This document is obsolete. Please
refer to RFC 6247.
Network Working Group R. Braden Request for Comments: 1379 ISI November 1992
Extending TCP for Transactions -- Concepts
Status of This Memo
This memo provides information for the Internet community. It does not specify an Internet standard. Distribution of this memo is unlimited.
Abstract
This memo discusses extension of TCP to provide transaction-oriented service, without altering its virtual-circuit operation. This extension would fill the large gap between connection-oriented TCP and datagram-based UDP, allowing TCP to efficiently perform many applications for which UDP is currently used. A separate memo contains a detailed functional specification for this proposed extension.
This work was supported in part by the National Science Foundation under Grant Number NCR-8922231.
TABLE OF CONTENTS
1. INTRODUCTION .................................................. 2 2. TRANSACTIONS USING STANDARD TCP ............................... 3 3. BYPASSING THE 3-WAY HANDSHAKE ................................. 6 3.1 Concept of TAO ........................................... 6 3.2 Cache Initialization ..................................... 10 3.3 Accepting <SYN,ACK> Segments ............................. 11 4. SHORTENING TIME-WAIT STATE .................................... 13 5. CHOOSING A MONOTONIC SEQUENCE ................................. 15 5.1 Cached Timestamps ........................................ 16 5.2 Current TCP Sequence Numbers ............................. 18 5.3 64-bit Sequence Numbers .................................. 20 5.4 Connection Counts ........................................ 20 5.5 Conclusions .............................................. 21 6. CONNECTION STATES ............................................. 24 7. CONCLUSIONS AND ACKNOWLEDGMENTS ............................... 32 APPENDIX A: TIME-WAIT STATE AND THE 2-PACKET EXCHANGE ............ 34 REFERENCES ....................................................... 37 Security Considerations .......................................... 38 Author's Address ................................................. 38
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The TCP protocol [STD-007] implements a virtual-circuit transport service that provides reliable and ordered data delivery over a full-duplex connection. Under the virtual circuit model, the life of a connection is divided into three distinct phases: (1) opening the connection to create a full-duplex byte stream; (2) transferring data in one or both directions over this stream; and (3) closing the connection. Remote login and file transfer are examples of applications that are well suited to virtual-circuit service.
Distributed applications, which are becoming increasingly numerous and sophisticated in the Internet, tend to use a transaction-oriented rather than a virtual circuit style of communication. Currently, a transaction-oriented Internet application must choose to suffer the overhead of opening and closing TCP connections or else build an application-specific transport mechanism on top of the connectionless transport protocol UDP. Greater convenience, uniformity, and efficiency would result from widely-available kernel implementations of a transport protocol supporting a transaction service model [RFC- 955].
The transaction service model has the following features:
* The fundamental interaction is a request followed by a response.
* An explicit open or close phase would impose excessive overhead.
* At-most-once semantics is required; that is, a transaction must not be "replayed" by a duplicate request packet.
* In favorable circumstances, a reliable request/response handshake can be performed with exactly one packet in each direction.
* The minimum transaction latency for a client is RTT + SPT, where RTT is the round-trip time and SPT is the server processing time.
We use the term "transaction transport protocol" for a transport- layer protocol that follows this model [RFC-955].
The Internet architecture allows an arbitrary collection of transport protocols to be defined on top of the minimal end-to-end datagram service provided by IP [Clark88]. In practice, however, production systems implement only TCP and UDP at the transport layer. It has proven difficult to leverage a new transport protocol into place, to be widely enough available to be useful for application builders.
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This memo explores an alternative approach to providing a transaction transport protocol: extending TCP to implement the transaction service model, while continuing to support the virtual circuit model. Each transaction will then be a single instance of a TCP connection. The proposed transaction extension is effectively implementable within current TCPs and operating systems, and it should also scale to the much faster networks, interfaces, and CPUs of the future.
The present memo explains the theory behind the extension, in somewhat exquisite detail. Despite the length and complexity of this memo, the TCP extensions required for transactions are in fact quite limited and simple. Another memo [TTCP-FS] provides a self-contained functional specification of the extensions.
Section 2 of this memo describes the limitations of standard TCP for transaction processing, to motivate the extensions. Sections 3, 4, and 5 explore the fundamental extensions that are required for transactions. Section 6 discusses the changes required in the TCP connection state diagram. Finally, Section 7 presents conclusions and acknowledgments. Familiarity with the standard TCP protocol [STD-007] is assumed.
Reliable transfer of data depends upon sequence numbers. Before data transfer can begin, both parties must "synchronize" the connection, i.e, agree on common sequence numbers. The synchronization procedure must preserve at-most-once semantics, i.e., be free from replay hazards due to duplicate packets. The TCP developers adopted a synchronization mechanism known as the 3-way handshake.
Consider a simple transaction in which client host A sends a single- segment request to server host B, and B returns a single-segment response. Many current TCP implementations use at least ten segments (i.e., packets) for this sequence: three for the 3-way handshake opening the connection, four to send and acknowledge the request and response data, and three for TCP's full-duplex data-conserving close sequence. These ten segments represent a high relative overhead for two data-bearing segments. However, a more important consideration is the transaction latency seen by the client: 2*RTT + SPT, larger than the minimum by one RTT. As CPU and network speeds increase, the relative significance of this extra transaction latency also increases.
Proposed transaction transport protocols have typically used a "timer-based" approach to connection synchronization [Birrell84]. In this approach, once end-to-end connection state is established in the client and server hosts, a subset of this state is maintained for
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some period of time. A new request before the expiration of this timeout period can then reestablish the full state without an explicit handshake. Watson pointed out that the timer-based approach of his Delta-T protocol [Watson81] would encompass both virtual circuits and transactions. However, the TCP group adopted the 3-way handshake (because of uncertainty about the robustness of enforcing the packet lifetime bounds required by Delta-T, within a general Internet environment). More recently, Liskov, Shrira, and Wroclawski [Liskov90] have proposed a different timer-based approach to connection synchronization, requiring loosely-synchronized clocks in the hosts.
The technique proposed in this memo, suggested by Clark [Clark89], depends upon cacheing of connection state but not upon clocks or timers; it is described in Section 3 below. Garlick, Rom, and Postel also proposed a connection synchronization mechanism using cached state [Garlick77]. Their scheme required each host to maintain connection records containing the highest sequence number on each connection. The technique suggested here retains only per-host state, not per-connection state.
During TCP development, it was suggested that TCP could support transactions with data segments containing both SYN and FIN bits. (These "Kamikaze" segments were not supported as a service; they were used mainly to crash other experimental TCPs!) To illustrate this idea, Figure 1 shows a plausible application of the current TCP rules to create a minimal transaction. (In fact, some minor adjustments in the standard TCP spec would be required to make Figure 1 fully legal [STD-007]).
Figure 1, like many of the examples shown in this memo, uses an abbreviated form to illustrate segment sequences. For clarity and brevity, it omits explicit sequence and acknowledgment numbers, assuming that these will follow the well-known TCP rules. The notation "ACK(x)" implies a cumulative acknowledgment for the control bit or data "x" and everything preceding "x" in the sequence space. The referent of "x" should be clear from the context. Also, host A will always be the client and host B will be the server in these diagrams.
