TCP Maintenance (TCPM) D. Borman
Internet-Draft Quantum Corporation
Intended status: Standards Track B. Braden
Expires: February 07, 2014 University of Southern California
V. Jacobson
Google, Inc.
R. Scheffenegger, Ed.
NetApp, Inc.
August 06, 2013

TCP Extensions for High Performance


This document specifies a set of TCP extensions to improve performance over paths with a large bandwidth * delay product and to provide reliable operation over very high-speed paths. It defines TCP options for scaled windows and timestamps. The timestamps are used for two distinct mechanisms, RTTM (Round Trip Time Measurement) and PAWS (Protection Against Wrapped Sequences).

This document obsoletes RFC 1323 and describes changes from it.

Status of This Memo

This Internet-Draft is submitted 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

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 February 07, 2014.

Copyright Notice

Copyright (c) 2013 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 ( 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 described in the Simplified BSD License.

Table of Contents

1. Introduction

The TCP protocol [RFC0793] was designed to operate reliably over almost any transmission medium regardless of transmission rate, delay, corruption, duplication, or reordering of segments. Over the years, advances in networking technology has resulted in ever-higher transmission speeds, and the fastest paths are well beyond the domain for which TCP was originally engineered.

This document defines a set of modest extensions to TCP to extend the domain of its application to match the increasing network capability. It is an update to and obsoletes [RFC1323], which in turn is based upon and obsoletes [RFC1072] and [RFC1185].

Changes between [RFC1323] and this document are detailed in Appendix H.

For brevity, the full discussions of the merits and history behind the TCP options defined within this document have been omitted. [RFC1323] should be consulted for reference. It is recommended that a modern TCP stack implements and make use of the extensions described in this document.

1.1. TCP Performance

TCP performance problems arise when the bandwidth * delay product is large. A network having such paths is referred to as "long, fat network" (LFN).

There are two fundamental performance problems with basic TCP over LFN paths:

Window Size Limit

The TCP header uses a 16 bit field to report the receive window size to the sender. Therefore, the largest window that can be used is 2^16 = 64 KiB. For LFN paths where the bandwidth * delay product exceeds 64 KiB, the receive window limits the maximum throughput of the TCP connection over the path, i.e., the amount of unacknowledged data that TCP can send in order to keep the pipeline full.

To circumvent this problem, Section 2 of this memo defines a TCP option, "Window Scale", to allow windows larger than 2^16. This option defines an implicit scale factor, which is used to multiply the window size value found in a TCP header to obtain the true window size.
Recovery from Losses

Packet losses in an LFN can have a catastrophic effect on throughput.

To generalize the Fast Retransmit / Fast Recovery mechanism to handle multiple packets dropped per window, Selective Acknowledgments are required. Unlike the normal cumulative acknowledgments of TCP, Selective Acknowledgments give the sender a complete picture of which segments are queued at the receiver and which have not yet arrived.

Selective acknowledgements and their use are specified in separate documents, "TCP Selective Acknowledgment Options" [RFC2018], "An Extension to the Selective Acknowledgement (SACK) Option for TCP" [RFC2883], and "A Conservative Selective Acknowledgment (SACK)-based Loss Recovery Algorithm for TCP" [RFC6675], and not further discussed in this document.

1.2. TCP Reliability

An especially serious kind of error may result from an accidental reuse of TCP sequence numbers in data segments. TCP reliability depends upon the existence of a bound on the lifetime of a segment: the "Maximum Segment Lifetime" or MSL.

Duplication of sequence numbers might happen in either of two ways:

Sequence number wrap-around on the current connection

A TCP sequence number contains 32 bits. At a high enough transfer rate, the 32-bit sequence space may be "wrapped" (cycled) within the time that a segment is delayed in queues.
Earlier incarnation of the connection

Suppose that a connection terminates, either by a proper close sequence or due to a host crash, and the same connection (i.e., using the same pair of port numbers) is immediately reopened. A delayed segment from the terminated connection could fall within the current window for the new incarnation and be accepted as valid.

Duplicates from earlier incarnations, case (2), are avoided by enforcing the current fixed MSL of the TCP specification, as explained in Section 4.8 and Appendix B. However, case (1), avoiding the reuse of sequence numbers within the same connection, requires an upper bound on MSL that depends upon the transfer rate, and at high enough rates, a dedicated mechanism is required.

A possible fix for the problem of cycling the sequence space would be to increase the size of the TCP sequence number field. For example, the sequence number field (and also the acknowledgment field) could be expanded to 64 bits. This could be done either by changing the TCP header or by means of an additional option.

Section 4 presents a different mechanism, which we call PAWS (Protection Against Wrapped Sequence numbers), to extend TCP reliability to transfer rates well beyond the foreseeable upper limit of network bandwidths. PAWS uses the TCP Timestamps option defined in Section 3.2 to protect against old duplicates from the same connection.

1.3. Using TCP options

The extensions defined in this document all use TCP options.

When [RFC1323] was published, there was concern that some buggy TCP implementation might be crashed by the first appearance of an option on a non-<SYN> segment. However, bugs like that can lead to DOS attacks against a TCP. Research has shown that most TCP implementations will properly handle unknown options on non-<SYN> segments ([Medina04], [Medina05]). But it is still prudent to be conservative in what you send, and avoiding buggy TCP implementation is not the only reason for negotiating TCP options on <SYN> segments.

The window scale option negotiates fundamental parameters of the TCP session. Therefore, it is only sent during the initial handshake. Furthermore, the window scale option will be sent in a <SYN,ACK> segment only if the corresponding option was received in the initial <SYN> segment.

The Timestamps option may appear in any data or <ACK> segment, adding 12 bytes to the 20-byte TCP header. It is required that this TCP option will be sent on all non-<SYN> segments after an exchange of options on the <SYN> segments has indicated that both sides understand this extension.

Research has shown that the use of the Timestamps option to arrive at an optimal retransmission timeout value has only limited benefit ([Allman99]. However, there are other uses of the Timestamps option, such as the Eifel mechanism [RFC3522], [RFC4015], and PAWS (see Section 4) which improve overall TCP security and performance. The extra header bandwidth used by this option should be evaluated for the gains in performance and security in an actual deployment.

Appendix A contains a recommended layout of the options in TCP headers to achieve reasonable data field alignment.

Finally, we observe that most of the mechanisms defined in this document are important for LFN's and/or very high-speed networks. For low-speed networks, it might be a performance optimization to NOT use these mechanisms. A TCP vendor concerned about optimal performance over low-speed paths might consider turning these extensions off for low- speed paths, or allow a user or installation manager to disable them.

1.4. Terminology

The key words "MUST", "MUST NOT", "REQUIRED", "SHALL", "SHALL NOT", "SHOULD", "SHOULD NOT", "RECOMMENDED", "MAY", and "OPTIONAL" in this document are to be interpreted as described in [RFC2119].

