ATM Switch. Traffic Shaper. ATM Switch. Traffic Shaper. ATM Switch. Terminal Equipment. Terminal Equipment UNI UNI

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1 Mult-rate Trac Shapng and End-to-End Performance Guarantees n ATM Debanjan Saha Department of Computer Scence Unversty of Maryland College Park, MD E-mal: debanjan@cs.umd.edu Sart Mukherjee y Dept. of Computer Scence & Engg. Unversty of Nebraska Lncoln, NE E-mal: sart@cse.unl.edu Satsh K. Trpath Department of Computer Scence Unversty of Maryland College Park, MD E-mal: trpath@cs.umd.edu Abstract Ths paper proposes a trac control scheme for ntegrated servces ATM networks. The control strategy comprses of two components: a shapng mechansm at the network entry pont and a frame based servce dscplne at the swtches. The shaper enforces a short trem peak rate, and a long term average rate. The multplexng scheme at a swtch allocates a guaranteed bandwdth to a connecton. A connecton may get more than the guaranteed amount, up to a connecton specc maxmum, f slack bandwdth s avalable. By mposng an upper bound on the allocated bandwdth, we secure a better handle on the delay jtter. Unlke most frame-based schemes, our scheme allows allocaton of bandwdth at any arbtrary granularty. We suggest a smple admsson control polcy and derve determnstc bounds on end-to-end delay and jtter. An outlne of a hardware realzaton of the scheme s also presented. Supported n part by ARPA and Phlps Labs under contract DASG and NSF under contract CCR y Supported n part by the Oce of Research Councl, UNL, under the grant LWT/

2 1 Introducton An ntegrated servces ATM network s expected to oer servces to a wde range of real-tme applcatons. The ntegraton of real-tme trac n ATM networks places strngent requrements on the worst-case end-to-end delay and jtter 1 bounds assocated wth such trac. At the same tme, to accommodate dverse trac from these applcatons, network servces need to be very exble. In ths paper we develop a framework for trac control that can provde determnstc guarantees on end-to-end performance, and can support dverse throughput requrements of a wde mx of applcatons. The trac control schemes [10, 3, 4, 7, 5, 9, 6, 2] proposed for ntegrated servces packet swtched networks comprse of two components: trac shapng at the network entry pont and a packet servce dscplne at the swtches. Besdes smoothenng, the shaper helps characterze the trac enterng the network. A succnct characterzaton s extremely mportant n obtanng tght performance bounds. However, trac shapng does not necessarly guarantee these bounds snce the characterstcs of the shaped trac may get moded at the ntermedate swtches. In order to provde predctable end-to-end performance, we need to characterze the trac enterng every swtch on the path of a connecton. Hence, t s mportant to use a sutable servce dscplne at swtches so that trac from a connecton extng a swtch and enterng the next swtch on ts path can be easly and precsely characterzed. In most of the proposed schemes, the shaper enforces a specc rate of packet arrval, typcally a declared peak or average rate. However, most applcatons generate nherently bursty trac. Hence, enforcng an average rate results n hgher delay, and on the other hand, a peak rate allocaton leads to lower network utlzaton. To strke a balance between network delay/jtter and bandwdth utlzaton, t s desrable that the network enforces multple rates a peak rate to allow bursty trac over short ntervals, and an average rate (much lower than the peak rate) at all other tme. Packet servce dscplne at swtchng nodes can be broadly classed as ether work-conservng or non-work-conservng. In a work-conservng servce dscplne an output lnk s never dle so long as there are packets to send. A non-work-conservng dscplne can keep a lnk dle even f there are packets watng. Although non-work-conservng polces oer hgher average delay than workconservng polces, typcally the former enforces lower delay jtter (hence smaller buer space). For guaranteed worst case performance, a non-work-conservng servce dscplne at swtches s a preferred choce. In our scheme, the shaper enforces a maxmum allowed \average" rate over a long nterval whle gvng the connectons an opton to volate ths constrant by choosng a \peak" rate over shorter 1 The maxmum derence over all packets between departure and arrval tmes. 1

