Performance analysis of TCP/IP over high bandwidth delay product networks

Performance analysis of TCP/IP over high bandwidth delay product networks

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Performance analysis of TCP/IP over high bandwidth delay product networks
Kerkar, Subodh
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University of South Florida
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ABSTRACT: In today's Internet scenario, the current TCP has performed reasonably well. As the Internet has scaled up in load, speed, size and connectivity by the order of six over the past fifteen years, the TCP has consistently avoided severe congestion throughout this same period. Applications involving high performance computings such as bulk-data transfer, multimedia Web streaming, and computational grids demand high bandwidth. These applications usually operate over wide-area networks and, hence, performance over wide-area networks has become a critical issue. Future applications will need steady transfer rates in the order of gigabits per second to support collaborative work. TCP, which is the most widely used protocol, is expected to be used in these scenarios. It has been shown that TCP doesn't work well in this new environment, and several new TCP versions have been developed in recent years to address this issue. To date, there has not been a performance evaluation of various TCP protocols. In this thesis, various TCP versions 3/4 Tahoe, Reno, Newreno, Vegas, Westwood, Sack, Highspeed TCP, Scalable TCP 3/4 have been evaluated for their performance over high bandwidth delay product networks. It was found that the flow and congestion control mechanism used in TCP was unable to reach full utilization on high-speed links. Also discussed in this Thesis are fairness issues related to these new protocols with respect to themselves and with others.
Thesis (M.S.C.S.)--University of South Florida, 2004.
Includes bibliographical references.
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by Subodh Kerkar.

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Performance analysis of TCP/IP over high bandwidth delay product networks
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by Subodh Kerkar.
[Tampa, Fla.] :
University of South Florida,
Thesis (M.S.C.S.)--University of South Florida, 2004.
Includes bibliographical references.
Text (Electronic thesis) in PDF format.
System requirements: World Wide Web browser and PDF reader.
Mode of access: World Wide Web.
Title from PDF of title page.
Document formatted into pages; contains 75 pages.
ABSTRACT: In today's Internet scenario, the current TCP has performed reasonably well. As the Internet has scaled up in load, speed, size and connectivity by the order of six over the past fifteen years, the TCP has consistently avoided severe congestion throughout this same period. Applications involving high performance computings such as bulk-data transfer, multimedia Web streaming, and computational grids demand high bandwidth. These applications usually operate over wide-area networks and, hence, performance over wide-area networks has become a critical issue. Future applications will need steady transfer rates in the order of gigabits per second to support collaborative work. TCP, which is the most widely used protocol, is expected to be used in these scenarios. It has been shown that TCP doesn't work well in this new environment, and several new TCP versions have been developed in recent years to address this issue. To date, there has not been a performance evaluation of various TCP protocols. In this thesis, various TCP versions 3/4 Tahoe, Reno, Newreno, Vegas, Westwood, Sack, Highspeed TCP, Scalable TCP 3/4 have been evaluated for their performance over high bandwidth delay product networks. It was found that the flow and congestion control mechanism used in TCP was unable to reach full utilization on high-speed links. Also discussed in this Thesis are fairness issues related to these new protocols with respect to themselves and with others.
Adviser: Miguel Labrador.
Dissertations, Academic
x Computer Science
t USF Electronic Theses and Dissertations.
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Performance Analysis Of Tcp/Ip Over High Bandwidth Delay Product Networks by Subodh Kerkar A thesis submitted in partial fulfillment of the requirements for the degree of Master of Science in Computer Science Department of Computer Science & Engineering College of Engineering University of South Florida Major Professor: Miguel Labrador, Ph.D. Srinivas Katkoori, Ph.D. Dewey Rundus, Ph.D. Date of Approval: July 6, 2004 Keywords: highspeed, fast-tcp, optical, evaluation, reno Copyright 2004 Subodh Kerkar


i TABLE OF CONTENTS LIST OF FIGURES ii ABSTRACT iv CHAPTER 1 INTRODUCTION 1 1.1 Motivation for Present Work 1 1.2 Contribution of this Thesis 5 1.3 Outline of this Thesis 5 CHAPTER 2 LITERATURE REVIEW 7 2.1 The Transmission Control Protocol (TCP) 7 2.2 TCP Versions 10 2.2.1 Old Versions 10 2.2.2 TCP for High Bandwidth-Delay Product Networks 23 2.3 Router Queueing Techniques 26 2.3.1 Droptail 26 2.3.2 Random Early Detection (RED) 28 CHAPTER 3 METHODOLOGY 31 3.1 Simulation Topology for a HBDP Link ( Single source/single sink ) 31 3.2 Simulation Topology for Fairness to Itself and Others Scenario (Two Sources/ Two Sinks) 33 CHAPTER 4 PERFORMANCE EVALUATION 35 4.1 Single Source – Single Sink Topology 35 4.2 Two Sources – Two Sinks Topology ( Fairness to Others ) 46 4.3 Two Sources – Two Sinks Topology ( Fairness to Itself ) 57 CHAPTER 5 CONCLUSIONS 64 REFERENCES 66


ii LIST OF FIGURES Figure 1. Variation of TCP’s Congestion Window With Time 10 Figure 2. Scalable TCP Congestion Window Properties 26 Figure 3. Response Curves of Scalable TCP and Regular TCP 27 Figure 4. Drop Probabilities of Different RED Modes 29 Figure 5. Simulation Topology for One Source – One Sink 32 Figure 6. Simulation Topology for 2 Sources – 2 Sinks (Fairness to Itself) 33 Figure 7. Simulation Topology for 2 Sources – 2 Sinks (Fairness to Others) 34 Figure 8. Normalized Throughput of Protocols Under Consideration 37 Figure 9. TCP Sequence Numbers for Link Bandwidth 1Gbps 37 Figure 10. Congestion Windows of All the Protocols When Link Speed 1Gbps 39 Figure 11. Slowstart Times 42 Figure 12. Packet Loss Rate 43 Figure 13. Recovery Time 44 Figure 14. Performance of HSTCP Against TCP Protocols 48 Figure 15. HSTCP Fairness to Vegas and Westwood 49 Figure 16. Performance of TCP Westwood 51 Figure 17. Behavior of TCP Westwood’s Cwnd and Ssthresh when the Bottleneck Link is Set to 1Gbps 51


iii Figure 18. Performance of Vegas 52 Figure 19. Performance of Scalable TCP With Other Protocols 55 Figure 20. Performances of Tahoe, Newreno, Reno and SACK TCP 57 Figure 21. Fairness to Itself (Newreno, Reno, SACK, Scalable, Tahoe) 62 Figure 22. Fairness to Itself (Vegas) 63


iv PERFORMANCE ANALYSIS OF TCP/IP OVER HIGH BANDWIDTH DELAY PRODUCT NETWORKS Subodh Kerkar ABSTRACT In todayÂ’s Internet scenario, the current TCP has performed reasonably well. As the Internet has scaled up in load, speed, size and connectivity by the order of six over the past fifteen years, the TCP has consistently avoided severe congestion throughout this same period. Applications involving high performance computings such as bulk-data transfer, multimedia Web streaming, and computational grids demand high bandwidth. These applications usually operate over wide-area networks and, hence, performance over wide-area networks has become a critical issue. Future applications will need steady transfer rates in the order of gigabits per second to support collaborative work. TCP, which is the most widely used protocol, is expected to be used in these scenarios. It has been shown that TCP doesnÂ’t work well in this new environment, and several new TCP versions have been developed in recent years to address this issue. To date, there has not been a performance evaluation of various TCP protocols. In this thesis, various TCP versions Tahoe, Reno, Newreno, Vegas, Westwood, Sack, Highspeed TCP, Scalable TCP have been evaluated for their performance over high


v bandwidth delay product networks. It was found that the flow and congestion control mechanism used in TCP was unable to reach full ut ilization on high-speed links. Also discussed in this Thesis are fairness issues related to these new protocols with respect to themselves and with others.


1 CHAPTER 1 INTRODUCTION 1.1 Motivation of Present Work Todays Internet scenario is on th e brink of testing the widely used communication protocols, the Transmission Control Protocol (TCP). Applications needing connections with steady transfer rates are consistently showing up on the horizon. In the first International Work shop on Protocols for Fast Long-Distance Networks [1], several presentations put forwar d the point of future network applications requiring steady transfer rates in the order of gigabits per second to suppor t collaborative work. Although the raw transmission bit rate of next generation networks will definitely support these high speeds, it is unknown whet her the communication protocols will become the performance bottleneck. This is the case of TCP when running over optical networks using the Dynamic Bandwidth on Demand (DBOD) service or high bandwidthdelay product networks. Many communities use such networks and need to distribute a substantial amount of data over them. The large datasets collect ed by the High Energy Physics, Bioinformatics and Radio Astronomy communities require global distribution for the data to be analyzed effectively. This is one example of su ch a network. Internet paths operating in this region are usually re ferred to as Long Fat Pipes. High capacity packet satellite channels e.g. DARPAs wideband net are called LFNs. Terrestrial


