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RFC 2760

Ongoing TCP Research Related to Satellites

Pages: 46
Informational
Part 2 of 3 – Pages 12 to 29
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3.3 Loss Recovery

3.3.1 Non-SACK Based Mechanisms

3.3.1.1 Mitigation Description
Several similar algorithms have been developed and studied that improve TCP's ability to recover from multiple lost segments in a window of data without relying on the (often long) retransmission timeout. These sender-side algorithms, known as NewReno TCP, do not depend on the availability of selective acknowledgments (SACKs) [MMFR96]. These algorithms generally work by updating the fast recovery algorithm to use information provided by "partial ACKs" to trigger retransmissions. A partial ACK covers some new data, but not all data outstanding when a particular loss event starts. For instance, consider the case when segment N is retransmitted using the fast
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   retransmit algorithm and segment M is the last segment sent when
   segment N is resent.  If segment N is the only segment lost, the ACK
   elicited by the retransmission of segment N would be for segment M.
   If, however, segment N+1 was also lost, the ACK elicited by the
   retransmission of segment N will be N+1.  This can be taken as an
   indication that segment N+1 was lost and used to trigger a
   retransmission.

3.3.1.2 Research
Hoe [Hoe95,Hoe96] introduced the idea of using partial ACKs to trigger retransmissions and showed that doing so could improve performance. [FF96] shows that in some cases using partial ACKs to trigger retransmissions reduces the time required to recover from multiple lost segments. However, [FF96] also shows that in some cases (many lost segments) relying on the RTO timer can improve performance over simply using partial ACKs to trigger all retransmissions. [HK99] shows that using partial ACKs to trigger retransmissions, in conjunction with SACK, improves performance when compared to TCP using fast retransmit/fast recovery in a satellite environment. Finally, [FH99] describes several slightly different variants of NewReno.
3.3.1.3 Implementation Issues
Implementing these fast recovery enhancements requires changes to the sender-side TCP stack. These changes can safely be implemented in production networks and are allowed by RFC 2581 [APS99].
3.3.1.4 Topology Considerations
It is expected that these changes will work well in all environments outlined in section 2.
3.3.1.5 Possible Interaction and Relationships with Other Research
See section 3.3.2.2.5.

3.3.2 SACK Based Mechanisms

3.3.2.1 Fast Recovery with SACK
3.3.2.1.1 Mitigation Description
Fall and Floyd [FF96] describe a conservative extension to the fast recovery algorithm that takes into account information provided by selective acknowledgments (SACKs) [MMFR96] sent by the receiver. The algorithm starts after fast retransmit triggers the resending of a
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   segment.  As with fast retransmit, the algorithm cuts cwnd in half
   when a loss is detected.  The algorithm keeps a variable called
   "pipe", which is an estimate of the number of outstanding segments in
   the network.  The pipe variable is decremented by 1 segment for each
   duplicate ACK that arrives with new SACK information.  The pipe
   variable is incremented by 1 for each new or retransmitted segment
   sent.  A segment may be sent when the value of pipe is less than cwnd
   (this segment is either a retransmission per the SACK information or
   a new segment if the SACK information indicates that no more
   retransmits are needed).

   This algorithm generally allows TCP to recover from multiple segment
   losses in a window of data within one RTT of loss detection.  Like
   the forward acknowledgment (FACK) algorithm described below, the SACK
   information allows the pipe algorithm to decouple the choice of when
   to send a segment from the choice of what segment to send.

   [APS99] allows the use of this algorithm, as it is consistent with
   the spirit of the fast recovery algorithm.

