ=SND.NXT><ACK=RCV.NXT><CTL=ACK>
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This Acknowledgement segment is referred to as a "challenge ACK"
as, in the event the RST segment that elicited it had been
legitimate (but silently dropped as a result of enforcing the
above checks), the challenge ACK would elicit a new reset segment
that would fall exactly on the left edge of the window and would
thus pass all the above checks, finally resetting the connection.
We recommend the implementation of this countermeasure. However,
we are aware of patent claims on this counter-measure, and suggest
vendors to research the consequences of the possible patents that
may apply.
[US-CERT, 2003a] is an advisory of a firewall system that was
found particularly vulnerable to resets attack because of not
validating the TCP Sequence Number of RST segments. Clearly, all
TCPs (including those in middle-boxes) should validate RST
segments as discussed in this section.
11.1.2. SYN flag
Processing of SYN segments received for connections in the
synchronized states SHOULD occur as follows:
o If a SYN segment is received for a connection in any synchronized
state other than TIME-WAIT, respond with an ACK, applying rate-
throttling. [Ramaiah et al, 2008]
o If the corresponding connection is in the TIME-WAIT state, then
process the incomming SYN as specified in
[I-D.ietf-tcpm-tcp-timestamps].
DISCUSSION:
Section 3.9 (page 71) of RFC 793 [Postel, 1981c] states that if a
SYN segment is received with a valid (i.e., "in window") Sequence
Number, an RST segment should be sent in response, and the
connection should be aborted.
The IETF has published an RFC, "Improving TCP's Resistance to
Blind In-Window Attacks" [Ramaiah et al, 2008] which addresses,
among others, this variant of TCP-based connection-reset attack.
This section describes the counter-measure proposed by the IETF, a
problem that may arise from the implementation of that solution,
and a workaround to it.
In order to mitigate this attack vector, [Ramaiah et al, 2008]
proposes to change TCP's reaction to SYN segments as follows.
When a SYN segment is received for a connection in any of the
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synchronized states, an Acknowledgement (ACK) segment is sent in
response.
As discussed in [Ramaiah et al, 2008], there is a corner-case that
would not be properly handled by this mechanism. If a host (TCP
A) establishes a TCP connection with a remote peer (TCP B), and
then crashes, reboots and tries to initiate a new incarnation of
the same connection (i.e., a connection with the same four-tuple
as the previous connection) using an Initial Sequence Number equal
to the RCV.NXT value at the remote peer (TCP B), the ACK segment
sent by TCP B in response to the SYN segment would contain an
Acknowledgement number that would be considered valid by TCP A,
and thus an RST segment would not be sent in response to the
Acknowledgement (ACK) segment. As this ACK would not have the SYN
bit set, TCP A (being in the SYN-SENT state) would silently drop
it (as stated on page 68 of RFC 793). After a Retransmission
Timeout (RTO), TCP A would retransmit its SYN segment, which would
lead to the same sequence of events as before. Eventually, TCP A
would timeout, and the connection would be aborted. This is a
corner case in which the introduced change would lead to a non-
desirable behavior. However, we consider this scenario to be
extremely unlikely and, in the event it ever took place, the
connection would nevertheless be aborted after retrying for a
period of USER TIMEOUT seconds.
However, when this change is implemented exactly as described in
[Ramaiah et al, 2008], the potential of interoperability problems
is introduced, as a heuristic widely incorporated in many TCP
implementations is disabled.
In a number of scenarios a socket pair may need to be reused while
the corresponding four-tuple is still in the TIME-WAIT state in a
remote TCP peer. For example, a client accessing some service on
a host may try to create a new incarnation of a previous
connection, while the corresponding four-tuple is still in the
TIME-WAIT state at the remote TCP peer (the server). This may
happen if the ephemeral port numbers are being reused too quickly,
either because of a bad policy of selection of ephemeral ports, or
simply because of a high connection rate to the corresponding
service. In such scenarios, the establishment of new connections
that reuse a four-tuple that is in the TIME-WAIT state would fail.
In order to avoid this problem, RFC 1122 [Braden, 1989] states (in
Section 4.2.2.13) that when a connection request is received with
a four-tuple that is in the TIME-WAIT state, the connection
request could be accepted if the sequence number of the incoming
SYN segment is greater than the last sequence number seen on the
previous incarnation of the connection (for that direction of the
data transfer).
