Abstract. Introduction. B. Integrated Services and RSVP. A. Classical Admission Control

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1 A SURVEY ON ADMISSION-CONTROL SCHEMES AND SCHEDULING ALGORITHMS Masaru Ouda, Murray State Unversty Abstract There has been a sustaned nterest among researchers and networ operators n provdng qualty of servce (QoS) over the Internet. As an essental tool for supportng QoS, development of effectve and scalable admsson control s an mportant topc of research. Over the years, varous admsson-control schemes have been proposed that clam to scale well n a networ envronment where the networ core s ept relatvely smple and processng burdens are pushed to the edges of the networ. Ths study surveyed selected admsson-control schemes of ths type. The contrbuton of ths study s an ntroducton of new classfcatons of admsson-control schemes, whch s based on locatons where the ey admsson-control mechansms are mplemented n a networ. The survey of the lterature was conducted n lght of locaton-based classfcaton of admsson controls and detals the worngs of schemes, dscusses ther contrbutons, and dentfes areas of further development. Introducton Admsson control s a process through whch a networ node determnes whether to accept a new flow request or deny t. It s a traffc management tool through whch the load on a networ s controlled. The admsson decson s made based on several crtera: ) the current and future avalablty of networ resources, ) the mpact of admsson decsons on the exstng flows, and 3) the polcy control mplemented by the networ admnstrator. Admsson control s essental when the networ promses servce guarantees or levels of servce assurances. The goals of the admsson control are to protect the performance obectves of the exstng flows, deny any requests the networ s unable to provde for, and accept as many new flows as the networ can commt to. A. Classcal Admsson Control Admsson control has been a topc of strong nterest among researchers for many years. Research actvtes were partcularly actve when ATM standards were emergng. ATM employs a connecton-orented, hop-by-hop admsson-control scheme as follows: A call s requested from a user to the networ by means of sgnalng. The sgnalng message carres a profle of the requested call, referred to as a traffc descrptor, whch detals the characterstcs of the generated traffc such as pea rate and delay requrement. Upon recevng the call request, the networ node executes an admsson test by examnng the traffc descrptor aganst the current state of the node. If enough resource s avalable, the node admts the new call and forwards the request to the downstream node. The downstream node, n turn, executes an admsson test and decdes whether to admt the requested call or not. Ths process s repeated untl the call request reaches the destnaton. In order to mae a sound admsson decson, each node mantans the state of all calls establshed through the node. Ths nformaton s updated every tme a new call s added or an exstng call s termnated. Due to a large amount of state nformaton requred at each node, concerns have been rased regardng resource usage effcency and the scalablty of such admsson processes. B. Integrated Servces and RSVP Successful deployment of ATM networs nspred researchers and engneers to buld IP networs capable of QoS support, smlar to that of ATM. Much of the nowledge and experence ganed from ATM has been ncorporated nto the desgn of new IP networs. Integrated Servces [] and RSVP [] are the outcomes of ther effort and defne the core specfcatons for QoS-enabled IP networs. The combnaton of IntServ and RSVP gave hope for the QoS support on IP networs. As wth ATM, IntServ archtecture ams to provde servce guarantees through resource reservaton. Through RSVP, t employs end-to-end sgnallng to communcate QoS parameters for the reservaton of resources. However, on ths type of archtecture, each reservaton of resource requres a state to be mantaned at every node along the path of an end-to-end flow. It has been sad that such archtectures may not scale well due to heavy processng overhead and large memory consumpton requred to mantan those flow states. Consderng the rate at whch the sze of Internet s growng and the number of hosts beng added, the concentraton of flows wthn the core routers can be a real ssue and the management of ndvdual flow states wll become ncreasngly dffcult. A mechansm that smplfes the operatons of the networ core s desred. A SURVEY ON ADMISSION-CONTROL SCHEMES AND SCHEDULING ALGORITHMS 9

2 C. Dfferentated Servces Archtecture To remedy the scalablty problem of IntServ wth an RSVP approach, dfferentated servces [3] have been proposed. DffServ acheves the scalablty by relevng the networ core from resource-ntensve operatons and placng the complexty at the edge routers. Specfcally, classfcaton and condtonng of pacets s performed only at the edges of the networ. DffServ does not employ hop-by-hop sgnalng n order to avod the mantenance of per-flow state swthn the core of the networ. Instead, flows wth smlar profles are aggregated at the edge routers so that the core routers only need to handle bundles of flows. DffServ supports class-of-servce dfferentatons. In order to mantan the promsed level of servce, the amount of traffc accepted at each class, especally at hgher levels of classes, must be lmted. Otherwse, the Servce Level Agreement between the user and the networ wll be volated [4]. Thus, there s a need for admsson control. In order for the edge nodes to mae sound admsson decsons, they must receve feedbac from other parts of the networ. DffServ specfcaton maes no menton of how ths s to be done. D. Multprotocol Label Swtchng Multprotocol Label Swtchng (MPLS) [5] s an evolvng and expandng set of protocols developed by IETF. MPLS can be seen as a combnaton of dfferent feature sets from ATM, IntServ, and DffServ. Ths s acheved through the creaton of a undrectonal sgnaled path, nown as Label Swtched Path. Label Swtched Path s establshed by RSVP-based call control, nown as RSVP-TE. MPLS ams to provde QoS-enabled transmsson paths over the Internet. MPLS employs encapsulaton of pacets wth short descrptors nown as labels at the entry nodes of the MPLS networ. The label determnes whch QoS class the pacet belongs to and where t wll be forwarded to. The same label wll