3.3.2.1.2 Research
[FF96] shows that the above described SACK algorithm performs better than several non-SACK based recovery algorithms when 1--4 segments are lost from a window of data. [AHKO97] shows that the algorithm improves performance over satellite links. Hayes [Hay97] shows the in certain circumstances, the SACK algorithm can hurt performance by generating a large line-rate burst of data at the end of loss recovery, which causes further loss.
3.3.2.1.3 Implementation Issues
This algorithm is implemented in the sender's TCP stack. However, it relies on SACK information generated by the receiver. This algorithm is safe for shared networks and is allowed by RFC 2581 [APS99].
3.3.2.1.4 Topology Considerations
It is expected that the pipe algorithm will work equally well in all scenarios presented in section 2.
3.3.2.1.5 Possible Interaction and Relationships with Other Research
See section 3.3.2.2.5.
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3.3.2.2 Forward Acknowledgments
3.3.2.2.1 Mitigation Description
The Forward Acknowledgment (FACK) algorithm [MM96a,MM96b] was developed to improve TCP congestion control during loss recovery. FACK uses TCP SACK options to glean additional information about the congestion state, adding more precise control to the injection of data into the network during recovery. FACK decouples the congestion control algorithms from the data recovery algorithms to provide a simple and direct way to use SACK to improve congestion control. Due to the separation of these two algorithms, new data may be sent during recovery to sustain TCP's self-clock when there is no further data to retransmit. The most recent version of FACK is Rate-Halving [MM96b], in which one packet is sent for every two ACKs received during recovery. Transmitting a segment for every-other ACK has the result of reducing the congestion window in one round trip to half of the number of packets that were successfully handled by the network (so when cwnd is too large by more than a factor of two it still gets reduced to half of what the network can sustain). Another important aspect of FACK with Rate-Halving is that it sustains the ACK self-clock during recovery because transmitting a packet for every-other ACK does not require half a cwnd of data to drain from the network before transmitting, as required by the fast recovery algorithm [Ste97,APS99]. In addition, the FACK with Rate-Halving implementation provides Thresholded Retransmission to each lost segment. "Tcprexmtthresh" is the number of duplicate ACKs required by TCP to trigger a fast retransmit and enter recovery. FACK applies thresholded retransmission to all segments by waiting until tcprexmtthresh SACK blocks indicate that a given segment is missing before resending the segment. This allows reasonable behavior on links that reorder segments. As described above, FACK sends a segment for every second ACK received during recovery. New segments are transmitted except when tcprexmtthresh SACK blocks have been observed for a dropped segment, at which point the dropped segment is retransmitted. [APS99] allows the use of this algorithm, as it is consistent with the spirit of the fast recovery algorithm.
3.3.2.2.2 Research
The original FACK algorithm is outlined in [MM96a]. The algorithm was later enhanced to include Rate-Halving [MM96b]. The real-world performance of FACK with Rate-Halving was shown to be much closer to
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   the theoretical maximum for TCP than either TCP Reno or the SACK-
   based extensions to fast recovery outlined in section 3.3.2.1
   [MSMO97].

3.3.2.2.3 Implementation Issues
In order to use FACK, the sender's TCP stack must be modified. In addition, the receiver must be able to generate SACK options to obtain the full benefit of using FACK. The FACK algorithm is safe for shared networks and is allowed by RFC 2581 [APS99].
3.3.2.2.4 Topology Considerations
FACK is expected to improve performance in all environments outlined in section 2. Since it is better able to sustain its self-clock than TCP Reno, it may be considerably more attractive over long delay paths.
3.3.2.2.5 Possible Interaction and Relationships with Other Research
Both SACK based loss recovery algorithms described above (the fast recovery enhancement and the FACK algorithm) are similar in that they attempt to effectively repair multiple lost segments from a window of data. Which of the SACK-based loss recovery algorithms to use is still an open research question. In addition, these algorithms are similar to the non-SACK NewReno algorithm described in section 3.3.1, in that they attempt to recover from multiple lost segments without reverting to using the retransmission timer. As has been shown, the above SACK based algorithms are more robust than the NewReno algorithm. However, the SACK algorithm requires a cooperating TCP receiver, which the NewReno algorithm does not. A reasonable TCP implementation might include both a SACK-based and a NewReno-based loss recovery algorithm such that the sender can use the most appropriate loss recovery algorithm based on whether or not the receiver supports SACKs. Finally, both SACK-based and non-SACK-based versions of fast recovery have been shown to transmit a large burst of data upon leaving loss recovery, in some cases [Hay97]. Therefore, the algorithms may benefit from some burst suppression algorithm.

3.3.3 Explicit Congestion Notification

3.3.3.1 Mitigation Description
Explicit congestion notification (ECN) allows routers to inform TCP senders about imminent congestion without dropping segments. Two major forms of ECN have been studied. A router employing backward ECN (BECN), transmits messages directly to the data originator
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   informing it of congestion.  IP routers can accomplish this with an
   ICMP Source Quench message.  The arrival of a BECN signal may or may
   not mean that a TCP data segment has been dropped, but it is a clear
   indication that the TCP sender should reduce its sending rate (i.e.,
   the value of cwnd).  The second major form of congestion notification
   is forward ECN (FECN).  FECN routers mark data segments with a
   special tag when congestion is imminent, but forward the data
   segment.  The data receiver then echos the congestion information
   back to the sender in the ACK packet.  A description of a FECN
   mechanism for TCP/IP is given in [RF99].