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This requirement aims at avoiding the sequence number space of the
new and old incarnations of the connection to overlap, thus
avoiding old segments from the previous incarnation of the
connection to be accepted as valid by the new connection.
The requirement in [Ramaiah et al, 2008] to disregard SYN segments
received for connections in any of the synchronized states forbids
the implementation of the heuristic described above. As a result,
we argue that the processing of SYN segments proposed in [Ramaiah
et al, 2008] should apply only for connections in any of the
synchronized states other than the TIME-WAIT state.
11.1.3. Security/Compartment
If the security/compartment field of an incoming TCP segment does not
match the value recorded in the corresponding TCB, TCP SHOULD NOT
abort the connection, but simply discard the corresponding packet.
Additionally, this whole event SHOULD be logged as a security
violation.
DISCUSSION:
Section 3.9 (page 71) of RFC 793 [Postel, 1981c] states that if
the IP security/compartment of an incoming segment does not
exactly match the security/compartment in the TCB, a RST segment
should be sent, and the connection should be aborted.
A discussion of the IP security options relevant to this section
can be found in Section 3.13.2.12, Section 3.13.2.13, and Section
3.13.2.14 of [CPNI, 2008].
This certainly provides another attack vector for performing
connection-reset attacks, as an attacker could forge TCP segments
with a security/compartment that is different from that recorded
in the corresponding TCB and, as a result, the attacked connection
would be reset.
It is interesting to note that for connections in the ESTABLISHED
state, this check is performed after validating the TCP Sequence
Number and checking the RST bit, but before validating the
Acknowledgement field. Therefore, even if the stricter validation
of the Acknowledgement field (described in Section 3.4) was
implemented, it would not help to mitigate this attack vector.
This attack vector can be easily mitigated by relaxing the
reaction to TCP segments with "incorrect" security/compartment
values as specified in this section.
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11.1.4. Precedence
If the Precedence field of an incomming TCP segment does not match
the value recorded in the corresponding TCB, TCP MUST NOT abort the
connection, and MUST instead continue processing the segment as
specified by RFC 793.
DISCUSSION:
Section 3.9 (page 71) of RFC 793 [Postel, 1981c] states that if
the IP Precedence of an incoming segment does not exactly match
the Precedence recorded in the TCB, a RST segment should be sent,
and the connection should be aborted.
This certainly provides another attack vector for performing
connection-reset attacks, as an attacker could forge TCP segments
with a IP Precedence that is different from that recorded in the
corresponding TCB and, as a result, the attacked connection would
be reset.
It is interesting to note that for connections in the ESTABLISHED
state, this check is performed after validating the TCP Sequence
Number and checking the RST bit, but before validating the
Acknowledgement field. Therefore, even if the stricter validation
of the Acknowledgement field (described in Section 3.4) were
implemented, it would not help to mitigate this attack vector.
This attack vector can be easily mitigated by relaxing the
reaction to TCP segments with "incorrect" IP Precedence values.
That is, even if the Precedence field does not match the value
recorded in the corresponding TCB, TCP should not abort the
connection, and should instead continue processing the segment as
specified by RFC 793.
It is interesting to note that resetting a connection due to a
change in the Precedence value might have a negative impact on
interoperability. For example, the packets that correspond to the
connection could temporarily take a different internet path, in
which some middle-box could re-mark the Precedence field (due to
administration policies at the network to be transited). In such
a scenario, an implementation following the advice in RFC 793
would abort the connection, when the connection would have
probably survived.
While the IPv4 Type of Service field (and hence the Precedence
field) has been redefined by the Differentiated Services (DS)
field specified in RFC 2474 [Nichols et al, 1998], RFC 793
[Postel, 1981c] was never formally updated in this respect. We
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note that both legacy systems that have not been upgraded to
implement the differentiated services architecture described in
RFC 2475 [Blake et al, 1998] and current implementations that have
extrapolated the discussion of the Precedence field to the
Differentiated Services field may still be vulnerable to the
connection reset vector discussed in this section.
11.1.5. Illegal options
TCP MUST silently drop those TCP segments that contain TCP options
with illegal option lengths.
DISCUSSION:
Section 4.2.2.5 of RFC 1122 [Braden, 1989] discusses the
processing of TCP options. It states that TCP must be able to
receive a TCP option in any segment, and must ignore without error
any option it does not implement. Additionally, it states that
TCP should be prepared to handle an illegal option length (e.g.,
zero) without crashing, and suggests handling such illegal options
by resetting the corresponding connection and logging the reason.