be placed on all pacets that belong to the same QoS class and the same forwardng destnaton. MPLS s ganng wde acceptance as a WAN protocol-ofchoce and replacng Frame Relay and ATM-based WANs. It s used to transport Voce Over IP (VOIP) traffc and extend Ethernet LANs over the Internet. The strengths of MPLS nclude the seamless support of IP pacets wth QoS support, ablty to operate through segments of networ that do not support MPLS, and scalablty afforded by the mplementaton of labels and smple operatons at the core of the networ. MPLS s protocol agonstc n that the payload of the labeled pacets may be of any type, such as Ethernet frames or ATM cells. MPLS s desgned specfcally for those protocols that do not support QoS natvely, such as IP. MPLS allows multple layers of label encapsulatons to allow tunnelng through dfferent admnstratve MPLS domans. Because MPLS uses RSVP-based call control, t nherts the same strengths and weanesses of RSVP. There s a need for admsson control that scales well n an envronment where core routers are ept relatvely smple and processng burdens are pushed to the edges of the networ. Ths study surveyed admsson-control schemes recently proposed, all of whch clam to offer some level of scalablty. The remanng sectons of ths paper are organzed as follows: Secton classfes the admsson-control schemes and schedulng algorthms n several categores. Secton 3 surveys the admsson-control schemes beng proposed n recent years and descrbes the goals, approaches, contrbutons, and shortcomngs of each scheme; Secton 4 concludes the survey. Classfcatons of Admsson Control and Schedulng Algorthms Ths secton descrbes the classfcatons of admssoncontrol schemes and schedulng algorthms. A. Parameter-Based vs. Measurement- Based Admsson Control Admsson-control schemes are generally classfed as ether parameter-based or measurement-based approaches. In ether case, users request servce from the networ by sendng flow specfcatons. Flow specfcatons descrbe the nature of pacet flows (e.g. pea rate) and requrements for pacet handlng wthn the networ (e.g. loss rate). The networ uses parameters specfed n the flow specfcatons to compute how much resource t must set asde n order to support the requested flow. The admsson decson wll be made by comparng the requred resource aganst what s avalable on a node. The dfferences between the two approaches exst n the way the allocated resource on a node s beng estmated. In parameter-based admsson control, the node computes ts reserved resource by eepng trac of parameter values n flow specfcatons at each flow establshment and termnaton. Wth ths approach, the amount of allocated resource s a dscrete functon and the networ node nows exactly how much resource s used or reserved at any gven tme. The strength of ths approach les n ts ablty to provde hard guarantees to each flow beng accepted. One of the shortcomngs s that t does not use the resource effcently. The worst-case scenaro s typcally used to compute the resource reservaton requrements to assure hard guarantees. Once the 9 INTERNATIONAL JOURNAL OF MODERN ENGINEERING VOLUME 0, NUMBER, SPRING/SUMMER 00

3 resource s mared as reserved, t s no longer avalable for new flows that request guaranteed servce. Under the measurement-based approach, the resource consumed by exstng flows on a networ s estmated by measurng the actual traffc flow. It apples statstcal prncples to assess the current and very-near-future state of the networ. Expressed by way of confdence level, t can predct a lelhood of beng able to support a requested level of servce based on the traffc pattern of the past. Usng ths nformaton, a networ node decdes whether to admt a flow or reect t. Ths approach s shown to have much better utlzaton of networ resources than the parameter-based one. However, measurement-based admsson control does not provde hard guarantees. The level of assurance ths approach gves s based on the past hstory; the applcablty of confdence level depends on whether the traffc pattern wll reman smlar to that of the past. Snce the networ s not mmune to sudden changes n ts envronment (e.g. traffc pattern changes, ln falures), the measurement-based approach may be effectve only on stable networs. Another shortcomng of ths approach s that t requres an accumulaton of a long hstory. In order to yeld a hgh utlzaton, the confdence nterval at a gven confdence level must be ept short. Ths requres many samples. Wthout a long hstory, admsson decsons must be made wth a very conservatve vew of the unused resources. In recent years, admsson-control schemes that are hybrd between parameter-based and measurement-based approaches have been nvestgated [4], [6], [], [4]. They ncorporate past hstory (.e., measurements) to adust the reserved bandwdth (.e., parameters) of flows. Due to ther dualty, the strengths of ether approach may mtgate the weanesses of the other. Because of ths unque property, a hybrd approach to admsson control s ganng nterest. B. Stateful vs. Stateless Schedulng Algorthms The manner n whch the arrvng pacets are queued and processed at each networ node, referred to as schedulng, can have a sgnfcant mpact on the way the admsson control s carred out. Schedulng algorthms are generally classfed as stateful or stateless for the purpose of scalablty dscusson. Stateful algorthms requre mantenance of ndvdual flow state at every node along the path of a flow. Examples of stateful-schedulng algorthms nclude Far Queueng [7], [8], Vrtual Cloc [9], and ther varants such as Weghted Far Queueng [8] and Jtter-Vrtual Cloc [4]. These algorthms have been developed for the support of guaranteed servce as ther prmary obectve. They gve precse control over the treatment of ndvdual flow and can provde bounds on bandwdth allocaton and end-to-end delay. The maor drawbac of stateful schedulers s that they requre mantenance of per flow QoS state of all flows at each networ node. Due to the sze of and the complexty nvolved n the management of QoS state nformaton, the scalablty of ths approach has been challenged. When a stateful scheduler s deployed n a networ, admsson control maes use of ndvdual QoS state mantaned at each node and determnes whether the node has suffcent resources to meet the demands of the newly-requested flow. Stateless-schedulng algorthms, on the other hand, mantan no QoS state