In this document, these words will appear with that interpretation only when in UPPER CASE. Lower case uses of these words are not to be interpreted as carrying [RFC2119] significance.

2. TCP Window Scale Option

2.1. Introduction

The window scale extension expands the definition of the TCP window to 30 bits and then uses an implicit scale factor to carry this 30-bit value in the 16-bit Window field of the TCP header (SEG.WND in [RFC0793]). The exponent of the scale factor is carried in a TCP option, Window Scale. This option is sent only in a <SYN> segment (a segment with the SYN bit on), hence the window scale is fixed in each direction when a connection is opened.

The maximum receive window, and therefore the scale factor, is determined by the maximum receive buffer space. In a typical modern implementation, this maximum buffer space is set by default but can be overridden by a user program before a TCP connection is opened. This determines the scale factor, and therefore no new user interface is needed for window scaling.

2.2. Window Scale Option

The three-byte Window Scale option MAY be sent in a <SYN> segment by a TCP. It has two purposes: (1) indicate that the TCP is prepared to do both send and receive window scaling, and (2) communicate the exponent of a scale factor to be applied to its receive window. Thus, a TCP that is prepared to scale windows SHOULD send the option, even if its own scale factor is 1 and the exponent 0. The scale factor is limited to a power of two and encoded logarithmically, so it may be implemented by binary shift operations. The maximum scale exponent is limited to 14 for a maximum permissible receive window size of 1 GiB (2^(14+16)).

TCP Window Scale Option (WSopt):

Kind: 3

Length: 3 bytes

       | Kind=3  |Length=3 |shift.cnt|
            1         1         1

This option is an offer, not a promise; both sides MUST send Window Scale options in their <SYN> segments to enable window scaling in either direction. If window scaling is enabled, then the TCP that sent this option will right-shift its true receive-window values by 'shift.cnt' bits for transmission in SEG.WND. The value 'shift.cnt' MAY be zero (offering to scale, while applying a scale factor of 1 to the receive window).

This option MAY be sent in an initial <SYN> segment (i.e., a segment with the SYN bit on and the ACK bit off). It MAY also be sent in a <SYN,ACK> segment, but only if a Window Scale option was received in the initial <SYN> segment. A Window Scale option in a segment without a SYN bit MUST be ignored.

The window field in a segment where the SYN bit is set (i.e., a <SYN> or <SYN,ACK>) is never scaled.

2.3. Using the Window Scale Option

SND.WND = SEG.WND << Snd.Wind.Shift              

SEG.WND = RCV.WND >> Rcv.Wind.Shift              

A model implementation of window scaling is as follows, using the notation of [RFC0793]:

max window < 2^30

TCP determines if a data segment is "old" or "new" by testing whether its sequence number is within 2^31 bytes of the left edge of the window, and if it is not, discarding the data as "old". To insure that new data is never mistakenly considered old and vice versa, the left edge of the sender's window has to be at most 2^31 away from the right edge of the receiver's window. Similarly with the sender's right edge and receiver's left edge. Since the right and left edges of either the sender's or receiver's window differ by the window size, and since the sender and receiver windows can be out of phase by at most the window size, the above constraints imply that two times the maximum window size must be less than 2^31, or

Since the max window is 2^S (where S is the scaling shift count) times at most 2^16 - 1 (the maximum unscaled window), the maximum window is guaranteed to be < 2^30 if S <= 14. Thus, the shift count MUST be limited to 14 (which allows windows of 2^30 = 1 GiB). If a Window Scale option is received with a shift.cnt value larger than 14, the TCP SHOULD log the error but MUST use 14 instead of the specified value. This is safe as a sender can always choose to only partially use any signaled receive window.

The scale factor applies only to the Window field as transmitted in the TCP header; each TCP using extended windows will maintain the window values locally as 32-bit numbers. For example, the "congestion window" computed by Slow Start and Congestion Avoidance (see [RFC5681]) is not affected by the scale factor, so window scaling will not introduce quantization into the congestion window.

2.4. Addressing Window Retraction

When a non-zero scale factor is in use, there are instances when a retracted window can be offered - see Appendix F for a detailed example. The end of the window will be on a boundary based on the granularity of the scale factor being used. If the sequence number is then updated by a number of bytes smaller than that granularity, the TCP will have to either advertise a new window that is beyond what it previously advertised (and perhaps beyond the buffer), or will have to advertise a smaller window, which will cause the TCP window to shrink. Implementations MUST ensure that they handle a shrinking window, as specified in section of [RFC1122].

For the receiver, this implies that:

The receiver MUST honor, as in-window, any segment that would have been in-window for any <ACK> sent by the receiver.
When window scaling is in effect, the receiver SHOULD track the actual maximum window sequence number (which is likely to be greater than the window announced by the most recent <ACK>, if more than one segment has arrived since the application consumed any data in the receive buffer).

On the sender side:

The initial transmission MUST be within the window announced by the most recent <ACK>.
On first retransmission, or if the sequence number is out-of-window by less than 2^Rcv.Wind.Shift then do normal retransmission(s) without regard to receiver window as long as the original segment was in window when it was sent.
Subsequent retransmissions MAY only be sent, if they are within the window announced by the most recent <ACK>.

3. TCP Timestamps option

3.1. Introduction

TCP measures the round trip time (RTT), primarily for the purpose of arriving at a reasonable value for the Retransmission Timeout (RTO) timer interval. Accurate and current RTT estimates are necessary to adapt to changing traffic conditions, while a conservative estimate of the RTO interval is necessary to minimize spurious RTOs.

When [RFC1323] was originally written, it was perceived that taking RTT measurements for each segment, and also during retransmissions, would contribute to reduce spurious RTOs, while maintaining the timeliness of necessary RTOs. At the time, RTO was also the only mechanism to make use of the measured RTT. It has been shown, that taking more RTT samples has only a very limited effect to optimize RTOs [Allman99].

This document makes a clear distinction between the round trip time measurement (RTTM) mechanism, and subsequent mechanisms using the RTT signal as input, such as RTO (see Section 3.4).

The Timestamps option is important when large receive windows are used, to allow the use of the PAWS mechanism (see Section 4). Furthermore, the option is useful for all TCP's, since it simplifies the sender and allows the use of additional optimizations such as Eifel ([RFC3522], [RFC4015]) and others.

3.2. Timestamps option

TCP is a symmetric protocol, allowing data to be sent at any time in either direction, and therefore timestamp echoing may occur in either direction. For simplicity and symmetry, we specify that timestamps always be sent and echoed in both directions. For efficiency, we combine the timestamp and timestamp reply fields into a single TCP Timestamps option.

TCP Timestamps option (TSopt):

Kind: 8

Length: 10 bytes

       |Kind=8 |  10   |   TS Value (TSval)  |TS Echo Reply (TSecr)|
           1       1              4                     4

The Timestamps option carries two four-byte timestamp fields. The Timestamp Value field (TSval) contains the current value of the timestamp clock of the TCP sending the option.