However, this suggested behavior could be exploited to perform
connection-reset attacks. Therefore, as discussed in Section 3.10
of this document, we advise TCP implementations to silently drop
those TCP segments that contain illegal option lengths.
11.2. Blind data-injection attacks
An attacker could try to inject data in the stream of data being
transferred on the connection. As with the other attacks described
in Section 11 of this document, in order to perform a blind data
injection attack the attacker would need to know or guess the four-
tuple that identifies the TCP connection to be attacked.
Additionally, he should be able to guess a valid ("in window") TCP
Sequence Number, and a valid Acknowledgement Number.
As discussed in Section 3.4 of this document, [Ramaiah et al, 2008]
proposes to enforce a more strict check on the Acknowledgement Number
of incoming segments than that specified in RFC 793 [Postel, 1981c].
Implementation of the proposed check requires more packets on the
side of the attacker to successfully perform a blind data-injection
attack. However, it should be noted that applications concerned with
any of the attacks discussed in Section 11 of this document should
make use of proper authentication techniques, such as those specified
for IPsec in RFC 4301 [Kent and Seo, 2005].
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12. Information leaking
12.1. Remote Operating System detection via TCP/IP stack fingerprinting
Clearly, remote Operating System (OS) detection is a useful tool for
attackers. Tools such as nmap [Fyodor, 2006b] can usually detect the
operating system type and version of a remote system with an
amazingly accurate precision. This information can in turn be used
by attackers to tailor their exploits to the identified operating
system type and version.
Evasion of OS fingerprinting can prove to be a very difficult task.
Most systems make use of a variety of protocols, each of which have a
large number of parameters that can be set to arbitrary values.
Thus, information on the operating system may be obtained from a
number of sources ranging from application banners to more obscure
parameters such as TCP's retransmission timer.
Nmap [Fyodor, 2006b] is probably the most popular tool for remote OS
detection via active TCP/IP stack fingerprinting. p0f [Zalewski,
2006a], on the other hand, is a tool for performing remote OS
detection via passive TCP/IP stack fingerprinting. SinFP [SinFP,
2006] can perform both active and passive fingerprinting. Finally,
TBIT [TBIT, 2001] is a TCP fingerprinting tool that aims at
characterizing the behavior of a remote TCP peer based on active
probes, and which has been widely used in the research community.
TBIT [TBIT, 2001] implements a number of tests not present in other
tools, such as characterizing the behavior of a TCP peer with respect
to TCP congestion control.
[Fyodor, 1998] and [Fyodor, 2006a] are classic papers on the subject.
[Miller, 2006] and [Smith and Grundl, 2002] provide an introduction
to passive TCP/IP stack fingerprinting. [Smart et al, 2000] and
[Beck, 2001] discuss some techniques for evading OS detection through
TCP/IP stack fingerprinting.
The following subsections discuss TCP-based techniques for remote OS
detection via and, where possible, propose ways to mitigate them.
12.1.1. FIN probe
TCP MUST silently drop TCP any segments received for a connection in
the LISTEN state that do not have the SYN, RST, or ACK flags set. In
the rest of the cases, the processing rules in RFC 793 MUST be
applied.
DISCUSSION:
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The attacker sends a FIN (or any packet without the SYN or the ACK
flags set) to an open port. RFC 793 [Postel, 1981c] leaves the
reaction to such segments unspecified. As a result, some
implementations silently drop the received segment, while others
respond with a RST.
12.1.2. Bogus flag test
TCP MUST ignore any flags not supported, and MUST NOT reflect them if
a TCP segment is sent in response to the one just received.
DISCUSSION:
The attacker sends a TCP segment setting at least one bit of the
Reserved field. Some implementations ignore this field, while
others reset the corresponding connection or reflect the field in
the TCP segment sent in response.
12.1.3. TCP ISN sampling
The attacker samples a number of Initial Sequence Numbers by sending
a number of connection requests. Many TCP implementations differ on
the ISN generator they implement, thus allowing the correlation of
ISN generation algorithm to the operating system type and version.
This document advises implementing an ISN generator that follows the
behavior described in RFC 1948 [Bellovin, 1996]. However, it should
be noted that even if all TCP implementations generated their ISNs as
proposed in RFC 1948, there is still a number of implementation
details that are left unspecified, which would allow remote OS
fingerprinting by means of ISN sampling. For example, the time-
dependent parameter of the hash could have a different frequency in
different TCP implementations.