at any part of the networ. FIFO and LIFO queueng are examples of stateless algorthms. Snce t requres no state mantenance, t s scalable. However, t does not provde the control necessary to support varous QoS requrements. The Internet, for the most part, s composed of networ nodes supportng stateless-queueng algorthms. In recent years, a new type of scheduler has been added to the above. It s called core-stateless schedulng. Corestateless schedulng ams to provde a smlar level of QoS control offered by stateful algorthms, yet tres to acheve networ scalablty comparable to one offered by stateless algorthms. In core-stateless algorthms, the edge nodes mantan QoS states of ndvdual flows, but the core routers do not. The core routers may mantan aggregate-level nformaton that asssts n controllng flows, dependng on the mplementaton. The elmnaton of ndvdual flow states from the core routers s made possble by embeddng the QoS states n each pacet header. There have been some novel deas proposed usng ths schedulng mechansm. Core-Stateless Far Queueng [0], Core-Jtter VC [4], and Vrtual Tme Reference System [] are examples of corestateless algorthms. They are further explaned later n the survey secton of ths paper. C. Locaton Based Classfcaton of Admsson Control A contrbuton of ths study s the ntroducton of locatonbased classfcaton of admsson-control schemes. It s a new classfcaton based on locatons at whch the ey admsson-control algorthms are appled. Accordng to locaton-based classfcaton, admsson-control algorthms proposed n recent years are classfed nto the followng fve categores: admsson control at ) edge nodes (Edges), ) central node (Central), 3) ngress node (Ingress), 4) egress node (Egress), and 5) end-user staton (End-to-End). The taxonomy of admsson-control schemes s gven n Fgure. A SURVEY ON ADMISSION-CONTROL SCHEMES AND SCHEDULING ALGORITHMS 93

4 Admsson Control Schemes returns the summary statstcs. Upon recevng the summary nformaton, the sender decdes whether the networ s capable of carryng the requested load. Scalable Reservaton Protocol [5] and others [6] [8] belong to ths category. Parameter Based Hybrd Measurement Based Survey of Admsson-Control Schemes Edges Central Ingress Egress End-to-End Fgure. Taxonomy of admsson-control schemes Admsson control at edge nodes (.e., Edges) lessons the processng requrements of core routers through flow aggregaton at networ edges. Core routers process and mantan only the aggregate flow reservaton nformaton. Through aggregaton, overhead reducton s made possble by fewer sgnalng-message exchanges and less call-state mantenance. Aggregaton of RSVP [3] belongs to ths category. Admsson control at central node (.e., Central) employs a master server that performs admsson-control functons on behalf of all routers n a networ. By off-loadng resourcentensve servces, core routers become lghtweght. Bandwdth Broer [] uses ths approach. Admsson control at ngress node (.e., Ingrss) enables core routers to mae admsson decsons wthout needng to mantan ndvdual flow states. Ingress node measures the rate of pacet arrvals for each ndvdual flow and nserts ths nformaton n each pacet header. Core routers read ths nformaton and accumulate them per aggregate flow. Thus, the core routers only mantan aggregate flow states and are able to mae admsson decsons at an ndvdual flow bass. Dynamc Pacet Sate (DPS) [4] uses ths approach. Admsson control at egress node (.e., Egress) pushes the complexty to the egress routers so that no per-flow states need to be mantaned n the core of the networ. Egress routers construct profles of flows by montorng the pacet arrvals and departures. By measurng delay experenced by each pacet, egress routers estmate the dynamcally changng networ load. Based on ths nformaton, the egress routers mae admsson decsons. Egress admsson control [4] belongs to ths category. Admsson control at the end-user staton (.e., End-to- End) uses a form of n-band sgnalng to estmate the avalablty of networ resources. The admsson decson s typcally made by end users, rather than the networ. Pror to sendng data traffc, an orgnatng end user sends a stream of pacets at a constant rate for a short perod of tme. The recever measures the arrval pattern of probng pacets and In lght of locaton-based classfcatons of admssoncontrol schemes descrbed above, ths secton surveys selected admsson-control schemes. A. Admsson Control at Ingress Node Dynamc Pacet State (DPS) [4] s an ngress-node based admsson-control scheme and employees a core-stateless scheduler. Its goal s to mae admsson decsons for new flows wthout mantanng ndvdual flow states n the core of the networ. DPS also ams to acheve end-to-end perflow delay and bandwdth guarantees on a networ, where only the edge routers perform per-flow management. To meet these goals, DPS uses a pacet-header marng technque, where the ngress node encodes state nformaton on the header of each pacet. The core nodes apply control to pacets accordng to ther header marngs. DPS proposes two nnovatve schemes, one n admsson control and the other n schedulng. How these schemes wor s descrbed n subsequent sectons. DPS s admsson-control scheme s comprsed of two algorthms: ) per-hop admsson control and ) aggregate reservaton estmaton. The former s parameter-based, whle the latter s measurement-based. Each algorthm ndependently computes an estmated reserved bandwdth of aggregated flows. These estmates should be very close, f not the same. However, under certan condtons, devatons from the true reserved bandwdth are observed on each of the two algorthms n opposte drectons. One algorthm estmates at a hgher rate than the true reserved bandwdth and the other estmates at a lower rate. The frst algorthm does not account for the duplcate reservaton requests. Ths can lead to an under-utlzaton of a ln due to nflated estmaton of the reserved bandwdth. The second algorthm does not nclude the effects of new calls beng admtted n the mddle of an estmate cycle. Ths results n estmatng the reserved bandwdth at a lower rate than the actual rate. The results from these two algorthms are reconcled at the end of a fxed nterval and arrve at one value that better reflects the true reserved bandwdth. The goal of admsson control n DPS s to estmate a close upper bound on a reserved aggregate rate so that a determnstc guarantee can be made to those calls beng accepted, whle 94 INTERNATIONAL JOURNAL OF MODERN ENGINEERING VOLUME 0, NUMBER, SPRING/SUMMER 00