   As described in [RF99], senders transmit segments with an "ECN-
   Capable Transport" bit set in the IP header of each packet.  If a
   router employing an active queueing strategy, such as Random Early
   Detection (RED) [FJ93,BCC+98], would otherwise drop this segment, an
   "Congestion Experienced" bit in the IP header is set instead.  Upon
   reception, the information is echoed back to TCP senders using a bit
   in the TCP header.  The TCP sender adjusts the congestion window just
   as it would if a segment was dropped.

   The implementation of ECN as specified in [RF99] requires the
   deployment of active queue management mechanisms in the affected
   routers.  This allows the routers to signal congestion by sending TCP
   a small number of "congestion signals" (segment drops or ECN
   messages), rather than discarding a large number of segments, as can
   happen when TCP overwhelms a drop-tail router queue.

   Since satellite networks generally have higher bit-error rates than
   terrestrial networks, determining whether a segment was lost due to
   congestion or corruption may allow TCP to achieve better performance
   in high BER environments than currently possible (due to TCP's
   assumption that all loss is due to congestion).  While not a solution
   to this problem, adding an ECN mechanism to TCP may be a part of a
   mechanism that will help achieve this goal.  See section 3.3.4 for a
   more detailed discussion of differentiating between corruption and
   congestion based losses.

3.3.3.2 Research
[Flo94] shows that ECN is effective in reducing the segment loss rate which yields better performance especially for short and interactive TCP connections. Furthermore, [Flo94] also shows that ECN avoids some unnecessary, and costly TCP retransmission timeouts. Finally, [Flo94] also considers some of the advantages and disadvantages of various forms of explicit congestion notification.
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3.3.3.3 Implementation Issues
Deployment of ECN requires changes to the TCP implementation on both sender and receiver. Additionally, deployment of ECN requires deployment of some active queue management infrastructure in routers. RED is assumed in most ECN discussions, because RED is already identifying segments to drop, even before its buffer space is exhausted. ECN simply allows the delivery of "marked" segments while still notifying the end nodes that congestion is occurring along the path. ECN is safe (from a congestion control perspective) for shared networks, as it maintains the same TCP congestion control principles as are used when congestion is detected via segment drops.
3.3.3.4 Topology Considerations
It is expected that none of the environments outlined in section 2 will present a bias towards or against ECN traffic.
3.3.3.5 Possible Interaction and Relationships with Other Research
Note that some form of active queueing is necessary to use ECN (e.g., RED queueing).

3.3.4 Detecting Corruption Loss

Differentiating between congestion (loss of segments due to router buffer overflow or imminent buffer overflow) and corruption (loss of segments due to damaged bits) is a difficult problem for TCP. This differentiation is particularly important because the action that TCP should take in the two cases is entirely different. In the case of corruption, TCP should merely retransmit the damaged segment as soon as its loss is detected; there is no need for TCP to adjust its congestion window. On the other hand, as has been widely discussed above, when the TCP sender detects congestion, it should immediately reduce its congestion window to avoid making the congestion worse. TCP's defined behavior, as motivated by [Jac88,Jac90] and defined in [Bra89,Ste97,APS99], is to assume that all loss is due to congestion and to trigger the congestion control algorithms, as defined in [Ste97,APS99]. The loss may be detected using the fast retransmit algorithm, or in the worst case is detected by the expiration of TCP's retransmission timer. TCP's assumption that loss is due to congestion rather than corruption is a conservative mechanism that prevents congestion collapse [Jac88,FF98]. Over satellite networks, however, as in many wireless environments, loss due to corruption is more common than on terrestrial networks. One common partial solution to this problem is
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   to add Forward Error Correction (FEC) to the data that's sent over
   the satellite/wireless link.  A more complete discussion of the
   benefits of FEC can be found in [AGS99].  However, given that FEC
   does not always work or cannot be universally applied, other
   mechanisms have been studied to attempt to make TCP able to
   differentiate between congestion-based and corruption-based loss.