12.1.4. TCP initial window
Many TCP implementations differ on the initial TCP window they use.
There are a number of factors that should be considered when
selecting the TCP window to be used for a given system. A number of
implementations that use static windows (i.e., no automatic buffer
tuning mechanisms are implemented) default to a window of around 32
KB, which seems sensible for the general case. On the other hand, a
window of 4 KB seems to be common practice for connections servicing
critical applications such as BGP. It is clear that the window size
is a tradeoff among a number of considerations. Section 3.7
discusses some of the considerations that should be made when
selecting the window size for a TCP connection.
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If automatic tuning mechanisms are implemented, we suggest the
initial window to be at least 4 * RMSS segments. We note that a
remote OS fingerprinting tool could still sample the advertised TCP
window, trying to correlate the advertised window with the potential
automatic buffer tuning algorithm and Operating System.
12.1.5. RST sampling
If an RST must be sent in response to an incoming segment, then if
the ACK bit of an incoming TCP segment is off, a Sequence Number of
zero MUST be used in the RST segment sent in response. That is,
<SEQ=0><ACK=SEG.SEQ+SEG.LEN><CTL=RST, ACK>
It should be noted that the SEG.LEN value used for the
Acknowledgement Number MUST be incremented once for each flag set in
the original segment that makes use of a byte of the sequence number
space. That is, if only one of the SYN or FIN flags were set in the
received segment, the Acknowledgement Number of the response should
be set to SEG.SEQ+SEG.LEN+1. If both the SYN and FIN flags were set
in the received segment, the Acknowledgement Number should be set to
SEG.SEQ+SEG.LEN+2.
We also RECOMMEND that TCP sets ACK bit (and the Acknowledgement
Number) in all outgoing RST segments, as it allows for additional
validation checks to be enforced at the system receiving the segment.
DISCUSSION:
[Fyodor, 1998] reports that many implementations differ in the
Acknowledgement Number they use in response to segments received
for connections in the CLOSED state. In particular, these
implementations differ in the way they construct the RST segment
that is sent in response to those TCP segments received for
connections in the CLOSED state.
RFC 793 [Postel, 1981c] describes (in pages 36-37) how RST
segments are to be generated. According to this RFC, the ACK bit
(and the Acknowledgment Number) is set in a RST only if the
incoming segment that elicited the RST did not have the ACK bit
set (and thus the Sequence Number of the outgoing RST segment must
be set to zero). However, we recommend TCP implementations to set
the ACK bit (and the Acknowledgement Number) in all outgoing RST
segments, as it allows for additional validation checks to be
enforced at the system receiving the segment.
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12.1.6. TCP options
Different implementations differ in the TCP options they enable by
default. Additionally, they differ in the actual contents of the
options, and in the order in which the options are included in a TCP
segment. There is currently no recommendation on the order in which
to include TCP options in TCP segments.
12.1.7. Retransmission Timeout (RTO) sampling
TCP uses a retransmission timer for retransmitting data in the
absence of any feedback from the remote data receiver. The duration
of this timer is referred to as "retransmission timeout" (RTO). RFC
2988 [Paxson and Allman, 2000] specifies the algorithm for computing
the TCP retransmission timeout (RTO).
The algorithm allows the use of clocks of different granularities, to
accommodate the different granularities used by the existing
implementations. Thus, the difference in the resulting RTO can be
used for remote OS fingerprinting. [Veysset et al, 2002] describes
how to perform remote OS fingerprinting by sampling and analyzing the
RTO of the target system. However, this fingerprinting technique has
at least the following drawbacks:
o It is usually much slower than other fingerprinting techniques, as
it may require considerable time to sample the RTO of a given
target.
o It is less reliable than other fingerprinting techniques, as
latency and packet loss can lead to bogus results.
While in principle it would be possible to defeat this fingerprinting
technique (e.g., by obfuscating the granularity of the clock used for
computing the RTO), we consider that a more important step to defeat
remote OS detection is for implementations to address the more
effective fingerprinting techniques described in Sections 12.1.1
through 12.1.7 of this document.