3 Termnal Equpment Traffc Shaper UNI ATM Swtch ATM Swtch ATM Swtch UNI Traffc Shaper Termnal Equpment Fgure 1: ATM network archtecture. ntervals. We use a frame based non-work-conservng packet schedulng scheme at the swtches. The scheme guarantees all connectons a mnmum bandwdth (atleast the average rate) negotated at the connecton set up tme. A connecton may use more than the allocated mnmum, up to a connecton specc maxmum, f slack bandwdth s avalable. By mposng an upper bound on the allocated bandwdth we secure a better handle on delay jtter. Our scheme also overcomes a common shortcomng of most frame based servce dscplnes,.e., allocaton of bandwdth only n multples of a xed quantum. In a hgh-speed network supportng a wde mx of applcatons, ths nexblty can lead to severe under-utlzaton of resources. We solve ths problem by multplexng several connectons on a sngle slot of a frame (to be dscussed n detal n secton 2). The rest of the paper s organzed as follows. Secton 2 descrbes trac shapng and schedulng. Secton 3 s devoted to the propertes of the control scheme and secton 4 to the mpact of the control parameters. In secton 5 we present a hardware realzaton of the scheme. We brey dscuss related work n secton 6 and conclude n secton 7. 2 Trac Shapng and Schedulng We consder a network archtecture as shown n gure 1. Trac from all sources are regulated at the user network nterface (UNI), whch s the network entry/access pont. Insde the network, trac generated by a source/stream s constraned to follow the same path called a connecton. From now onwards, we wll use C nterchangeably for a source, a stream as well as a connecton. We dvde the tme axs nto successve perods of equal length, called a frame. Every nternal or access lnk of the network can use an arbtrary tme orgn to dene frames correspondng to that lnk. Every lnk can also choose an arbtrary frame sze. However, to keep the dscusson smple we wll assume a global tme orgn and a unform frame sze F (n tme unts) for all lnks. Snce the lnks may be of derent speeds, the number of slots 2 a frame can accommodate could der 2 Slots are of xed sze, expressed n number of bts. A slot can hold a packet/cell. From now onwards, we wll use 2

4 on derent lnks. Over each lnk we vew the correspondng frames of the lnk as travelng wth xed sze slots from one end of the lnk to the other end (refer to gure 2). We denote by P l the number of slots per frame on lnk l. We assume that propagaton and transmsson delays are xed and known. In the rest of the dscusson we focus on queung delays only. Departng Frames 0 F 2F 3F Propagaton Delay Arrvng Frames Fgure 2: Departng and arrvng frames on a lnk. 2.1 Mult-rate Shapng A trac shaper (gure 1) shapes the trac from a source enterng the network. The advantage of usng a shaper s that t allows decomposton of packet delay nto two components delay n the shaper and delay n the network. The rst of these components can be estmated from the statstcal characterzaton of the source. The shaper can shape the trac n a manner that can be succnctly characterzed. Ths makes t easer for the network to bound the second component. From ths pont onwards, we wll restrct our attenton to the network component of the delay. Trac s shaped accordng to a smoothness crteron that lmts the number of packet entres nto the network over an nterval (consstng of an ntegral number of frames). Ths essentally enforces a certan packet arrval rate over that nterval. A mult-rate shaper allows two rates of packet arrval. A peak rate s enforced over short ntervals. However, over a long nterval, number of packet entry s regulated n such a fashon so that an average rate (much lower than the peak rate) s obtaned. To formalze the trac shapng we dene below the \smoothness" of a trac stream. Denton 2.1 A trac stream C s sad to be (W; M) smooth f, slot, packet and cell nterchangeably. 3