   TCP segments that have been corrupted are most often dropped by
   intervening routers when link-level checksum mechanisms detect that
   an incoming frame has errors.  Occasionally, a TCP segment containing
   an error may survive without detection until it arrives at the TCP
   receiving host, at which point it will almost always either fail the
   IP header checksum or the TCP checksum and be discarded as in the
   link-level error case.  Unfortunately, in either of these cases, it's
   not generally safe for the node detecting the corruption to return
   information about the corrupt packet to the TCP sender because the
   sending address itself might have been corrupted.

3.3.4.1 Mitigation Description
Because the probability of link errors on a satellite link is relatively greater than on a hardwired link, it is particularly important that the TCP sender retransmit these lost segments without reducing its congestion window. Because corrupt segments do not indicate congestion, there is no need for the TCP sender to enter a congestion avoidance phase, which may waste available bandwidth. Simulations performed in [SF98] show a performance improvement when TCP can properly differentiate between between corruption and congestion of wireless links. Perhaps the greatest research challenge in detecting corruption is getting TCP (a transport-layer protocol) to receive appropriate information from either the network layer (IP) or the link layer. Much of the work done to date has involved link-layer mechanisms that retransmit damaged segments. The challenge seems to be to get these mechanisms to make repairs in such a way that TCP understands what happened and can respond appropriately.
3.3.4.2 Research
Research into corruption detection to date has focused primarily on making the link level detect errors and then perform link-level retransmissions. This work is summarized in [BKVP97,BPSK96]. One of the problems with this promising technique is that it causes an effective reordering of the segments from the TCP receiver's point of view. As a simple example, if segments A B C D are sent across a noisy link and segment B is corrupted, segments C and D may have already crossed the link before B can be retransmitted at the link
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   level, causing them to arrive at the TCP receiver in the order A C D
   B.  This segment reordering would cause the TCP receiver to generate
   duplicate ACKs upon the arrival of segments C and D.  If the
   reordering was bad enough, the sender would trigger the fast
   retransmit algorithm in the TCP sender, in response to the duplicate
   ACKs.  Research presented in [MV98] proposes the idea of suppressing
   or delaying the duplicate ACKs in the reverse direction to counteract
   this behavior.  Alternatively, proposals that make TCP more robust in
   the face of re-ordered segment arrivals [Flo99] may reduce the side
   effects of the re-ordering caused by link-layer retransmissions.

   A more high-level approach, outlined in the [DMT96], uses a new
   "corruption experienced" ICMP error message generated by routers that
   detect corruption.  These messages are sent in the forward direction,
   toward the packet's destination, rather than in the reverse direction
   as is done with ICMP Source Quench messages.  Sending the error
   messages in the forward direction allows this feedback to work over
   asymmetric paths.  As noted above, generating an error message in
   response to a damaged packet is problematic because the source and
   destination addresses may not be valid.  The mechanism outlined in
   [DMT96] gets around this problem by having the routers maintain a
   small cache of recent packet destinations; when the router
   experiences an error rate above some threshold, it sends an ICMP
   corruption-experienced message to all of the destinations in its
   cache.  Each TCP receiver then must return this information to its
   respective TCP sender (through a TCP option).  Upon receiving an ACK
   with this "corruption-experienced" option, the TCP sender assumes
   that packet loss is due to corruption rather than congestion for two
   round trip times (RTT) or until it receives additional link state
   information (such as "link down", source quench, or additional
   "corruption experienced" messages).  Note that in shared networks,
   ignoring segment loss for 2 RTTs may aggravate congestion by making
   TCP unresponsive.

3.3.4.3 Implementation Issues
All of the techniques discussed above require changes to at least the TCP sending and receiving stacks, as well as intermediate routers. Due to the concerns over possibly ignoring congestion signals (i.e., segment drops), the above algorithm is not recommended for use in shared networks.
3.3.4.4 Topology Considerations
It is expected that corruption detection, in general would be beneficial in all environments outlined in section 2. It would be particularly beneficial in the satellite/wireless environment over which these errors may be more prevalent.
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3.3.4.5 Possible Interaction and Relationships with Other Research
SACK-based loss recovery algorithms (as described in 3.3.2) may reduce the impact of corrupted segments on mostly clean links because recovery will be able to happen more rapidly (and without relying on the retransmission timer). Note that while SACK-based loss recovery helps, throughput will still suffer in the face of non-congestion related packet loss.