12.2. System uptime detection
The "uptime" of a system may prove to be valuable information to an
attacker. For example, it might reveal the last time a security
patch was applied. Information about system uptime is usually leaked
by TCP header fields or options that are (or may be) time-dependent,
and are usually initialized to zero when the system is bootstrapped.
As a result, if the attacker knows the frequency with which the
corresponding parameter or header field is incremented, and is able
to sample the current value of that parameter or header field, the
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system uptime will be easily obtained. Two fields that can
potentially reveal the system uptime is the Sequence Number field of
a SYN or SYN/ACK segment (i.e., when it contains an ISN) and the
TSval field of the timestamp option. Section 3.3.1 of this document
discusses the generation of TCP Initial Sequence Numbers. Section
4.7.1 of this document discusses the generation of TCP timestamps.
13. Covert channels
As virtually every communications protocol, TCP can be exploited to
establish covert channels. While an exhaustive discussion of covert
channels is out of the scope of this document, for completeness of
the document we simply note that it is possible for a (probably
malicious) user to establish a covert channel by means of TCP, such
that data can be surreptitiously passed to a remote system, probably
unnoticed by a monitoring system, and with the possibility of
concealing the location of the source system.
In most cases, covert channels based on manipulation of TCP fields
can be eliminated by protocol scrubbers and other middle-boxes. On
the other hand, "timing channels" may prove to be more difficult to
eliminate.
[Rowland, 1996] contains a discussion of covert channels in the
TCP/IP protocol suite, with some TCP-based examples. [Giffin et al,
2002] describes the use of TCP timestamps for the establishment of
covert channels. [Zander, 2008] contains an extensive bibliography
of papers on covert channels, and a list of freely-available tools
that implement covert channels with the TCP/IP protocol suite.
14. TCP Port scanning
TCP port scanning aims at identifying TCP port numbers on which there
is a process listening for incoming connections. That is, it aims at
identifying TCPs at the target system that are in the LISTEN state.
The following subsections describe different TCP port scanning
techniques that have been implemented in freely-available tools.
These subsections focus only on those port scanning techniques that
exploit features of TCP itself, and not of other communication
protocols.
For example, the following subsections do not discuss the
exploitation of application protocols (such as FTP) or the
exploitation of features of underlying protocols (such as the IP
Identification field) for port-scanning purposes.
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14.1. Traditional connect() scan
The most trivial scanning technique consists in trying to perform the
TCP three-way handshake with each of the port numbers at the target
system (e.g. by issuing a call to the connect() function of the
Sockets API). The three-way handshake will complete for port numbers
that are "open", but will fail for those port numbers that are
"closed".
As this port-scanning technique can be implemented by issuing a call
to the connect() function of the Sockets API that normal applications
use, it does not require the attacker to have superuser privileges.
The downside of this port-scanning technique is that it is less
efficient than other scanning methods (e.g., the "SYN scan" described
in Section 14.2), and that it can be easily logged by the target
system.
14.2. SYN scan
The SYN scan was introduced as a "stealth" port-scanning technique.
It aims at avoiding the target system from logging the port scan by
not completing the TCP three-way handshake. When a SYN/ACK segment
is received in response to the initial SYN segment, the system
performing the port scan will respond with an RST segment, thus
preventing the three-way handshake from completing. While this port-
scanning technique is harder to detect and log than the traditional
connect() scan described in Section 14.1, most current NIDS (Network
Intrusion Detection Systems) can detect and log it.
SYN scans are sometimes mistakenly reported as "SYN flood" attacks by
NIDS, though.
The main advantage of this port scanning technique is that it is much
more efficient than the traditional connect() scan.
In order to implement this port-scanning technique, port-scanning
tools usually bypass the TCP API, and forge the SYN segments they
send (e.g., by using raw sockets). This typically requires the
attacker to have superuser privileges to be able to run the port-
scanning tool.
14.3. FIN, NULL, and XMAS scans
TCP SHOULD respond with an RST when a TCP segment is received for a
connection in the LISTEN state, and the incoming segment has neither
the SYN bit nor the RST bit set.
DISCUSSION:
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RFC 793 [Postel, 1981c] states, in page 65, that an incoming
segment that does not have the RST bit set and that is received
for a connection in the fictional state CLOSED causes an RST to be
sent in response. Pages 65-66 of RFC 793 describes the processing
of incoming segments for connections in the state LISTEN, and
implicitly states that an incoming segment that does not have the
ACK bit set (and is not a SYN or an RST) should be silently
dropped.