5 '.. INITIALIZE(C ) /* Intalzaton for connecton C */ n 0 /* Current queue length of C */ w p 0; wa 0 /* Current peak and average wndows */ m p M p ; ma M a /* Current peak and average number of releasable packets */ $ TIMER(C ) /* Operates on all actve connectons. */ At the begnnng of a frame: f (w a (wa a + 1) mod W M a end f f (w p (wp p + 1) mod W then mp M p end f & SHAPER() 1. On arrval of a packet from C : n n + 1; hold the packet n the buffer 2. Release packets from C : whle m a > 0 ^ mp > 0 ^ n > 0 do release the next packet n the network m a ma? 1; mp mp? 1; n n? 1 end whle % Table 1: Functons performed at the shaper (j+1)w?1 X k=jw m k M; where j = 1; 2; : : :, and m k s the number of packets from stream C that travel n the k th frame. Informally, a trac stream s (W; M) smooth f t carres no more than M packets n a tme wndow of W frames. We assume that the shaper smooths the trac from the stream C nto a (W a; M a) and (W p ; M p a ) smooth ow, where W W p and M a=w a M p =W p. Essentally, the shaper enforces a maxmum average rate of M a cells over an nterval (or wndow) of W a frames and a maxmum peak rate of M p cells over an nterval (or wndow) of W p frames. The shaper functons n the followng way. It has a buer space to hold the ncomng packets. We assume t to be of nnte sze. When a packet arrves from connecton C, t s stored n the 4

6 buer. If releasng the packet nto the network does not volate the mposed rate constrants (both peak and average), the packet s admtted. In table 1 we brey outlne a functonal descrpton of the shaper. It uses several counters. The counter n keeps track of the number of buered packets belongng to connecton C. The counters m p and ma represent the number of packets that can be admtted nto the network n the current peak and average wndows, respectvely. If both of them are greater than zero, a packet s released. The counters w p and w a keep runnng counts of the number of frames expred n the current peak and average wndows, respectvely. The INITIALIZE() functon s nvoked when a new stream s admtted. It ntalzes the counters for the connecton. The TIMER() functon s nvoked at the begnnng of every frame. It ncrements w p and w a at every nvocaton, and checks f the peak and/or average wndow(s) expres n ths frame. It sets m p and m a to M p and M a at the begnnng of a new peak and average wndow, respectvely. The SHAPER() functon determnes f releasng ncomng packets n the network volates the smoothness crtera of the correspondng connecton. admtted to the network. Otherwse, they are held n the shaper buer. If they are not volated, packets are In order to compute end-to-end performance metrcs, shapng trac at the network entry pont s not sucent. The same trac characterstcs may not hold at ntermedate swtches unless a sutable multplexng scheme s used there. In the next secton we present a servce dscplne that mantans the orgnal shape of the trac throughout the network. 2.2 Servce Dscplne From the modelng perspectve, a swtch can be vewed as a queue wth rate of servce equal to the output lnk speed. If a number of connectons are multplexed over the same output lnk, the eectve rate of servce receved by ndvdual connectons depends on the servce dscplne used. We use a frame based non-work-conservng servce dscplne to be dscussed shortly. The justcaton behnd usng such a scheme s to mantan trac smoothness. In our scheme, each connecton C s guaranteed a mnmum rate of R (hgher than the average rate, and preferably less than the peak rate) slots per frame, negotated at the tme of connecton set up. In a frame based allocaton scheme, bandwdth s always allocated [6, 5] n multples of a xed quanta 3. However, ths allocaton strategy tends to under-utlze the lnk capacty. To allevate the problem, n our scheme, we allow R to be non-ntegral. The ntegral part of R can be easly allocated by schedulng br c slots per frame. However, fracton of a slot cannot be allocated to a connecton. In order to meet the fractonal requrement, the easest approach s to allocate one slot to that connecton. Ths s equvalent to allocaton of dr e slots per frame. However, ths 3 Ths quanta s usually one slot per frame. 5