The first three segments in Figure 1 implement the standard TCP three-way handshake. If segment #1 had been an old duplicate, the client side would have sent an RST (Reset) bit in segment #3, terminating the sequence. The request data included on the initial SYN segment cannot be delivered to user B until segment #3 completes the 3-way handshake. Loading control bits onto the segments has reduced the total number of segments to 5, but the client still observes a transaction latency of 2*RTT + SPT. The 3-way handshake
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The TCP close sequence also poses a performance problem for transactions: one or both end(s) of a closed connection must remain in "TIME-WAIT" state until a 4 minute timeout has expired [STD-007]. The same connection (defined by the host and port numbers at both ends) cannot be reopened until this delay has expired. Because of TIME-WAIT state, a client program should choose a new local port number (i.e., a different connection) for each successive transaction. However, the TCP port field of 16 bits (less the "well-known" port space) provides only 64512 available user ports. This limits the total rate of transactions between any pair of hosts to a maximum of 64512/240 = 268 per second. This is much too low a rate for low-delay paths, e.g., high-speed LANs. A high rate of short connections (i.e., transactions) could also lead to excessive consumption of kernel memory by connection control blocks in TIME- WAIT state.
In summary, to perform efficient transaction processing in TCP, we need to suppress the 3-way handshake and to shorten TIME-WAIT state.
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Protocol mechanisms to accomplish these two goals are discussed in Sections 3 and 4, respectively. Both require the choice of a monotonic sequence-like space; Section 5 analyzes the choices and makes a selection for this space. Finally, the TCP connection state machine must be extended as described in Section 6.
Transaction processing in TCP raises some other protocol issues, which are discussed in the functional specification memo [TTCP-FS]. These include:
(1) augmenting the user interface for transactions,
(2) delaying acknowledgment segments to allow maximum piggy-backing of control bits with data,
(3) measuring the retransmission timeout time (RTO) on very short connections, and
(4) providing an initial server window.
A recently proposed set of enhancements [RFC-1323] defines a TCP Timestamps option that carries two 32-bit timestamp values. The Timestamps option is used to accurately measure round-trip time (RTT). The same option is also used in a procedure known as "PAWS" (Protect Againsts Wrapped Sequence) to prevent erroneous data delivery due to a combination of old duplicate segments and sequence number reuse at very high bandwidths. The particular approach to transactions chosen in this memo does not require the RFC-1323 enhancements; however, they are important and should be implemented in every TCP, with or without the transaction extensions described here.
To avoid 3-way handshakes for transactions, we introduce a new mechanism for validating initial SYN segments, i.e., for enforcing at-most-once semantics without a 3-way handshake. We refer to this as the TCP Accelerated Open, or TAO, mechanism.
The basis of TAO is this: a TCP uses cached per-host information to immediately validate new SYNs [Clark89]. If this validation fails, e.g., because there is no current cached state or the segment is an old duplicate, the procedure falls back to a normal 3-way handshake to validate the SYN. Thus, bypassing a 3-way handshake is considered to be an optional optimization.
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The proposed TAO mechanism uses a finite sequence-like space of values that increase monotonically with successive transactions (connections) between a given (client, server) host pair. Call this monotonic space M, and let each initial SYN segment carry an M value SEG.M. If M is not the existing sequence (SEG.SEQ) field, SEG.M may be carried in a TCP option.
When host B receives from host A an initial SYN segment containing a new value SEG.M, host B compares this against cache.M[A], the latest M value that B has cached for host A. This comparison is the "TAO test". Because the M values are monotonically increasing, SEG.M > cache.M[A] implies that the SYN must be new and can be accepted immediately. If not, a normal 3-way handshake is performed to validate the initial SYN segment. Figure 2 illustrates the TAO mechanism; cached M values are shown enclosed in square brackets. The M values generated by host A satisfy x0 < x1, and the M values generated by host B satisfy y0 < y1.
An appropriate choice for the M value space is discussed in Section 5. M values are drawn from a finite number space, so inequalities must be defined in the usual way for sequence numbers [STD-007]. The M space must not wrap so quickly that an old duplicate SYN will be erroneously accepted. We assume that some maximum segment lifetime (MSL) is enforced by the IP layer.
RFC 1379 Transaction TCP -- Concepts November 1992
Figure 2 shows the simplest case: each side has cached the latest M value of the other, and the SEG.M value in the client's SYN segment is greater than the value in the cache at the server host. As a result, B can accept the client A's request data1 immediately and pass it to the server application. B's reply data2 is shown piggybacked on the <SYN,ACK> segment. As a result of this 2-way exchange, the cached M values are updated at both sites; the client side becomes relevant only if the client/server roles reverse. Validation of the <SYN,ACK> segment at host A is discussed later.
Figure 3 shows the TAO test failing but the consequent 3-way handshake succeeding. B updates its cache with the value x2 >= x1 when the initial SYN is known to be valid.
Figure 3. TAO Test Fails but 3-Way Handshake Succeeds.
There are several possible causes for a TAO test failure on a legitimate new SYN segment (not an old duplicate).
(1) There may be no cached M value for this particular client host.
(2) The SYN may be the one of a set of nearly-simultaneous SYNs for different connections but from the same host, which
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arrived out of order.
(3) The finite M space may have wrapped around between successive transactions from the same client.
(4) The M values may advance too slowly for closely-spaced transactions.
None of these TAO failures will cause a lockout, because the resulting 3-way handshake will succeed. Note that the first transaction between a given host pair will always require a 3-way handshake; subsequent transactions can take advantage of TAO.
The per-host cache required by TAO is highly desirable for other reasons, e.g., to retain the measured round trip time and MTU for a given remote host. Furthermore, a host should already have a per-host routing cache [HR-COMM] that should be easily extensible for this purpose.
Figure 4 illustrates a complete TCP transaction sequence using the TAO mechanism. Bypassing the 3-way handshake leads to new connection states; Figure 4 shows three of them, "SYN-SENT*", "CLOSE-WAIT*", and "LAST-ACK*". Explanation of these states is deferred to Section 6.
TCP A (Client) TCP B (Server) _______________ ______________
CLOSED LISTEN
1. SYN-SENT* --> <SYN,data1,FIN,M=x1> --> CLOSE-WAIT* (TAO test OK=> data1->user_B)
The first connection between hosts A and B will find no cached state at one or both ends, so both M caches must be initialized. This requires that the first transaction carry a specially marked SEG.M value, which we call SEG.M.NEW. Receiving a SEG.M.NEW value in an initial SYN segment, B will cache this value and send its own M back to initialize A's cache. When a host crashes and restarts, all its cached M values cache.M[*] must be invalidated in order to force a re-synchronization of the caches at both ends.