As a result, an attacker can exploit this situation to perform a
port scan by sending TCP segments that do not have the ACK bit set
to the target system. When a port is "open" (i.e., there is a TCP
in the LISTEN state on the corresponding port), the target system
will respond with an RST segment. On the other hand, if the port
is "closed" (i.e., there is a TCP in the fictional state CLOSED)
the attacker will not get any response from the target system.
Since the only requirement for exploiting this port scanning
vector is that the probe segments must not have the ACK bit set,
there are a number of different TCP control-bits combinations that
can be used for the probe segments.
When the probe segment sent to the target system is a TCP segment
that has only the FIN bit set, the scanning technique is usually
referred to as a "FIN scan". When the probe packet is a TCP
segment that does not have any of the control bits set, the
scanning technique is usually known as a "NULL scan". Finally,
when the probe packet sent to the target system has only the FIN,
PSH, and the URG bits set, the port-scanning technique is known as
a "XMAS scan".
It should be clear that while the aforementioned control-bits
combinations are the most popular ones, other combinations could
be used to exploit this port-scanning vector. For example, the
CWR, ECE, and/or any of the Reserved bits could be set in the
probe segments.
The advantage of this port-scanning technique is that in can
bypass some stateless firewalls. However, the downside is that a
number of implementations do not comply strictly with RFC 793
[Postel, 1981c], and thus always respond to the probe segments
with an RST, regardless of whether the port is open or closed.
This port-scanning vector can be easily defeated as rby responding
with an RST when a TCP segment is received for a connection in the
LISTEN state, and the incoming segment has neither the SYN bit nor
the RST bit set.
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14.4. Maimon scan
If a TCP that is in the CLOSED or LISTEN states receives a TCP
segment with both the FIN and ACK bits set, it MUST respond with a
RST.
DISCUSSION:
This port scanning technique was introduced in [Maimon, 1996] with
the name "StealthScan" (method #1), and was later incorporated
into the nmap tool [Fyodor, 2006b] as the "Maimon scan".
This port scanning technique employs TCP segments that have both
the FIN and ACK bits sets as the probe segments. While according
to RFC 793 [Postel, 1981c] these segments should elicit an RST
regardless of whether the corresponding port is open or closed, a
programming flaw found in a number of TCP implementations has
caused some systems to silently drop the probe segment if the
corresponding port was open (i.e., there was a TCP in the LISTEN
state), and respond with an RST only if the port was closed.
Therefore, an RST would indicate that the scanned port is closed,
while the absence of a response from the target system would
indicate that the scanned port is open.
While this bug has not been found in current implementations of
TCP, it might still be present in some legacy systems.
14.5. Window scan
When sending an RST segment, TCP SHOULD set the Window field to zero.
DISCUSSION:
This port-scanning technique employs ACK segments as the probe
packets. ACK segments will elicit an RST from the target system
regardless of whether the corresponding TCP port is open or
closed. However, as described in [Maimon, 1996], some systems set
the Window field of the RST segments with different values
depending on whether the corresponding TCP port is open or closed.
These systems set the Window field of their RST segments to zero
when the corresponding TCP port is closed, and set the Window
field to a non-zero value when the corresponding TCP port is open.
As a result, an attacker could exploit this situation for
performing a port scan by sending ACK segments to the target
system, and examining the Window field of the RST segments that
his probe segments elicit.
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In order to defeat this port-scanning technique, we recommend TCP
implementations to set the Window field to zero in all the RST
segments they send. Most popular implementations of TCP already
implement this policy.
14.6. ACK scan
The so-called "ACK scan" is not really a port-scanning technique
(i.e., it does not aim at determining whether a specific port is open
or closed), but rather aims at determining whether some intermediate
system is filtering TCP segments sent to that specific port number.
The probe packet is a TCP segment with the ACK bit set which,
according to RFC 793 [Postel, 1981c] should elicit an RST from the
target system regardless of whether the corresponding TCP port is
open or closed. If no response is received from the target system,
it is assumed that some intermediate system is filtering the probe
packets sent to the target system.
It should be noted that this "port scanning" techniques exploits
basic TCP processing rules, and therefore cannot be defeated at an
end-system.
15. Processing of ICMP error messages by TCP
TCP SHOULD silently ignore received ICMP Source Quench messages.