The Timestamp Echo Reply field (TSecr) is valid if the ACK bit is set in the TCP header; if it is valid, it echoes a timestamp value that was sent by the remote TCP in the TSval field of a Timestamps option. When TSecr is not valid, its value MUST be zero. However, a value of zero does not imply TSecr being invalid. The TSecr value will generally be from the most recent Timestamps option that was received; however, there are exceptions that are explained below.

A TCP MAY send the Timestamps option (TSopt) in an initial <SYN> segment (i.e., segment containing a SYN bit and no ACK bit), and MAY send a TSopt in other segments only if it received a TSopt in the initial <SYN> or <SYN,ACK> segment for the connection.

Once TSopt has been successfully negotiated (sent and received) during the <SYN>, <SYN,ACK> exchange, TSopt MUST be sent in every non- <RST> segment for the duration of the connection, and SHOULD be sent in an <RST> segment (see Section 4.2 for details). If a non-<RST> segment is received without a TSopt, a TCP SHOULD silently drop the segment. A TCP MUST NOT abort a TCP connection because any segment lacks an expected TSopt.

Implementations are strongly encouraged to follow the above rules for handling a missing Timestamps option, and the order of precedence mentioned in Section 4.3 when deciding on the acceptance of a segment.

If a receiver chooses to accept a segment without an expected Timestamps option, it must be clear that undetectable data corruption may occur.

Such a TCP receiver may experience undetectable wrapped- sequence effects, such as data (payload) corruption or session stalls. In order to maintain the integrity of the payload data, in particular on high speed networks, it is paramount to follow the described processing rules.

However, it has been mentioned that under some circumstances, the above guidelines are too strict, and some paths sporadically suppress the Timestamps option, while maintaining payload integrity. A path behaving in this manner should be deemed unacceptable, but it has been noted that some implementations relax the acceptance rules as a workaround, and allow TCP to run across such paths.

If a TSopt is received on a connection where TSopt was not negotiated in the initial three-way handshake, the TSopt MUST be ignored and the packet processed normally.

In the case of crossing <SYN> segments where one <SYN> contains a TSopt and the other doesn't, both sides MAY send a TSopt in the <SYN,ACK> segment.

TSopt is required for the two mechanisms described in sections 3.3 and 4.2. There are also other mechanisms that rely on the presence of the TSopt, e.g. [RFC3522]. If a TCP stopped sending TSopt at any time during an established session, it interferes with these mechanisms. This update to [RFC1323] describes explicitly the previous assumption (see Section 4.2), that each TCP segment must have TSopt, once negotiated.

3.3. The RTTM Mechanism

RTTM places a Timestamps option in every segment, with a TSval that is obtained from a (virtual) "timestamp clock". Values of this clock MUST be at least approximately proportional to real time, in order to measure actual RTT.

These TSval values are echoed in TSecr values in the reverse direction. The difference between a received TSecr value and the current timestamp clock value provides an RTT measurement.

When timestamps are used, every segment that is received will contain a TSecr value. However, these values cannot all be used to update the measured RTT. The following example illustrates why. It shows a one-way data flow with segments arriving in sequence without loss. Here A, B, C... represent data blocks occupying successive blocks of sequence numbers, and ACK(A),... represent the corresponding cumulative acknowledgments. The two timestamp fields of the Timestamps option are shown symbolically as <TSval=x,TSecr=y>. Each TSecr field contains the value most recently received in a TSval field.

TCP  A                                     TCP B

                <A,TSval=1,TSecr=120> ----->

     <---- <ACK(A),TSval=127,TSecr=1>

                <B,TSval=5,TSecr=127> ----->

     <---- <ACK(B),TSval=131,TSecr=5>

  . . . . . . . . . . . . . . . . . . . . . .

                <C,TSval=65,TSecr=131> ---->

     <---- <ACK(C),TSval=191,TSecr=65>


The dotted line marks a pause (60 time units long) in which A had nothing to send. Note that this pause inflates the RTT which B could infer from receiving TSecr=131 in data segment C. Thus, in one-way data flows, RTTM in the reverse direction measures a value that is inflated by gaps in sending data. However, the following rule prevents a resulting inflation of the measured RTT:

RTTM Rule:
A TSecr value received in a segment MAY be used to update the averaged RTT measurement only if the segment advances the left edge of the send window, i.e. SND.UNA is increased.

Since TCP B is not sending data, the data segment C does not acknowledge any new data when it arrives at B. Thus, the inflated RTTM measurement is not used to update B's RTTM measurement.

3.4. Updating the RTO value

[Ludwig00] and [Floyd05] have highlighted the problem that an unmodified RTO calculation, which is updated with per- packet RTT samples, will truncate the path history too soon. This can lead to an increase in spurious retransmissions, when the path properties vary in the order of a few RTTs, but a high number of RTT samples are taken on a much shorter timescale.

Implementers should note that with timestamps multiple RTTMs can be taken per RTT. The [RFC6298] RTO estimator has weighting factors, alpha and beta, based on an implicit assumption that at most one RTTM will be sampled per RTT. When multiple RTTMs per RTT are available to update the RTO estimator, this implicit assumption must be considered. An implementation suggestion is detailed in Appendix G.

3.5. Which Timestamp to Echo

If more than one Timestamps option is received before a reply segment is sent, the TCP must choose only one of the TSvals to echo, ignoring the others. To minimize the state kept in the receiver (i.e., the number of unprocessed TSvals), the receiver should be required to retain at most one timestamp in the connection control block.

There are three situations to consider:

Delayed ACKs.

Many TCP's acknowledge only every second segment out of a group of segments arriving within a short time interval; this policy is known generally as "delayed ACKs". The data-sender TCP must measure the effective RTT, including the additional time due to delayed ACKs, or else it will retransmit unnecessarily. Thus, when delayed ACKs are in use, the receiver SHOULD reply with the TSval field from the earliest unacknowledged segment.
A hole in the sequence space (segment(s) have been lost).

The sender will continue sending until the window is filled, and the receiver may be generating <ACK>s as these out-of-order segments arrive (e.g., to aid "fast retransmit").

The lost segment is probably a sign of congestion, and in that situation the sender should be conservative about retransmission. Furthermore, it is better to overestimate than underestimate the RTT. An <ACK> for an out-of-order segment SHOULD therefore contain the timestamp from the most recent segment that advanced RCV.NXT.

The same situation occurs if segments are re-ordered by the network.
A filled hole in the sequence space.

The segment that fills the hole and advances the window represents the most recent measurement of the network characteristics. An RTT computed from an earlier segment would probably include the sender's retransmit time-out, badly biasing the sender's average RTT estimate. Thus, the timestamp from the latest segment (which filled the hole) MUST be echoed.