This cache synchronization procedure is illustrated in Figure 5, where client host A has crashed and restarted with its cache entries undefined, as indicated by "??". Since cache.TS[B] is undefined, A sends a SEG.M.NEW value instead of SEG.M in the <SYN> segment of its first transaction request to B. Receiving this SEG.M.NEW, the server host B invalidates cache.TS[A] and performs a 3-way handshake. SEG.M in segment #2 updates A's cache, and when the handshake completes successfully, B updates its cached M value to x2 >= x1.
Suppose that the 3-way handshake failed, presumably because
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segment #1 was an old duplicate. Then segment #3 from host A would be an RST segment, with the result that both side's caches would be left undefined.
Figure 6 shows the procedure when the server crashes and restarts. Upon receiving a <SYN> segment from a host for which it has no cached M value, B initiates a 3-way handshake to validate the request and sends its own M value to A. Again the result is to update cached M values on both sides.
Transactions introduce a new hazard of erroneously accepting an old duplicate <SYN,ACK> segment. To be acceptable, a <SYN,ACK> segment must arrive in SYN-SENT state, and its ACK field must acknowledge something that was sent. In current TCPs the effective send window in SYN-SENT state is exactly one octet, and an acceptable <SYN,ACK> must exactly ACK this one octet. The clock-driven selection of Initial Sequence Number (ISN) makes an erroneous acceptance exceedingly unlikely. An old duplicate SYN could be accepted erroneously only if successive connection attempts occurred more often than once every 4 microseconds, or if the segment lifetime exceeded the 4 hour wraparound time for ISN
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selection.
However, when TCP is used for transactions, data sent with the initial SYN increases the range of sequence numbers that have been sent. This increases the danger of accepting an old duplicate <SYN,ACK> segment, and the consequences are more serious. In the example in Figure 7, segments 1-3 form a normal transaction sequence, and segment 4 begins a new transaction (incarnation) for the same connection. Segment #5 is a duplicate of segment #2 from the preceding transaction. Although the new transaction has a larger ISN, the previous ACK value 402 falls into the new range [200,700) of sequence numbers that have been sent, so segment #5 could be erroneously accepted and passed to the client as the response to the new request.
_T_C_P__A _T_C_P__B
CLOSED LISTEN
1. --> <seq=100,SYN,data=300,FIN,M=x1> --> (TAO test OK)
(New Request) 4. --> <seq=200,SYN,data=500,FIN,M=x2> --> ...
(Duplicate of segment #2) 5. <-- <seq=800,ack=402,SYN,data=300,FIN,M=y1> <--... (Acceptable!!)
Figure 7: Old Duplicate <SYN,ACK> Causing Error
Unfortunately, we cannot simply use TAO on the client side to detect and reject old duplicate <SYN,ACK> segments. A TAO test at the client might fail for a valid <SYN,ACK> segment, due to out- of-order delivery, and this could result in permanent non-delivery of a valid transaction reply.
Instead, we include a second M value, an echo of the client's M value from the initial <SYN> segment, in the <SYN,ACK> segment. A
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specially-marked M value, SEG.M.ECHO, is used for this purpose. The client knows the value it sent in the initial <SYN> and can therefore positively validate the <SYN,ACK> using the echoed value. This is illustrated in Figure 12, which is the same as Figure 4 with the addition of the echoed value on the <SYN,ACK> segment #2.
It should be noted that TCP allows a simultaneous open sequence in which both sides send and receive an initial <SYN> (see Figure 8 of [STD-007]. In this case, the TAO test must be performed on both sides to preserve the symmetry. See [TTCP-FS] for an example.
Once a transaction has been initiated for a particular connection (pair of ports) between a given host pair, a new transaction for the same connection cannot take place for a time that is at least:
RTT + SPT + TIME-WAIT_delay
Since the client host can cycle among the 64512 available port numbers, an upper bound on the transaction rate between a particular host pair is:
[1] TRmax = 64512 /(RTT + TIME-WAIT_Delay)
in transactions per second (Tps), where we assumed SPT is negligible. We must reduce TIME-WAIT_Delay to support high-rate TCP transaction processing.
TIME-WAIT state performs two functions: (1) supporting the full- duplex reliable close of TCP, and (2) allowing old duplicate segments from an earlier connection incarnation to expire before they can cause an error (see Appendix to [RFC-1185]). The first function impacts the application model of a TCP connection, which we would not want to change. The second is part of the fundamental machinery of TCP reliable delivery; to safely truncate TIME-WAIT state, we must provide another means to exclude duplicate packets from earlier incarnations of the connection.
To minimize the delay in TIME-WAIT state while performing both functions, we propose to set the TIME-WAIT delay to:
[2] TIME-WAIT_Delay = max( K*RTO, U )
where U and K are constants and RTO is the dynamically-determined retransmission timeout, the measured RTT plus an allowance for the
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RTT variance [Jacobson88]. We choose K large enough so that there is high probability of the close completing successfully if at all possible; K = 8 seems reasonable. This takes care of the first function of TIME-WAIT state.
In a real implementation, there may be a minimum RTO value Tr, corresponding to the precision of RTO calculation. For example, in the popular BSD implementation of TCP, the minimum RTO is Tr = 0.5 second. Assuming K = 8 and U = 0, Eqns [1] and [2] impose an upper limit of TRmax = 16K Tps on the transaction rate of these implementations.
It is possible to have many short connections only if RTO is very small, in which case the TIME-WAIT delay [2] reduces to U. To accelerate the close sequence, we need to reduce U below the MSL enforced by the IP layer, without introducing a hazard from old duplicate segments. For this purpose, we introduce another monotonic number sequence; call it X. X values are required to be monotonic between successive connection incarnations; depending upon the choice of the X space (see Section 5), X values may also increase during a connection. A value from the X space is to be carried in every segment, and a segment is rejected if it is received with an X value smaller than the largest X value received. This mechanism does not use a cache; the largest X value is maintained in the TCP connection control block (TCB) for each connection.
The value of U depends upon the choice for the X space, discussed in the next section. If X is time-like, U can be set to twice the time granularity (i.e, twice the minimum "tick" time) of X. The TIME-WAIT delay will then ensure that current X values do not overlap the X values of earlier incarnations of the same connection. Another consequence of time-like X values is the possibility that an open but idle connection might allow the X value to wrap its sign bit, resulting in a lockup of the connection. To prevent this, a 24-day idle timer on each open connection could bypass the X check on the first segment following the idle period, for example. In practice, many implementations have keep-alive mechanisms that prevent such long idle periods [RFC-1323].
Referring back to Figure 4, our proposed transaction extension results in a minimum exchange of 3 packets. Segment #3, the final ACK segment, does not increase transaction latency, but in combination with the TIME-WAIT delay of K*RTO it ensures that the server side of the connection will be closed before a new transaction is issued for this same pair of ports. It also provides an RTT measurement for the server.
We may ask whether it would be possible to further reduce the TIME-
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WAIT delay. We might set K to zero; alternatively, we might allow the client TCP to start a new transaction request while the connection was still in TIME-WAIT state, with the new initial SYN acting as an implied acknowledgment of the previous FIN. Appendix A summarizes the issues raised by these alternatives, which we call "truncating" TIME-WAIT state, and suggests some possible solutions. Further study would be required, but these solutions appear to bend the theory and/or implementations of the TCP protocol farther than we wish to bend them.