2 fiber-optical paths also fall into the LFN cl ass, which is moving out of the domain for which TCP was originally crafted. Since TCP is expected to be used in this scenario, its most important aspect Flow and Congestion control mechanism should perform as the network evolves. As the bandwidth or the delay goes on increasing, this mechanism causes problems as TCP reacts adversely in such cases. According to Sally Floyd [3], TCP faces three main difficulties: Bit Error Rate (BER), slow-start mechanism, and congestion avoidance mechanism. First, the Bit Error Rate (BER) of the links on which high data transfer is expected should be very small, much smalle r than the current BER. Secondly, The slowstart phase of the TCPs congestion cont rol mechanism sees the congestion window increase exponentially. In high bandwidt h-delay product scenar ios, the congestion window increases to a very la rge value and a large number of packets is dropped once the channel capacity is filled. And finally, TC P has been shown to waste a considerable amount of bandwidth in its congestion a voidance phase when the window increases linearly. Hence, in high capacity link and l ong propagation delays, it will take TCP a very long time to fill up the whole pipe. As the high-bandwidth network becomes more widely used, the problems of the Additive Increase Multiplicative Decrease (AIM D) algorithm of the TCP becomes more apparent. For many years, network resear ch has been seeing improvements in TCP efficiency and stability. As a result, different versions of TCP including Tahoe, Reno, Vegas [6], Sack [7,8] and Westwood [9] ha ve been developed. These variants brought about a significant amount of improvement, as a result of their improved congestion control mechanism, selected acknowledgement, and fast recovery; bu t all these variants


3 had the same unchanged window-based algorithm as is specified in RFC 2581 [10]. There have been many other TCP variants that employ techniques other than the windowbased algorithm. Mechanisms that employ ra te-based techniques involve the controlling of the congestion window, based on feedback re ceived from routers. But these kinds of techniques are unlikely to be in corporated in the future, as they require the modifications of routers by the ISPs. In th e last two to three years, researchers have come upon many proposals for the modification of TCP on the senders side for its use on High Bandwidth-Delay Product links. A few exampl es are Sally Floyds High-speed TCP [3], Kellys Scalable TCP [4], Ca ltechs FAST protocol [5]. Today, there have been many TCP vers ions showing significant performance improvement over the original versions. These versions have been te sted on a stand-alone basis or, at most, with the classic TCP Reno version. The performance of recent TCP versions like TCP Westwood and FAST TCP has only been compared with TCP Reno [4]. Research on TCP Westwood has explor ed its improvement on high-bandwidth networks and its friendliness with just tw o TCP versions, Reno and Vegas. The Caltech group at UCLA has conducted research on FAST TCP and has compared its results with respect to TCP Reno. Scalable TCP, which is based on Highspeed TCP, is still wide open for exploration [6]. Explicit Control Protocol (XCP) [11] also compares its performance with TCP Reno. In our discussions, we will ca tegorize the protocols under study as old and new protocols. The new protoc ols for the High Bandwidth-Delay Product networks include the HSTCP and the Scalable TCP; the rest fall under the old category. Apart from the research mentioned above, ther e has not been a perf ormance evaluation of any kind conducted where the old protocols a nd the new have been compared together


4 under similar conditions. So far the matter of how a set of new and old protocols would behave separately if run on a link of varying capacity has not been tested. Features like Packet Loss Ratio, slow-start time, sequence number, congestion window, and throughput have not been compared or analyzed In this thesis, a performance evaluation of these old and new protocols is conducted and the results analyzed. It is also known that TCPs throughput is inversely proportional to the round trip time (RTT). Hence, fairness issues come into play. Connections with larger RTT take a longer time to fulfill the avai lable bandwidth over high-speed links; connections with shorter RTT, which share the same segment of link, obtain more bandwidth resources. In addition, the TCP congestion control is depe ndent on the number of flows. Suppose there are N connections sharing the same link, all th e connections will increa se the sending rate by one segment every RTT, so the overall incr ease of all TCP flows is a function of N. As a greater number of flows compete for a fair share of bandwidth, fairness to each flow becomes an important factor. Unfairness is bound to result when more than one flow having different RTTs are competing for the same bottleneck link. These were the problems that existed in the older versions of TCP. It has yet to be seen whether these problems still persist with the new protocols. As mentioned earlier, these protocols have been tested and compared mostly with TCP Re no; they have not been tried against their new highspeed counterparts. So far, no analysis has been do ne on the friendliness issues of these protocols. An analysis for fairness of these protocols in various combinations like Highspeed TCP Scalable TCP, FAST TCP XCP, etc., would result in a very interesting discussion. It is not known how th ese protocols react w ith one another over high bandwidth delay product links; this is still an area to be explore d. In this thesis, the


5 friendliness of these protocols when conte nded with the old and new ones over a common channel will be analyzed. 1.2 Contribution of this Thesis This thesis makes the following contribut ion, meant to address the above-mentioned aspects. 1. A performance evaluation of TCP Ta hoe, Reno, Newreno, Sack, Westwood and Highspeed TCP, and Scalable TCP over High Bandwidth Delay product networks. 2. A study of the fairness of new protocol s when sharing the same bottleneck link with its peer protocols as well as with themselves. 3. Analysis of the effect of Droptail and RED queuing techniques on performance and fairness. 1.3 Outline of this Thesis This thesis is organized as follows: Chapter 2 offers an elaborate view of the various TCP protocols under consideration. It is categorized into 4 parts: 1). The first part describes the basic operation of the TCP protocol and the phases of a congestion window; 2). The second part offers a detailed description of the eight protocols is given with respect to their congestion control mechanisms and the research conducted in relatio n to them so far; 3). These protocols are classified into two parts: the older versi on of TCP and the more recent versions for HBDPN. TCP Tahoe, Reno, Newreno, SACK, VEGAS and Westwood have been


6 explained under the former, while HSTCP and Scalable TCP have been explained under the latter section; 4). The fourth part e xplains the two router queue management techniques, Droptail and RED, whic h are used in the simulations. Chapter 3 explains the three simulati on topologies used in the experiments conducted to evaluate the perf ormance of these protocols wh en running on a stand alone basis, when running against ot her protocols, and when runni ng against each other for issues concerning fairness. The router queuing techniques and the parameters used in the simulation scenarios are explained in detail. Chapter 4 discusses the results of the simulations conducted along with graphs for throughput, congestion window, slow start time, packet loss ratio, recovery time, and throughput ratios. The behavior of these prot ocols over the network topologies mentioned above is explained in detail. Performance of these protocols when competing with other protocols over a bottleneck li nk is analyzed. These protoc ols are also evaluated and analyzed when competing with themselves over a bottleneck link. Chapter 5 concludes the thesis, pointing out the best protocol for the HBDPN. It also discusses future work and additional e xperiments that could be conducted in this direction.


7 CHAPTER 2 LITERATURE REVIEW Congestion avoidance, slow-sta rt, fast retransmit and fast recovery mechanisms of TCP are very well known and studied in vast detail [7],[10],[11],[13],[15], they are considered the building bloc ks of regular TCP Reno, Ne wreno, Tahoe and SACK [11], [12], [13], [14]. In this di scussion, TCP Vegas, Westwood, High-speed TCP and Scalable TCP are also put forward for comparison. In this section, the m echanism of TCP is discussed, and a few terms are explained in brief. The above-mentioned TCP protocols are discussed with respect to their congesti on control mechanisms and the research done on them to date. Later in the chapter the performance of the above-mentioned protocols on high bandwidth delay product networks is studied. 2.1 The Transmission Control Protocol (TCP) In order that effectiv e communication take place between the sender and the receiver, TCP uses error, flow and congesti on control algorithms. Th ese include the Slow Start, Congestion Avoidance, Fast Retransm it and Fast Recovery mechanisms. Tahoe is the oldest and the simplest of all the TCP versions. The rest of this section details the mechanisms on which a TCP Tahoe version performs. A TCP receiver uses cumulative acknowledgements to specify the sequence number of the next packet the receiver expects. The generation of acknowledgements


8 allows the sender to get continuous feedback from the receiver. Every time a sender sends a segment, the sender starts a timer a nd waits for the acknowledgement. If the timer expires before the acknowledgment is receive d, TCP assumes that the segment is lost and retransmits it. This expiration of the timer is referred to as a timeout. If the acknowledgement is received, however, TCP re cords the time at which the segment was received and calculates the Round Trip Time (RTT). A weighted moving average of the RTT is maintained and used to calculat e the timeout value for each segment. TCP uses a sliding window mechanism to achieve flow control that allows multiple packets to be present in flight so th at the available bandwidth can be used more efficiently. This keeps the sender from overw helming the receivers buffers. However, the most important variation of TCPs s liding window mechanis m over other sliding window mechanisms is the variation of the window size in TCP with respect to time. If the receiver is unable to send acknowledgements at the ra te at which the sender is sending data, the sender reduces its sending wi ndow. The sender and receiver agree upon the number of packets that a sender can send without being acknowledged, and upon number of packets the receiver is able to re ceive, before its buffers become overwhelmed. This is accomplished by the Advertised Window (AWND) parameter, which is the receiver side estimate of the number of pack ets it is able to receive without overflowing its buffer queues. TCP also includes several variables fo r performing congestion control. The CWND variable defines the number of consecutive packets that a sender is able to send before receiving an acknowledgement and the variable is changed based on network conditions. At any given point in time the sender is allowed to send as many consecutive


9 packets as provided by the minimum of CW ND and AWND, there by considering the condition of the receiver and th e network simultaneously. At the connection startup time, CWND is started at 1 and incremented by 1 for every acknowledgment received thereafter. This leads to an exponential growth of the transmission rate and is referred to as the Slow Start algorithm The growth continues until the Slow Start Threshold (SSTHRESH) is reached. After that, the CWND is increased by 1 for every RTT, presenting a linear growth characteristic in the Congestion Avoidance Phase. This is the additive increase mechanism of congestion c ontrol in TCP, as the transmission rate additively increases for every successful packet transmission. The Congestion Avoidance phase continues increasing the CWND until a packet is lost, in which case the congestion window is reduced to 1 and TCP enters the Slow Start phase. This is a multiplicative decrease since CWND reduces to a value of 1 as shown in Figure 1. The loss of a packet in the congestion avoidance state leads to a timeout in Tahoe. TCP includes error control mechanisms to provide a reliable service. TCP detects packet losses by means of the retransmissi on time out or the reception of 3 duplicate acknowledgements (DUPACKS). Upon the receipt of 3 DUPACKS asking for the retransmission of the same packet, TCP assumes that the segment is lost due to congestion. At this point, TCP retransmits the missing packet instead of waiting for a timeout to occur. This is called the Fast Retransmit algorithm. A TCP Tahoe sender has these three main algorithms available to perf orm error and congestion control. Because of the drastic reduction of its CWND, TCP Tahoe has been shown to provide very low throughput.