3.4 Congestion Avoidance

3.4.1 Mitigation Description

During congestion avoidance, in the absence of loss, the TCP sender adds approximately one segment to its congestion window during each RTT [Jac88,Ste97,APS99]. Several researchers have observed that this policy leads to unfair sharing of bandwidth when multiple connections with different RTTs traverse the same bottleneck link, with the long RTT connections obtaining only a small fraction of their fair share of the bandwidth. One effective solution to this problem is to deploy fair queueing and TCP-friendly buffer management in network routers [Sut98]. However, in the absence of help from the network, other researchers have investigated changes to the congestion avoidance policy at the TCP sender, as described in [Flo91,HK98].

3.4.2 Research

The "Constant-Rate" increase policy has been studied in [Flo91,HK98]. It attempts to equalize the rate at which TCP senders increase their sending rate during congestion avoidance. Both [Flo91] and [HK98] illustrate cases in which the "Constant-Rate" policy largely corrects the bias against long RTT connections, although [HK98] presents some evidence that such a policy may be difficult to incrementally deploy in an operational network. The proper selection of a constant (for the constant rate of increase) is an open issue. The "Increase-by-K" policy can be selectively used by long RTT connections in a heterogeneous environment. This policy simply changes the slope of the linear increase, with connections over a given RTT threshold adding "K" segments to the congestion window every RTT, instead of one. [HK98] presents evidence that this policy, when used with small values of "K", may be successful in reducing the unfairness while keeping the link utilization high, when a small number of connections share a bottleneck link. The selection of the constant "K," the RTT threshold to invoke this policy, and performance under a large number of flows are all open issues.
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3.4.3 Implementation Issues

Implementation of either the "Constant-Rate" or "Increase-by-K" policies requires a change to the congestion avoidance mechanism at the TCP sender. In the case of "Constant-Rate," such a change must be implemented globally. Additionally, the TCP sender must have a reasonably accurate estimate of the RTT of the connection. The algorithms outlined above violate the congestion avoidance algorithm as outlined in RFC 2581 [APS99] and therefore should not be implemented in shared networks at this time.

3.4.4 Topology Considerations

These solutions are applicable to all satellite networks that are integrated with a terrestrial network, in which satellite connections may be competing with terrestrial connections for the same bottleneck link.

3.4.5 Possible Interaction and Relationships with Other Research

As shown in [PADHV99], increasing the congestion window by multiple segments per RTT can cause TCP to drop multiple segments and force a retransmission timeout in some versions of TCP. Therefore, the above changes to the congestion avoidance algorithm may need to be accompanied by a SACK-based loss recovery algorithm that can quickly repair multiple dropped segments.

3.5 Multiple Data Connections

3.5.1 Mitigation Description

One method that has been used to overcome TCP's inefficiencies in the satellite environment is to use multiple TCP flows to transfer a given file. The use of N TCP connections makes the sender N times more aggressive and therefore can improve throughput in some situations. Using N multiple TCP connections can impact the transfer and the network in a number of ways, which are listed below. 1. The transfer is able to start transmission using an effective congestion window of N segments, rather than a single segment as one TCP flow uses. This allows the transfer to more quickly increase the effective cwnd size to an appropriate size for the given network. However, in some circumstances an initial window of N segments is inappropriate for the network conditions. In this case, a transfer utilizing more than one connection may aggravate congestion.
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   2. During the congestion avoidance phase, the transfer increases the
      effective cwnd by N segments per RTT, rather than the one segment
      per RTT increase that a single TCP connection provides.  Again,
      this can aid the transfer by more rapidly increasing the effective
      cwnd to an appropriate point.  However, this rate of increase can
      also be too aggressive for the network conditions.  In this case,
      the use of multiple data connections can aggravate congestion in
      the network.

   3. Using multiple connections can provide a very large overall
      congestion window.  This can be an advantage for TCP
      implementations that do not support the TCP window scaling
      extension [JBB92].  However, the aggregate cwnd size across all N
      connections is equivalent to using a TCP implementation that
      supports large windows.