7 scheme under-utlzes the lnk capacty, and to be precse, reduces to the aforementoned allocaton by quanta scheme. We adopt the followng approach to acheve the fractonal allocaton wthout reducng the lnk utlzaton: We call a frame busy f all of ts slots are used to fulll the guaranteed rate of all the connectons beng served by t. If the frame has any slack slot after t has met the total requrement, t s called non-busy. A (non-)busy perod refers to the nterval (consstng of frames) where all the consecutve frames are (non-)busy. The allocaton algorthm s derent for a busy and a non-busy perod. In a non-busy perod, each connecton C s allocated at least dr e slots n a frame. If there are stll free slots left, they are dstrbuted randomly makng sure that C gets no more than M p slots n a wndow of W p frames 4. In a busy perod, a more complex schedulng s requred. Unlke n a non-busy perod, n a busy perod there are not enough slots to over-allocate all connectons wth the next hgher ntegral number of slots. In order to be as far as possble, the scheduler rst satses the ntegral porton of the slot requrements of all connectons. Of the remanng slots, some connectons wth fractonal requrements are allocated one slot each, whle the fractonal requrements of others are left unsatsed. Slot requrements for the next frame are sutably adjusted to compensate for ths unfar allocaton. The ratonale behnd ths allocaton scheme s the followng by sutable admsson control, n every frame the ntegral slot requrements of all connectons are always satsed. Of the connectons wth fractonal requrements, some may get an extra slot whle others may not. However, f a connecton s deprved of ts fractonal requrements n consecutve frames, they accumulate and exceed 1. At that stage a slot must be allocated to t. Once agan, the scheduler lmts the maxmum number of slots used by C to M p Table 2 gves a functonal descrpton of the scheduler. n a tme wndow of W p frames. In the descrpton of the scheduler, r denotes remanng slot requrement of C n the current frame, w p s the number of frames expred n the current peak wndow and m p counts the number of slots the connecton can get n the remander of the peak wndow (.e. next W p? w p frames, ncludng the current one). At the tme of connecton set up the INITIALIZE() functon sets r to R. It also sets m p to M p and w p to 0, followng the same reason of the shaper. The TIMER() functon resets m p back to M p every W p frames. The SCHEDULE() functon allocates the departng slots of lnk l. It rst satses ntegral requrements of all the connectons. Dependng on the number of slots remanng, t arbtrarly selects some connectons wth fractonal requrement and allocates one slot to each of them. If there are stll slots left 5, connectons may be gven addtonal bandwdth to clear backlogs of packets, f any. Snce the requrement of a connecton gets fully satsed n a non-busy perod, r for the next frame s re-ntalzed to the orgnal requrement R. 4 Ths s done to ensure trac smoothness. 5 Ths s a non-busy frame. 6

8 Otherwse, the leftover requrement s carred over to the next frame along wth R. In all frame based allocaton schemes known to us, bandwdth s allocated n ncremental steps of one slot. Ths means that f we choose a frame sze of 1ms, granularty of bandwdth allocaton s approxmately 425 Kb/s, assumng a packet sze of 53-bytes. Ths coarse granularty of bandwdth allocaton can lead to severe under-utlzaton of network resources. Our scheme overcomes ths problem by allocatng \non-ntegral" number of slots to a connecton and there by allowng allocaton of bandwdth at any arbtrary granularty. 3 Computng the QoS Envelope In ths secton we prove some mportant propertes of the trac control scheme and derve worst case end-to-end bounds on delay and jtter. We rst prove that under certan admsson control crtera, the multplexng at swtches does not destroy the orgnal smoothness of the trac. That s to say that the output stream from a swtch conforms to the same shapng envelope as the nput stream to the swtch. Consequently, end-to-end delay (jtter) n a mult-hop connecton reduces to the product of the number of hops and the delay (jtter) suered at a swtch n solaton. The rst step towards provng the smoothness property s to determne the admsson control crtera under whch a connecton C s guaranteed a bandwdth of R slots per frame. In the followng lemma, and n the subsequent theorem we nvestgate the behavor of the algorthm n a busy perod. Bandwdth guarantee s trvally satsed n a non-busy perod. Lemma 3.1 Let N l be the number of connectons beng multplexed on lnk l, r k be the slot requrement of connecton C at the begnnng of the k th frame and P N l =1 R P l. Then P N l =1 rk P l s satsed for all frames n a busy perod. Proof: Let us consder a busy perod consstng of frames K; K + 1; : : : ; K + m. We need to show P N 8k 2 [K; K + m]; l =1 rk P l. We prove the lemma by nducton. Base Case: From the allocaton algorthm we observe that r K = R. Clearly, XN l =1 r K = and hence, the asserton holds n the base case. XN l =1 R P l ; Inductve Hypothess: Assume that the premse holds for all k K + j. To prove that t holds for all k, we need to show that t holds for k = K + j + 1. From the allocaton algorthm, XN l =1 r K+j+1 = XN l XN l r K+j + =1 =1 7 R? P l P l :