Figure 1. Variation of TCPs Congestion Window with Time. 2.2 TCP Versions The various TCP versions that are to be studied are classified into 2 parts: old versions and new versions [TCP for High Bandwidth Delay Product Network (HBDPN)]. While the older versions include Tahoe, Reno, Newreno, SACK, Vegas and Westwood, the more recent ones, which address the performance issues of TCP over high bandwidth delay product networks, include High-speed TCP, and Scalable TCP. The details of each of these are investigated in the following sections. 2.2.1 Older Versions The TCP versions falling under this category have been performing very well over not-so-large bandwidth networks over a long period of time. These are the 10


11 versions that have faced many challenge s when run on very high bandwidth delay product networks. These categories of protoc ols are discussed one by one as follows. TCP Tahoe Early TCP implementations followed a go-back-n technique using cumulative positive acknowledgement, and required a retransmit timer expiration to resend data lost during the flight. These TCPs did very litt le to handle congestion. TCP Tahoe added a number of new algorithms and refinement s to earlier implementations. The new algorithms include slow-start, congestion avoi dance, and fast-retransmit [15]. One of the major refinements was the modification of the roundtrip time estim ator used to set retransmission timeout values. Initially, it wa s assumed that lost packets represented congestion. Therefore, it was assumed by Jacobs on that when a packet loss occurred, the sender should lower its share of the bandwidth. The mechanism of TCP Tahoe is the same as explained in section 2.1.TCP Tahoe does not deal well with multiple packet drops within a single window of data. The two phases in increasing the conge stion window, the slow-start and the congestion avoidance phases can be summed up with the following equations. Slow-start phase: cwnd = cwnd + 1 if cwnd < ssthresh Congestion avoidance phase: cwnd = cwnd + 1/cwnd if cwnd ssthresh


12 where ssthresh is the threshold value at whic h TCP changes its phase from slow-start to congestion avoidance. When a segment lo ss is detected, the cwnd and ssthresh are updated as follows. ssthresh = cwnd/2 cwnd = 1 During the time when TCP Tahoe came up, the network environment and the applications that were being used did not demand high ba ndwidth links. Hence, this variant of TCP did not have to face the challenge of s caling to the high bandwidth delay product network. Studies done in [16] reflect that TCP Tahoe has major drawbacks as a means of providing data services over a multimedia network, since random loss resulting from fluctuations in real-time traffic can lead to significant thr oughput deterioration in the high bandwidth delay product network. The result s of these studies conclude that the performance is degraded when the product of the loss probability and the square of the bandwidth-delay product is la rge. Also, for the high bandwid th delay product network, TCP is extremely unfair towards connec tions with higher propagation delays. TCP Reno The TCP Reno [19] implementation m odified the sender to incorporate a mechanism called fast recovery. Unlike Tahoe, Reno does not empty the pipe unnecessarily on the receipt of a few numbers of dupacks. Instead, with the mechanism of fast recovery the congestion window is set to ha lf its previous value. The idea is that the only way for a loss to be detected via a time out and not via the receipt of a dupack is when the flow of packets and ACKs has completely stopped, which would be an


13 indication of heavy congestion. But if the sender is still able to rece ive an ACK, then it should not fall back into slow-s tart, as it does in the case of TCP Tahoe. This case does not imply heavy congestion, since the flow s till exists, but the sender should send with relatively less vigor, utilizing a lower amount of resources. The mechanism of fast recovery comes into picture at this stage. After receiving a certain number of dupacks, the sender will retransmit the lost packet; but, unl ike Tahoe, it will not fall back into slowstart. It will rather take adva ntage of the fact that the curr ently existing flow should keep on sending, albeit using fewer resources. By using fast recovery the sender uses a congestion window that is half the size of the congestion wi ndow present just before the loss. This factor forces Reno to send less pack ets out until it knows that it is feasible to send more. Therefore, it has indeed reduced its utilization of th e network. Although Reno TCP is better than Tahoe in cases of single pa cket loss, Reno TCP is not much better than Tahoe when multiple packets are lost within a window of data [15], [17]. Fast recovery ensures that the pipe does not become empty. Therefore, slow-start is executed only when a packet is timed out. This is implemente d by setting ssthresh to half the current congestion window size and then setting the congestion window to 1 segment, causing the TCP connection to slow-start until the ss thresh is reached; then it goes into the congestion avoidance phase like in the case of Tahoe. Th e Reno TCP represented in equation form looks like this. Slow-start phase cwnd = cwnd + 1 When a segment is detected, the fast re transmission algorithm halves the congestion window.


14 ssthresh = cwnd/2 cwnd = ssthresh TCP Reno then enters fast recovery phase. In this phase, the window size is increased by one segment when a duplicate acknowledgement is received; and th e congestion window is restored to ssthresh when a non-dup licate acknowledgement corresponding to the retransmitted segments is received. The basic problem in TCP Reno is that fast retransmit assumes that only one segment was lost. This can result in loss of ACK clocking and timeout s if more than one segment is lost. Reno faces several problem s when multiple packet losses occur in a window of data. This usually occurs when fast retransmit and fast recovery is invoked. It is invoked several times in succession lead ing to multiplicative d ecreases of cwnd and ssthresh impacting the throughput of the conne ction. Another problem with Reno TCP is ACK starvation. This occurs due to the ambi guity of duplicate ACKs. The sender reduces the congestion window when it enters fast re transmit; it receives dupack s that inflate the congestion window so that it sends new pack ets until it fills its sending window. It then receives a non-dupack and exits fast recovery. However, due to multiple losses in the past, the ACK will be followed by 3 dupacks sign aling that another segment was lost; this way, fast retransmit is entered again after another reducing of sst hresh and cwnd. This happens several times in succession and duri ng this time the left edge of the sending window advances only after each successive fast retransmit; and the amount of data in flight eventually becomes more than the congestion window. When there are no more ACKs to be received, the sender stalls and recovers from this deadlock only through


15 timeout, which causes slow-start. There ar e two solutions available for the above problems: Newreno and TCP SACK. TCP Newreno The TCP Newreno [11] modifies the fast re transmit and fast recovery mechanisms of Reno TCP. These modifications are impleme nted to fix the drawbacks of TCP Reno. Here, the wait for the retransmit timer is elimin ated when multiple packets are lost from a window. Newreno is the same as Reno but applies more intelligence during fast recovery. It utilizes the idea of partial ACKs. When there are multiple packet losses, the ACK for the retransmitted packet will acknowledge some but not all the packets sent before the fast retransmit. In Newreno, a partial ACK is ta ken as an indication of another lost packet and as such the sender transmits the firs t unacknowledged packet. Unlike Reno, partial ACKs do not take Newreno out of fast recove ry. This way Newreno retransmits 1 packet per RTT until all lost packets are retransmitted, and avoids requiring multiple fast retransmits from a single window of data This Newreno modification of Reno TCP defines a fast recovery procedure that begi ns when three duplicate ACKs are received and ends when either a retransmission timeout occurs or an ACK arrives that acknowledges all of the data up to and including the data that was outstanding when the fast recovery procedure began [18]. The Newreno algorithm can be explained in the following steps: On the receipt of the third dupack, if the sender is not already in fast recovery procedure, then set ssthresh to no more than the value below [19]. ssthresh = max(flightsize/2, 2*MSS)


16 Also, remember the highest sequence number transmitted in a variable. Retransmit the lost packet and set cwnd to ssthresh + 3*MSS. This artificially inflates the congestion window by the numbe r of segments that have left the network and that the receiver has buffered. For each additional dupack received, increment the congestion window by MSS. Transmit a segment, if allowed by the new value of cwnd and the receivers advertised window. When an ACK arrives that acknowledges new data, this ACK could be the acknowledgement elicited by the retransmission from step 2, or one elicited by a later retransmission. TCP Vegas In 1994, Brakmo, O'Malley and Peterson came with a new TCP implementation called Vegas that achieves between 40% and 70% better throughput and 1/5 to 1/2 the losses when compared with TCP Reno. TCP Ve gas [20] also had all the changes and modifications on the sender side. In Reno, th e RTT is computed using a coarse-grained timer, which does not give an accurate estimat e of RTT. Tests conducted conclude that for losses that resulted in a timeout, it t ook Reno an average of 1100ms from the time it sent a segment that was lost, until it timed out and resent the segment; whereas less than 300ms would have been the correct timeout in terval had a more accurate clock been used. TCP Vegas fixes this problem using a finer co arse-grained timer. Vegas also changed the retransmission mechanism. The system clock is read and saved each time a segment is sent; when an ACK arrives, the clock is r ead again and the RTT calculation is computed