We therefore propose using formula [2] with K=8 and retaining the final ACK(FIN) transmission. To raise the transaction rate, therefore, we require small values of RTO and U.
5. CHOOSING A MONOTONIC SEQUENCE
For simplicity, we want the monotonic sequence X used for shortening TIME-WAIT state to be identical to the monotonic sequence M for bypassing the 3-way handshake. Calling the common space M, we will send an M value SEG.M in each TCP segment. Upon receipt of an initial SYN segment, SEG.M will be compared with a per-host cached value to authenticate the SYN without a 3-way handshake; this is the TAO mechanism. Upon receipt of a non-SYN segment, SEG.M will be compared with the current value in the connection control block and used to discard old duplicates.
Note that the situation with TIME-WAIT state differs from that of bypassing 3-way handshakes in two ways: (a) TIME-WAIT requires duplicate detection on every segment vs. only on SYN segments, and (b) TIME-WAIT applies to a single connection vs. being global across all connections. This section discusses possible choices for the common monotonic sequence.
The SEG.M values must satisfy the following requirements.
* The values must be monotonic; this requirement is defined more precisely below.
* Their granularity must be fine-grained enough to support a high rate of transaction processing; the M clock must "tick" at least once between successive transactions.
* Their range (wrap-around time) must be great enough to allow a realistic MSL to be enforced by the network.
The TCP spec calls for an MSL of 120 secs. Since much of the Internet does not carefully enforce this limit, it would be safer to have an MSL at least an order of magnitude larger. We set as an
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objective an MSL of at least 2000 seconds. If there were no TIME- WAIT delay, the ultimate limit on transaction rate would be set by speed-of-light delays in the network and by the latency of host operating systems. As the bottleneck problems with interfacing CPUs to gigabit LANs are solved, we can imagine transaction durations as short as 1 microsecond. Therefore, we set an ultimate performance goal of TRmax at least 10**6 Tps.
A particular connection between hosts A and B is identified by the local and remote TCP "sockets", i.e., by the quadruplet: {A, B, Port.A, Port.B}. Imagine that each host keeps a count CC of the number of TCP connections it has initiated. We can use this CC number to distinguish different incarnations of the same connection. Then a particular SEG.M value may be labeled implicitly by 6 quantities: {A, B, Port.A, Port.B, CC, n}, where n is the byte offset of that segment within the connection incarnation.
To bypass the 3-way handshake, we require thgt SEG.M values on successive SYN segments from a host A to a host B be monotone increasing. If CC' > CC, then we require that:
To delete old duplicates (allowing TIME-WAIT state to be shortened), we require that SEG.M values be disjoint across different incarnations of the same connection. If CC' > CC then
We now consider four different choices for the common monotonic space: RFC-1323 timestamps, TCP sequence numbers, the connection count, and 64-bit TCP sequence numbers. The results are summarized in Table I.
The PAWS mechanism [RFC-1323] uses TCP "timestamps" as monotonically increasing integers in order to throw out old duplicate segments within the same incarnation. Jacobson suggested the cacheing of these timestamps for bypassing 3-way handshakes [Jacobson90], i.e., that TCP timestamps be used for our common monotonic space M. This idea is attractive since it would allow the same timestamp options to be used for RTTM, PAWS, and transactions.
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To obtain at-most-once service, the criterion for immediate acceptance of a SYN must be that SEG.M is strictly greater than the cached M value. That is, to be useful for bypassing 3-way handshakes, the timestamp clock must tick at least once between any two successive transactions between the same pair of hosts (even if different ports are used). Hence, the timestamp clock rate would determine TRmax, the maximum possible transaction rate.
Unfortunately, the timestamp clock frequency called for by RFC- 1323, in the range 1 sec to 1 ms, is much too slow for transactions. The TCP timestamp period was chosen to be comparable to the fundamental interval for computing and scheduling retransmission timeouts; this is generally in the range of 1 sec. to 1 ms., and in many operating systems, much closer to 1 second. Although it would be possible to increase the timestamp clock frequency by several orders of magnitude, to do so would make implementation more difficult, and on some systems excessively expensive.
The wraparound time for TCP timestamps, at least 24 days, causes no problem for transactions.
The PAWS mechanism uses TCP timestamps to protect against old duplicate non-SYN segments from the same incarnation [RFC-1323]. It can also be used to protect against old duplicate data segments from earlier incarnations (and therefore allow shortening of TIME-WAIT state) if we can ensure that the timestamp clock ticks at least once between the end of one incarnation and the beginning of the next. This can be achieved by setting U = 2 seconds, i.e., to twice the maximum timestamp clock period. This value in formula [2] leads to an upper bound TRmax = 32K Tps between a host pair. However, as pointed out above, old duplicate SYN detection using timestamps leads to a smaller transaction rate bound, 1 Tps, which is unacceptable. In addition, the timestamp approach is imperfect; it allows old ACK segments to enter the new connection where they can cause a disconnect. This happens because old duplicate ACKs that arrive during TIME-WAIT state generate new ACKs with the current timestamp [RFC-1337].
We therefore conclude that timestamps are not adequate as the monotonic space M; see Table I. However, they may still be useful to effectively extend some other monotonic number space, just as they are used in PAWS to extend the TCP sequence number space. This is discussed below.
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It is useful to understand why the existing 32-bit TCP sequence numbers do not form an appropriate monotonic space for transactions.
The sequence number sent in an initial SYN is called the Initial Sequence Number or ISN. According to the TCP specification, an ISN is to be selected using:
[3] ISN = (R*T) mod 2**32
where T is the real time in seconds (from an arbitrary origin, fixed when the system is started) and R is a constant, currently 250 KBps. These ISN values form a monotonic time sequence that wraps in 4.55 hours = 16380 seconds and has a granularity of 4 usecs. For transaction rates up to roughly 250K Tps, the ISN value calculated by formula [3] will be monotonic and could be used for bypassing the 3-way handshake.
However, TCP sequence numbers (alone) could not be used to shorten TIME-WAIT state, because there are several ways that overlap of the sequence space of successive incarnations can occur (as described in Appendix to [RFC-1185]). One way is a "fast connection", with a transfer rate greater than R; another is a "long" connection, with a duration of approximately 4.55 hours. TIME-WAIT delay is necessary to protect against these cases. With the official delay of 240 seconds, formula [1] implies a upper bound (as RTT -> 0) of TRmax = 268 Tps; with our target MSL of 2000 sec, TRmax = 32 Tps. These values are unacceptably low.
To improve this transaction rate, we could use TCP timestamps to effectively extend the range of the TCP sequence numbers. Timestamps would guard against sequence number wrap-around and thereby allow us to increase R in [3] to exceed the maximum possible transfer rate. Then sequence numbers for successive incarnations could not overlap. Timestamps would also provide safety with an MSL as large as 24 days. We could then set U = 0 in the TIME-WAIT delay calculation [2]. For example, R = 10**9 Bps leads to TRmax <= 10**9 Tps. See 2(b) in Table I. These values would more than satisfy our objectives.