   4. The overall cwnd decrease in the face of dropped segments is
      reduced when using N parallel connections.  A single TCP
      connection reduces the effective size of cwnd to half when a
      single segment loss is detected.  When utilizing N connections
      each using a window of W bytes, a single drop reduces the window
      to:

        (N * W) - (W / 2)

   Clearly this is a less dramatic reduction in the effective cwnd size
   than when using a single TCP connection.  And, the amount by which
   the cwnd is decreased is further reduced by increasing N.

   The use of multiple data connections can increase the ability of
   non-SACK TCP implementations to quickly recover from multiple dropped
   segments without resorting to a timeout, assuming the dropped
   segments cross connections.

   The use of multiple parallel connections makes TCP overly aggressive
   for many environments and can contribute to congestive collapse in
   shared networks [FF99].  The advantages provided by using multiple
   TCP connections are now largely provided by TCP extensions (larger
   windows, SACKs, etc.).  Therefore, the use of a single TCP connection
   is more "network friendly" than using multiple parallel connections.
   However, using multiple parallel TCP connections may provide
   performance improvement in private networks.
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3.5.2 Research

Research on the use of multiple parallel TCP connections shows improved performance [IL92,Hah94,AOK95,AKO96]. In addition, research has shown that multiple TCP connections can outperform a single modern TCP connection (with large windows and SACK) [AHKO97]. However, these studies did not consider the impact of using multiple TCP connections on competing traffic. [FF99] argues that using multiple simultaneous connections to transfer a given file may lead to congestive collapse in shared networks.

3.5.3 Implementation Issues

To utilize multiple parallel TCP connections a client application and the corresponding server must be customized. As outlined in [FF99] using multiple parallel TCP connections is not safe (from a congestion control perspective) in shared networks and should not be used.

3.5.4 Topological Considerations

As stated above, [FF99] outlines that the use of multiple parallel connections in a shared network, such as the Internet, may lead to congestive collapse. However, the use of multiple connections may be safe and beneficial in private networks. The specific topology being used will dictate the number of parallel connections required. Some work has been done to determine the appropriate number of connections on the fly [AKO96], but such a mechanism is far from complete.

3.5.5 Possible Interaction and Relationships with Other Research

Using multiple concurrent TCP connections enables use of a large congestion window, much like the TCP window scaling option [JBB92]. In addition, a larger initial congestion window is achieved, similar to using [AFP98] or TCB sharing (see section 3.8).

3.6 Pacing TCP Segments

3.6.1 Mitigation Description

Slow-start takes several round trips to fully open the TCP congestion window over routes with high bandwidth-delay products. For short TCP connections (such as WWW traffic with HTTP/1.0), the slow-start overhead can preclude effective use of the high-bandwidth satellite links. When senders implement slow-start restart after a TCP connection goes idle (suggested by Jacobson and Karels [JK92]),
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   performance is reduced in long-lived (but bursty) connections (such
   as HTTP/1.1, which uses persistent TCP connections to transfer
   multiple WWW page elements) [Hei97a].

   Rate-based pacing (RBP) is a technique, used in the absence of
   incoming ACKs, where the data sender temporarily paces TCP segments
   at a given rate to restart the ACK clock.  Upon receipt of the first
   ACK, pacing is discontinued and normal TCP ACK clocking resumes.  The
   pacing rate may either be known from recent traffic estimates (when
   restarting an idle connection or from recent prior connections), or
   may be known through external means (perhaps in a point-to-point or
   point-to-multipoint satellite network where available bandwidth can
   be assumed to be large).

   In addition, pacing data during the first RTT of a transfer may allow
   TCP to make effective use of high bandwidth-delay links even for
   short transfers.  However, in order to pace segments during the first
   RTT a TCP will have to be using a non-standard initial congestion
   window and a new mechanism to pace outgoing segments rather than send
   them back-to-back.  Determining an appropriate size for the initial
   cwnd is an open research question.  Pacing can also be used to reduce
   bursts in general (due to buggy TCPs or byte counting, see section
   3.2.2 for a discussion on byte counting).

3.6.2 Research

Simulation studies of rate-paced pacing for WWW-like traffic have shown reductions in router congestion and drop rates [VH97a]. In this environment, RBP substantially improves performance compared to slow-start-after-idle for intermittent senders, and it slightly improves performance over burst-full-cwnd-after-idle (because of drops) [VH98]. More recently, pacing has been suggested to eliminate burstiness in networks with ACK filtering [BPK97].