SEG.TSval >= TS.recent and SEG.SEQ <= Last.ACK.sent

An algorithm that covers all three cases is described in the following rules for Timestamps option processing on a synchronized connection:

The connection state is augmented with two 32-bit slots:

TS.Recent holds a timestamp to be echoed in TSecr whenever a segment is sent, and Last.ACK.sent holds the ACK field from the last segment sent. Last.ACK.sent will equal RCV.NXT except when <ACK>s have been delayed.
When a TSopt is sent, its TSecr field is set to the current TS.Recent value.

<A, TSval=1> ------------------->
<B, TSval=2> ------------------->
<C, TSval=3> ------------------->
         <---- <ACK(C), TSecr=1>

<A, TSval=1> ------------------->
         <---- <ACK(A), TSecr=1>
<C, TSval=3> ------------------->
         <---- <ACK(A), TSecr=1>
<B, TSval=2> ------------------->
         <---- <ACK(C), TSecr=2>
<E, TSval=5> ------------------->
         <---- <ACK(C), TSecr=2>
<D, TSval=4> ------------------->
         <---- <ACK(E), TSecr=4>

The following examples illustrate these rules. Here A, B, C... represent data segments occupying successive blocks of sequence numbers, and ACK(A),... represent the corresponding acknowledgment segments. Note that ACK(A) has the same sequence number as B. We show only one direction of timestamp echoing, for clarity.

4. PAWS - Protection Against Wrapped Sequence Numbers

4.1. Introduction

Section 4.2 describes a simple mechanism to reject old duplicate segments that might corrupt an open TCP connection; we call this mechanism PAWS (Protection Against Wrapped Sequence numbers). PAWS operates within a single TCP connection, using state that is saved in the connection control block. Section 4.8 and Appendix H discuss the implications of the PAWS mechanism for avoiding old duplicates from previous incarnations of the same connection.

4.2. The PAWS Mechanism

PAWS uses the same TCP Timestamps option as the RTTM mechanism described earlier, and assumes that every received TCP segment (including data and <ACK> segments) contains a timestamp SEG.TSval whose values are monotonically non-decreasing in time. The basic idea is that a segment can be discarded as an old duplicate if it is received with a timestamp SEG.TSval less than some timestamp recently received on this connection.

s < t  if 0 < (t - s) < 2^31, 

In both the PAWS and the RTTM mechanism, the "timestamps" are 32-bit unsigned integers in a modular 32-bit space. Thus, "less than" is defined the same way it is for TCP sequence numbers, and the same implementation techniques apply. If s and t are timestamp values,

The choice of incoming timestamps to be saved for this comparison MUST guarantee a value that is monotonically increasing. For example, we might save the timestamp from the segment that last advanced the left edge of the receive window, i.e., the most recent in-sequence segment. Instead, we choose the value TS.Recent introduced in Section 3.5 for the RTTM mechanism, since using a common value for both PAWS and RTTM simplifies the implementation of both. As Section 3.5 explained, TS.Recent differs from the timestamp from the last in-sequence segment only in the case of delayed <ACK>s, and therefore by less than one window. Either choice will therefore protect against sequence number wrap-around.

RTTM was specified in a symmetrical manner, so that TSval timestamps are carried in both data and <ACK> segments and are echoed in TSecr fields carried in returning <ACK> or data segments. PAWS submits all incoming segments to the same test, and therefore protects against duplicate <ACK> segments as well as data segments. (An alternative non-symmetric algorithm would protect against old duplicate <ACK>s: the sender of data would reject incoming <ACK> segments whose TSecr values were less than the TSecr saved from the last segment whose ACK field advanced the left edge of the send window. This algorithm was deemed to lack economy of mechanism and symmetry.)

TSval timestamps sent on <SYN> and <SYN,ACK> segments are used to initialize PAWS. PAWS protects against old duplicate non- <SYN> segments, and duplicate <SYN> segments received while there is a synchronized connection. Duplicate <SYN> and <SYN,ACK> segments received when there is no connection will be discarded by the normal 3-way handshake and sequence number checks of TCP.

[RFC1323] recommended that <RST> segments NOT carry timestamps, and that they be acceptable regardless of their timestamp. At that time, the thinking was that old duplicate <RST> segments should be exceedingly unlikely, and their cleanup function should take precedence over timestamps. More recently, discussions about various blind attacks on TCP connections have raised the suggestion that if the Timestamps option is present, SEG.TSecr could be used to provide stricter acceptance tests for <RST> segments. While still under discussion, to enable research into this area it is now RECOMMENDED that when generating an <RST>, that if the segment causing the <RST> to be generated contained a Timestamps option, that the <RST> also contain a Timestamps option. In the <RST> segment, SEG.TSecr SHOULD be set to SEG.TSval from the incoming segment and SEG.TSval SHOULD be set to zero. If an <RST> is being generated because of a user abort, and Snd.TS.OK is set, then a Timestamps option SHOULD be included in the <RST>. When an <RST> segment is received, it MUST NOT be subjected to PAWS checks, and information from the Timestamps option MUST NOT be used to update connection state information. SEG.TSecr MAY be used to provide stricter <RST> acceptance checks.

4.3. Basic PAWS Algorithm

If the PAWS algorithm is used, the following processing MUST be performed on all incoming segments for a synchronized connection. Also, PAWS processing MUST take precedence over the regular TCP acceptablitiy check (Section 3.3 in [RFC0793]), which is performed after verification of the received Timestamps option:

If there is a Timestamps option in the arriving segment, SEG.TSval < TS.Recent, TS.Recent is valid (see later discussion) and the RST bit is not set, then treat the arriving segment as not acceptable:

If the segment is outside the window, reject it (normal TCP processing)
If an arriving segment satisfies: SEG.SEQ <= Last.ACK.sent (see Section 3.5), then record its timestamp in TS.Recent.
If an arriving segment is in-sequence (i.e., at the left window edge), then accept it normally.
Otherwise, treat the segment as a normal in-window, out-of-sequence TCP segment (e.g., queue it for later delivery to the user).

Steps R2, R4, and R5 are the normal TCP processing steps specified by [RFC0793].

It is important to note that the timestamp MUST be checked only when a segment first arrives at the receiver, regardless of whether it is in- sequence or it must be queued for later delivery.

Consider the following example.

This rule allows reasonable performance under loss. A full window of data is in transit at all times, and after a loss a full window less one segment will show up out-of-sequence to be queued at the receiver (e.g., up to ~2^30 bytes of data); the Timestamps option must not result in discarding this data.

In certain unlikely circumstances, the algorithm of rules R1-R5 could lead to discarding some segments unnecessarily, as shown in the following example:

This case is very unlikely to occur. If the retransmission was triggered by a timeout, some of the segments C.1, ... Z.1 must have been delayed longer than the RTO time. This is presumably an unlikely event, or there would be many spurious timeouts and retransmissions. If B's retransmission was triggered by the "fast retransmit" algorithm, i.e., by duplicate <ACK>s, then the queued segments that caused these <ACK>s must have been received already.