9 Ths completes the proof. From now on, we would assume that a connecton s admtted only f P N l =1 R P l s satsed for all lnks l t passes through. Ths condton can be used as a smple admsson control test for a new connecton. The next theorem proves the mnmum number of slots that C gets n the worst case. Theorem 3.1 If N l be the number of connectons beng multplexed on lnk l and P N l =1 R P l, all C 's get at least (bmr c? 1) slots n m consecutve frames startng at the begnnng of a busy perod. Proof: Let us assume that a busy perod starts from frame K. From the allocaton algorthm we observe that, n each busy perod C gets atleast br c slots. Therefore, t gets atleast mbr c slots n m consecutve frames. Let us denote the fractonal part of R by fr g(= R? br c). In the worst case C wll not get any extra slot for ts fractonal requrement. Therefore, the fractonal requrement bulds up to mfr g. When mfr g 1, C s gven one extra slot, and C gets bmr c slots. If mfr g < 1, C does not get any extra slot n the worst case, and the number of slots allocated to t becomes (bmr c? 1). Theorem 3.2 If a connecton C s (W a; M a p ) and (W ; M p ) smooth before enterng a swtchng node and M a (bw ar c? 1), then the output stream from the swtch s also (W a; M a) and (W p ; M p ) smooth. Proof: The allocaton algorthm enforces an upper bound of M p p packets n a perod of W frames. Hence the output stream s (W p ; M p ) smooth. Now, snce C s (W a; M a ) smooth, the trac that enters a swtchng node n a perod of W a frames s at most M a. Snce M a bw ar c? 1, the trac that enters the swtchng node n a perod of W a frames s served wthn that tme, leavng no backlog. Hence, the output s also (W a; M a) smooth. Theorem 3.2 s extremely mportant n dervng end-to-end performance. Snce a trac stream conforms to ts orgnal shape after passng through a swtch, performance bounds for a mult-hop connecton can be computed by smply addng the correspondng bounds at the ndvdual swtches on the path. In the subsequent dscusson we wll assume that M a bw ar c? 1 for all streams at all nodes. Ths condton denes a lower bound on the servce rate (R ) requred by a connecton. The choce of R depends on other factors such as maxmum tolerable delay, delay jtter etc. Next, we derve the worst case bound on the delay and jtter suered at a swtch n solaton. In order to compute delay and jtter we rst derve the worst case queue length of a connecton at a swtch. 8

10 No Of Packets a W Max Backlog M a p W Servce Curve M p Arrval Curve Tme Fgure 3: Delay bound computaton. Lemma 3.2 The worst case queue length (backlog) of a connecton C at a swtch cannot exceed, B max max(m p ; M a? b((m a =M p? 1)W p? 1)R c + 1): Proof: Consder a busy perod. The worst case arses when a burst of sze M p cells comes every W p frames (refer to gure 3). If the server can clear the burst before the arrval of the next one, the backlog can be at most M p cells. If the next burst comes before the last one s cleared, backlog accumulates and reaches the maxmum at the pont shown n gure 3. For smplcty, we assume that M p dvdes M a. The worst case backlog can be computed by subtractng the number of packets transmtted from the number of packet that entered the swtchng node from the begnnng of a busy perod to the pont of maxmum backlog. Clearly, trac nput s at most M a packets. To nd the number of packets transmtted, we observe that number of frames expred from the begnnng of the busy perod to the pont of maxmum backlog s (M a=m p? 1)W p. Out of these, at most one frame tme may be wasted because of asynchrony between the arrvng and departng frames. Hence, at least (b((m a =M p the proof.? 1)W p? 1)R c? 1) packets are transmtted. Ths completes 9