We should make clear how this proposal, sequence numbers plus timestamps, differs from the timestamps alone discussed (and rejected) in the previous section. The difference lies in what is cached and tested for TAO; the proposal here is to cache and test BOTH the latest TCP sequence number and the latest TCP timestamp. In effect, we are proposing to use timestamps to logically extend
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the sequence space to 64 bits. Another alternative, presented in the next section, is to directly expand the TCP sequence space to 64 bits.
Unfortunately, the proposed solution (TCP sequence numbers plus timestamps) based on equation [3] would be difficult or impossible to implement on many systems, which base their TCP implementation upon a very low granularity software clock, typically O(1 sec). To adapt the procedure to a system with a low granularity software clock, suppose that we calculate the ISN as:
[4] ISN = ( R*Ts*floor(T/Ts) + q*CC) mod 2**32
where Ts is the time per tick of the software clock, CC is the connection count, and q is a constant. That is, the ISN is incremented by the constant R*Ts once every clock tick and by the constant q for every new connection. We need to choose q to obtain the required monotonicity.
For monotonicity of the ISN's themselves, q=1 suffices. However, monotonicity during the entire connection requires q = R*Ts. This value of q can be deduced as follows. Let S(T, CC, n) be the sequence number for byte offset n in a connection with number CC at time T:
For any T1 > T2, we require that: S(T2, CC+1, 0) - S(T1, CC, n) > 0 for all n. Since R is assumed to be an upper bound on the transfer rate, we can write down:
R > n/(T2 - T1), or T2/Ts - T1/Ts > n/(R*Ts)
Using the relationship: floor(x)-floor(y) > x-y-1 and a little algebra leads to the conclusion that using q = R*Ts creates the required monotonic number sequence. Therefore, we consider:
[5] ISN = R*Ts*(floor(T/Ts) + CC) mod 2**32
(which is the algorithm used for ISN selection by BSD TCP).
For error-free operation, the sequence numbers generated by [5] must not wrap the sign bit in less than MSL seconds. Since CC cannot increase faster than TRmax, the safe condition is:
R* (1 + Ts*TRmax) * MSL < 2**31.
We are interested in the case: Ts*TRmax >> 1, so this relationship
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reduces to:
[6] R * Ts * TRmax * MSL < 2**31.
This shows a direct trade-off among the maximum effective bandwidth R, the maximum transaction rate TRmax, and the maximum segment lifetime MSL. For reasonable limiting values of R, Ts, and MSL, formula [6] leads to a very low value of TRmax. For example, with MSL= 2000 secs, R=10**9 Bps, and Ts = 0.5 sec, TRmax < 2*10**-3 Tps.
To ease the situation, we could supplement sequence numbers with timestamps. This would allow an effective MSL of 2 seconds in [6], since longer times would be protected by differing timestamps. Then TRmax < 2**30/(R*Ts). The actual enforced MSL would be increased to 24 days. Unfortunately, TRmax would still be too small, since we want to support transfer rates up to R ~ 10**9 Bps. Ts = 0.5 sec would imply TRmax ~ 2 Tps. On many systems, it appears infeasible to decrease Ts enough to obtain an acceptable TRmax using this approach.
Another possibility would be to simply increase the TCP sequence space to 64 bits as suggested in [RFC-1263]. We would also increase the R value for clock-driven ISN selection, beyond the fastest transfer rate of which the host is capable. A reasonable upper limit might be R = 10**9 Bps. As noted above, in a practical implementation we would use:
ISN = R*Ts*( floor(T/Ts) + CC) mod 2**64
leading to:
R*(1 + Ts * TRmax) * MSL < 2**63
For example, suppose that R = 10**9 Bps, Ts = 0.5, and MSL = 16K secs (4.4 hrs); then this result implies that TRmax < 10**6 Tps. We see that adding 32 bits to the sequence space has provided feasible values for transaction processing.
The Connection Count CC is well suited to be the monotonic sequence M, since it "ticks" exactly once for each new connection incarnation and is constant within a single incarnation. Thus, it perfectly separates segments from different incarnations of the same connection and would allow U = 0 in the TIME-WAIT state delay
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formula [2]. (Strictly, U cannot be reduced below 1/R = 4 usec, as noted in Section 4. However, this is of little practical consequence until the ultimate limits on TRmax are approached).
Assume that CC is a 32-bit number. To prevent wrap-around in the sign bit of CC in less than MSL seconds requires that:
TRmax * MSL < 2**31
For example, if MSL = 2000 seconds then TRmax < 10**6 Tp. These are acceptable limits for transaction processing. However, if they are not, we could augment CC with TCP timestamps to obtain very far-out limits, as discussed below.
It would be an implementation choice at the client whether CC is global for all destinations or private to each destination host (and maintained in the per-host cache). In the latter case, the last CC value assigned for each remote host could also be maintained in the per-host cache. Since there is not typically a large amount of parallelism in the network connection of a host, there should be little difference in the performance of these two different approaches, and the single global CC value is certainly simpler.
To augment CC with TCP timestamps, we would bypass a 3-way handshake if both SEG.CC > cache.CC[A] and SEG.TSval >= cache.TS[A]. The timestamp check would detect a SYN older than 2 seconds, so that the effective wrap-around requirement would be:
TRmax * 2 < 2**31
i.e., TRmax < 10**9 Tps. The required MSL would be raised to 24 days. Using timestamps in this way, we could reduce the size of CC. For example, suppose CC were 16 bits. Then the wrap-around condition TRmax * 2 < 2**15 implies that TRmax is 16K.
Finally, note that using CC to delete old duplicates from earlier incarnations would not obviate the need for the time-stamp-based PAWS mechanism to prevent errors within a single incarnation due to wrapping the 32-bit TCP sequence space at very high transfer rates.
The alternatives for monotonic sequence are summarized in Table I. We see that there are two feasible choices for the monotonic space: the connection count and 64-bit sequence numbers. Of these two, we believe that the simpler is the connection count.
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Implementation of 64-bit sequence numbers would require negotiation of a new header format and expansion of all variables and calculations on the sequence space. CC can be carried in an option and need be examined only once per packet.
We propose to use a simple 32-bit connection count CC, without augmentation with timestamps, for the transaction extension. This choice has the advantages of simplicity and directness. Its drawback is that it adds a third sequence-like space (in addition to the TCP sequence number and the TCP timestamp) to each TCP header and to the main line of packet processing. However, the additional code is in fact very modest.
We now have a general outline of the proposed TCP extensions for transactions.
o A host maintains a 32-bit global connection counter variable CC.
o The sender's current CC value is carried in an option in every TCP segment.
o CC values are cached per host, and the TAO mechanism is used to bypass the 3-way handshake when possible.
o In non-SYN segments, the CC value is used to reject duplicates from earlier incarnations. This allows TIME-WAIT state delay to be reduced to K*RTO (i.e., U=0 in Eq. [2]).