5 mnmzng over-reservaton. DPS proposes a schedulng algorthm that provdes servce guarantees at levels comparable to IntServ on DffServ-le envronments. Ths schedulng algorthm s called Core-Jtter Vrtual Cloc (Core- Jtter VC). It s a non-wor conservng schedulng algorthm. Core-Jtter Vrtual Cloc s a varant of Jtter Vrtual Cloc (Jtter VC). The prmary dfference between the two algorthms s that Core-Jtter VC s a core-stateless-based scheduler, whle Jtter VC s a stateful scheduler. Core-Jtter VC provdes the same delay guarantee as Jtter VC at an end-to-end path, but not at ntermedate routers. Jtter VC has been proven to provde the same level of guarantee as Weghted Far Queueng (WFQ) [9]. Thus, Core-Jtter VC also provdes the same guarantee as WFQ at the end-to-end path. Jtter VC and Core-Jtter VC, are based on a pacet-header marng and queueng archtecture, where each router n a path of a flow reads and re-mars pacet-header nformaton for queueng and schedulng purposes. They employ a delaytter-rate-controller unt [0] for queueng purposes and a Vrtual-Cloc scheduler for schedulng purposes. A pacet enterng nto a Jtter-VC router or a Core-Jtter-VC router wll be held n a watng room by the delay-tter-ratecontroller untl t becomes elgble for transmsson. Once the pacet s released from the watng room, Vrtual-Cloc scheduler servces them n order of ther earlest deadlne. Each pacet s gven a deadlne by whch t must leave the Jtter-VC server or the Core-Jtter-VC server., of a prevously receved pacet when t computes the elgblty tme, e,, of an arrvng pacet from the same flow. In order to better explan the worngs of Core-Jtter VC, Jtter VC s descrbed frst. For the th pacet of flow, ts Core-Jtter VC mproves upon Jtter VC and maes the elgble tme e, and deadlne d, at the th node on ts scheme stateless. It does so by removng the term d, from path under the Jtter-VC algorthm are computed as follows: the equaton (4) and ntroducng a new term, δ, a slac varable, nstead, whch holds the followng property: e = a e = max( a + g, d ),, > (4),,,,,, l d, = e, +,,, (5) r where a, s the arrval tme,, l s the pacet length, and g s the amount of tme between the pacet s deadlne and the actual departure tme. At every pacet departure at every router, ths g value s computed and recorded n ts pacet header and read at the subsequent router. A sample tme dagram of pacets gong through a seres of Jtter-VC servers s depcted n Fgure (a). The shaded area depcts delays experenced by the second pacet at each node. node (ngress) node node 3 node 4 (egress) node (ngress) node node 3 node 4 (egress) e, d, e, e, d, e, e, d, e,3 (a) e, d, e,3 (b) d, e, d, d, d, 3 e,3 d, 3 e, e,4 d, 3 d, e,4 e,3 d, 4 d, 3 d, 4 e,4 e,4 d, 4 tme d, 4 tme Fgure. The tme dagram of pacets through (a) Jtter- VC servers and (b) Core-Jtter-VC servers Jtter VC s a stateful servce because, by equaton (4), each router must mantan the deadlne,,, δ,, d, d + g + g > (6) Wth the above defnton, the elgblty tme of a pacet at the th node can be computed as follows compare t to equaton (4): e = a + g + > (7),,, δ, The detals of how the actual value for the slac varable δ s determned are gven by Stoca and Zhang [4]. A sample tme dagram of pacets gong through a seres of Core-Jtter-VC servers s depcted n Fgure (b). Observe that the slac tme, δ, s a fxed value for all partcpatng nodes n flow for the th pacet. A SURVEY ON ADMISSION-CONTROL SCHEMES AND SCHEDULING ALGORITHMS 95

6 The strengths of DPS nclude ts ablty to guarantee bandwdth and delay bound-through Core-Jtter VC. The proposed admsson-control algorthm s robust n the presence of networ falures and partal reservaton snce the algorthm to estmate the reservaton rate does not remember the past beyond the perod, T W. DPS s largest contrbuton s that t s the frst of ts nd to demonstrate that there s a way to provde a hard guarantee on bandwdth and delay requrements wthout mantanng ndvdual flow states n the core of the networ. DPS proposed a noble dea of nsertng ndvdual flow states n the header of pacets so that the core nodes don t have to mantan them. DPS nspred others n developng new schemes based on ths premse [], []. Whle DPS offers guaranteed servces wthout ndvdual flow-state mantenance at core routers, the overall scalablty ganed from ths archtecture remans a queston. Inserton and nterpretaton of state nformaton n every data pacet can be an expensve operaton. Indeed, there s a concern that DPS may be transformng all data pacets nto control pacets such that core nodes must pay extra attenton to every pacet they receve regardless of ts type. Core-Jtter Vrtual Cloc scheduler requres both ngress and core nodes to montor and alter the header of every data pacet that travels through. For admsson control, only the ngress router wrtes to the pacet header, yet core routers must read and process every data pacet. Snce the control nformaton s embedded n the data pacet headers, all pacets become essental to the healthy operaton of the networ. Consderng the parameter-based QoS model, where only the control pacets need extra attenton from the routers, DPS s new approach could potentally add hgher processng demands on networ routers. Another drawbac of DPS s admsson control s that t requres nserton of dummy pacets at the ngress router when there s no data flow. Dummy pacets must be nected n the networ every tme there s a gap between data pacets, whch s larger than the maxmum nter-pacet arrval tme T I. T I s typcally a small wndow compared to the perod T W used to compute the aggregate reserved rate. Ths type of approach wors well for those applcatons that generate traffc at a constant bt rate and always termnate the reservaton as soon as the transmsson s over, such as telephony. DPS s admsson-control scheme may not appeal strongly to other types of networ applcatons. If the source s slent for an extended perod of tme, constant bt-rate dummy pacets must be nserted nto a networ at / T I rate. Ths could result n wasted bandwdth because even the besteffort traffc cannot tae full advantage of unused bandwdth. B. Admsson Control at Central Node Bandwdth Broer (BB) [] belongs to the centrallycontrolled admsson-control approach that ams to provde scalablty n the networ by