11 Theorem 3.3 The maxmum queung delay (expressed n number of frames) suered by a packet from connecton C at a swtch s, Proof: D max d max(m p ; M a? b((m a=m p? 1)W p? 1)R c + 1) e + 1 R We derve the bounds for worst case delay suered by a packet from a connecton C by rst computng the maxmum backlog and then ndng the tme t takes to clear t. From lemma 3.2 we get the worst case backlog. In the worst case, the tme t takes to clear ths backlog s db max =R e frames. Addng to that the worst case delay due to asynchrony n arrvng and departng frames, we get the expresson on the rght hand sde. Corollary 3.1 The buer space requred at each swtch to guarantee no packet loss for connecton C s B max. Proof: Keepng a buer of the sze of maxmum possble buer sze s sucent to guarantee no pack loss. The proof follows drectly from theorem 3.2 Corollary 3.2 The maxmum jtter suered by packets from a connecton C at a swtch s, Proof: J max D max The delay jtter s dened to be the derence between the maxmum and mnmum delay. Snce the mnmum queung delay suered by a packet s zero the result s obvous. Theorem 3.4 The maxmum end-to-end queung delay and jtter suered by a packet from an H-hop connecton C may not exceed, HD max. Proof: From theorem 3.3 and corollary 3.2 we get, maxmum delay and delay jtter suered by a packet from a connecton C s D max. From theorem 3.2 we get, trac from connecton C mantans ts smoothness all through the network. Hence, the delay and the jtter bound for an H-hop connecton s the sum of the bounds at ndvdual swtches. Ths concludes the proof. The results n ths secton are derved under the assumpton of a unform frame sze all over the network. They can be extended when derent swtches/lnks use derent frame szes. In that case the shape of the trac would not be the same all through the network. However, knowng the nput characterzaton of a stream enterng a swtch, the output can be characterzed. Hence, to compute the end-to-end performance bounds, bounds at each swtch on the path has to be determned ndvdually. 10

12 4 Choosng the Control Parameters In the followng dscusson we study the mpact of derent control parameters on worst case delay usng numercal examples. Note that the averagng nterval W a for delay bound. However, R mean = M a =W a lower bound on the allocated bandwdth. On the other hand R peak does not appear n the expresson denes the average rate of arrval and determnes the = M p =W p denes the peak arrval rate. To get a quanttatve feelng, n gure 4 we plot the maxmum delay D max aganst. As expected, we observe longer worst case delay for hgher mean to peak rato. We R peak =R mean also observe a sharp mprovement n delay bound wth an ncrease n rate of servce Servce Rate/Mean Arrval Rate=1 Servce Rate/Mean Arrval Rate=2 Servce Rate/Mean Arrval Rate=3 140 Delay Bound Peak to Mean Rato Fgure 4: Eect of burstness on delay bound. In ths example W a = 200,M a = 1000 and W p = 125. The worst case delay not only depends on the peak and mean arrval rates and rate of servce, t also depends on the averagng wndow sze. The followng plot examnes the mpact of M p and W p for the same value of R peak = M p =W p. It s apparent that for the same peak rate of arrval, longer averagng wndow sze and hence larger bursts lead to longer delays. As mentoned earler, trac smoothng at the network entry pont allows one to separate delay nto two components delay n the shaper buer and delay n the network. The rst component of the delay depends on the statstcal characterstcs of the source and the shapng envelope chosen by the user. The second component s ndependent of the actual characterstcs of the source and depends on the shapng parameters and network load and can be bounded by the network. The user, dependng on the cost and performance tradeos, can choose the most approprate parameters. 11