Even if a segment were delayed past the RTO, the Fast Retransmit mechanism [Jacobson90c] will cause the delayed segments to be retransmitted at the same time as B.2, avoiding an extra RTT and therefore causing a very small performance penalty.

We know of no case with a significant probability of occurrence in which timestamps will cause performance degradation by unnecessarily discarding segments.

4.4. Timestamp Clock

It is important to understand that the PAWS algorithm does not require clock synchronization between sender and receiver. The sender's timestamp clock is used to stamp the segments, and the sender uses the echoed timestamp to measure RTTs. However, the receiver treats the timestamp as simply a monotonically increasing serial number, without any necessary connection to its clock. From the receiver's viewpoint, the timestamp is acting as a logical extension of the high-order bits of the sequence number.

The receiver algorithm does place some requirements on the frequency of the timestamp clock.

The timestamp clock must not be "too slow".

It MUST tick at least once for each 2^31 bytes sent. In fact, in order to be useful to the sender for round trip timing, the clock SHOULD tick at least once per window's worth of data, and even with the window extension defined in Section 2.2, 2^31 bytes must be at least two windows.

To make this more quantitative, any clock faster than 1 tick/sec will reject old duplicate segments for link speeds of ~8 Gbps. A 1 ms timestamp clock will work at link speeds up to 8 Tbps (8*10^12) bps!
The timestamp clock must not be "too fast".

The recycling time of the timestamp clock MUST be greater than MSL seconds. Since the clock (timestamp) is 32 bits and the worst-case MSL is 255 seconds, the maximum acceptable clock frequency is one tick every 59 ns.

However, it is desirable to establish a much longer recycle period, in order to handle outdated timestamps on idle connections (see Section 4.5), and to relax the MSL requirement for preventing sequence number wrap-around. With a 1 ms timestamp clock, the 32-bit timestamp will wrap its sign bit in 24.8 days. Thus, it will reject old duplicates on the same connection if MSL is 24.8 days or less. This appears to be a very safe figure; an MSL of 24.8 days or longer can probably be assumed in the internet without requiring precise MSL enforcement.

Based upon these considerations, we choose a timestamp clock frequency in the range 1 ms to 1 sec per tick. This range also matches the requirements of the RTTM mechanism, which does not need much more resolution than the granularity of the retransmit timer, e.g., tens or hundreds of milliseconds.

The PAWS mechanism also puts a strong monotonicity requirement on the sender's timestamp clock. The method of implementation of the timestamp clock to meet this requirement depends upon the system hardware and software.

4.5. Outdated Timestamps

If a connection remains idle long enough for the timestamp clock of the other TCP to wrap its sign bit, then the value saved in TS.Recent will become too old; as a result, the PAWS mechanism will cause all subsequent segments to be rejected, freezing the connection (until the timestamp clock wraps its sign bit again).

With the chosen range of timestamp clock frequencies (1 sec to 1 ms), the time to wrap the sign bit will be between 24.8 days and 24800 days. A TCP connection that is idle for more than 24 days and then comes to life is exceedingly unusual. However, it is undesirable in principle to place any limitation on TCP connection lifetimes.

We therefore require that an implementation of PAWS include a mechanism to "invalidate" the TS.Recent value when a connection is idle for more than 24 days. (An alternative solution to the problem of outdated timestamps would be to send keep-alive segments at a very low rate, but still more often than the wrap-around time for timestamps, e.g., once a day. This would impose negligible overhead. However, the TCP specification has never included keep-alives, so the solution based upon invalidation was chosen.)

Note that a TCP does not know the frequency, and therefore, the wraparound time, of the other TCP, so it must assume the worst. The validity of TS.Recent needs to be checked only if the basic PAWS timestamp check fails, i.e., only if SEG.TSval < TS.Recent. If TS.Recent is found to be invalid, then the segment is accepted, regardless of the failure of the timestamp check, and rule R3 updates TS.Recent with the TSval from the new segment.

To detect how long the connection has been idle, the TCP MAY update a clock or timestamp value associated with the connection whenever TS.Recent is updated, for example. The details will be implementation-dependent.

4.6. Header Prediction

"Header prediction" [Jacobson90a] is a high-performance transport protocol implementation technique that is most important for high-speed links. This technique optimizes the code for the most common case, receiving a segment correctly and in order. Using header prediction, the receiver asks the question, "Is this segment the next in sequence?" This question can be answered in fewer machine instructions than the question, "Is this segment within the window?"

Adding header prediction to our timestamp procedure leads to the following recommended sequence for processing an arriving TCP segment:

Check timestamp (same as step R1 above)
Do header prediction: if segment is next in sequence and if there are no special conditions requiring additional processing, accept the segment, record its timestamp, and skip H3.
Process the segment normally, as specified in RFC 793. This includes dropping segments that are outside the window and possibly sending acknowledgments, and queuing in-window, out-of-sequence segments.

Another possibility would be to interchange steps H1 and H2, i.e., to perform the header prediction step H2 *first*, and perform H1 and H3 only when header prediction fails. This could be a performance improvement, since the timestamp check in step H1 is very unlikely to fail, and it requires unsigned modulo arithmetic. To perform this check on every single segment is contrary to the philosophy of header prediction. We believe that this change might produce a measurable reduction in CPU time for TCP protocol processing on high-speed networks.

However, putting H2 first would create a hazard: a segment from 2^32 bytes in the past might arrive at exactly the wrong time and be accepted mistakenly by the header-prediction step. The following reasoning has been introduced in [RFC1185] to show that the probability of this failure is negligible.

However, this probabilistic argument is not universally accepted, and the consensus at present is that the performance gain does not justify the hazard in the general case. It is therefore recommended that H2 follow H1.

4.7. IP Fragmentation

At high data rates, the protection against old segments provided by PAWS can be circumvented by errors in IP fragment reassembly (see [RFC4963]). The only way to protect against incorrect IP fragment reassembly is to not allow the segments to be fragmented. This is done by setting the Don't Fragment (DF) bit in the IP header. Setting the DF bit implies the use of Path MTU Discovery as described in [RFC1191], [RFC1981], and [RFC4821], thus any TCP implementation that implements PAWS MUST also implement Path MTU Discovery.

4.8. Duplicates from Earlier Incarnations of Connection

The PAWS mechanism protects against errors due to sequence number wrap-around on high-speed connections. Segments from an earlier incarnation of the same connection are also a potential cause of old duplicate errors. In both cases, the TCP mechanisms to prevent such errors depend upon the enforcement of a maximum segment lifetime (MSL) by the Internet (IP) layer (see Appendix of RFC 1185 for a detailed discussion). Unlike the case of sequence space wrap-around, the MSL required to prevent old duplicate errors from earlier incarnations does not depend upon the transfer rate. If the IP layer enforces the recommended 2 minute MSL of TCP, and if the TCP rules are followed, TCP connections will be safe from earlier incarnations, no matter how high the network speed. Thus, the PAWS mechanism is not required for this case.