3.6.3 Implementation Issues

RBP requires only sender-side changes to TCP. Prototype implementations of RBP are available [VH97b]. RBP requires an additional sender timer for pacing. The overhead of timer-driven data transfer is often considered too high for practical use. Preliminary experiments suggest that in RBP this overhead is minimal because RBP only requires this timer for one RTT of transmission [VH98]. RBP is expected to make TCP more conservative in sending bursts of data after an idle period in hosts that do not revert to slow start after an idle period. On the other hand, RBP makes TCP more aggressive if the sender uses the slow start algorithm to start the ACK clock after a long idle period.
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3.6.4 Topology Considerations

RBP could be used to restart idle TCP connections for all topologies in Section 2. Use at the beginning of new connections would be restricted to topologies where available bandwidth can be estimated out-of-band.

3.6.5 Possible Interaction and Relationships with Other Research

Pacing segments may benefit from sharing state amongst various flows between two hosts, due to the time required to determine the needed information. Additionally, pacing segments, rather than sending back-to-back segments, may make estimating the available bandwidth (as outlined in section 3.2.4) more difficult.

3.7 TCP Header Compression

The TCP and IP header information needed to reliably deliver packets to a remote site across the Internet can add significant overhead, especially for interactive applications. Telnet packets, for example, typically carry only a few bytes of data per packet, and standard IPv4/TCP headers add at least 40 bytes to this; IPv6/TCP headers add at least 60 bytes. Much of this information remains relatively constant over the course of a session and so can be replaced by a short session identifier.

3.7.1 Mitigation Description

Many fields in the TCP and IP headers either remain constant during the course of a session, change very infrequently, or can be inferred from other sources. For example, the source and destination addresses, as well as the IP version, protocol, and port fields generally do not change during a session. Packet length can be deduced from the length field of the underlying link layer protocol provided that the link layer packet is not padded. Packet sequence numbers in a forward data stream generally change with every packet, but increase in a predictable manner. The TCP/IP header compression methods described in [DNP99,DENP97,Jac90] reduce the overhead of TCP sessions by replacing the data in the TCP and IP headers that remains constant, changes slowly, or changes in a predictable manner with a short "connection number". Using this method, the sender first sends a full TCP/IP header, including in it a connection number that the sender will use to reference the connection. The receiver stores the full header and uses it as a template, filling in some fields from the limited
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   information contained in later, compressed headers.  This compression
   can reduce the size of an IPv4/TCP headers from 40 to as few as 3 to
   5 bytes (3 bytes for some common cases, 5 bytes in general).

   Compression and decompression generally happen below the IP layer, at
   the end-points of a given physical link (such as at two routers
   connected by a serial line).  The hosts on either side of the
   physical link must maintain some state about the TCP connections that
   are using the link.

   The decompresser must pass complete, uncompressed packets to the IP
   layer.  Thus header compression is transparent to routing, for
   example, since an incoming packet with compressed headers is expanded
   before being passed to the IP layer.

   A variety of methods can be used by the compressor/decompressor to
   negotiate the use of header compression.  For example, the PPP serial
   line protocol allows for an option exchange, during which time the
   compressor/decompressor agree on whether or not to use header
   compression.  For older SLIP implementations, [Jac90] describes a
   mechanism that uses the first bit in the IP packet as a flag.

   The reduction in overhead is especially useful when the link is
   bandwidth-limited such as terrestrial wireless and mobile satellite
   links, where the overhead associated with transmitting the header
   bits is nontrivial.  Header compression has the added advantage that
   for the case of uniformly distributed bit errors, compressing TCP/IP
   headers can provide a better quality of service by decreasing the
   packet error probability.  The shorter, compressed packets are less
   likely to be corrupted, and the reduction in errors increases the
   connection's throughput.

   Extra space is saved by encoding changes in fields that change
   relatively slowly by sending only their difference from their values
   in the previous packet instead of their absolute values.  In order to
   decode headers compressed this way, the receiver keeps a copy of each
   full, reconstructed TCP header after it is decoded, and applies the
   delta values from the next decoded compressed header to the
   reconstructed full header template.