13 Servce Rate/Mean Arrval Rate=1 Servce Rate/Mean Arrval Rate=2 Servce Rate/Mean Arrval Rate=3 140 Delay Bound Peak Wndow Sze Fgure 5: Eect of M p on delay bound. In ths example W a = 200,M a = 1000 and R peak =R mean = 2. 5 Realzaton of the Scheme In ths secton we brey outlne the desgn of a buer manager (gure 6), the key component used to mplement the schedulng mechansms at the swtch. The archtecture of the buer manager s based on the desgn suggested n [2]. However, our scheme does not requre the sequencer, the most complex component used n [2]. The buer manager keeps track of the requrements of each actve connecton C n the current frame. It also stores ponters to the Head and Tal of each connecton queue. At any tme a packet at the head of connecton queue s elgble to be scheduled for transmsson f t satses the elgblty crtera descrbed n the schedulng algorthm (refer to secton 2). The buer manager conssts of a cell pool, an dle address FIFO, a table-look-up memory and a processor. The cells are stored n the cell pool. The table-look-up-memory, referred to as connecton table, stores necessary control nformaton. The dle address FIFO contans addresses of the current empty cell locatons n the pool. The buer manager mantans a vrtual queue for each connecton by lnkng the packets through ponters. Cells stored n the cell buer are tagged wth the ponter to the next cell from the same connecton, f any. When a cell arrves, t s stored n the cell pool at the address gven by the dle address FIFO. Whle the cell s beng wrtten nto the cell pool, ts vrtual channel denter s extracted by the processor and the correspondng entry n the connecton table s moded. If the connecton queue s empty, the buer address of the cell s wrtten nto both Head and Tal elds. 12

14 Otherwse, the buer address s wrtten nto the Tal eld only. Ponters are sutably adjusted to add the arrvng cell at the end of the connecton queue. Connecton Table VCI Head Tal r m p E1 E2 E3 Processor Idle-Address FIFO Waddr Buffer Space Tag Raddr Arrvng Cell In Out Departng Cell Cell Pool Fgure 6: Archtecture of the buer manager. The connecton table stores three elgblty ags, E ; = 1; 2; 3, for all actve connectons. The ag E 1 s set when a stream has an outstandng ntegral requrement, the ag E 2 s set when t has an outstandng fractonal requrement and ag E 3 s set f the stream has not exceeded ts maxmum allowed lmt 6. Flags are set only f a connecton has a packet to send. Ths ag can be set/reset by a local logc and does not need processor cycles. Durng transmsson of a cell, an elgble connecton s selected. The buer address of the cell at the head of the connecton s stored n the connecton table and s extracted by the processor and fed to the Raddr lnes of the cell pool. The buer address of the next cell n the queue, f any, s wrtten nto the Head eld of the correspondng entry n the connecton table. The processor also makes necessary changes to the control nformaton stored n the connecton table. 6 A connecton C can send at most M p cells n a tme wndow of W p frames. 13

15 6 Related Work Trac control n ntegrated servces packet swtched networks has receved tremendous attenton n the last few years. For a very good survey on ths topc refer to [1]. In the rest of the secton we brey revew some of the related lterature on rate based trac control. The potental benet of trac regulaton at the swtchng nodes has been analyzed by Cruz n [3, 4]. Usng a two parameter characterzaton of trac he derves bounds on burstness and end-to-end delay. For some cases they are shown to ncrease exponentally wth the number of hops n a connecton. Golestan [5] proposes a frame based scheme for peak rate constraned sources and stop and go queung at the nodes. Kalmenak et. al. [6] propose a conceptually smlar scheme usng a herarchy of frame szes. Though these schemes are very smple, they requre network wde standardzaton of frame szes. Besdes, allocatng bandwdth at the peak rate can lead to consderable wastage of network resources. Another rate-based approach to trac management s the vrtual clock scheme proposed by Zhang [10]. The vrtual clock scheme can be vewed as a generalzaton of the round robn servce dscplne. Its dstngushng feature s the provson of farness and mnmum throughput guarantees to competng connectons. However, no bound on delay or jtter was obtaned n [10]. Parekh et. al. [7] analyze the network behavor for leaky bucket constraned sources and headof-the-lne processor sharng servce dscplne at the swtchng nodes. Zhang et. al. [9] propose a statc prorty based rate controlled scheme. In ther scheme, a connecton s assgned a lnk deadlne and a correspondng schedulng prorty on each lnk along ts path. Ths scheme can guarantee tght delay and jtter bounds. However, t s not clear how to choose the lnk deadlnes to acheve a specc qualty of servce. The schemes closest to ours are [5, 6, 9, 8]. Our work ders from [5, 6] n that we consder multrate constraned trac as opposed to peak rate constraned trac used n these schemes. Also, our scheme does not requre network wde standardzaton of frame szes and s more exble n accommodatng dverse throughput requrements of derent applcatons. The derence between our approach and the one suggested n [9, 8] s n the servce dscplne at the swtches. We use a frame based scheme n contrast to calendar based approach used n [9, 8]. Unlke [9, 8], we use a very smple bandwdth allocaton algorthm to acheve a specc qualty of servce. 7 Concludng Remarks We present a framework for trac control n an ntegrated servces ATM network. The proposed scheme comprses of two components: a shapng mechansm at the network entry pont, and a servce dscplne at the swtches. The shaper enforces dual rates, namely a mean and a peak rate. 14