17 using this time and the timestamp recorded for the relevant segment. With the use of this accurate RTT, retransmission is decided as follows: When a dupack is received, Vegas checks to see if the new RTT is greater than RTO. If it is, Vegas retransmits the segment without having to wait for the 3rd dupack. Wh ereas, when a non-dupack is received, if it is the first or second one afte r a retransmission, Vegas checks again to see if RTT > RTO; if so, then the segment is retransmitted. This process catches any other segment that may have been lost previous to the retransm ission without requiring a waiting period to receive a dupack. Vegas treats the receipt of certain ACKs as a tr igger to check if a timeout should happen, but still contain Renos timeout code in case this mechanism fails to recognize a lost segment. Vegas' congestion avoidance actions ar e based on changes in the estimated amount of extra data in the network. Vegas defines the RTT of a connection as its BaseRTT when the connection is not congeste d. In practice, it is the minimum of all measured roundtrip times and mostly it is the RTT of the first segment sent by the connection before the router queues increase. Vegas uses this value to calculate the expected throughput. Secondly, it ca lculates the current actual sending rate. This is done by recording the sending time for a segment, recording how many bytes are transmitted between the time that segment is sent and its acknowledgement is received, computing the RTT for the segment when its acknowledge ment arrives, and dividing the number of bytes transmitted by the sample RTT. This calculation is done once per round trip time. Thirdly Vegas compares actual to exp ected throughput and adjusts the window accordingly. Difference between the actual a nd expected throughput is recorded. Vegas defines two thresholds, and which roughly correspond to having too little and too


18 much extra data in the network, respectiv ely. Following is the mechanism of the congestion control in equation form. Diff is the difference between actual and expected throughput. Diff < 0 : change BaseRTT to the latest sampled RTT Diff < : increase the congestion window linearly Diff > : decrease the congestion window linearly < Diff < : do nothing To be able to detect and avoid congesti on during slow-start, Ve gas allows exponential growth only every other RTT. In between, th e congestion window stays fixed so a valid comparison of the expected and actual rates can be made. When th e actual rate falls below the expected rate by the equivalent of one router buffer, Vegas changes from slowstart mode to linear increase/decrease mode. A couple of problems with TCP Vegas that could have a serious impact on its performance, are the issues of rerouting and stability. Rero uting a path may change the propagation delay of the connect ion; Vegas uses the connection to adjust the window size and it can affect the throughput considerably. Another issue of TCP Vegas is its stability. Since each TCP connection attempts to keep a few packets in the network when their estimation of the propagation delay is off, this could lead the connection to inadvertently keep many more packets in the netw ork causing a persistent congestion. Research on TCP Vegas to date consists primarily of analys es of the protocol, improving its congestion avoidance and detec tion techniques. [21] [22]. Most of the studies involving TCP Vegas consist of its performance evaluation with respect TCP Reno. [23][24]. Recent research at Caltech is exploring a new Vegas version, which


19 Caltech claims is a stabilized version of Ve gas [25]. This stabilized version of Vegas is completely source-based and requires no networ k support. They further suggest that this stabilized Vegas be deployed in an incremen ting fashion when a network contains a mix of links (some with active queue management and some without). Also, the performance of TCP Vegas is compared against that of TCP Reno on high performance computation grids [26] by Eric Weidge and Wu-chon Feng at Ohio State University. With the help of real traffic distributions We idge and Feng show that Vegas performs well over modern high performance links and better than TCP Reno, provided th at the TCP Vegas parameters and are properly selected. TCP SACK TCP throughput can be affected consider ably by multiple packets lost from a window of data. TCPs cumulative acknowledgeme nt scheme causes the sender to either wait for a round trip time to find out about a lo st packet, or to unne cessarily retransmit segments that have been corre ctly received. With this type of scheme, multiple dropped segments generally cause TCP to lose its ACK-based clock, which reduces the overall throughput. Selective Acknowledgement (SACK) [ 27] is a strategy th at rectifies this behavior. With selective acknowledgement, th e data receiver can inform the sender about all segments that have arrived successfully, so that the sender need retransmit only those segments that have actually been lost. This mechanism uses two TCP options: the first is an enabling option, SACK-permitted which can be sent in a SYN se gment to indicate that the SACK option can be used once the connection is establishe d; the second is the SACK option itself, which may be sent once permission has been given by SACK-


20 permitted. In other words, a selective ac knowledgement (SACK) mechanism combined with a selective repeat retransmission policy can help to overcome these limitations. The receiving TCP sends back SACK packets to the sender TCP indicati ng to the sender data that has been received. The sender can then retransmit only the missing segments [28]. The congestion control algorithms present in the standard TCP implementations must be preserved. In particular, to pr eserve robustness in the presen ce of packets reordered by the network, recovery is not triggered by a single ACK repor ting out-of-order packets at the receiver. Further, during the recovery, the data sender limits the number of segments sent in response to each ACK. Existing impl ementations limit the data sender to sending one segment during Reno-style fast recover y, or two segments during slow-start. Other aspects of congestion control, such as reducing the conge stion window in response to congestion, must similarly be preserved. Th e use of time-outs as a fallback mechanism for detecting dropped packets is unchange d by the SACK option. Because the data receiver is allowed to discard SACKed data when a retransmit timeout occurs the data sender must ignore prior SACK information, wh en determining which data to retransmit. Studies regarding TCP SACK include issues concerning aggressiveness of the protocol in the presence of congestion in comparison to other TCP implementations. Also, the issues concerning current TCP im plementation performance in a congested environment when competing against TCP implementations with SACK have been explored. [29]. TCP SACK has also been used to enhance performance of TCP in satellite environments. In [40], TCP with selective acknowledgement is examined and compared to traditional TCP implementations.


21 TCP Westwood TCP Westwood is a scheme [30] empl oyed by the TCP source to estimate the available bandwidth and use the bandwidth estimation to recover faster, thus achieving higher throughput. It is based on two concepts: the end-to-end estimati on of the available bandwidth and the way such an estimation is us ed to set the slow-sta rt threshold and the congestion window. Also, it is im portant to note that the feed back is merely end-to-end and does not depend on any intermediate nodes at the network level. The TCP Westwood (TCPW) source continuously estimates the packet rate of the connection by properly averaging the rate of returni ng ACKs. This estimate is used to compute the allowable congestion window and slow-start threshold to be used af ter a congestion episode is detected that is after three duplicate acknow ledgements or a timeout. Unlike TCP Reno, which simply halves the congestion window after three dupacks, TCPW attempts to make a more intelligent decision. It selects a slow-sta rt threshold and a congestion window that are consistent with the effective connection rate at the time of congestion. These types of techniques for bandwidth esti mation have been proposed before, (packet pair [31] and TCP Vegas [32]) but, due to tech nical reasons they have not been deployed onto the network. The key thing about TCPW is that it probes the network for the actual rate that a connection is ach ieving during the data transfer not the available bandwidth before the connection is starte d. TCPW offers a number of fe atures that are not available in TCP Reno or SACK. The knowledge of the av ailable bandwidth can be used to adjust the rate of a variable rate source. In the TCPW the se nder continuously computes the connection Bandwidth Estimate (BWE) that is defined as the share of bottleneck


22 bandwidth used by the connection. After a p acket loss indication, th e sender resets the congestion window and the slow-sta rt threshold based on BWE as cwnd = BWE x RTT. Another important element of this procedur e is the RTT estimati on. RTT is required to compute the window that supports the estimated rate BWE. Ideally, RTT should be measured when the bottleneck is empty. In prac tice, it is set equal to the overall minimum roundtrip delay (RTTmin) measured so far on that connection. In TCPW, congestion window increments during the slow start and congestion avoidance remain the same as in Reno that is they are exponen tial and linear, re spectively. In case of 3 dupacks, TCPW sets the congestion window and slow-start threshold as follows: ssthresh = (BWE*RTTmin)/MSS if(cwnd > ssthresh) /*congestion avoidance*/ cwnd = ssthresh; endif In the case of a packet loss being indicat ed by timeout expiration, cwnd and ssthresh are set as follows: cwnd = 1; ssthresh = (BWE*RTTmin)/MSS; if(ssthresh < 2) ssthresh = 2; endif; Recent research on TCPWs performance over large bandwidth pipes includes modifications of TCPW to TCP Westwood w ith Bulk Repeat (TCPW BR) [33]. TCPW BR has three sender-side modifications, namely Bulk Repeat, fixed Retransmission timeout, and intelligent window adjustment to help a sender recover from multiple losses in the same congestion window and to keep window size reasonably large when there is


23 no congestion along the path. TCPW BR also uses a loss differentiation algorithm (LDA), which is based on two schemes: spike and rate gap threshold, which is used to differentiate between losses due to congesti on and losses due to erro r. In losses due to congestion, congestion loss mode TCPW BR works in the same way as TCPW. In cases of error losses, error loss mode TCPW BR relies on the three sender-side modifications discussed above. This protocol has shown significant performance improvement in heavy loss environments. TCPW has also been modified for its performance over large bandwidth networks. Techniques like Adaptive restart (Astart) [ 34], paced-Westwood [35], TCP-Westwood with easy-RED [36], and TCP Westwood with rate estimates [37] explore the fairness issues, efficiency, frie ndliness issues, and performance issues of TCPW over high bandwidth-delay product networks that have small buffers. 2.2.2 TCP for High Bandwidth-Delay Product Networks As the next generation of applications will require network links with steady transfer rates in the order of gi gabits per second to transfer hu ge data in a reliable amount of time, the widely used TCP protocol will become the bottleneck. Since the windowbased mechanism of current TCP implementatio ns is not suitable for achieving high link utilization, many researchers have proposed modifications to TCP to improve the performance that it presents in very high ba ndwidth-delay product links. Here we discuss two protocols proposed to be efficient ove r high bandwidth links, namely High speed TCP and Scalable TCP.