   A disadvantage to using this delta encoding scheme where values are
   encoded as deltas from their values in the previous packet is that if
   a single compressed packet is lost, subsequent packets with
   compressed headers can become garbled if they contain fields which
   depend on the lost packet.  Consider a forward data stream of packets
   with compressed headers and increasing sequence numbers.  If packet N
   is lost, the full header of packet N+1 will be reconstructed at the
   receiver using packet N-1's full header as a template.  Thus the
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   sequence number, which should have been calculated from packet N's
   header, will be wrong, the checksum will fail, and the packet will be
   discarded.  When the sending TCP times out and retransmits a packet
   with a full header is forwarded to re-synchronize the decompresser.

   It is important to note that the compressor does not maintain any
   timers, nor does the decompresser know when an error occurred (only
   the receiving TCP knows this, when the TCP checksum fails).  A single
   bit error will cause the decompresser to lose sync, and subsequent
   packets with compressed headers will be dropped by the receiving TCP,
   since they will all fail the TCP checksum. When this happens, no
   duplicate acknowledgments will be generated, and the decompresser can
   only re-synchronize when it receives a packet with an uncompressed
   header.  This means that when header compression is being used, both
   fast retransmit and selective acknowledgments will not be able
   correct packets lost on a compressed link.  The "twice" algorithm,
   described below, may be a partial solution to this problem.

   [DNP99] and [DENP97] describe TCP/IPv4 and TCP/IPv6 compression
   algorithms including compressing the various IPv6 extension headers
   as well as methods for compressing non-TCP streams.  [DENP97] also
   augments TCP header compression by introducing the "twice" algorithm.
   If a particular packet fails to decompress properly, the twice
   algorithm modifies its assumptions about the inferred fields in the
   compressed header, assuming that a packet identical to the current
   one was dropped between the last correctly decoded packet and the
   current one.  Twice then tries to decompress the received packet
   under the new assumptions and, if the checksum passes, the packet is
   passed to IP and the decompresser state has been re-synchronized.
   This procedure can be extended to three or more decoding attempts.
   Additional robustness can be achieved by caching full copies of
   packets which don't decompress properly in the hopes that later
   arrivals will fix the problem.  Finally, the performance improvement
   if the decompresser can explicitly request a full header is
   discussed.  Simulation results show that twice, in conjunction with
   the full header request mechanism, can improve throughput over
   uncompressed streams.

3.7.2 Research

[Jac90] outlines a simple header compression scheme for TCP/IP. In [DENP97] the authors present the results of simulations showing that header compression is advantageous for both low and medium bandwidth links. Simulations show that the twice algorithm, combined with an explicit header request mechanism, improved throughput by 10-15% over uncompressed sessions across a wide range of bit error rates.
Top   ToC   RFC2760 - Page 29
   Much of this improvement may have been due to the twice algorithm
   quickly re-synchronizing the decompresser when a packet is lost.
   This is because the twice algorithm, applied one or two times when
   the decompresser becomes unsynchronized, will re-sync the
   decompresser in between 83% and 99% of the cases examined.  This
   means that packets received correctly after twice has resynchronized
   the decompresser will cause duplicate acknowledgments.  This re-
   enables the use of both fast retransmit and SACK in conjunction with
   header compression.

3.7.3 Implementation Issues

Implementing TCP/IP header compression requires changes at both the sending (compressor) and receiving (decompresser) ends of each link that uses compression. The twice algorithm requires very little extra machinery over and above header compression, while the explicit header request mechanism of [DENP97] requires more extensive modifications to the sending and receiving ends of each link that employs header compression. Header compression does not violate TCP's congestion control mechanisms and therefore can be safely implemented in shared networks.

3.7.4 Topology Considerations

TCP/IP header compression is applicable to all of the environments discussed in section 2, but will provide relatively more improvement in situations where packet sizes are small (i.e., overhead is large) and there is medium to low bandwidth and/or higher BER. When TCP's congestion window size is large, implementing the explicit header request mechanism, the twice algorithm, and caching packets which fail to decompress properly becomes more critical.

3.7.5 Possible Interaction and Relationships with Other Research

As discussed above, losing synchronization between a sender and receiver can cause many packet drops. The frequency of losing synchronization and the effectiveness of the twice algorithm may point to using a SACK-based loss recovery algorithm to reduce the impact of multiple lost segments. However, even very robust SACK- based algorithms may not work well if too many segments are lost.


(page 29 continued on part 3)

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