16 At the swtches, each connecton s guaranteed a mnmum bandwdth but, may receve a better servce f there s unused capacty. We propose a frame based bandwdth allocaton scheme. Unlke other frame based schemes, our scheme allows allocaton of bandwdth at any arbtrary granularty. We outlne a hardware realzaton of the control mechansm and derve worst case end-to-end delay and jtter bounds. References [1] C. M. Aras, J. F. Kurose, D. S. Reeves, and H. Schulzrnne. Real-Tme Communcaton n Packet-Swtched Networks. Proceedngs of the IEEE, 82(1), January [2] H. J. Chao. Archtecture Desgn for Regulatng and Schedulng User's Trac n ATM Networks. In Proceedngs, SIGCOMM, August [3] R. L. Cruz. A Calculus for Network Delay, Part I: Network Elements n Isolaton. IEEE Transsactons on Informaton Theory, 37, [4] R. L. Cruz. A Calculus for Network Delay, Part II: Network Analyss. IEEE Transsactons on Informaton Theory, 37, [5] S. J. Golestan. A Framng Strategy for Congeston Management. IEEE Journal on Selected Areas n Communcaton, 9(7), September [6] C. R. Kalmanek, H. Kanaka, and S. Keshav. Rate Controlled Servers for Very Hgh Speed Networks. In Proceedngs, GLOBECOMM, December [7] A. K. Parekh and R. G. Gallager. A Generalzed Processor Sharng Approach to Flow Control n Integrated Servces Network: The Sngle Node Case. IEEE/ACM Transactons on Networkng, 1(3), June [8] H. Zhang. Servce Dscplnes for Packet-Swtched Integrated-Servces Networks. Ph.D. Thess, Unversty of Calforna, Berkeley, [9] H. Zhang and D. Ferrar. Rate Controlled Statc Prorty Queung. In Proceedngs, INFOCOM, [10] L. Zhang. Vrtual Clock: A New Trac Control Algorthm for Packet Swtchng Networks. In Proceedngs, SIGCOMM,

17 '.. INITIALIZE(C ) /* Intalzaton for connecton C. */ r R /* Slot requrement n current frame */ w p 0; mp M p /* To enforce peak rate constrant */ $ & TIMER(C ) /* Operates on all actve connectons */ At the begnnng of a frame: f (w p (wp p + 1) mod W ) = 0 then mp M p end f SCHEDULE(l) /* Schedulng over lnk l */ At the begnnng of a frame: 1. Intalze N P l /* Number of slots n a frame */ 2. Satsfy ntegral requrements: whle (9C :: r 1 ^ m p > 0) ^ N > 0 do allocate next slot to C r r? 1; m p mp? 1; N N? 1 end whle 3. Satsfy fractonal requrements: whle (9C :: r > 0 ^ m p > 0) ^ N > 0 do allocate next slot to C r r? 1; m p mp? 1; N N? 1 end whle f N > 0 then EndofBusyPerod true 4. Dstrbute slack slots: whle (9C :: m p > 0) ^ N > 0 do allocate next slot to C m p mp? 1; N N? 1; end whle 5. allocate remanng slots to non-real-tme traffc 6. Compute requrements for next frame: f EndofBusyPerod then else end f r R ; EndofBusyPerod r r + R false % Table 2: Swtch schedulng functons. 16