24 Highspeed TCP High-speed TCP [3] is a m odification to TCPs congestion control mechanism to be used with TCP connections that have large congestion windows. The congestion control mechanism of the current standard TCP constrains the c ongestion windows that can be achieved in realistic environments. For example, for a standard TCP connection with a 1500 byte packets and a 100msec round trip time, achieving a steady state throughput of 10Gbps would require an average congestion window of 83,333 segments and a packet drop rate of, at most one congestion event every 5000,000,000 packets, which is a very unrealistic constraint. High-speed TCP is designed to have a different response in environments with a very low c ongestion event rate. It is also designed to have the standard TCP response in environments with packet loss rates of, at most, 10 -3 Since HSTCP leaves TCPs behavior unchanged in environments with mild to heavy congestion, it does not increase the risk of congestion collaps e. In environments with very low packet loss rates HSTCP presents a more aggressive response function. The high-speed TCP response function is specifi ed using three parameters: low_window, high_window, and high_p; low_window is used to establish a point of transition. The HSTCP response function uses the same res ponse function as regular TCP, when the current congestion window is at most lo w_window; it uses the high-speed TCP response function when the current c ongestion window is greater than low_window; high_window and high_p are used to specify the upper e nd of the high-speed TCP response function [36]. The high-speed TCP response function is represented by new additive increase and multiplicative decrease parameters. In congestion avoidance phase the behavior of the congestion window can be give n by the following equations:


25 ACK: cwnd = cwnd + a(cwnd)/cwnd Drop: cwnd = cwnd b(cwnd) x cwnd Research related to HSTCP includes wo rk done by Evandro De Souza [36], Deb Agarwal, and Sally Floyd [37]. In [36], HSTCP is tested for deployment issues. According to De Souza, HSTCP is appropriate to bulk transfer appl ication because it is able to maintain high throughput in different ne twork conditions, and because it is easy to deploy when compared with other solutions already in use. Anot her study conducted by Sally Floyd in collaboration with Evandro de Souza and Deb Agarwal, which is a part of HSTCP proposal, involves the limited slow -start mechanism for TCP with large windows. [37] Scalable TCP Scalable TCP [38] is a simple change to the traditional TC P congestion control algorithm; it claims to improve TCP perfor mance in high-speed wide area networks. Scalable TCP changes the algorithm to update TCPs congestion window as follows. For each acknowledgement received in a round trip time during which congestion has not been detected: cwnd = cwnd + 0.01 And on the first detection of conge stion in a given round trip time: cwnd = cwnd [0.125 cwnd] Figure 2 shows the main difference in the scaling properties of traditional and scalable TCP. Traditional TCP probing times are propor tional to the sending rate and the round


trip time. Scalable TCP probing times are proportional only to round trip time making the scheme scalable to high-speed networks. Scalable TCP has been designed from a strong theoretical base to ensure resource sharing and stability while maintaining agility in conjunction with prevailing network conditions. The response curves for both a traditional TCP connection and a scalable TCP connection is shown below. The scalable TCP algorithm is only used for windows above a certain size. By choosing the point at which the response curves intersect, good resource sharing with traditional connections can be ensured. 26 Figure 2. Scalable TCP Congestion Window Properties [4]. This allows scalable TCP to be deployed incrementally. Scalable TCP builds directly on the high-speed TCP proposal and works on engineering stable and scalable TCP variants. 2.3 Router Queuing Techniques 2.3.1 Droptail


27 The drop tail [42] scheme is the traditional and the simplest technique for managing router queue lengths. Droptail does not selectively drop packets; it drops them Figure 3. Response Curves of Scalable TCP and Regular TCP [4]. when there is no buffer space available. After setting a maximum length to the queue, Droptail accepts packets for the queue until the maximum length is reached, and then drops subsequent incoming packets until the queue decreases (as a packet from the queue has been transmitted). This technique is also known as "tail drop", since most recently arrived packet, which is at the tail of the queue, is dropped when the queue is full. Connections sending more traffic will get more system resources (although not necessarily better performance). This method has served the Internet well for years, but it has two important drawbacks. In some situations tail drop allows a single connection or a few flows to monopolize queue space, preventing other connections from finding room in the queue. This "lock-out" phenomenon is often the result of synchronization or other timing effects.


28 The tail drop discipline allows queues to main tain a full status for long periods of time, since tail drop signals congestion only when the queue has become full. It is important to reduce the steady-state queue size, and this is perhaps queue management's most important goal. 2.3.2 Random Early Detection (RED) There are two basic parts to RED [43]: detecting congestion, and responding to congestion. The algorithms for bot h of these tasks are simple, efficient in terms of both time and space, and easy to implement. To track the congestion level, an Exponentially Weighted Moving Average (EWMA) of the que ue length is kept. RED recalculates the average queue length avg each time a packet arrives in order to have an up to date estimate of the current conges tion level when determining what to do with the incoming packet. If the queue is not empty, then the avg is calculated using the following equation: avg = (1-w q )avg + w q q where q is the instantaneous queue length given by the number of packets currently enqueued, and w q is an operator-set parameter called the queue weight, which determines how quickly avg can change. If the queue is empt y, then the equation used to update avg depends upon the amount of time the queue was idle before the packet arrived, q time and the number of small packets that could have been transmitted by the gateway during that time. Thus, m = (time q time )/s avg = (1 w q ) m avg


29 where time is the current time, and m is simply a temporary value representing the number of packets that could have been transmitted during the idle time (time q time ), and s is the time needed to transfer a typical small packet. Once an estimation of the congestion level has been calculated, the gateway uses the value to determine what to do with the incoming packet. For this purpose RED queues are configured with two values, min th and max th which represent minimum and maximum thresholds for calculating a random drop probability. At avg values below min th the incoming packet will simply be en-queued, while at values above max th it will be marked. This marking can consist either of dropping the packet or performing some action such as setting a bit in the packets encapsulated transport header to indicate a congestion event to the flow. If avg falls between min th and max th however, the gateway randomly marks the packet with a probability p, which it generates internally. Further RED experiments [44] have led Floyd to recommend using what is known as the gentle RED variation, in which the only difference from standard RED is the probability that a packet will be dropped that varies from max p (maximum drop probability) to 1, when avg is between max th and 2*max th (a) (b) Figure 4. Drop Probabilities of Different RED Modes. (a) Normal Drop Probability (b) Gentle Drop Probability.


30 The goal of this algorithmic adjustment is to allow more leeway away from optimal values in the network op erators selection of max p and max th without severely influencing performance. The differences between normal and gentle modes are illustrated by Figures 4(a) and 4(b), respectively.


31 CHAPTER 3 METHODOLOGY This chapter explains in detail, the s imulation scenarios used in the performance evaluation of these eight protocols. This chap ter is divided into two sections. The first section describes the simulation setup and the parameters used for evaluating the performance of the TCP protocols on a sta nd-alone basis over a si ngle link. The second section describes the simulation setup for evaluating the fairness of these protocols to themselves as well as to ot her variants. The simulation t opology for evaluating these two types of fairness is the same. All the simu lations are conducted using both DropTail and RED ( Gentle ) router queuing techniques. The packet si ze in all the simulations is taken as 1000 bytes. Also, the application, which is used to send data, is FTP in all scenarios. All the simulations use the NS-2 [39] simulator. The parameters for RED queuing technique are set to default in the tcl script so that NS-2 configures them automatically. 3.1 Simulation Topology for a HBDP Link ( Single source/single sink ) This section includes the de scription of the simulati on topology and parameters used to evaluate the performance of TCP Tahoe, Reno, Newreno, SACK, Vegas, Westwood, HSTCP and Scalable TCP over high bandwidth-delay product networks. The network topology used consists of one TCP source, one TCP sink node (destination), and


two routers connected by a bottleneck link as shown in Figure 5. The simulations were carried out using network simulator NS-2 [39]. The maximum values of the congestion window for the TCP versions are set such that he connections could achieve full link utilization. The propagation delay of the bottleneck link is fixed and set to 25msec (two-way). Figure 5. Simulation Topology for One Source One Sink. The bandwidth of the link is varied from 1.5Mbps to 1000Mbps with values of 10, 100, 250, 550 and 800Mbps in between. In this way the bandwidth-delay product increases; this increase is used to show the performance of these protocols. The queue limit of the bottleneck link is set to 200 packets to absorb part of the sudden congestion. The experiments performed are used to show the bandwidth utilization, congestion window behavior, packet loss rate during the slow-start phase, and recovery time after a packet drop event. 32


3.2 Simulation Topology for Fairness to Itself and Others Scenario (Two Sources/Two Sinks) The simulation topology used is a dumbbell with a single bottleneck link as shown in Figure 6. The bottleneck link is connected on either sides by two TCP sources and two TCP sink nodes. The queue limit is kept at 200 packets for both scenarios; fairness to itself and fairness to other protocols. The link bandwidth is varied from 10Mbps to 1000Mbps with values of 100, 250, 550 and 800 Mbps in between. The link delay is 25msec both ways and is kept fixed. The maximum window size is kept large enough so as not to impose any limits. Figure 6. Simulation Topology for 2 Source 2 Sinks (Fairness to Itself) In the case of experiments regarding fairness to itself of a TCP protocol, the propagation delay of one TCP source is varied and its value is kept as the multiples of the propagation delay of the other TCP source. For example, if S1s delay is 10msec, then S2s delay is taken as 20msec through 60msec with multiples of 10 for the various 33


scenarios. All TCP protocols are tested for performance evaluation on this Figure 7. Simulation Topology for 2 Sources 2 Sinks (Fairness to Others) topology to observe how they perform when competing with themselves. The topology for the performance of a TCP protocol, when competing with another TCP protocol, is shown in Figure 7. The only difference in these simulations is that S1 and S2 are two different TCP sources and the propagation delays of both sources are the same (in our case 10msec). All eight protocols are tested against one another under this scenario and their performance is evaluated with respect to the throughput ratios they achieve. 34