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TABLE I: Summary of Monotonic Sequences
APPROACH TRmax (Tps) Required MSL COMMENTS __________________________________________________________________
1. Timestamp & PAWS 1 24 days TRmax is too small __________________________________________________________________
2. Current TCP Sequence Numbers
(a) clock-driven ISN: eq. [3] 268 240 secs TRmax & MSL too small
(b) Timestamps& clock- driven ISN [3] & 10**9 24 days Hard to R=10**9 implement
(c) Timestamps & c-dr ISN: eq. [4] 2**30/(R*Ts) 24 days TRmax too small. __________________________________________________________________
3. 64-bit TCP Sequence Numbers
2**63/(MSL*R*Ts) MSL Significant TCP change e.g., R=10**9 Bps, MSL = 4.4 hrs, Ts = 0.5 sec=> TRmax = 10**6 __________________________________________________________________
4. Connection Counts
(a) no timestamps 2**31/MSL MSL 3rd sequence e.g., MSL=2000 sec space TRmax = 10**6
(b) with timestamps 2**30 24 days (ditto) and PAWS __________________________________________________________________
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TCP has always allowed a connection to be half-closed. TAO makes a significant addition to TCP semantics by allowing a connection to be half-synchronized, i.e., to be open for data transfer in one direction before the other direction has been opened. Thus, the passive end of a connection (which receives an initial SYN) can accept data and even a FIN bit before its own SYN has been acknowledged. This SYN, data, and FIN may arrive on a single segment (as in Figure 4), or on multiple segments; packetization makes no difference to the logic of the finite-state machine (FSM) defining transitions among connection states.
Half-synchronized connections have several consequences.
(a) The passive end must provide an implied initial data window in order to accept data. The minimum size of this implied window is a parameter in the specification; we suggest 4K bytes.
(b) New connection states and transitions are introduced into the TCP FSM at both ends of the connection. At the active end, new states are required to piggy-back the FIN on the initial SYN segment. At the passive end, new states are required for a half-synchronized connection.
This section develops the resulting FSM description of a TCP connection as a conventional state/transition diagram. To develop a complete FSM, we take a constructive approach, as follows: (1) write down all possible events; (2) write down the precedence rules that govern the order in which events may occur; (3) construct the resulting FSM; and (4) augment it to support TAO. In principle, we do this separately for the active and passive ends; however, the symmetry of TCP results in the two FSMs being almost entirely coincident.
Figure 8 lists all possible state transitions for a TCP connection in the absence of TAO, as elementary events and corresponding actions. Each transition is labeled with a letter. Transitions a-g are used by the active side, and c-i are used by the passive side. Without TAO, transition "c" (event "rcv ACK(SYN)") synchronizes the connection, allowing data to be accepted for the user.
By definition, the first transition for an active (or passive) side must be "a" (or "i", respectively). During a single instance of a connection, the active side will progress through some permutation of the complete sequence of transitions {a b c d e f } or the sequence {a b c d e f g}. The set of possible permutations is determined by precedence rules governing the order in which transitions can occur.
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Label Event / Action _____ ________________________ a OPEN / snd SYN
b rcv SYN [No TAO]/ snd ACK(SYN)
c rcv ACK(SYN) /
d CLOSE / snd FIN
e rcv FIN / snd ACK(FIN)
f rcv ACK(FIN) /
g timeout=2MSL / delete TCB ___________________________________________________ h passive OPEN / create TCB
i rcv SYN [No TAO]/ snd SYN, ACK(SYN) ___________________________________________________
Figure 8. Basic TCP Connection Transitions
Using the notation "<." to mean "must precede", the precedence rules are:
(1) Logical ordering: must open connection before closing it:
b <. e
(2) Causality -- cannot receive ACK(x) before x has been sent:
a <. c and i <. c and d <. f
(3) Acknowledgments are cumulative
c <. f
(4) First packet in each direction must contain a SYN.
b <. c and b <. f
(5) TIME-WAIT state
Whenever d precedes e in the sequence, g must be the last transition.
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Applying these rules, we can enumerate all possible permutations of the events and summarize them in a state transition diagram. Figure 9 shows the result, with boxes representing the states and directed arcs representing the transitions.
________ ________ | | h | | | CLOSED |--------->| LISTEN | |________| |________| | | | a | i ____V____ ____V___ ________ | | b | | e | | | |--------->| |-------------->| | |________| |________| |________| / / | / | / / | c d / | c / / __V_____ | ____V___ / / | | e | | | d | d / | |------------>| | | | |________| | |________| | | | | | | | | ___V____ | | | | | | | | | | | | | | | | |________| | | | | | | ____V___ ______V_ | ________ | | | | b | | e | | | | | | |------->| |--------->| | | | |________| |________| | |________| | | | / | | | c | / d c | c | d | | / | | | _V___V__ ____V___ V_____V_ | | e | | | | | |---->| | | | |________| |________| |________| | | | | f | f | f ____V___ ____V___ ___V____ | | e | TIME- | g | | | |---->| WAIT |-->| CLOSED | |________| |________| |________|
Figure 9: Basic State Diagram
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Although Figure 9 gives a correct representation of the possible event sequences, it is not quite correct for the actions, which do not compose as shown. In particular, once a control bit X has been sent, it must continue to be sent until ACK(X) is received. This requires new transitions with modified actions, shown in the following list. We use the labeling convention that transitions with the same event part all have the same letter, with different numbers of primes to indicate different actions.
Label Event / Action _____ _______________________________________ b' (=i) rcv SYN [No TAO] / snd SYN,ACK(SYN) b'' rcv SYN [No TAO] / snd SYN,FIN,ACK(SYN) d' CLOSE / snd SYN,FIN e' rcv FIN / snd FIN,ACK(FIN) e'' rcv FIN / snd SYN,FIN,ACK(FIN)
Figure 10 shows the state diagram of Figure 9, with the modified transitions and with the states used by standard TCP [STD-007] identified. Those states that do not occur in standard TCP are numbered 1-5.
Standard TCP has another implied restriction: a FIN bit cannot be recognized before the connection has been synchronized, i.e., c <. e. This eliminates from standard TCP the states 1, 2, and 5 shown in Figure 10. States 3 and 4 are needed if a FIN is to be piggy-backed on a SYN segment (note that the states shown in Figure 1 are actually wrong; the states shown as SYN-SENT and ESTABLISHED are really states 3 and 4). In the absence of piggybacking the FIN bit, Figure 10 reduces to the standard TCP state diagram [STD-007].
The FSM described in Figure 10 is intended to be applied cumulatively; that is, parsing a single packet header may lead to more than one transition. For example, the standard TCP state diagram includes a direct transition from SYN-SENT to ESTABLISHED:
rcv SYN,ACK(SYN) / snd ACK(SYN).
This is transition b followed immediately by c.