TCP SHOULD process ICMP "hard errors" as "soft errors" when they are
received for connections that are in any of he synchronized states.
TCP SHOULD process ICMP "fragmentation needed and DF bit set" and
ICMPv6 "Packet Too Big" error messages as described in [RFC5927].
DISCUSSION:
[RFC5927] analyzes a number of vulnerabilities based on crafted
ICMP messages, along with possible counter-measures.
16. TCP interaction with the Internet Protocol (IP)
16.1. TCP-based traceroute
The traceroute tool is used to identify the intermediate systems the
local system and the destination system. It is usually implemented
by sending "probe" packets with increasing IP Time to Live values
(starting from 0), without maintaining any state with the final
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destination.
Some traceroute implementations use ICMP "echo request" messages as
the probe packets, while others use UDP packets or TCP SYN segments.
In some cases, the state-less nature of the traceroute tool may
prevent it from working correctly across stateful devices such as
Network Address Translators (NATs) or firewalls.
In order to by-pass this limitation, an attacker could establish a
TCP connection with the destination system, and start sending TCP
segments on that connection with increasing IP Time to Live values
(starting from 0) [Zalewski, 2007] [Zalewski, 2008]. Provided ICMP
error messages are not blocked by any intermediate system, an
attacker could exploit this technique to map the network topology
behind the aforementioned stateful devices in scenarios in which he
could not have achieved this goal using the traditional traceroute
tool.
NATs [Srisuresh and Egevang, 2001] and other middle-boxes could
defeat this network-mapping technique by overwriting the Time to Live
of the packets they forward to the internal network. For example,
they could overwrite the Time to Live of all packets being forwarded
to an internal network with a value such as 128. We strongly
recommend against overwriting the IP Time to Live field with the
value 255 or other similar large values, as this could allow an
attacker to bypass the protection provided by the Generalized TTL
Security Mechanism (GTSM) described in RFC 5087 [Gill et al, 2007].
[Gont and Srisuresh, 2008] discusses the security implications of
NATs, and proposes mitigations for this and other issues.
16.2. Blind TCP data injection through fragmented IP traffic
As discussed in Section 11.2, TCP data injection attacks usually
require an attacker to guess or know a number of parameters related
with the target TCP connection, such as the connection-id {Source
Address, Source Port, Destination Address, Destination Port}, the TCP
Sequence Number, and the TCP Acknowledgement Number. Provided these
values are obfuscated as recommended in this document, the chances of
an off-path attacker of successfully performing a data injection
attack against a TCP connection are fairly low for many of the most
common scenarios.
As discussed in this document, randomization of the values contained
in different TCP header fields is not a replacement for cryptographic
methods for protecting a TCP connection, such as IPsec (specified in
RFC 4301 [Kent and Seo, 2005]).
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However, [Zalewski, 2003b] describes a possible vector for performing
a TCP data injection attack that does not require the attacker to
guess or know the aforementioned TCP connection parameters, and could
therefore be successfully exploited in some scenarios with less
effort than that required to exploit the more traditional data-
injection attack vectors.
The attack vector works as follows. When one system is transferring
information to a remote peer by means of TCP, and the resulting
packet gets fragmented, the first fragment will usually contain the
entire TCP header which, together with the IP header, includes all
the connection parameters that an attacker would need to guess or
know to successfully perform a data injection attack against TCP. If
an attacker were able to forge all the fragments other than the first
one, his forged fragments could be reassembled together with the
legitimate first fragment, and thus he would be relieved from the
hard task of guessing or knowing connection parameters such as the
TCP Sequence Number and the TCP Acknowledgement Number.
In order to successfully exploit this attack vector, the attacker
should be able to guess or know both of the IP addresses involved in
the target TCP connection, the IP Identification value used for the
specific packet he is targeting, and the TCP Checksum of that target
packet. While it would seem that these values are hard to guess, in
some specific scenarios, and with some security-unwise implementation
approaches for the TCP and IP protocols, these values may be feasible
to guess or know. For example, if the sending system uses
predictable IP Identification values, the attacker could simply
perform a brute force attack, trying each of the possible
combinations for the TCP Checksum field. In more specific scenarios,
the attacker could have more detailed knowledge about the data being
transferred over the target TCP connection, which might allow him to
predict the TCP Checksum of the target packet. For example, if both
of the involved TCP peers used predictable values for the TCP
Sequence Number and for the IP Identification fields, and the
attacker knew the data being transferred over the target TCP
connection, he could be able to carefully forge the IP payload of his
IP fragments so that the checksum of the reassembled TCP segment
matched the Checksum included in the TCP header of the first (and
legitimate) IP fragment.