35 CHAPTER 4 PERFORMANCE EVALUATION As mentioned in the previous chapte r, the simulations explained here are categorized into three sections. First, each protocol is evaluated separately on a single source single sink topology with a varying ba ndwidth. Second, each protocol is evaluated against every other protocol fo r fairness issues. And third, each protocol is evaluated for issues concerning fairness when sharing a bottleneck link with itself. Simulations in each section have been performed using Droptail as well as RED (gentle) router queuing techniques. 4.1 Single Source Single Sink Topology In this scenario, while evaluating the result s, it is important to have the knowledge that Droptail and RED queuing techniques will virtually have the same effect. RED in this case with a single source (connection) will mark all the packets for dropping when the average value of the queue goes above th e maximum threshold. Since it is a single source scenario, the whole queue (buffer) is for this connection, and hence RED will drop packets as the buffer is full, which is similar to the mechanism Droptail follows. The effective buffer in this case will be less than 200 packets, which is also used in Droptail. This buffer value is automatically configured in the NS-2s [39] RED mechanism. Hence


in the following section, we see that the results of droptail and RED queuing techniques are similar. Figure 8 shows the utilization achieved by the protocols under consideration as a function of the bandwidth of the bottleneck link. As the figure shows, there are considerable differences among these protocols. As a general trend, it can be seen that the performance of most protocols degrades as the bandwidth increases, revealing clear scalability problems. This is expected behavior, and it confirms what other researchers have found [41]. Only TCP Vegas, Scalable TCP, and HighSpeed TCP seem to perform well and scale better to higher speeds. The importance of this graph is the addition of other TCP versions as well as Vegas and Westwood, which had not been compared previously. The TCP versions also perform as expected with Tahoe presenting the worst performance followed by Reno, Newreno, and SACK, in that order. This sequence reflects the behavior of these protocols according to their reaction to packet losses and multiple packet losses from the same congestion window. Finally, TCP Westwood improves over the regular TCP versions but still below HighSpeed TCP, Scalable TCP and Vegas. (a) 36


(b) 37 Figure 8. Normalized Throughput of Protocols Under Consideration. (a) Using Droptail (b) Using RED. (a) (b) Figure 9. TCP Sequence Numbers for Link Bandwidth 1Gbps. (a) Using Droptail (b) Using RED


Figure 9 also shows the throughput achieved by all these protocols while they are using the TCP sequence numbers in the case of a 1 Gbps link. As it can be seen from the figures, the results relate very well to each other. The throughput performance of the protocols can be explained by looking at the behavior of the congestion window variable. Figure 10 plots the cwnd of the protocols over time, where the bottleneck bandwidth is set to 1 Gbps. With the exception of Vegas, the figure shows the expected sawtooth pattern of TCP. It can be seen that TCP Tahoe is the only protocol reducing its cwnd to 1, while the other TCP versions only reduce it to half or less than half of the current value. Reno presents deeper and longer reactions, while Newreno and SACK are very similar. Interesting behaviors are experienced by TCP Vegas, Westwood and HighSpeed TCP. TCP Westwood achieves better throughput because its cwnd does not drop as deep as the regular TCP versions, guided by the Fair Share Estimate (FSE). It will be seen later that TCP Westwood goes through a rather long Congestion Avoidance phase. The case is the same (a) (b) Figure 10. Congestion Windows of All the Protocols when Link Speed 1Gbps. (a) Using Droptail (b) Using RED (Continued.) 38


(a) (b) Figure 10. Congestion Windows of All the Protocols when Link Speed 1Gbps. (a) Using Droptail (b) Using RED 39


40 with the congestion window of Scalable TCP. Its window behavior is similar to that of TCP Westwood, but Scalable TCP achieves hi gher throughput because it does not take the initial computation tim e that TCP Westwood takes to compute the estimated bandwidth. The cwnd of HighSpeed TCP takes values similar to TCP Westwood, but it achieves better throughput since it manages to transmit more packets, particularly at the beginning of the connection. It also achiev es better throughput becau se it increases the congestion window faster. HighSpeed TCP pr esents the oscillatory behavior also experienced in the simulation results in [36] when only one source is in the system. TCP Vegass behavior is the best as the cwnd is rather steady after the Slow Start. Two conclusions are important at this point First, it is definitively impossible to achieve full bandwidth utiliz ation using the window-based approach utilized by current TCP implementations. The behavior of the cwnd shows that it takes TCP too much time to reach the maximum window size, and too little time to reduce its si ze in the presence of packet losses. Furtherm ore, the reduction of the cwnd is very drastic. The second conclusion is more important and has to do with Vegas behavior. If full utilization is to be achieved, the mechanism used by Vegas needs to be improved upon further. In the case of the bottleneck bandwidth being set to 1 Gbps, for example, it is known that the theoretical value of the congestion window re quired to achieve full link utilization is approximately 3325 packets, given by the ba ndwidth-delay product of the network and the buffer size. It can be observed from Figure 11 that this is the maximum value achieved by all protocols, and that TCP Ve gas congestion window is very steady and close to 3325 after the Slow Start phase, indicating that TCP Vegas is very good at estimating the available bandwidth. The main problem with Vegas lies in the first


Congestion Avoidance phase; it takes Vegas a rather long amount of time to initially reach the 3325 value. Next, the performance of the protocols were evaluated during the Slow Start phase. Here, the Slow Start time is a primary area of interest, as is the Packet Loss Rate (PLR) during that period of time. The Slow Start time is important because the longer it takes, the more wasted capacity results. The PLR is an indication of how efficient the Slow Start mechanism is. Obviously, the higher the PLR, the worse. The PLR was measured as the number of packets lost, divided by the total number of packets sent during the Slow Start phase. From Figure 11 it can be seen that all protocols have a similar and very short Slow Start duration. This is expected since they all employ the same exponential mechanism. Vegas has a slightly longer duration because it increases the cwnd exponentially every other RTT; however, its slow-start time of 0.4 seconds is still a very short time. Figure 12 shows the PLR achieved by the different protocols during the Slow Start phase. As can be seen, most TCP versions show a similar and steady PLR as the bandwidth is increased. This is expected because the buffer at the bottleneck link fills out at the same time regardless of the 41 (a)


42 (b) Figure 11. Slowstart Times. (a) Using Droptail (b) Using RED link capacity. This is in contradiction to other studies that say that one of the problems of current TCP versions is the very large value of cwnd that is achieved during Slow Start and resulting the high PLR. The explanation is in the buffer size of the bottleneck link. If the buffer size is set to the bandwidth-delay product of the link, the cwnd will in fact grow to very large values (in the order of 10000 in this studys 1 Gbps case). In reality, however, the system can only absorb around 6250 packets; but if the buffer size is set to more realistic values, as in this studys example, the cwnd will grow to modest values and the number of packets dropped will not be substantial. For instance, the PLR was in the order of 6% for this study. An interesting point to mention here is the fact that TCP Vegas and HighSpeed TCP were the only protocols with zero PLR. While TCP Vegas Slow Start phase takes a little longer than the other protocols, its Slow Start procedure is rather effective in avoiding packet losses during this time. This is in complete alignment with the design goals of Vegas as explained in [20]. The performance of the protocols over the


Congestion Avoidance phase is also investigated in the current study. Here, interestis mostly in the Congestion Avoidance phase time, called the recovery time, or the time that it normally (a) (b) Figure 12. Packet Loss Rate. (a) Using Droptail (b) Using RED takes the congestion window to reach its maximum value after a drastic reduction resulting from a packet drop. Figure 13 shows this time in seconds for the different 43


protocols as the bandwidth of the channel is increased. As was expected, the recovery time takes longer as the channel capacity increases. From the graph, it can be seen that the recovery time for regular TCP versions is around 70 seconds, or 2800 RTTs, while the recovery time of TCP Westwood and Vegas is around 10 seconds longer. Also, it can be observed that the recovery time of HighSpeed TCP is very small. This is due to the oscillatory behavior presented by this protocol as observed in Figure 10. For this experiment, only 44 (a) (b) Figure 13. Recovery Time (a) Using Droptail (b) Using RED


45 the first Congestion Avoidance phase was utilized. Another interesting point is related to TCP Westwood. It was found that the bandwidth calculation during th e initial phase is not very accurate, and therefor e, after the initial loss of packets, TCP Westwood sets the cwnd and ssthresh at very low values. Figure 17 shows the values of cwnd and ssthresh over time in the case where the bottleneck li nk is set to 1 Gbps. Th e figure clearly shows the bandwidth estimation problems of Westwood experiences during th e initial phase of the connection and how the Congestion Avoi dance phase starts with very small cwnd and ssthresh values. As a result, Westwood stays in that phase for a very long time, wasting a lot of bandwidth. In fact, the cwnd was set equal to 41 and grew to 3325 in a linear manner. A similar case was found in Vegas where the first Congestion Avoidance phase started at a cwnd of 72. At the beginning of the Slow Start phase, the expected bandwidth was a high value because the network is empt y. However, the actual bandwidth decreased substantially, since the e xponential increase of the cwnd quickly filled the buffers. At this point, Vegas lost some packets, reduced its cwnd, and then entered into Congestion Avoidance with a very low value of cwnd. Under realistic network conditions with normal buffer sizes this situation is rather un avoidable. Since the expected bandwidth is very close to the link speed, the cwnd starts increasing linearly until the actual bandwidth equals the expected bandwidth, and, at that time, the cwnd remains steady until the end of the simulation, achieving full utilization. The problem is that this initial Congestion Avoidance phase is very long and increases with the bandwidth of the bottleneck link. Modifying the slow-start phase procedures or the algorithms that drive the Congestion Avoidance phase can solve this problem.