We may still ask whether the PAWS mechanism can provide additional security against old duplicates from earlier connections, allowing us to relax the enforcement of MSL by the IP layer. Appendix B explores this question, showing that further assumptions and/or mechanisms are required, beyond those of PAWS. This is not part of the current extension.

5. Conclusions and Acknowledgements

This memo presented a set of extensions to TCP to provide efficient operation over large bandwidth * delay product paths and reliable operation over very high-speed paths. These extensions are designed to provide compatible interworking with TCP stacks that do not implement the extensions.

These mechanisms are implemented using TCP options for scaled windows and timestamps. The timestamps are used for two distinct mechanisms: RTTM (Round Trip Time Measurement) and PAWS (Protection Against Wrapped Sequences).

The Window Scale option was originally suggested by Mike St. Johns of USAF/DCA. The present form of the option was suggested by Mike Karels of UC Berkeley in response to a more cumbersome scheme defined by Van Jacobson. Lixia Zhang helped formulate the PAWS mechanism description in [RFC1185].

Finally, much of this work originated as the result of discussions within the End-to-End Task Force on the theoretical limitations of transport protocols in general and TCP in particular. Task force members and other on the end2end-interest list have made valuable contributions by pointing out flaws in the algorithms and the documentation. Continued discussion and development since the publication of [RFC1323] originally occurred in the IETF TCP Large Windows Working Group, later on in the End-to-End Task Force, and most recently in the IETF TCP Maintenance Working Group. The authors are grateful for all these contributions.

6. Security Considerations

The TCP sequence space is a fixed size, and as the window becomes larger it becomes easier for an attacker to generate forged packets that can fall within the TCP window, and be accepted as valid segments. While use of timestamps and PAWS can help to mitigate this, when using PAWS, if an attacker is able to forge a packet that is acceptable to the TCP connection, a timestamp that is in the future would cause valid segments to be dropped due to PAWS checks. Hence, implementers should take care to not open the TCP window drastically beyond the requirements of the connection.

A naive implementation that derives the timestamp clock value directly from a system uptime clock may unintentionally leak this information to an attacker. This does not directly compromise any of the mechanisms described in this document. However, this may be valuable information to a potential attacker. An implementer should evaluate the potential impact and mitigate this accordingly (i.e. by using a random offset for the timestamp clock on each connection, or using an external, real-time derived timestamp clock source).

Expanding the TCP window beyond 64 KiB for IPv6 allows Jumbograms [RFC2675] to be used when the local network supports packets larger than 64 KiB. When larger TCP segments are used, the TCP checksum becomes weaker.

Mechanisms to protect the TCP header from modification should also protect the TCP options.

Middleboxes and TCP options:

Implementations that depend on PAWS could provide a mechanism for the application to determine whether or not PAWS is in use on the connection, and chose to terminate the connection if that protection doesn't exist. This is not just to protect the connection against middleboxes that might remove the Timestamps option, but also against remote hosts that do not have Timestamp support.

7. IANA Considerations

This document has no actions for IANA.

8. References

8.1. Normative References

[RFC0793] Postel, J., "Transmission Control Protocol", STD 7, RFC 793, September 1981.
[RFC1191] Mogul, J. and S. Deering, "Path MTU discovery", RFC 1191, November 1990.
[RFC2119] Bradner, S., "Key words for use in RFCs to Indicate Requirement Levels", BCP 14, RFC 2119, March 1997.

8.2. Informative References

[RFC0896] Nagle, J., "Congestion control in IP/TCP internetworks", RFC 896, January 1984.
[RFC1072] Jacobson, V. and R. Braden, "TCP extensions for long-delay paths", RFC 1072, October 1988.
[RFC1110] McKenzie, A., "Problem with the TCP big window option", RFC 1110, August 1989.
[RFC1122] Braden, R., "Requirements for Internet Hosts - Communication Layers", STD 3, RFC 1122, October 1989.
[RFC1185] Jacobson, V., Braden, B. and L. Zhang, "TCP Extension for High-Speed Paths", RFC 1185, October 1990.
[RFC1323] Jacobson, V., Braden, B. and D. Borman, "TCP Extensions for High Performance", RFC 1323, May 1992.
[RFC1981] McCann, J., Deering, S. and J. Mogul, "Path MTU Discovery for IP version 6", RFC 1981, August 1996.
[RFC2018] Mathis, M., Mahdavi, J., Floyd, S. and A. Romanow, "TCP Selective Acknowledgment Options", RFC 2018, October 1996.
[RFC2581] Allman, M., Paxson, V. and W. Stevens, "TCP Congestion Control", RFC 2581, April 1999.
[RFC2675] Borman, D., Deering, S. and R. Hinden, "IPv6 Jumbograms", RFC 2675, August 1999.
[RFC2883] Floyd, S., Mahdavi, J., Mathis, M. and M. Podolsky, "An Extension to the Selective Acknowledgement (SACK) Option for TCP", RFC 2883, July 2000.
[RFC3522] Ludwig, R. and M. Meyer, "The Eifel Detection Algorithm for TCP", RFC 3522, April 2003.
[RFC4015] Ludwig, R. and A. Gurtov, "The Eifel Response Algorithm for TCP", RFC 4015, February 2005.
[RFC4821] Mathis, M. and J. Heffner, "Packetization Layer Path MTU Discovery", RFC 4821, March 2007.
[RFC4963] Heffner, J., Mathis, M. and B. Chandler, "IPv4 Reassembly Errors at High Data Rates", RFC 4963, July 2007.
[RFC5681] Allman, M., Paxson, V. and E. Blanton, "TCP Congestion Control", RFC 5681, September 2009.
[RFC6298] Paxson, V., Allman, M., Chu, J. and M. Sargent, "Computing TCP's Retransmission Timer", RFC 6298, June 2011.
[RFC6675] Blanton, E., Allman, M., Wang, L., Jarvinen, I., Kojo, M. and Y. Nishida, "A Conservative Loss Recovery Algorithm Based on Selective Acknowledgment (SACK) for TCP", RFC 6675, August 2012.
[RFC6691] Borman, D., "TCP Options and Maximum Segment Size (MSS)", RFC 6691, July 2012.
[Garlick77] Garlick, L., Rom, R. and J. Postel, "Issues in Reliable Host-to-Host Protocols", Proc. Second Berkeley Workshop on Distributed Data Management and Computer Networks, May 1977.
[Hamming77] Hamming, R., "Digital Filters", Prentice Hall, Englewood Cliffs, N.J. ISBN 0-13-212571-4, 1977.
[Jacobson88a] Jacobson, V., "Congestion Avoidance and Control", SIGCOMM '88, Stanford, CA., August 1988.
[Jacobson90a] Jacobson, V., "4BSD Header Prediction", ACM Computer Communication Review, April 1990.
[Jacobson90c] Jacobson, V., "Modified TCP congestion avoidance algorithm", Message to the end2end-interest mailing list, April 1990.
[Allman99] Allman, M. and V. Paxson, "On Estimating End-to-End Network Path Properties", Proc. ACM SIGCOMM Technical Symposium, Cambridge, MA, September 1999.
[Jain86] Jain, R., "Divergence of Timeout Algorithms for Packet Retransmissions", Proc. Fifth Phoenix Conf. on Comp. and Comm., Scottsdale, Arizona, March 1986.
[Karn87] Karn, P. and C. Partridge, "Estimating Round-Trip Times in Reliable Transport Protocols", Proc. SIGCOMM '87, August 1987.
[Martin03] Martin, D., "[Tsvwg] RFC 1323.bis", Message to the tsvwg mailing list, September 2003.
[Medina04] Medina, A., Allman, M. and S. Floyd, "Measuring Interactions Between Transport Protocols and Middleboxes", Proc. ACM SIGCOMM/USENIX Internet Measurement Conference. October 2004, August 2004.
[Medina05] Medina, A., Allman, M. and S. Floyd, "Measuring the Evolution of Transport Protocols in the Internet", ACM Computer Communication Review 35(2), April 2005.
[Ludwig00] Ludwig, R. and K. Sklower, "The Eifel Retransmission Timer", ACM SIGCOMM Computer Communication Review Volume 30 Issue 3, July 2000.
[Ekstroem04] Ekstroem, H. and R. Ludwig, "The Peak-Hopper: A New End-to-End Retransmission Timer for Reliable Unicast Transport", INFOCOM 2004 IEEE, March 2004.
[Mathis08] Mathis, M., "[tcpm] Example of 1323 window retraction problem", Message to the tcpm mailing list, March 2008.
[Floyd05] Floyd, S., "[tcpm] How the RTO should be estimated with timestamps", Message from 26.Jan.2007 to the tcpm mailing list, August 2005.
[Honda11] Honda, M., Nishida, Y., Raiciu, C., Greenhalgh, A., Handley, M. and H. Tokuda, "Is it still possible to extend TCP?", Proc. of ACM Internet Measurement Conference (IMC) '11, November 2011.
[Watson81] Watson, R., "Timer-based Mechanisms in Reliable Transport Protocol Connection Management", Computer Networks, Vol. 5, 1981.
[Zhang86] Zhang, L., "Why TCP Timers Don't Work Well", Proc. SIGCOMM '86, Stowe, VT, August 1986.