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Next we introduce TAO. If the TAO test succeeds, the connection becomes half-synchronized. This requires a new set of states, mirroring the states of Figure 10, beginning with acceptance of a SYN (transition "b" or "i"), and ending when ACK(SYN) arrives (transition
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"c"). Figure 11 shows the result of augmenting Figure 10 with the additional states for TAO. The transitions are defined in the following table:
Key for Figure 11: Complete State Diagram with TAO
a OPEN / create TCB, snd SYN b' rcv SYN [no TAO]/ snd SYN,ACK(SYN) b'' rcv SYN [no TAO]/ snd SYN,FIN,ACK(SYN) c rcv ACK(SYN) / d CLOSE / snd FIN d' CLOSE / snd SYN,FIN e rcv FIN / snd ACK(FIN) e' rcv FIN / snd SYN,ACK(FIN) e'' rcv FIN / snd FIN,ACK(FIN) e''' rcv FIN / snd SYN,FIN,ACK(FIN) f rcv ACK(FIN) / g timeout=2MSL / delete TCB h passive OPEN / create TCB i (= b') rcv SYN [no TAO]/ snd SYN,ACK(SYN) j rcv SYN [TAO OK] / snd SYN,ACK(SYN) k rcv SYN [TAO OK] / snd SYN,FIN,ACK(SYN)
Each new state in Figure 11 bears a very simple relationship to a standard TCP state. We indicate this by naming the new state with the standard state name followed by a star. States SYN-SENT* and SYN-RECEIVED* differ from the corresponding unstarred states in recording the fact that a FIN has been sent. The other new states with starred names differ from the corresponding unstarred states in being half-synchronized (hence, a SYN bit needs to be transmitted).
The state diagram of Figure 11 is more general than required for transaction processing. In particular, it handles simultaneous connection synchronization from both sides, allowing one or both sides to bypass the 3-way handshake. It includes other transitions that are unlikely in normal transaction processing, for example, the server sending a FIN before it receives a FIN from the client (ESTABLISHED* -> FIN-WAIT-1* in Figure 11).
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The relationship between starred and unstarred states is very regular. As a result, the state extensions can be implemented very simply using the standard TCP FSM with the addition of two "hidden" boolean flags, as described in the functional specification memo
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[TTCP-FS].
As an example of the application of Figure 11, consider the minimal transaction shown in Figure 12.
TCP A (Client) TCP B (Server) _______________ ______________
CLOSED LISTEN
1. SYN-SENT* --> <SYN,data1,FIN,CC=x1> --> CLOSE-WAIT* (TAO test OK=> data1->user_B)
LAST-ACK* <-- <SYN,ACK(FIN),data2,FIN,CC=y1,CC.ECHO=x1> <-- 2. TIME-WAIT (TAO test OK, data2->user_A)
3. TIME-WAIT --> <ACK(FIN),CC=x2> --> CLOSED
(timeout) CLOSED
Figure 12: Minimal Transaction Sequence
Sending segment #1 leaves the client end in SYN-SENT* state, which differs from SYN-SENT state in recording the fact that a FIN has been sent. At the server end, passing the TAO test enters ESTABLISHED* state, which passes the data to the user as in ESTABLISHED state and also records the fact that the connection is half synchronized. Then the server processes the FIN bit of segment #1, moving to CLOSE-WAIT* state.
Moving to CLOSE-WAIT* state should cause the server to send a segment containing SYN and ACK(FIN). However, transmission of this segment is deferred so the server can piggyback the response data and FIN on the same segment, unless a timeout occurs first. When the server does send segment #2 containing the response data2 and a FIN, the connection advances from CLOSE-WAIT* to LAST-ACK* state; the connection is still half-synchronized from B's viewpoint.
Processing segment #2 at the client again results in multiple transitions:
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These correspond respectively to receiving a SYN, a FIN, an ACK for A's SYN, and an ACK for A's FIN.
Figure 13 shows a slightly more complex example, a transaction sequence in which request and response data each require two segments. This figure assumes that both client and server TCP are well-behaved, so that e.g., the client sends the single segment #5 to acknowledge both data segments #3 and #4. SEG.CC values are omitted for clarity.
TCP was designed to be a highly symmetric protocol. This symmetry is evident in the piggy-backing of acknowledgments on data and in the common header format for data segments and acknowledgments. On the other hand, the examples and discussion in this memo are in general highly unsymmetrical; the actions of a "client" are clearly distinguished from those of a "server". To explain this apparent discrepancy, we note the following. Even when TCP is used for virtual circuit service, the data transfer phase is symmetrical but the open and close phases are not. A minimal transaction, consisting of one segment in each direction, compresses the open, data transfer, and close phases together, and making the asymmetry of the open and
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close phases dominant. As request and response messages increase in size, the virtual circuit model becomes increasingly relevant, and symmetry again dominates.
TCP's 3-way handshake precludes any performance gain from including data on a SYN segment, while TCP's full-duplex data-conserving close sequence ties up communication resources to the detriment of high- speed transactions. Merely loading more control bits onto TCP data segments does not provide efficient transaction service. To use TCP as an effective transaction transport protocol requires bypassing the 3-way handshake and shortening the TIME-WAIT delay. This memo has proposed a backwards-compatible TCP extension to accomplish both goals. It is our hope that by building upon the current version of TCP, we can give a boost to community acceptance of the new facilities. Furthermore, the resulting protocol implementations will retain the algorithms that have been developed for flow and congestion control in TCP [Jacobson88].
O'Malley and Peterson have recently recommended against backwards- compatible extensions to TCP, and suggested instead a mechanism to allow easy installation of alternative versions of a protocol [RFC- 1263]. While this is an interesting long-term approach, in the shorter term we suggest that incremental extension of the current TCP may be a more effective route.
Besides the backward-compatible extension proposed here, there are two other possible approaches to making efficient transaction processing widely available in the Internet: (1) a new version of TCP or (2) a new protocol specifically adapted to transactions. Since current TCP "almost" supports transactions, we favor (1) over (2). A new version of TCP that retained the semantics of STD-007 but used 64 bit sequence numbers with the procedures and states described in Sections 3, 4, and 6 of this memo would support transactions as well as virtual circuits in a clean, coherent manner.
A potential application of transaction-mode TCP might be SMTP. If commands and responses are batched, in favorable cases complete SMTP delivery operations on short messages could be performed with a single minimal transaction; on the other hand, the body of a message may be arbitrarily large. Using a TCP extended as in this memo could significantly reduce the load on large mail hosts.
This work began as an elaboration of the concept of TAO, due to Dave Clark. I am grateful to him and to Van Jacobson, John Wroclawski, Dave Borman, and other members of the End-to-End Research group for helpful ideas and critiques during the long development of this work. I also thank Liming Wei, who tested the initial implementation in Sun OS.
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APPENDIX A -- TIME-WAIT STATE AND THE 2-PACKET EXCHANGE
This appendix considers the implications of reducing TIME-WAIT state delay below that given in formula [2].
An immediate consequence of this would be the requirement for the server host to accept an initial SYN for a connection in LAST-ACK state. Without the transaction extensions, the arrival of a new <SYN> in LAST-ACK state looks to TCP like a half-open connection, and TCP's rules are designed to restore correspondence by destroying the state (through sending a RST segment) at one end or the other. We would need to thwart this action in the case of transactions.