off-loadng the routers controlplane functonaltes to a master server nown as a Bandwdth Broer. Bandwdth Broer mantans QoS-state nformaton for all flows of every router wthn a desgnated doman. Networ routers perform only the data-plane functonaltes (.e., pacet forwardng), n addton to the exchange of QoS-related nformaton of each flow wth Bandwdth Broer. Bandwdth-Broer archtecture s bult upon the Vrtual Tme Reference System (VTRS) []. It s classfed as a core-stateless schedulng scheme, where the core routers n the networ do not mantan ndvdual flow states. VTRS s a framewor on whch guaranteed servces can be offered n a networ wthout mandatng that a specfc schedulng algorthm (e.g., Core-Jtter VC) be employed. It conssts of three logcal components: a pacet state carred by pacets, edgetraffc condtonng at the networ edge, and a per-hop vrtual-tme reference and update mechansm at the core routers. VTRS was nspred by the wor presented n Dynamc Pacet State (DPS) [4], where the core-stateless approach was frst ntroduced. VTRS s an extenson to DPS; however, VTRS has unque and sgnfcant contrbutons beyond what DPS proposed. Frst, t establshed generalzed mathematcal expressons that bound end-to-end delay and bandwdth requrement for the support of flows that travel through core-stateless routers. Second, framewor defned n VTRS s generc enough that t not only expresses delay bounds and sustanable rates of a flow through core-stateless schedulers, but also through stateful schedulers (e.g., WFQ) as well as stateless (e.g., FIFO) schedulers. Thrd, the framewor allows mxng of rate-based and delay-based schedulers n the path of a flow. Fourth, t ntroduced two new wor-conservng core-stateless schedulng algorthms: Core Stateless Vrtual Cloc (CSVC), whch s rate-based and Vrtual Tme Earlest Deadlne Frst (VT-EDF), whch s delay-based. Consder the path of a flow on a networ, traversng h hops of routers. Suppose that q routers execute rate-based schedulng and h q routers employ delay-based schedulers. Pacets enterng the networ wll be shaped at the edge node and move through a seres of core nodes. Delays that the pacets experence wll be at the shaper and at each core node. Then, the total delay of end-to-end path of flow, d, s ee d = d + d (8) ee shaper core 96 INTERNATIONAL JOURNAL OF MODERN ENGINEERING VOLUME 0, NUMBER, SPRING/SUMMER 00

7 shaper For smplcty, we do not consder the delay experenced at the shaper n ths paper. It suffces to say that d vares by the type of shaper beng used and the maxmum delay that can be bounded. Zhang gves an example of delay experenced at the shaper usng a dual toen-bucet regulator []. The delay experenced at core nodes s bounded by L d q h q d r,max h h ee = + ( ) + π + ψ = (9) = The term,max L q represents the delay experenced at r rate-based routers and ( h q) d represents the delay observed at delay-based routers. h π = s the total propagaton delay and ψ s the error term of node, whch has the followng property: ˆ fˆ vˆ + ψ (0),,, where f s the vrtual fnsh tme of pacet n flow at node. It means that the targeted pacet-departure tme n a vrtual tme lne s the latest tme the pacet may leave the, node and stll meet the delay requrement. v ˆ s the actual fnsh tme (.e., actual pacet departure tme) of pacet n flow at node. Havng acheved the bandwdth and delay guarantees through VTRS on core-stateless networ, the desgners of VTRS enhanced ts scalablty further by movng the QoSrelated control functons out of core routers to a master server nown as a Bandwdth Broer. Bandwdth Broer s composed of three servce components: polcy control, QoS routng, and admsson control. Polcy control determnes whch hosts and applcatons are allowed to access the networ. QoS routng selects a path that fulflls the requrements of a requested flow. Admsson control determnes the elgblty and feasblty of the requested flow by consultng the polcy control and QoS routng control. Bandwdth Broer (BB) suggests that by movng the admsson-control functon from the core routers to a central server, several postve outcomes can be expected. Frst, t further allevates the core routers from burdensome processng and mang them potentally more effcent. Second, servce guarantees can be made for both per-flow and aggregate flows. Thrd, by decouplng the QoS-related functonaltes of control plane from core routers, t may be possble to ntroduce new guaranteed servces wthout requrng software or hardware upgrades at core routers. Fourth, t allows the executon of sophstcated and optmzed admsson control for the entre networ, whch mght have been dffcult under the hop-by-hop admsson control. Ffth, the problem of nconsstent QoS states observed n the hop-byhop reservaton mechansm can be lessened. Sxth, through the physcal separaton of control- and data-plane functonaltes, ssues n control plane (e.g., scalablty of Bandwdth Broer) can be dealt separately from the ssues n data plane. Seventh, admsson control can be performed at an entre path level, as opposed to a local level as done by the hop-byhop approach, and could reduce the complexty of admsson-control algorthms. Fnally, BB addresses the effects of dynamc on and leave of ndvdual flows to and from an aggregate flow and ncorporates such effects nto the admsson-control algorthm. There are several open ssues wth the desgn of Bandwdth Broer. Whle t addresses the core routers scalablty ssues well, t does not elaborate much on the Bandwdth Broer s scalablty ssues. The amount of flow state nformaton the Bandwdth Broer must manage could ncrease dramatcally as the sze of the networ grows. There s a menton [] that ths problem can be allevated by employng multple Bandwdth Broers n dstrbuted fashon. Ths s contrary to one of the orgnal motves of BB, where t tres to avod the problem of nconsstent networ vew, whch s often ntroduced by the dstrbuted approach. There also may be a potental delay ncurred when the concentraton of communcatons to and from the Bandwdth Broer becomes severe. Though t s convenent to have polcy and QoS-routng nformaton on-hand for admsson decsons, performng all three tass for the entre networ can be demandng and t warrants a careful feasblty study. Fnally, there s always a danger of a sngle pont of falure, whch results not only n an nablty to mae admsson decsons, but also n loss of all QoS-related control-plane functonaltes, whch Bandwdth Broer provdes. C. Admsson Control at Edge Nodes Aggregaton of RSVP reservatons [3] belongs to the edges-based admsson-control approach and t ams to provde scalablty n the networ core by aggregatng reservaton requests of ndvdual flows at the edge nodes. Indvdual flows between the same par of source and destnaton nodes can form an aggregate. It s an extenson to RSVP specfcatons. The prmary focus of ths approach s on the reducton of RSVP message exchanges, whch leads to conservaton of memory and processng power at those locatons where the volume of ndvdual flows may be heavest. A SURVEY ON ADMISSION-CONTROL SCHEMES AND SCHEDULING ALGORITHMS 97