46 4.2 Two Sources Two Sinks ( Fairness to Others ) For the current study, the four recent and ve ry important TCP modules are analyzed separately: Highspeed TCP, TCP Westwood, TC P Vegas, and Scalable TCP. Currently these TCP versions seem to have the poten tial to make their way through the high bandwidth delay product barrier. The follo wing sections analyze the behavior of Highspeed TCP, TCP Westwood, TCP Vega s and Scalable TCP when they are competing with one another and with regula r TCP protocols. Finally, the regular TCP protocols are discussed together. For the si mulations, a dumbbell topology is employed as shown in Figure 7. Two different TCP versions are attached to two sources to evaluate their performance over the bottle neck link. The analysis is d one by taking the throughput ratio of these protocols. All the graphs plot the throughput ratio against the bandwidth for these protocols. For example, if the tw o protocols under study are HSTCP and TCP Reno, then their throughputs are calculated ov er the varying bandwidth and their ratios are taken and represented as High/Reno. Highspeed TCP (HSTCP) In Figure 14, as was seen in the previous section, HSTCP is shown to be very aggressive grabbing a high percentage of ba ndwidth share when competing with other TCP protocols. HSTCP increases its congestion window quickly in the slow-start phase because it transmits more packets at the start of the connection. This behavior of HSTCP makes it grab more bandwidth; hence, it achieves on average, 25-50 times higher throughput than other protocols. The fairness tends to improve as the bandwidth continues to increase.


This is due to the fact that as the bandwidth of the link increases, the buffer size of 200 tends to become small for this bottleneck link; hence the congestion status changes to heavy from moderate. The packet loss ratio increases which causes the HSTCP to follow the highspeed response function with fairness being improved according to [41]. Figure 15 shows the throughput ratios of HSTCP when competing with TCP Westwood and TCP Vegas. These two protocols suffer drastically with respect to HSTCP, as a result of their congestion control mechanisms and bandwidth estimation process, which will be considered in more detail in the following section. (a) 47


(b) Figure 14. Performance of HSTCP Against TCP Protocols (a) Using Droptail (b) RED These protocols always tend to be dominated by other protocols and never perform well in competent environments although as seen in the previous section, these protocols achieve a very high throughput when run on their own. (a) 48


49 (b) Figure 15. HSTCP Fairness to Vegas and Westwood (a) Using Droptail (b) Using RED The effect of the router queuing techniques was minimal and existed only when the link capacity was small. The relative fairness between HSTCP and other protocols is less in the case of Droptail, than in the case of RED. This is because with RED the fraction of packet drops received by each flow should be roughly proportional to that flows share of link bandwidth, while this property is no longer true in case of Droptail queue management. As the link capacity continues to increase, the router queue management technique does not have effect on the results. In todays Internet scenario, there are not a lot of TCP connections that are operating effectively with congestion windows containing thousands of packets. Therefore, the benefits of the HSTCP would outweigh the unfairness that would be experienced by regular TCP protocols.


TCP Westwood (TCPW) The TCPW uses a bandwidth estimation mechanism with which it manipulates its ssthresh and congestion window after every RTT. Unlike other TCP protocols, TCPW does not reduce the ssthresh to half during a congestion event but converges it to a steady value as shown in Figure 17. This takes a fair amount of time and during this time, another TCP version with which TCPW is competing (as it was in this study) grabs more than its fair share of bandwidth. As Figure 16 shows, all the protocols dominate over TCP Westwood; and, as the bandwidth increases, this dominance tends to increase. The router queuing techniques do not have a significant effect on the issue of fairness. 50 (a)


(b) Figure 16. Performance of TCP Westwood Figure 17. Behavior of TCP Westwoods Cwnd and Ssthresh when the Bottleneck Link is Set to 1 Gbps. TCP Vegas According to Figure 11, TCP Vegas has a slightly longer slowstart time than do other protocols. The reason for this is that Vegas increases its congestion window exponentially every other RTT. The performance of Vegas can be seen in Figure 18. The 51


plots suggest that TCP Vegas does suffer when it shares a bottleneck link with other TCP protocols, due to its bandwidth estimation technique, which prolongs the slowstart phase. This is because the other protocols use most of the buffer space, causing the TCP Vegas connection to back off, since it interprets this as a signal of network congestion. 52 (a) (b) Figure 18. Performance of Vegas (a) Using Droptail (b) Using RED.


53 In Figure 18, the prominent difference be tween the two graphs during the small bandwidth phase can be seen. The reason for this is based on the assumption that the buffer occupancy decreases as the bandwidth increases. This study analyzed the Droptail and RED cases. When Droptail is used, Vegas due to its less aggressive mechanism of increasing its congestion window, can not occu py much buffer space; thus, the other protocols sharing the bottleneck link with Vegas will occupy more buffer space and a larger share of the bottleneck link. As the bandwidth increases, however, the buffer tends to become less occupied; hence, Vegas receiv es a much better share of the buffer. In addition to the buffer share, the bandwidth estimation technique of Vegas outweighs the effect of regular AIMD techniques used in traditional TCP. In the case of RED, Vegas receives a better share of the bottleneck link wh en compared to Droptail. This is because of the mechanism of RED. RED gives fair share of the buffer to each flow. Hence, Vegas, during the small bandwidth links receives its share of buffer space; and more packets from the other TCP protocols that ar e sharing the bottleneck link with Vegas are dropped in this stage, increasi ng fairness. As the bandwidth increases, the buffer tends to become empty and Vegas always receives its sh are of the buffer (if required), and so does not end up suffering. Scalable TCP Scalable TCPs performance when compe ting with the rest of the protocols is shown in the Figure 19. The gra phs clearly indicate that ex cept for the Scalable/Vegas combination, all other protocols receive a fair share of bandwidth when run with Scalable TCP, irrespective of the router queuing technique used; th e Scalable/Westwood


combination is not represented in the graph because of its high magnitude. Scalable TCP uses up most of the bandwidth when sharing a bottleneck link with Westwood, due to TCP Westwoods congestion window mechanism( see Figure 17). The difference in the (a) and (b) graphs of Figure 19 is the Scalable/Vegas combination. Again, here, the effect of router queuing management is seen only on the low-bandwidth links, and RED increases the fairness towards Vegas when the channel capacity is low. The reason for this was explained in the previous section. The Scalable TCP, being more aggressive, occupies much of the bottleneck link in the smaller bandwidth links (40 times more than Vegas). As the link speed increases, as seen in the previous section, Vegas receives a fair share of the bottleneck link. Here, in this case, because of the aggressive nature of Scalable TCP, it dominates over Vegas fractionally more than the other TCP protocols. 54 (a)


55 (b) Figure 19. Performance of Scalable TCP with Other Protocols (a) Using Droptail (b) Using RED. Other TCP protocols [TCP Tahoe, Reno, Newreno and SACK] These protocols do not show any drastic differences in fairness when competing with one another. Their behavior is shown in Figure 20: all protocols act fair to one another without any noticeable aggressiveness. The reason for this being all these protocols use the same traditional window-based algorithm. (a)


56 (b) (c)


57 (d) Figure 20. Performances of Tahoe, Newreno, Reno and SACK TCP. 4.3 Two Sources Two Sinks (Fairness to Itself) This section explains the results obtained from the simulations conducted with the topology shown in Figure 6. These simulations concern the issues regarding fairness shown when each of these protocols shares a bottleneck link with itself. The two sources S1 and S2 in Figure 6 are two TCP sources of the same version. While the propagation delay of one source is kept constant at 10ms, the propagation delay of the second source varies in multiples of 10 60ms. The graphs show the RTT scale plotted on the x-axis and Throughput plotted on the y-axis. The throughputs of the S1 and S2 flows are plotted on the graphs for the link capacities of 10,100,550 and 1000Mbps. TCP Tahoe, Reno, Newreno, SACK, and Scalable TCP are analyzed


58 together because of the similar results they present. Analysis of HSTCP, TCP Westwood and TCP Vegas is done separately. Analysis of Tahoe, Reno, Newr eno, SACK and Scalable TCP These protocols present some common behavior when they share a common bottleneck link, where one source has a larger propagation delay than the other. The first important and basic thing that can be observed from the graphs in Figure 21 is that, as the propagation delay of one source in creases, that source tends to occupy less bandwidth. The reason for this is that, as the propagation delay of a TCP sender increases, it takes longer for it to rece ive its acknowledgment and increase its congestion window by sending more packets in to the bottleneck link. Hence, it fills up a lesser percentage of the shared link th an do TCP senders with lesser propagation delay. As the bandwidth of the bottleneck li nk increases, the Droptail and RED queuing techniques produce similar results because, at the low bandwidth links, the buffer is utilized the most and, as the bandwidth incr eases, the buffer tends to be less occupied. Hence, during the Droptail scenario for all the protocols, when the bandwidth of the bottleneck link is 10Mbps, a prominent fairne ss issue exists, which is absent in the RED technique. This is because in the cas e of Droptail, the TCP source with the shorter propagation delay fills up the buffer faster than the connection with the longer propagation delay, which means packets are dropped of the latter connection and less of the link is utilized by this source. Another observation is that, as the bandwidth of the link goes on increasing, the total link u tilization of the bottl eneck link decreases.