There are two different possible ways to further reduce TIME-WAIT delay.
(1) Explicit Truncation of TIME-WAIT state
TIME-WAIT state could be explicitly truncated by accepting a new sendto() request for a connection in TIME-WAIT state.
This would allow the ACK(FIN) segment to be delayed and sent only if a timeout occurs before a new request arrives. This allows an ideal 2-segment exchange for closely-spaced transactions, which would restore some symmetry to the transaction exchange. However, explicit truncation would represent a significant change in many implementations.
It might be supposed that even greater symmetry would result if the new request segment were a <SYN,ACK> that explicitly acknowledges the previous reply, rather than a <SYN> that is only an implicit acknowledgment. However, the new request segment might arrive at B to find the server side in either LAST-ACK or CLOSED state, depending upon whether the ACK(FIN) had arrived. In CLOSED state, a <SYN,ACK> would not be acceptable. Hence, if the client sent an initial <SYN,ACK> instead of a <SYN> segment, there would be a race condition at the server.
(2) No TIME-WAIT delay
TIME-WAIT delay could be removed entirely. This would imply that the ACK(FIN) would always be sent (which does not of course guarantee that it will be received). As a result, the arrival of a new SYN in LAST-ACK state would be rare.
This choice is much simpler to implement. Its drawback is that the server will get a false failure report if the ACK(FIN) is
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lost. This may not matter in practice, but it does represent a significant change of TCP semantics. It should be noted that reliable delivery of the reply is not an issue. The client enter TIME-WAIT state only after the entire reply, including the FIN bit, has been received successfully.
The server host B must be certain that a new request received in LAST-ACK state is indeed a new SYN and not an old duplicate; otherwise, B could falsely acknowledge a previous response that has not in fact been delivered to A. If the TAO comparison succeeds, the SYN must be new; however, the server has a dilemma if the TAO test fails.
In Figure A.1, for example, the reply segment from the first transaction has been lost; since it has not been acknowledged, it is still in B's retransmission queue. An old duplicate request, segment #3, arrives at B and its TAO test fails. B is in the position of having old state it cannot discard (the retransmission queue) and needing to build new state to pursue a 3-way handshake to validate the new SYN. If the 3-way handshake failed, it would need to restore the earlier LAST-ACK* state. (Compare with Figure 15 "Old Duplicate SYN Initiates a Reset on Two Passive Sockets" in STD-007). This would be complex and difficult to accomplish in many implementations.
TCP A (Client) TCP B (Server) _______________ ______________
CLOSED LISTEN
1. SYN-SENT* --> <SYN,data1,FIN> --> CLOSE-WAIT* (TAO test OK; data1->server)
The only practical action A can taken when the TAO test fails on a new SYN received in LAST-ACK state is to ignore the SYN, assuming it is really an old duplicate. We must pursue the possible consequences
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of this action.
Section 3.1 listed four possible reasons for failure of the TAO test on a legitimate SYN segment: (1) no cached state, (2) out-of-order delivery of SYNs, (3) wraparound of CCgen relative to the cached value, or (4) the M values advance too slowly. We are assuming that there is a cached CC value at B (otherwise, the SYN cannot be acceptable in LAST-ACK state). Wrapping the CC space is very unlikely and probably impossible; it is difficult to imagine circumstances which would allow the new SYN to be delivered but not the ACK(FIN), especially given the long wraparound time of CCgen.
This leaves the problem of out-of-order delivery of two nearly- concurrent SYNs for different ports. The second to be delivered may have a lower CC option and thus be locked out. This can be solved by using a new CCgen value for every retransmission of an initial SYN.
Truncation of TIME-WAIT state and acceptance of a SYN in LAST-ACK state should take place only if there is a cached CC value for the remote host. Otherwise, a SYN arriving in LAST-ACK state is to be processed by normal TCP rules, which will result in a RST segment from either A or B.
This discussion leads to a paradigm for rejecting old duplicate segments that is different from TAO. This alternative scheme is based upon the following:
(a) Each retransmission of an initial SYN will have a new value of CC, as described above.
This provision takes care of reordered SYNs.
(b) A host maintains a distinct CCgen value for each remote host. This value could easily be maintained in the same cache used for the received CC values, e.g., as cache.CCgen[].
Once the caches are primed, it should always be true that cache.CCgen[B] on host A is equal to cache.CC[A] on host B, and the next transaction from A will carry a CC value exactly 1 greater. Thus, there is no problem of wraparound of the CC value.
(c) A new SYN is acceptable if its SEG.CC > cache.CC[client], otherwise the SYN is ignored as an old duplicate.
This alternative paradigm was not adopted because it would be a somewhat greater perturbation of TCP rules, because it may not have the robustness of TAO, and because all of its consequences may not be
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understood.
REFERENCES
[Birrell84] Birrell, A. and B. Nelson, "Implementing Remote Procedure Calls", ACM TOCS, Vo. 2, No. 1, February 1984.
[Clark88] Clark, D., "The Design Philosophy of the Internet Protocols", ACM SIGCOMM '88, Stanford, CA, August 1988.
[Clark89] Clark, D., Private communication, 1989.
[Garlick77] Garlick, L., R. Rom, and J. Postel, "Issues in Reliable Host-to-Host Protocols", Proc. Second Berkeley Workshop on Distributed Data Management and Computer Networks, May 1977.
[HR-COMM] Braden, R., Ed., "Requirements for Internet Hosts -- Communication Layers", STD-003, RFC-1122, October 1989.
[Jacobson88] Jacobson, V., "Congestion Avoidance and Control", SIGCOMM '88, Stanford, CA., August 1988.
[Liskov90] Liskov, B., Shrira, L., and J. Wroclawski, "Efficient At-Most-Once Messages Based on Synchronized Clocks", ACM SIGCOMM '90, Philadelphia, PA, September 1990.
[RFC-955] Braden, R., "Towards a Transport Service Transaction Protocol", RFC-955, September 1985.
[RFC-1185] Jacobson, V., Braden, R., and Zhang, L., "TCP Extension for High-Speed Paths", RFC-1185, October 1990.
[RFC-1263] O'Malley, S. and L. Peterson, "TCP Extensions Considered Harmful", RFC-1263, University of Arizona, October 1991.
[RFC-1323] Jacobson, V., Braden, R., and Borman, D., "TCP Extensions for High Performance, RFC-1323, February 1991.
[RFC-1337] Braden, R., "TIME-WAIT Assassination Hazards in TCP", RFC-1337, May 1992.
[STD-007] Postel, J., "Transmission Control Protocol - DARPA Internet Program Protocol Specification", STD-007, RFC-793, September 1981.
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[TTCP-FS] Braden, R., "Transaction TCP -- Functional Specification", Work in Progress, September 1992.
[Watson81] Watson, R., "Timer-based Mechanisms in Reliable Transport Protocol Connection Management", Computer Networks, Vol. 5, 1981.
Security Considerations
Security issues are not discussed in this memo.
Author's Address
Bob Braden University of Southern California Information Sciences Institute 4676 Admiralty Way Marina del Rey, CA 90292