As discussed in Section 4.1 of [CPNI, 2008], IP fragmentation
provides a vector for performing a variety of attacks against an IP
implementation. Therefore, we discourage the reliance on IP
fragmentation by end-systems, and recommend the implementation of
mechanisms for the discovery of the Path-MTU, such as that described
in Section 15.7.3 of this document and/or that described in RFC 4821
[Mathis and Heffner, 2007]. We nevertheless recommend randomization
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of the IP Identification field as described in Section 3.5.2 of
[CPNI, 2008]. While randomization of the IP Identification field
does not eliminate this attack vector, it does require more work on
the side of the attacker to successfully exploit it.
16.3. Broadcast and multicast IP addresses
TCP connection state is maintained between only two endpoints at a
time. As a result, broadcast and multicast IP addresses should not
be allowed for the establishment of TCP connections. Section 4.3 of
[CPNI, 2008] provides advice about which specific IP address blocks
should not be allowed for connection-oriented protocols such as TCP.
17. Security Considerations
This document provides a thorough security assessment of the
Transmission Control Protocol (TCP), identifies a number of
vulnerabilities, and specifies possible counter-measures.
Additionally, it provides implementation guidance such that the
resilience of TCP implementations is improved.
18. Acknowledgements
The author would like to thank (in alphabetical order) David Borman,
Wesley Eddy, and Alfred Hoenes, for providing valuable feedback on
earlier versions of thi document.
This document is heavily based on the document "Security Assessment
of the Transmission Control Protocol (TCP)" [CPNI, 2009] written by
Fernando Gont on behalf of CPNI (Centre for the Protection of
National Infrastructure).
The author would like to thank (in alphabetical order) Randall
Atkinson, Guillermo Gont, Alfred Hoenes, Jamshid Mahdavi, Stanislav
Shalunov, Michael Welzl, Dan Wing, Andrew Yourtchenko, Michal
Zalewski, and Christos Zoulas, for providing valuable feedback on
earlier versions of the UK CPNI document.
Additionally, the author would like to thank (in alphabetical order)
Mark Allman, David Black, Ethan Blanton, David Borman, James Chacon,
John Heffner, Jerrold Leichter, Jamshid Mahdavi, Keith Scott, Bill
Squier, and David White, who generously answered a number of
questions that araised while the aforementioned document was being
written.
Finally, the author would like to thank CPNI (formely NISCC) for
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their continued support.
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20. References
20.1. Normative References
[I-D.ietf-tcpm-tcp-timestamps]
Gont, F., "Reducing the TIME-WAIT state using TCP
timestamps", draft-ietf-tcpm-tcp-timestamps-03 (work in
progress), December 2010.
[I-D.ietf-tsvwg-port-randomization]
Larsen, M. and F. Gont, "Transport Protocol Port
Randomization Recommendations",
draft-ietf-tsvwg-port-randomization-09 (work in progress),
August 2010.
[RFC6093] Gont, F. and A. Yourtchenko, "On the Implementation of the
TCP Urgent Mechanism", RFC 6093, January 2011.
20.2. Informative References
[I-D.gont-timestamps-generation]
Gont, F. and A. Oppermann, "On the generation of TCP
timestamps", draft-gont-timestamps-generation-00 (work in
progress), June 2010.
[RFC5927] Gont, F., "ICMP Attacks against TCP", RFC 5927, July 2010.
Appendix A. TODO list
A Number of formatting issues still have to be fixed in this
document. Among others are:
o The ASCII-art corresponding to some figures are still missing. We
still have to convert the nice JPGs of the UK CPNI document into
ugly ASCII-art.
o The references have not yet been converted to xml, but are
hardcoded, instead. That's why they may not look as expected
Appendix B. Change log (to be removed by the RFC Editor before
publication of this document as an RFC)
B.1. Changes from draft-ietf-tcpm-tcp-security-01
A Number of formatting issues still have to be fixed in this
document. Among others are:
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o The whole document was reformatted with RFC 1122 style.
Author's Address
Fernando Gont
UK Centre for the Protection of National Infrastructure
Email: fernando@gont.com.ar
URI: http://www.cpni.gov.uk
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