Figure 21. Fairness to Itself (Newreno, Reno, SACK, Scalable, Tahoe) Analysis of TCP Vegas TCP Vegas shows a comparatively higher total link utilization than the previous protocols, but the fairness factor suffers. The TCP Vegas that has a longer propagation delay is severely affected in this scenario as is shown by the growing gap in Figure 22. This wider gap can be explained via the TCP Vegas congestion window mechanism, which increases its congestion window every other RTT. As the propagation delay of one Vegas source increases this source will take effectively double the time to increase its congestion window than other TCP protocols, which increases congestion window every RTT would take. Hence the broader gap. 62


63 CHAPTER 5 Figure 22. Fairness to Itself (Vegas)


64 CHAPTER 5 CONCLUSIONS This chapter is divided based on the experiments that were described in the previous chapter. The results and the observations of each experiment are summarized and the protocol is pointed, with advice as to which is best suited for the Internet, which is a large-scale combination of the experiments performed. From the experiments conducted over a single source single sink topology, TCP Vegas Scalable TCP, HSTCP and TCP We stwood performed fairly well over the TCP Reno, Newreno, Tahoe and SACK, as the band width of the bottleneck link increased. This experiment offers insight into the perf ormance of these protocols on an individual basis, although the environment in which these protocols were simulated was only remotely similar to the real world environm ent. A conclusion, which can be made from this set of experiments, is that, as the bandwidth of the link increases, the throughput decreases irrespective of the ro uter queuing technique used. From the experiments conducted for fai rness to others it can be summarized that the traditional TCP protocols Reno, Tahoe, SACK, Newreno are the fairest to one another. HSTCP is the most a ggressive protocol and it is not recommended that it be incorporated into the Internet. TCP Vegas, al though it is the best when run alone, suffers when it shares a bottleneck link with others, as does Westwood. Scalable TCP, on the


65 other hand, is a protocol that when sharing a bottleneck link with other protocols is comparatively less aggressive than HSTCP, a nd slightly more aggressive than regular TCP protocols. Additionally, when these protocols sh are a bottleneck link with themselves, Scalable TCP is not very aggressive, a nd doesnt allow one source to grab more bandwidth, like Vegas and HSTCP do. Scal able TCP acts similar to traditional TCP protocols in this scenario. C onsidering the observations of the three scenarios mentioned above, we come to a conclusion that Scalab le TCP, among the protocols under study, is the one with a better performance, with regards to throughput and overall fairness.


66 REFERENCES [1] First International Workshop on Prot ocols for Fast Long-Distance Networks, PFLDnet 2003. February 2003. [2] S. Low, F. Paganini, J. Wang, S. Adla kha, and J. Doyle. Dynamics of TCP/AQM and a Scalable control. In Proceedings of IEEE INFOCOM pages 239-248, June 2002. [3] S. Floyd, Highspeed TCP for Large Congestion Windows, IETF Internet draft floyd/papers/rfc3649.txt August 2002. [4] T. Kelly, Scalable TCP: Improving Performance in Highspeed Wide Area Networks. In Proceedings of First International Workshop on Protocols for Fast Long-Distance Networks February 2003. [5] C. Jin, D. Wei and S. Low. FAST TC P for High-speed long distance networks. IETF Internet draft (draft-jwl-tc p-fast-01.txt ) June 2003. [6] J. Mo, R. La, V. Anantharam, J. Warland Analysis & Comparison of TCP Reno and Vegas. In Proceedings of IEEE INFOCOM pages 1556-1563, March 1999. [7] M. Mathis, and J. Madhavi. Fo rward Acknowledgement: Redefining TCP Congestion Control. In Proceedings of SIGCOMM pages 281-291, August 1996. [8] K. Fall, and S. Floyd. Simulation base d Comparisons of Tahoe, Reno and SACK TCP. ACM Computer Communications Review Vol. 26 No. 3, pages 5-21, July 1996. [9] S. Mascolo, C. Casetti, M. Gerla, S. Lee, and M. Sanadidi. TCP Westwood: Congestion Control with Faster Recovery, UCLA CSD Technical Report #200017, 2000. [10] M. Allman, V. Paxson, and W. St evens. TCP Congestion Control, RFC 2581. April 1999. [11] S. Floyd and T. Henderson. The NewRe no Modification to TCPs Fast Recovery Algorithm, RFC 2582. April 1999.


67 [12] S. Floyd and V. Jacobson. Random Ea rly Detection Gateways for Congestion Avoidance. IEEE/ACM Transactions on Networking Vol.1 No.4, pages 397, August 1993. [13] V. Jacobson. Modified TCP Congestion Avoidance Algorithm. Technical Report April 1990. [14] M. Mathis, S. Floyd, J. Madhavi, and M. Podolsky. An Extension to the Selective Acknowledgement (SACK) Option for TCP, RFC 2883. July 2000. [15] V. Jacobson. Congestion Avoidance and Control. In Proceedings of SIGCOMM, pages 314-329, August 1988. [16] T. Lakshman, and U. Madhow. The Performance of TCP/IP for Networks with High Bandwidth-Delay Pr oducts and Random Loss. IEEE/ACM Transcations on Networking Vol.5 No. 3, pages 336-350, June 1997. [17] W. Stevens. TCP/IP Illustrated Volume 1 Professional Computing Series. Addison Wesley, 1st Ed., 1994. [18] S. Floyd and T. Henderson. The NewRe no Modification to TCP's Fast Recovery Algorithm, RFC 2582. April 1999. [19] M. Allman, V. Paxson, and W. St evens. TCP Congestion Control, RFC 2581, April 1999. [20] L. Brakmo and L. Peterson. TCP Vegas: End to End Congestion Avoidance on a Global Internet IEEE Journal on Selected Areas in Communication Vol. 13 No. 8, pages 1465-1480, October 1995. [21] L. Brakmo, S. O'Malley and L. Peterson. TCP Vegas: New Techniques for Congestion Detection and Avoidance. In Proceedings of SIGCOMM pages 2435, August 1994. [22] J. Ahn, P. Danzig, Z. Liu, and L. Ya n. Evaluation of TCP Vegas: Emulation and Experiment. ACM SIGCOMM Computer Communication Review Vol. 25, No 4, pages 185-195, October 1995. [23] T. Bonald. Comparison of TCP Reno and TCP Vegas via fluid approximation. Technical Report RR-3563 INRIA, November 1998. [24] J. Mo, R. La, V. Anantharam, and J. Warland. Analysis a nd Comparison of TCP Reno and Vegas. In Proceedings of IEEE INFOCOM pages 1556-1563, March 1999.


68 [25] H. Choe and S. Low. Stablized Vegas, In Proceedings of IEEE INFOCOM pages 2290 2300, April 2003. [26] E. Weigle and W. Feng. A Case for TCP Vegas in High-Performance Computational Grids In Proceedings of the 10th IEEE International Symposium on High Performance Distributed Computing page 158, August 2001. [27] M. Mathis, J. Mahdavi, S. Floyd and A. Romanow. TCP Selective Acknowledgment Options (SACK) RFC 2018, October 1996. [28] L. Brakmo, S. O'Malley and L. Peterson. TCP Vegas: New Techniques for Congestion Detection and Avoidance. In Proceedings of SIGCOMM pages 2435, August 1994. [29] S. Floyd. Issues of TCP with SACK Technical report January 1996. [30] M. Gerla, M. Sanadidi, R. Wang, A. Zanella, C. Casetti and S. Mascolo. TCP Westwood: Congestion Wi ndow Control Using Bandwidth Estimation. In Proceedings of IEEE GLOBECOM Vol. 3, pp 1698-1702, November 2001. [31] M. Allman and V. Paxson. On Estima ting End-to-End Network Path Properties. In Proceedings of SIGCOMM Vol. 29, No. 4, October 1999. [32] K. Lai and M. Baker. Measuring Li nk Bandwidths Using a Deterministic Model of Packet Delay. In Proceedings of ACM SIGCOMM August 2000. [33] G. Yang, R. Wang, F. Wang, M. Sanadi di, and M. Gerla. TCP Westwood with Bulk Repeat for Heavy Loss environments, UCLA CSD Technical Report #020023, 2002. [34] R. Wang, G. Pau, K. Yamada, M. Sanadidi, M. Gerla. TCP Startup Performance in Large Bandwidth Delay Networks. To appear in INFOCOM 2004 March 2004. [35] A. Razdan, A. Nandan, R. Wang, M. Sanadidi, and M. Gerla. Enhancing TCP Performance in Networks with Small Buffers. In Proceedings of 11 th International Conference on Com puter Communications and Networks October 2002. [36] E. Souza and D. Agarwal. A HighSpeed TCP Study: Characteristics and Deployment Issues. LBNL Technical Report Number LBNL-53215. Available at:


69 [37] S. Floyd. Limited Slow-Start for TCP with Large Congestion Windows, RFC 3742, March 2004. [38] T. Kelly. Scalable TCP: Improvi ng Performance in Highspeed Wide Area Networks. In Proceedings of First International Workshop on Protocols for Fast Long-Distance Networks February 2003. [39] NS-2 Network Simulator. In 2000. [40] M. Allman, C. Hayes, H. Kruse, and S. Ostermann. TCP Performance over Satellite Links In Proceedings of the 5th In ternational Conference on Telecommunication Systems March 1997. [41] S. Floyd. HighSpeed TCP for Large Congestion Windows, RFC 3649, December 2003. [42] B. Braden. et. al Recommendations on Queue Management and Congestion Avoidance in the Internet, RFC 2309, April 1998. [43] S. Floyd and V. Jacobson. Random early detection gatewa ys for congestion avoidance. IEEE/ACM Transactions on Networking Vol. 1 No.4, pages 397-413, August 1993. [44] S. Floyd. Description of gen tle Mode in NS. web page, notes/test-suite-red.txt


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