Appendix A. Implementation Suggestions

|   NOP  |  NOP   |  TSopt |   10   |
|          TSval timestamp          |
|          TSecr timestamp          |

TCP Option Layout

Interaction with the TCP Urgent Pointer

Appendix B. Duplicates from Earlier Connection Incarnations

There are two cases to be considered: (1) a system crashing (and losing connection state) and restarting, and (2) the same connection being closed and reopened without a loss of host state. These will be described in the following two sections.

B.1. System Crash with Loss of State

TCP's quiet time of one MSL upon system startup handles the loss of connection state in a system crash/restart. For an explanation, see for example "When to Keep Quiet" in the TCP protocol specification [RFC0793]. The MSL that is required here does not depend upon the transfer speed. The current TCP MSL of 2 minutes seemed acceptable as an operational compromise, when many host systems used to take this long to boot after a crash. Current host systems can boot considerably faster.

The Timestamps option may be used to ease the MSL requirements (or to provide additional security against data corruption). If timestamps are being used and if the timestamp clock can be guaranteed to be monotonic over a system crash/restart, i.e., if the first value of the sender's timestamp clock after a crash/restart can be guaranteed to be greater than the last value before the restart, then a quiet time is unnecessary.

To dispense totally with the quiet time would require that the host clock be synchronized to a time source that is stable over the crash/restart period, with an accuracy of one timestamp clock tick or better. We can back off from this strict requirement to take advantage of approximate clock synchronization. Suppose that the clock is always re-synchronized to within N timestamp clock ticks and that booting (extended with a quiet time, if necessary) takes more than N ticks. This will guarantee monotonicity of the timestamps, which can then be used to reject old duplicates even without an enforced MSL.

B.2. Closing and Reopening a Connection

When a TCP connection is closed, a delay of 2*MSL in TIME-WAIT state ties up the socket pair for 4 minutes (see Section 3.5 of [RFC0793]. Applications built upon TCP that close one connection and open a new one (e.g., an FTP data transfer connection using Stream mode) must choose a new socket pair each time. The TIME-WAIT delay serves two different purposes:

Implement the full-duplex reliable close handshake of TCP.

The proper time to delay the final close step is not really related to the MSL; it depends instead upon the RTO for the FIN segments and therefore upon the RTT of the path. (It could be argued that the side that is sending a FIN knows what degree of reliability it needs, and therefore it should be able to determine the length of the TIME-WAIT delay for the FIN's recipient. This could be accomplished with an appropriate TCP option in FIN segments.)

Although there is no formal upper-bound on RTT, common network engineering practice makes an RTT greater than 1 minute very unlikely. Thus, the 4 minute delay in TIME-WAIT state works satisfactorily to provide a reliable full-duplex TCP close. Note again that this is independent of MSL enforcement and network speed.

The TIME-WAIT state could cause an indirect performance problem if an application needed to repeatedly close one connection and open another at a very high frequency, since the number of available TCP ports on a host is less than 2^16. However, high network speeds are not the major contributor to this problem; the RTT is the limiting factor in how quickly connections can be opened and closed. Therefore, this problem will be no worse at high transfer speeds.
Allow old duplicate segments to expire.

To replace this function of TIME-WAIT state, a mechanism would have to operate across connections. PAWS is defined strictly within a single connection; the last timestamp (TS.Recent) is kept in the connection control block, and discarded when a connection is closed.

An additional mechanism could be added to the TCP, a per-host cache of the last timestamp received from any connection. This value could then be used in the PAWS mechanism to reject old duplicate segments from earlier incarnations of the connection, if the timestamp clock can be guaranteed to have ticked at least once since the old connection was open. This would require that the TIME-WAIT delay plus the RTT together must be at least one tick of the sender's timestamp clock. Such an extension is not part of the proposal of this RFC.

Note that this is a variant on the mechanism proposed by Garlick, Rom, and Postel [Garlick77], which required each host to maintain connection records containing the highest sequence numbers on every connection. Using timestamps instead, it is only necessary to keep one quantity per remote host, regardless of the number of simultaneous connections to that host.

Appendix C. Summary of Notation

The following notation has been used in this document.


Option Fields

Option Fields in Current Segment

Clock Values

Per-Connection State Variables


Appendix D. Event Processing Summary



SEG.WND = (RCV.WND >> Rcv.Wind.Shift).