8 The scheme employs two technques to acheve ts goals: suppresson and aggregaton of reservaton messages. Indvdual-flow RSVP requests are suppressed at the ngress node by alterng ther protocol ID. The subsequent nodes n the routng path wll not see these pacets as reservaton messages, except at the egress node. When the pacets reach the egress node, they wll be restored to ther orgnal IDs. As the egress node alters the protocol ID of reservaton pacets for ndvdual flow, t computes the total bandwdth requested from each flow. Once the requests reach a certan total bandwdth, the ngress node ntates an aggregaton and sends an AGGREGATE PATH message to downstream nodes. Upon recevng ths message, the egress node returns an AGGREGATE RESERVE message and nodes n the path commt the reserved resource for the aggregate flow. Each node n the downstream path mars an approprate amount of resources for reservaton. RSVP Reservaton Aggregaton s a logcal and natural extenson to the exstng RSVP. The man contrbuton s that RSVP wll be able to sgnal AGGREGATE PATH and RESERVE messages and that the core routers need not mantan per-flow states any longer. Antcpated resource savngs can be large when the number of aggregated flows s substantally fewer than the ndvdual flows. It wll wor well n an envronment where there are many end statons that are networed together wth a few edge routers, such as VPN. On the other hand, when the scheme s appled to a networ, where the number of edge routers s large and the dstrbuton of flows s evenly spread among all edges, resource savngs may not be as large as the prevously-descrbed envronment due to lac of concentratons. Servce-provder networs typcally belong to the latter type. Furthermore, when the number of aggregate flows ncreases to a substantal volume, they face smlar problems to havng many ndvdual flows. The scheme presented n RFC 375 allows only those ndvdual flows wth the same source and destnaton par to form an RSVP aggregate. Ths s dfferent from DffServ aggregaton, where any flow can form an aggregate, regardless of addresses, so long as they are mared wth the same DS Code Pont. RFC 375 ponts out that frequent modfcatons to the bandwdth reservaton of aggregate flows due to addtons and termnatons of ndvdual flow can lead to a large number of reservaton updates. Ths s contrary to the base assumpton that fewer reservaton messages are generated when ndvdual flow requests are aggregated. On the other hand, nfrequent updates to the reserved bandwdth of an aggregate flow can result n wasted bandwdth, snce a large bloc of resources wll need to be reserved to absorb temporal bandwdth fluctuatons. Thus, there s a trade-off between scalablty of the scheme and effcent use of bandwdth. D. Admsson Control at Egress Node Egress Admsson Control [4] performs data collecton and admsson decsons at the egress router. It processes reservaton messages only at the networ edge (egress router) and uses contnual passve montorng of a path to assess ts avalable resources. It models the networ as a blac-box system, where a flow of pacets arrves at one end of the box (ngress node), goes through the box (core nodes), and comes out at the other sde of the box (egress node). All other flows on the networ are modeled as nterferng crosstraffc of the measured flow. Usng ths bloc-box model, Egress Admsson Control ams to develop envelopes that accurately characterze the upper bounds on arrval and servce processes through measurement at the egress node. A unque characterstc of these envelopes s that they mplctly nclude the effects of cross traffc that are not drectly measured at the egress pont and mplctly prevent other egress ponts from admttng flows beyond an acceptable range. By applyng the extreme theory [] to the measured envelopes, t estmates the end-to-end servce avalablty of a certan traffc class. Ths estmate s used for mang admsson decsons. Egress Admsson Control constructs envelopes for arrval process and servce process. All edge nodes are synchronzed wth each other usng Networ Tme Protocol [3] and they tme stamp every pacet enterng the networ. When pacets reach the egress node, the tme stamp n the pacet header s read and an arrval envelope, nown as pea rate envelope [4], whch captures the behavor of the pea rate of the arrval process, s constructed. The pea rate envelope s constantly updated at a short fxed nterval. At a longer tme scale, changes n the envelope are measured, expressed as varance, and used to compute the confdence nterval of the pea-rate envelope. The servce envelope descrbes the behavor of the worst rate of servce process. When pacets arrve at an egress node, t examnes each pacet s header and computes the delay t experenced. Usng ths nformaton, the egress node constructs the trace of maxmum tme requred to servce a certan number of bts, called mnmum servce envelope. The varance observed by the changes n the servce envelop n a longer tme scale s used to compute the confdence nterval. When a new flow s requested, usng ts declared pea rate and delay bound, addng t to the measured pea rate arrval envelope, the admsson test wll compare ths value aganst the measured servce envelop, tang nto account varances, 98 INTERNATIONAL JOURNAL OF MODERN ENGINEERING VOLUME 0, NUMBER, SPRING/SUMMER 00

9 and determne f statstcally enough bandwdth exsts through the networ. Consder a blac-box system that has a measured pea arrval envelope wth mean R( t ) and varance σ ( t). Assume t has a mnmum servce envelope wth mean S ( t ) and varance ψ ( t). Suppose a new flow request arrves wth the pea-rate envelope r(t). Then, through the extreme theory [4], R( t ) and S ( t ) are Gumbel dstrbuted and the flow can be admtted wth delay bound D at confdence level of Φ(α), f tr t tr t S t D a t a t t D ( ) + ( ) ( + ) + ( ) + ψ ( + ) < 0 0 t T (7) S( t) lm t R( t) + r( t) lmt (8) t Egress Admsson Control has several noteworthy propertes. Frst, snce t employs a measurement-based algorthm, there s a potental for an effcent use of networ bandwdth. Second, t does not requre core nodes to process resource reservaton messages or store any nformaton assocated wth flows. Thrd, t does not assume or requre any specfc schedulng mechansm n the networ and that multple queueng dscplnes can co-exst. Fourth, route pnng, a ey ngredent for determnstc servce, s not fundamentally requred. Ffth, egress routers can perform admsson control on traffc aggregates and do not need to store or montor perflow traffc condtons. Whle the approach s novel and elegant, the scheme s vulnerable to sudden traffc-pattern changes. Snce the technque used to mae admsson decsons s based on statstcal nference through measurements, the scheme wll wor best n an envronment where the networ s stable, the pattern of traffc s relatvely unchangng, the amount of traffc added or subtracted at each flow admsson or termnaton s much smaller than the overall traffc beng carred, and the sze of the networ s large. On the other hand, f ths scheme s appled to a networ composed of few nodes wth hap-hazard traffc patterns, t can result n an unpredctable and unacceptable outcome. Snce the accuracy of ths scheme s closely ted to the networ condton, t would be dffcult to establsh contractual agreements between the user and the servce provder. Furthermore, the scheme wll not wor on the very frst flow on any gven par of edge nodes because t does not have past hstory to construct envelopes for admsson tests. It wll not hold up well when there are sudden changes n the traffc flow such as node and ln falures. A lengthy convergence perod may be observed after sgnfcant changes n the state of the networ occur. Snce the scheme does not explctly de-allocate the resources at flow termnaton, t s dffcult for the networ to dstngush whether a flow has been termnated or the source of a flow s beng slent temporarly. Once a source becomes slent or sends traffc below the declared sustaned rate for a perod of tme, t may need to re-ntate a flow request or send some type of control pacet to restore ts state. In order to do ths, the source must mantan a tmer and the tmer must be set wth some understandng of the behavor of the networ. Ths could add further complexty not only on the networ nodes, but also on the end systems. The scheme does not provde any graceful or ntellgent way to drop pacets when the load exceeds the antcpated lmt. No provsonng for correctng the ntal assessment of traffc n an explct manner s gven ether. It also mplctly assumes that the traffc sources always generate some pacets for the duraton of the reservaton. If a flow s admtted at a certan pea rate but s slent for a long tme, the scheme wll admt other flows durng the slent perod, resultng n overboong, and pacet drops could be observed when the slent source restarts the traffc generaton. E. Admsson Control at End-User Statons There are several versons of end-to-end measurementbased schemes proposed thus far [5] [8]. In ths secton, a system proposed by Karlsson and Ronngren [8] s revewed. The goal of ths end-to-end measurement-admsson scheme s to bound the loss probablty of pacets n hghprorty flows. A host wshng to establsh a low pacet-loss flow probes the networ pror to sendng data. Informaton gathered through probng s used to mae an admsson decson at the source host. Probng s performed as follows: a source host transmts blocs of pacets for a perod of tme at the pea rate of the flow t wshes to establsh. Each pacet contans nformaton regardng the probng, such as probe duraton and transmsson rate. Upon expraton of the probe duraton, the destnaton host returns a pacet, whch contans a measurement report such as the number of probng pacets receved. Based on the measurement report, the source host maes the admsson decson. The proposed servce archtecture employs smple queueng and schedulng mechansms at each node. Data and probng pacets belong to the controlled-load servce and are A SURVEY ON ADMISSION-CONTROL SCHEMES AND SCHEDULING ALGORITHMS 99

10 allocated a certan porton of the ln capacty. Wthn the controlled-load servce, there are two parttons: hghprorty queue and low-prorty queue. Data pacets are queued at the hgh-prorty queue and always servced pror to the low-prorty queue pacets. Probng pacets are queued at the low-prorty queue. All remanng pacets belong to the best-effort traffc and are queued at the besteffort queue. Ths queue s servced only when there are no pacets n the controlled-load servce. The end-to-end measurement-based approach s by far the smplest of all the admsson-control schemes surveyed n ths study. The processng requred by end system for the probng s lght. The queueng and schedulng mechansms necessary at each node are straghtforward, and no flow state needs to be mantaned n the networ. Due to ts smplcty, the scheme s unable to provde sophstcated servces. The proposed scheme can only gve a statstcal bound on the pacet loss; delay s not consdered by the admsson control as t offers no guarantee snce the source maes no requests to the networ and the networ maes no reservatons for the probed flow. Bandwdth blocng could result n a hghly contentous envronment, where probng pacets are generated at a hgh bandwdth rate compared to the remanng bandwdth of controlled-load servce. Suppose there are multple hosts wshng to establsh sessons. Some hosts may request flows at a hgher rate than the remanng bandwdth of controlled-load servce, whle others may wsh to establsh flows at a lower rate. Obvously, the attempts to establsh flows at hgher rates than the avalable bandwdth wll not succeed. The slower rate flows should be accepted, so long as enough bandwdth remans n the controlled-load servce. Under ths type of condton, even the slower rate probng pacets can be affected due to congestons n the controlled-load servce and may not be able to receve the requested servce. In terms of bandwdth-use effcency, t s desrable to eep the probng perod as short as possble. However, a short probng perod may not capture the average state of the networ and may result n overboong or under utlzaton. There s also an uncertanty n the probablty of pacet loss f ths scheme were appled on a networ wthout a route pnnng, yet there was no menton of t n the lterature revewed. Concluson There has been a sustaned nterest and effort n desgnng mechansms to offer guaranteed servces on IP networs. Admsson control s one of the essental tools n supportng QoS. MPLS has receved much attenton as a promsng transport archtecture that could offer servce dfferentatons n large-scale networs. The development of effectve and scalable admsson-control schemes and accompanyng schedulng algorthms sutable for MPLS networs has become an mportant topc of research. In ths paper, the author ntroduced a locaton-based classfcaton of admsson-control schemes, a new taxonomy for admsson-control schemes that complments the tradtonal parameter-based and measurement-based admsson-control categorzaton. In ths new classfcaton system, parameterbased and measurement-based admsson-control schemes are further categorzed nto ) edges-based, ) centrallycontrolled, 3) ngress-based, 4) egress-based, and 5) end-toend-based. The author also surveyed varous admssoncontrol schemes n lght of locaton-based classfcaton systems n order to better understand ther sutablty for MPLS networs. The analyss summary s gven n Table. Table. Summary of analyss of admsson-control schemes based on locaton classfcatons Locaton Admsson Schedulng Pro / Con Edges Parameter Stateful Guarantee / Lmted n scope Central Hybrd Stateful, Stateless, Core- Flexble schedulng / Sngle pont of falure Stateless Ingress Hybrd Core- Stateless Guarantee / May not scale Egress Hybrd Any Mathematcally modeled / Not proven to wor End-to- End Measurement Any Smple / No guarantee Each scheme asserts some level of scalablty. Indeed, there are some novel deas proposed and elegant approaches presented. Yet, every one of them has at least one sgnfcant shortcomng that prevents t from beng deployed over the Internet. Admsson controls that are hybrd, whose control mechansms are placed at an ngress node or central node, have characterstcs that are more promsng for future development than others. In subsequent studes, the author plans to desgn an admsson-control scheme that bulds upon those foundatons, yet exceeds n ts scalablty when compared wth those evaluated n ths study. 00 INTERNATIONAL JOURNAL OF MODERN ENGINEERING VOLUME 0, NUMBER, SPRING/SUMMER 00

11 References [] R. Braden, D. Clar, and S. Shener, RFC 633: Integrated servces n the Internet archtecture: an overvew, Jun [] R. Braden, Ed., L. Zhang, S. Berson, S. Herzog, and S. Jamn, RFC 05: Resource ReSerVaton Protocol (RSVP) verson functonal specfcaton, Sep [3] S. Blae, D. Blac, M. Carlson, E. Daves, Z. Wang, and W. Wess, RFC 475: An archtecture for dfferentated servces, Dec [4] I. Stoca and H. Zhang, Provdng guaranteed servces wthout per flow management, n SIGCOMM, 999, pp [5] E. Rosen, A. Vswanathan, and R. Callon, Multprotocol Label Swtchng Archtecture, RFC 303 (Proposed Standard), Internet Engneerng Tas Force, Jan. 00. [6] J. Mlbrandt, M. Menth, and J. Juner, Improvng experence-based admsson control through traffc type awareness, Journal of Networs, vol., no., pp., Aprl 007. [7] A. K. Pareh and R. G. Gallager, A generalzed processor sharng approach to flow control n ntegrated servces networs: The sngle-node case, IEEE/ACM Transactons on Networng, vol., no. 3, pp , June 993. [8] A. Demers, S. Keshav, and S. Shener, Analyss and smulaton of a far queueng algorthm, n Proc. of ACM SIGCOMM 89, 989, pp. 3. [9] L. Zhang, Vrtual cloc: A new traffc control algorthm for pacet swtchng networs, n Proc. of SIGCOMM 90, 990, pp [0] I. Stoca, S. Shener, and H. Zhang, Core-stateless far queueng: Achevng approxmately far bandwdth allocatons n hgh speed networs, n SIGCOMM, 998, pp [] Z. Zhang, Z. Duan, and Y. Hou, Vrtual tme reference system: A unfyng schedulng framewor for scalable support of guaranteed servces, 000. [] Z.-L. Zhang, Decouplng qos control from core routers: A novel bandwdth broer archtecture for scalable support of guaranteed servces, n SIGCOMM, 000, pp [3] F. Baer, C. Iturralde, F. L. Faucheur, and B. Dave, RFC 375: Aggregaton of rsvp for pv4 and pv6 reservatons, Dec. 00. [4] C. Cetnaya, V. Kanoda, and E. Knghtly, Scalable servces va egress admsson control, IEEE Transacton on Multmeda, vol. 3, no., March 00. [5] W. Almesberger, T. Ferrar, and J. Y, Srp: a scalable resource reservaton protocol for the nternet, 998. [6] F. Kelly, P. Key, and S. Zachary, Dstrbuted admsson control, December 000. [7] G. Banch, F. Borgonovo, A. Capone, L. Fratta, and C. Petrol, Pcp-dv: An end-to-end admsson control mechansm for p telephony, n Tyrrhenan IWDC 00 Evolutonary Trends of the Internet, Taormna, Italy, September 00. [8] V. Ele, G. Karlsson, and R. Ronngren, Admsson control based on end-to-end measurements, n IN- FOCOM (), 000, pp [9] R. L. Cruz, Qualty of servce guarantees n vrtual crcut swtched networs, IEEE Journal on Selected Areas n Communcatons, vol. 3, no. 6, pp , 995. [0] H. Zhang and D. Ferrar, Rate-controlled servce dscplnes, 994. [] R. Svaumar, T. eun Km, N. Ventaraman, J.-R. L, and V. Bharghavan, Achevng per-flow weghted rate farness n a core stateless networ, n Internatonal Conference on Dstrbuted Computng Systems, 000, pp [] E. Castllo, Extreme Value Theory n Engneerng. New Yor: Academc, 988. [3] D. L. Mlls, RFC 305: Networ tme protocol (verson 3) specfcaton, mplementaton, Mar. 99. [4] J. Schlembach, A. Soe, P. Yuan, and E. W. Knghtly, Desgn and mplementaton of scalable admsson control, n QoS-IP, 00, pp. 6. Bography MASARU OKUDA receved the B.S. degree n Informaton System and Computer Scence from Brgham Young Unversty - Hawa, Lae, HI, n 989, and the M.S. degree n Telecommuncatons and the Ph.D. degree n Informaton Scences from the Unversty of Pttsburgh, Pttsburgh, PA, n 996 and 006 respectvely. Currently, he s an assstant professor of Telecommuncatons Systems Management at Murray State Unversty, Murray, KY. Hs teachng and research areas nclude computer and networ securty, US telecom polces, networ protocol analyss, networ archtecture desgn, QoS enabled networs, peer-to-peer networs, and vdeo dstrbuton networs. Dr. Ouda may be reached at masaru.ouda@murraystate.edu. A SURVEY ON ADMISSION-CONTROL SCHEMES AND SCHEDULING ALGORITHMS 0