TrueWindow = SEG.WND << Snd.Wind.Shift,




Appendix E. Timestamps Edge Cases

While the rules laid out for when to calculate RTTM produce the correct results most of the time, there are some edge cases where an incorrect RTTM can be calculated. All of these situations involve the loss of segments. It is felt that these scenarios are rare, and that if they should happen, they will cause a single RTTM measurement to be inflated, which mitigates its effects on RTO calculations.

  tc=1   <A, TSval=1> ------------------->

  tc=2   (lost) <---- <ACK(A), TSecr=1, win=n>
      (RTTM would have been 1)

         (receive window opens, window update is sent)
  tc=5        <---- <ACK(A), TSecr=1, win=m>
         (RTTM is calculated at 4)

[Martin03] cites two similar cases when the returning <ACK> is lost, and before the retransmission timer fires, another returning <ACK> segment arrives, which aknowledges the data. In this case, the RTTM calculated will be inflated:

One thing to note about this situation is that it is somewhat bounded by RTO + RTT, limiting how far off the RTTM calculation will be. While more complex scenarios can be constructed that produce larger inflations (e.g., retransmissions are lost), those scenarios involve multiple segment losses, and the connection will have other more serious operational problems than using an inflated RTTM in the RTO calculation.

Appendix F. Window Retraction Example

Consider an established TCP connection using a scale factor of 128, Snd.Wind.Shift=7 and Rcv.Wind.Shift=7, that is running with a very small window because the receiver is bottlenecked and both ends are doing small reads and writes.

SEG.ACK  SEG.WIN computed SND.WIN   receiver's actual window
1000     2       1256               1300

Consider the ACKs coming back:

1040     2       1296               1300

The sender writes 40 bytes and receiver ACKs:

1045     2       1301               1300   - BEYOND BUFFER

1045     1       1173               1300   - RETRACTED WINDOW

The sender writes 5 additional bytes and the receiver has a problem. Two choices:

This is a general problem and can happen any time the sender does a write which is smaller than the window scale factor.

In most stacks it is at least partially obscured when the window size is larger than some small number of segments because the stacks prefer to announce windows that are an integral number of segments, rounded up to the next scale factor. This plus silly window suppression tends to cause less frequent, larger window updates. If the window was rounded down to a segment size there is more opportunity to advance the window, the BEYOND BUFFER case above, rather than retracting it.

Appendix G. RTO calculation modification

Taking multiple RTT samples per window would shorten the history calculated by the RTO mechanism in [RFC6298], and the below algorithm aims to maintain a similar history as originally intended by [RFC6298].

It is roughly known how many samples a congestion window worth of data will yield, not accounting for ACK compression, and ACK losses. Such events will result in more history of the path being reflected in the final value for RTO, and are uncritical. This modification will ensure that a similar amount of time is taken into account for the RTO estimation, regardless of how many samples are taken per window:

Note that the factor 2 in ExpectedSamples is due to "Delayed ACKs".

Instead of using alpha and beta in the algorithm of [RFC6298], use alpha' and beta' instead:

Appendix H. Changes from RFC 1323

Several important updates and clarifications to the specification in RFC 1323 are made in these document. The technical changes are summarized below:

A wrong reference to SND.WND was corrected to SEG.WND in Section 2.3
Section 2.4 was added describing the unavoidable window retraction issue, and explicitly describing the mitigation steps necessary.
In Section 3.2 the wording how the Timestamps option negotiation is to be performed was updated with RFC2119 wording. Further, a number of paragraphs were added to clarify the expected behavior with a compliant implementation using TSopt, as RFC1323 left room for interpretation - e.g. potential late enablement of TSopt.
The description of which TSecr values can be used to update the measured RTT has been clarified. Specifically, with timestamps, the Karn algorithm [Karn87] is disabled. The Karn algorithm disables all RTT measurements during retransmission, since it is ambiguous whether the <ACK> is for the original segment, or the retransmitted segment. With timestamps, that ambiguity is removed since the TSecr in the <ACK> will contain the TSval from whichever data segment made it to the destination.
RTTM update processing explicitly excludes segments not updating SND.UNA. The original text could be interpreted to allow taking RTT samples when SACK acknowledges some new, non-continuous data.
In RFC1323, section 3.4, step (2) of the algorithm to control which timestamp is echoed was incorrect in two regards:
It failed to update TS.recent for a retransmitted segment that resulted from a lost <ACK>.
It failed if SEG.LEN = 0.

In the new algorithm, the case of SEG.TSval >= TS.recent is included for consistency with the PAWS test.

It is now recommended that the Timestamps option is included in <RST> segments if the incoming segment contained a Timestamps option.
<RST> segments are explicitly excluded from PAWS processing.
Added text to clarify the precedence between regular TCP [RFC0793] and this document Timestamps option / PAWS processing. Discussion about combined acceptability checks are ongoing.
Snd.TSoffset and Snd.TSclock variables have been added. Snd.TSclock is the sum of my.TSclock and Snd.TSoffset. This allows the starting points for timestamp values to be randomized on a per-connection basis. Setting Snd.TSoffset to zero yields the same results as [RFC1323].
Appendix A has been expanded with information about the TCP Urgent Pointer. An earlier revision contained text around the TCP MSS option, which was split off into [RFC6691].
One correction was made to the Event Processing Summary in Appendix D. In SEND CALL/ESTABLISHED STATE, RCV.WND is used to fill in the SEG.WND value, not SND.WND.
Appendix G was added to exemplify how an RTO calculation might be updated to properly take the much higher RTT sampling frequency enabled by the Timestamps option into account.

Editorial changes of the document, that don't impact the implementation or function of the mechanisms described in this document include:

Removed much of the discussion in Section 1 to streamline the document. However, detailed examples and discussions in Section 2, Section 3 and Section 4 are kept as guideline for implementers.
Removed references to "new" options, as the options were introduced in [RFC1323] already. Changed the text in Section 1.3 to specifically address TS and WS options.
Section 1.4 was added for [RFC2119] wording. Normative text was updated with the appropriate phrases.
Added < > brackets to mark specific types of segments, and replaced most occurances of "packet" with "segment", where TCP segments are referred to.
Updated the text in Section 3 to take into account what has been learned since [RFC1323].
Removed the list of changes between [RFC1323] and prior versions. These changes are mentioned in Appendix C of [RFC1323].
Moved Appendix Changes from RFC 1323 to the end of the appendices for easier lookup. In addition, the entries were split into a technical and an editorial part, and sorted to roughly correspond with the sections in the text where they apply.

Authors' Addresses

David Borman Quantum Corporation Mendota Heights, MN 55120 USA EMail:
Bob Braden University of Southern California 4676 Admiralty Way Marina del Rey, CA 90292 USA EMail:
Van Jacobson Google, Inc. 1600 Amphitheatre Parkway Mountain View, CA 94043 USA EMail:
Richard Scheffenegger (editor) NetApp, Inc. Am Euro Platz 2 Vienna, 1120 Austria EMail: