1658 JOURNAL OF LIGHTWAVE TECHNOLOGY, VOL. 18, NO. 12, DECEMBER 2000Fig. 1. Generic wavelength-routed passive optical network (WR-PON) architecture. Each node has a single tunable transmitter (TT), and a single fixed receiver(FR). All filtering and optical routing is performed by passive optical elements within the network.tuning times of less than 10 ns between any of 40 channels  presented previously in . The present paper extends this pre-have also been reported. vious work, including a more comprehensive description of the Therefore, WDM LANs based on WR-PONs and requiring design and performance issues, more detailed results and com-only tunable transmitters represent a very promising approach to parisons with previously proposed applicable protocols.meeting future high-capacity data network requirements. To re- The structure of this paper is as follows. In Section II, weduce the complexity and cost of the node hardware further, MAC provide an overview of the operation of WRAP, and describeprotocols which do not require any additional control channel or the WDM network architectures to which the protocol is appli-carrier-sense capability are highly desirable. cable. Simulations of the protocol demonstrating the impact of In this paper, we consider the design and performance of a different allocation algorithms are described in Section III. InMAC protocol, called simply the WDM Request/Allocation Section IV we present simulation results demonstrating the per-Protocol (WRAP), applicable to WR-PONs with fast-tuning formance of WRAP as a function of the network and protocoltransmitters and fixed receivers. Our protocol does not re- parameters, and compare it with alternative applicable proto-quire an additional control channel, performing all necessary cols.signaling in-band. Unlike the previously-proposed FatMACprotocol , which also uses in-band signaling, WRAP II. OVERVIEW OF THE PROTOCOL AND NETWORK TOPOLOGIESdoes not require an optical broadcast facility, and is therefore Fig. 1 illustrates schematically the “generic” wave-applicable to systems using DBR or GCSR tunable semicon- length-routed PON considered in this paper. Each node hasductor lasers which can emit only one wavelength at a time. a single tunable transmitter (TT) and a single fixed opticalFurthermore, WRAP delivers similar performance to the ideal receiver (FR). Note that no optical filtering is employed atinterleaved time-division multiplexed (I-TDMA) scheme  the receivers—all necessary “filtering” is inherently providedunder heavy, uniform loading, while offering superior perfor- by the wavelength-routing functionality of the PON. Lightmance and capacity sharing under light and/or uneven loads. entering the PON at an input port is routed by passive opticalWRAP is robust in the presence of significant propagation components within the network to one and only one outputdelays in transmission, and in many situations shows improved port. The particular output port selected may depend onlyperformance in the presence of high delay-bandwidth product upon the input wavelength, or it may depend upon both theas a result of inherent pipelining within the network. wavelength and the input port. In either case, a transmitting The initial motivation for the development of WRAP was to node has the ability to select a specific destination node byprovide a MAC protocol for use in the MAWSON WDM net- tuning its transmitter to the corresponding wavelength. Thework demonstrator currently being developed within the Aus- network has no “master” or “controller” nodes, and allocationtralian Photonics Cooperative Research Centre at the University of network resources is distributed amongst the nodes. Theof Melbourne . MAWSON is a ring-based network, however lack of centralized control reduces the system complexityas we discuss in Section II, WRAP is also applicable to other and eliminates the possibility of there being a single criticalwavelength-routed topologies such as the AWG-based star. This congestion point, thus making the network simpler to installis in contrast to most previously proposed protocols for WDM and maintain.LANs which are applicable only to star topologies. The basic The MAC protocol described in this paper operates in adesign of WRAP, and initial performance simulations have been slotted mode to maximize efficiency. Each slot is divided into
SPENCER AND SUMMERFIELD: MAC PROTOCOL FOR WAVELENGTH-ROUTED PONS 1659Fig. 2. WR-PON queuing model.a number of shorter transmission periods known as minislots, followed immediately by the transmission of the correspondingsome of which are used for signaling, and others for data data (e.g., ); and those in which the reservations must betransmission, as described below. The protocol is applicable to received by all nodes before implicit slot allocations are madea range of wavelength-routed optical network topologies, the based on a common algorithm (e.g., , ). WRAP isonly requirement being that a mechanism for synchronizing based on a three-stage scheme in which a node wishing to sendtransmissions at the slot boundaries must be provided. The data first requests an allocation from the destination node. Themethod of synchronization, and the effective point of synchro- destination receives these requests from all source nodes, andnization, depends upon the details of the physical topology. For then sends back specific allocations to each source (hence thesome topologies this may have significant implementation and name WDM Request/Allocation Protocol). Finally, the sourceperformance implications, as we show later when discussing nodes send their data using the allocated slots.the application of WRAP in a ring topology. In this section we The WRAP scheme has a number of advantages over previ-first present an overview of the operation of WRAP, followed ously-proposed protocols:by further details of its implementation in star and ring topolo- • It does not require reservation requests from all sourcegies. This paper is concerned only with the performance of the nodes to be distributed globally within the network, norprotocol in the simplest case of a single-hop topology in which does it require a centralized scheduler to handle requests.the number of nodes is equal to the number of wavelengths This means that it does not require a shared control , i.e., . Other, more complex, implementations are channel, or a broadcast facility.possible, including multihop topologies, and variations on the • It is inherently fair, under any well-defined fairness crite-protocol in which wavelengths may be shared (i.e., ); rion, since access to any specified destination can be reg-however, these are beyond the scope of the present paper. ulated completely by the allocation algorithm used by the destination node.A. Queuing Model and Protocol Description • With a suitable choice of allocation algorithm, perfor- In the absence of any carrier-sense capability there are two mance can be made equivalent to that of I-TDMA underbasic classes of MAC protocol: those that allow collisions such conditions of uniform heavy loading.as ALOHA and slotted ALOHA ; and those that avoid all The queuing model of the network is shown in Fig. 2. Eachcollisions by the careful scheduling of transmissions. Protocols node maintains a set of queues, one for each of the dis-which allow collisions are inherently inefficient and suffer tinct destinations. The tunable transmitter at node can accessfrom poor throughput as a result of the high levels of data loss the set of wavelengths which, in the present work, is as-which occur under moderate to heavy loading . A simple sumed to be sufficient to reach all other nodes in the network.scheme based on the static preallocation of all data transmission The transmitter can only tune to one wavelength in at aslots such as the I-TDMA protocol  provides a fair and time. The receiver at node receives signals on the set of wave-optimal use of network resources for uniform network loading lengths but makes no distinction between these wave-and does not suffer any data collisions. Thus it is desirable lengths. If two wavelengths in are active simultaneously,as the uniformly loaded limit for any collision-free protocol. or the same wavelength is active via two sources simultaneously,Dynamic sharing of transmission resources requires the use of a collision occurs at node . The objective of the WRAP protocolsome kind of reservation scheme, usually based on a separate is to avoid all such collisions.control channel for exchanging the reservation information. To implement WRAP, time is divided into a series of uni-Most previously proposed reservation-based protocols fall into form length slots, of duration . The slot structure is shownone of two basic categories: those in which the reservations are in Fig. 3(a). The nodes are synchronized so that the start of a
1660 JOURNAL OF LIGHTWAVE TECHNOLOGY, VOL. 18, NO. 12, DECEMBER 2000 N N0 MFig. 3. WDM Request/Allocation Protocol (WRAP) slot structure and WDM transmission schedule. (a) Slot structure as viewed from a single transmitter or 1 Request/Allocation (R/A) minislots, and data minislots. Transmitted minislots are t treceiver in an -node network consists of a synchronization minislot,separated by a guard time , and the slot period is . (b) Sample multichannel transmission schedule in a three-node network. Each row represents transmissionsto a single destination node; the numbers in the R/A minslots represent the source node; allocation of data minislots to source nodes is determined by the WRAPprotocol.slot occurs at the same time on all the wavelengths in the net- resynchronization. They may also be the result of small vari-work at the synchronization point. As described later, the syn- ations in propagation delays caused by the effects of dispersionchronization points are at the network hub for a star topology on different wavelengths and by temperature changes and otherand at the WDM ADMs for a ring topology. As mentioned pre- environmental fluctuations in the fibers. The framing overheadviously, WRAP is based on a three-stage bandwidth allocation enables nodes to perform clock and data recovery on each in-scheme. In the first stage source nodes with data to transmit send coming minislot. This is needed for correct reception of the datarequests to the destination nodes. In the second stage the desti- contained within the minislot. The framing is also needed tonation nodes independently make allocations of data minislots byte-align and delineate the data within each minislot.and send them to the source nodes. In the final stage the source The synchronization minislot is followed by re-nodes transmit the data to the destinations nodes in the allocated quest/allocation (R/A) minislots and data minislots. Thedata minislots. In the simplest and most efficient implementa- R/A minislots are preallocated on each wavelength in antion, these three procedures should occur in consecutive slots. It interleaved TDMA manner in such a way that:is thus apparent that if the slot period cannot be too short oth-erwise, for example, requests will not arrive in time for alloca- • each source node is allocated one R/A minislot to eachtions to be sent out in the next slot. This implies that, regardless destination node in each slot; andof the topology, there is a minimum slot length. As we shall see, • no source node is required to transmit R/A minislots tothe ring topology results in even more severe constraints on the two different destinations at the same time.allowable slot length. The constraints on the slot length are dis- This is illustrated for the simple case of a three-node networkcussed in more detail in the following sections. in Fig. 3(b). The length of the data minislots is a selectable pa- As shown in Fig. 3(a), each slot is further divided into a se- rameter of the protocol, and is denoted by (bits).ries of minislots, whose lengths depend on their functions. The Once per slot, a node examines the state of its queues, and de-synchronization minislot is used to maintain slot synchroniza- cides how many data minislots to request, up to the maximumtion between nodes. The length of the synchronization minislot of per destination. Thus, a request requires at mostis bits. A practical implementation also requires a guard bits. Each node also examines the requests received from alltime, , between minislots and an additional framing overhead other nodes during the previous slot, and makes specific, col-of bits per minislot. The guard time accommodates the laser lision-free allocations of data minislots which are effective fortuning time as well as any timing variations. These timing vari- the next slot. A number of different allocation algorithms areations may be caused by a drift in the node clocks between possible, as discussed in Section III. Allocations therefore re-
SPENCER AND SUMMERFIELD: MAC PROTOCOL FOR WAVELENGTH-ROUTED PONS 1661quire bits (one per data minislot). The minimum number ofbits required in each R/A minislot (excluding overheads) is (1) Finally, each node examines the allocations it received duringthe previous slot (based on requests it made in the slot prior tothat), and transmits data from each of its queues during the dataminislot periods allocated to it by the corresponding destinationnodes. Note that although there is a minimum three-slot delay be-tween the arrival of data in a queue, and its delivery to the des-tination, under heavier loading (when there is usually data inthe queues), the protocol is expected to remain efficient dueto the effects of pipelining, in which requests, allocations anddata transmissions take place continually over long sequences ofcontiguous slots. Furthermore, as we will demonstrate, in prac-tical implementations the slot length is often comparable to theunavoidable propagation delays within the network, implyingthat the built-in protocol delay is of the same order as the prop-agation time. For an -node, -wavelength network, the total theoretical Fig. 4. Four-node wavelength-routed ring network. WDM add–dropdata transmission capacity of a slot in bits is given by multiplexers (ADMs) drop one wavelength out of the ring at each access point, while allowing any wavelength to be added. the length of a slot is constrained to be a submultiple of the (2) ring round trip propagation delay, , to ensure simple node synchronization, high utilization and collision free operationThe first expression gives an upper bound on the transmission . As described in the previous section, it is necessary forcapacity obtained by subtracting the protocol and network over- the nodes to be able to send requests, return allocations andheads for the total slot capacity, while the second expression is transmit data in consecutive slots. This is possible so long asthe actual capacity achieved given a particular choice of data 1) the slot period is not less than the ring round trip propagationminislot length . Differences between these two quantities delay, and 2) the largest distance from a node to an access pointmay arise in topologies in which is constrained, resulting is less than half the distance around the ring. Thus the onlyin a period of “dead time” at the end of each slot. Equation (2) viable choice for a slot length is , i.e.describes the fixed relationships between slot length, number (3)of nodes, and number of data minislots. There are tradeoffsamongst these parameters, which are investigated in more detail Each node must transmit its minislots at such a time that theythrough simulations in Section IV. Briefly, protocol overheads arrive at the ADM and are inserted into the ring at exactlyare more significant for small slots and/or large and , while the right time (to within the accuracy allowed by the guardthe choice of the number of data minislots is a tradeoff between times) to avoid collisions with adjacent minislots on the samethe impact of protocol overheads for large and the wasted wavelength. This requires that each node be aware of all thecapacity of small if the traffic patterns lead to many partially propagation delays within the network, and maintain an internalfilled data minislots. Furthermore, as we discuss later, the imple- timer to control transmissions. Global synchronization of thementation of an efficient allocation algorithm may place further node timers can be maintained using the synchronization min-constraints on . islot provided at the beginning of each slot. We have developed detailed cold- and warm-start procedures for establishing andB. Ring Topology maintaining synchronization which have been verified by sim- The basic ring topology is illustrated for the case of a ulation, and which will be implemented within the MAWSONfour-node network in Fig. 4. It consists of a fiber optic ring with demonstrator. A full description of these procedures is beyonda WDM add/drop multiplexer (ADM) located at each access the scope of the present paper.point. In the simplest configuration each ADM drops a single, The constraint [see (3)] on the slot length in the ring topologyfixed wavelength that is different from those dropped at other implies that the maximum available slot capacity is dependentaccess points. The ADM also allows any one wavelength, or upon the ring length, which is not a dynamically-selectable pa-any number of wavelengths simultaneously, to be added. This is rameter. However, the approximate length can be selected whenthe architecture of the MAWSON network demonstration being the network is installed, and the ring length may be increaseddeveloped by the Australian Photonics Cooperative Research by the addition of spooled fiber if the geographical size is foundCentre at the University of Melbourne. In the ring topology, to be too small. Equation (2) shows that the maximum capacity
1662 JOURNAL OF LIGHTWAVE TECHNOLOGY, VOL. 18, NO. 12, DECEMBER 2000Fig. 5. Four-node wavelength routed star network based on an arrayed waveguide grating (AWG) hub router. Signals are routed uniquely from source to destinationbased on their wavelength, and the port at which they are input to the AWG.per slot increases linearly with the slot length, and decreases within the accuracy allowed by the guard times) to avoid colli-approximately linearly with the number of nodes and/or data sions with adjacent minislots on different wavelengths, destinedminislots. For practical values of the network and protocol over- for the same output port. In this way, transmissions on all wave-heads, a rule-of-thumb which can be derived from (2) is that at lengths are fully synchronized at the hub, as shown in Fig. 3(b).1 Gb/s data rates, a maximum capacity of greater than 90% (i.e., This requires that each node be aware of the propagation delay900 Mb/s per wavelength) can be maintained if the length of the between itself and the hub within the network. As in the casering is greater than m (i.e., 1 km per node). Note that of the ring topology, global synchronization of the node timersthe maximum capacity tends toward 100% as the physical net- can be maintained using the synchronization minislot providedwork size or the bit-rate increases, the key parameter being the at the beginning of each slot. A technique for determining thedata-storage capacity of the ring itself. node-hub propagation delay, and achieving global synchroniza- tion, known as Distributed Slot Synchronization (DSS) has beenC. Star Topology developed for use in the CORD network testbed at Stanford, which is based on a broadcast star . This technique may be The basic star topology is illustrated for a four-node network readily adapted for use in the wavelength-routed star topology,in Fig. 5. It consists of a passive optical wavelength-routing hub although a full description of the scheme is beyond the scope of(which may be implemented using a single AWG, or equiva- the present paper.lent structure) to which the nodes are connected by fiber pairs. Once again it is necessary for the nodes to be able to send re-Light entering the hub on any given wavelength and input fiber quests, return allocations and transmit data in consecutive slots.is routed to a specific corresponding output port. The routing For the star topology this implies that the slot length must befunction is such that light on the same wavelength but arriving more than twice the longest propagation delay between the hubfrom different input ports must always be routed to different and a node. Other than this requirement, is a freely-se-output ports. Under this condition, the routing function is unique lectable parameter of the protocol, along with the number of dataand introduces no optical loss in principle, although any prac- minislots, . These parameters may be established by negoti-tical implementation of the passive hub will exhibit some finite ation among the nodes in the network. Analysis and simulationinsertion loss. (For AWG-type devices, the insertion loss is typ- of the network performance is therefore important to assist inically 3–6 dB, and is independent of the number of ports .) the development of formulae and/or heuristics that enable theseFig. 5 illustrates the specific routing properties of a four-port parameters to be optimized during network initialization.AWG. Entering the hub from Node 1, light of wavelengthis routed back to Node 1; wavelength is routed to Node 2; III. INVESTIGATION OF ALLOCATION ALGORITHMSand so on. However, wavelength is routed to Node 2 whenentering the hub from Node 2; to Node 3 when entering from A node must consider the amount of transmission bandwidthNode 3; and so on. The routing permutations follow the same requested by other nodes once per slot and make appropriatepattern for the other wavelengths. allocations to the requesting nodes. The goal of the algorithms With a star topology a central synchronization point at the considered here is to make this allocation fair, in the sense that,hub of the star is used. Each node transmits its minislots at such all else being equal, a receiving node will on average allocatea time that they arrive at the hub at exactly the right time (to bandwidth to requesting nodes in proportion to their demands,
SPENCER AND SUMMERFIELD: MAC PROTOCOL FOR WAVELENGTH-ROUTED PONS 1663while providing some guaranteed bandwidth to all nodes. The TABLE Inode must also consider the independent allocations received SYMMETRIC TRAFFIC LOAD DISTRIBUTION FOR SIMULATIONS OF A THREE-NODE NETWORKfrom other nodes, as they may clash, and make a choice of whichone to use for each data minislot. Three allocation algorithmshave been developed and have been compared using simulationsof a 3-node symmetrically loaded network. The simulations were performed using event driven simu-lation code written in C. A slot period of 20 s, a data rate of1 Gb/s, and four data minislots per slotwere assumed. Finite buffers of 40 data minislots per nodeper destination were used. Realistic protocol overheads wereaccounted for in the simulation as follows. A conservativeguard-time of ns was assumed, which allows for therelative drift of node clocks between synchronization events,as well as differences in propagation delays on different wave- access to the network capacity. Furthermore, under symmetriclengths due to fiber dispersion and environmental fluctuations. heavy loading, the cyclic allocation algorithm may result in re-The guard time thus contributes 100 bits overhead per minislot. ceiving nodes making allocations which are identical to the slotThe total of all framing and protocol overheads, each of which allocations in the I-TDMA protocol, which is known to be op-is rounded to the next 8-bit multiple (i.e., assuming a byte timal under these loading conditions.is the minimum transmission unit), is 996 bits per slot. This Fig. 6(a)–(d) shows the distribution of delays experienced bycomprises 124 bits for the synchronization minislot (100 for the data arriving at the source nodes obtained from a number ofguard time, 16 for the framing and 8 for the protocol as a slot simulation runs at total offered loads of 50%, 65%, 80%, andidentifier), 156 bits for each R/A minislot (100 for the guard 125%, respectively. At offered loads less than the network ca-time, 16 for framing and 40 for protocol, including 8 each for pacity most data is queued only as a result of the protocol delay,the requests and allocations) and 140 bits for each data minislot thus resulting in short queue lengths, as shown in Fig. 6(a). The(100 for the guard time, 16 for framing and 24 for protocol variability of the data arrivals sometimes results in more instan-overheads). Inclusion of the overheads results in an effective taneous load than the network can handle, and this results in thereduction of capacity, as described by (2). Data arrivals to the excess being queued. As the offered load is increased the queuenetwork were modeled as packets of uniformly distributed lengths increase, as shown in Fig. 6(b) but they are limited bylength of 0 to 16 000 bits arriving under a Poisson process. the buffer capacity. When the buffers are full, further data ar-The network loading was varied from 50 to 150% by changing rivals result in loss. Thus, under overloaded conditions the queuethe mean inter-arrival time appropriately. Network overload lengths approach the buffer capacity. The average rate at which(greater than 100% loading) was considered to enable us to the queues are served (i.e., the average number of data minis-investigate the limiting behavior of the protocol, however we lots transmitted per second) is proportional to the throughput.recognize that real networks are unlikely to be operated in this Data arriving at the queue must wait for all the data already inregime due to the substantial data losses incurred. Simulation the queue to be served before it will be sent. Thus the meanruns of 100 000 slot periods (2 s) were repeated eight times at queuing delay experienced by data is proportional to the meaneach of these values of loading. queue length and inversely proportional to the mean throughput. The network was loaded symmetrically as shown in Table I, Also, the delay is bounded, as the buffers are of finite lengthwith 50% of the traffic at each node sent to and received from the and there is a nonzero worst case network throughput. Protocolother two nodes. Unity (100%) loading of a source/destination and transmission delays are incurred in addition to the queuingroute corresponds to a mean arrival rate equal to the raw capacity delay, however pipelining results in the removal of most of theof a wavelength neglecting overheads, i.e., 1 Gb/s. protocol delay for heavy loading conditions. Thus for the net- work conditions simulated the buffer size of 40 minislots and worst case throughput of one minislot per slot bound the delayA. Cyclic Allocation to just above 40 slots. This bound on the delay is clearly visible In the cyclic allocation algorithm, each receiving node con- in Fig. 6(b) and (c).siders the requests from transmitting nodes in a round robin The cyclic allocation algorithm is simple, and allows eachfashion . Individual data minislots are allocated sequen- node to make allocations independently of its location in thetially, to nodes that have made requests until either there are network. However, Fig. 6(c) and (d) demonstrate an undesir-no more requests or no more unallocated minislots. This cyclic able property of the algorithm which only became apparentround-robin allocation continues in the next slot, recommencing during these simulation runs. The graphs clearly show thefrom the node following the last one considered in the previous development of a bimodal delay distribution under conditionsslot. When a transmitting node receives more than one alloca- of moderate-to-heavy loading. The observed results are duetion for the same data minislot from different destination nodes, to an undesirable synchronization and lock-up property ofthe conflict is resolved in a similar way, i.e., by cyclic allocation the cyclic allocation algorithm. Fig. 6(d) shows that underof such minislots in a round-robin fashion. When executed iden- very heavy load, there is a tendency for the protocol totically at every node, these algorithms result in fair and flexible fall into one of two possible states, indicated by the gray
1664 JOURNAL OF LIGHTWAVE TECHNOLOGY, VOL. 18, NO. 12, DECEMBER 2000 (a) (b) (c) (d)Fig. 6. Distribution of packet delays for the cyclic allocation algorithm in a three-node network under symmetric loading, for average network loads of (a) 50%,(b) 65%, (c) 80%, and (d) 125%.histogram (State 1) and the black histogram (State 2) respec- Fig. 6(c), the protocol falls in and out of the lock-up states intively. This behavior may be explained as follows. As the response to short-term peaks in demand, and the observed delaynetwork loading increases, the amount of data in the queues distribution is the time-averaged effect of the protocol spendingwill tend to exceed the maximum capacity of a slot. This periods of time in both states.will result in more outstanding requests every slot than there Fig. 7 shows the average delay and packet loss under theare available data minislots, as a result of which the source cyclic allocation algorithm as a function of the mean networknodes will consistently make identical allocations to the des- loading. The key feature of these results is the “knee” in thetination nodes. The cyclic round robin allocation method is delay characteristic, which occurs at an average load of aboutcompletely deterministic, and generates the same allocations 60%. This is preceded by a dramatic rise in the packet loss. Thefor as long as this overloaded state exists. State 1, which minimum delay under light loads is 110 s (5.5 slots). The delayhas lower mean delay, corresponds to the fortuitous circum- results under heavy loading conditions have been averaged overstance in which the allocations received back at each source the length of the simulation runs, and they represent the averagenode from the two destination nodes do not clash, i.e., both effect of both states. Delay in State 1 limits to 19.1 slots, whiledestinations have allocated different data minislots, and the State 2 has an average delay of 30.4 slots (a result of the bimodaltransmitter is able to use all allocations. State 2, with higher distribution with peaks near 20 and 40 slots delay).mean delay, corresponds to a situation in which many of theallocations clash, and the source node is unable to use both B. Random Allocationsince it can only transmit to one destination at a time. In the One strategy for avoiding the undesirable synchronization of3-node network, this state corresponds to a drop in available the allocations which occurs with the cyclic algorithm is to allo-capacity to 2/3 of the maximum capacity ( in the case cate the minislots randomly instead of sequentially. Similarly, ifof an -node network). It therefore corresponds to a form of more than one allocation is received at the source node from twocongestion collapse for the cyclic allocation algorithm, and (or more) destination nodes, then a random selection is made.must be avoided. This algorithm is simple, it obviously avoids any deterministic Under overloaded conditions, the protocol tends to fall into lock-up states, and it retains the property that allocations canone state or the other and stay there, and the two histograms in be made independently of a node’s location within the network.Fig. 6(d) show the results of separate simulation runs in which However, it clearly does not limit to the ideal I-TDMA scheduleeach state was observed. Under moderate loading, shown in under uniform heavy loading.
SPENCER AND SUMMERFIELD: MAC PROTOCOL FOR WAVELENGTH-ROUTED PONS 1665Fig. 7. Average delay and packet loss as a function of offered load for the cyclic allocation algorithm in a three-node network under symmetric loading. (a) (b) (c) (d)Fig. 8. Distribution of packet delays for the random allocation algorithm in a three-node network under symmetric loading, for average network loads of (a) 50%,(b) 65%, (c) 80%, and (d) 125%. Fig. 8(a)–(d) shows the delay distributions for the random al- lots per slot) as a result of the random allocation clashes and thuslocation algorithm obtained from a number of simulation runs the delay limits to an average value of 520 s (26 slots—slightlyat total offered loads of 50%, 65%, 80%, and 125%, respec- less than the theoretical 26-2/3 as the buffers never fill com-tively. The queues evolve in a similar manner to the cyclic allo- pletely). The evolution to this condition is clearly visible incation algorithm, however in this case the bimodal distribution Fig. 8(c) and (d). Fig. 9 shows average delay and packet loss ofdoes not emerge. As a result, the upper tail of the distribution the random allocation algorithm as a function of mean networkat 65% loading shown in Fig. 8(b) is less heavy than the corre- loading. The knee of the delay curve occurs at a slightly highersponding distribution for the cyclic allocation algorithm. Under loading of about 65% compared with the cyclic allocation algo-heavy loading the peak average capacity is 3/4 (i.e., 1-1/2 minis- rithm, and there is a corresponding improvement in packet loss.
1666 JOURNAL OF LIGHTWAVE TECHNOLOGY, VOL. 18, NO. 12, DECEMBER 2000Fig. 9. Average delay and packet loss as a function of offered load for the random allocation algorithm in a three-node network under symmetric loading.C. Preferential/Random Allocation minslot will be preferentially allocated. For example, Node 0 will preferentially allocate the first data minislot to Node The cyclic allocation algorithm was designed to converge to 3, the second to Node 2, the third to Node 1, and so on.the I-TDMA schedule under uniform heavy loading, but was The set of preferential allocations actually made is shown infound to suffer from undesirable lock-up states. The random al- Fig. 10(c)—these allocations are deterministic, and are derivedlocation algorithm avoids these lock-up states, resulting in im- directly from Fig. 10(a) and (b). Each node then makes randomproved average performance, but still underutilizes the available allocations of any unallocated data minislots to nodes withbandwidth under heavy loading. The ideal allocation algorithm additional requests. One such random allocation is shown inwould avoid the lock-up states and converge to the I-TDMA Fig. 10(d). At Node 0, Data Minislots 0 and 3 are not neededschedule under uniform heavy loading. We have developed such for preferential allocations, and in this example, Data Minislotan algorithm, which we call the preferential/random (P/R) algo- 0 has been randomly allocated to Node 2. Likewise, at Noderithm. Under this scheme the fixed I-TDMA schedule is consid- 1, Node 0 has been randomly allocated two minislots chosenered to be the preferred allocation at the receiving nodes. When from Data Minislots 2, 4, and 5. Fig. 10(e) summarizes therequests arrive from the source nodes, allocations are made ini- allocations received by each node, with a cross representingtially from the I-TDMA schedule. Once this has been done, if an allocation. An example of the allocations which maythere are further requests from some source nodes and there are subsequently be used, after resolving any clashes, is given inremaining free data minislots, then these are allocated by con- Fig. 10(f). At Node 0, Data Minislot 1, the random allocationsidering the requests from each source node in a round-robin from Node 3 clashes with the preferential allocation from Nodemanner (to ensure that the allocation of the remaining capacity 2, thus the preferential allocation is used. The clash whichis fair on average), and assigning to each source one of the re- occurs at Node 0, Data Minislot 4, is between two randommaining data minislots chosen at random (to avoid any deter- allocations (from Nodes 1 and 3). One of these is chosen atministic lock-up effects). Similarly, if a source node receives al- random—in this example it is the allocation from Node 1.locations for the same data minislot from two (or more) different Similar choices apply to all the other allocation clashes present.destinations, if one of these is the preferred I-TDMA allocation, The main disadvantage of the P/R algorithm is that it requiresthat one is used. If all are the result of random allocations, one each node to “know” its place in the I-TDMA schedule, and thusis chosen at random. it is not independent of the node’s location in the network. In An example of the P/R allocation scheme for a 4-node practice, the labeling scheme used for nodes and I-TDMA slots network with six data minislots is given in is arbitrary (i.e., the physical ordering of nodes in the network isFig. 10. Fig. 10(a) shows a table of all the requests made for the not significant), and can easily be assigned as part of the cold-slot under consideration. The numbers in the table indicate the and warm-start procedures which are required in any case. Fornumber of data minislots requested for transmission by each the preferred I-TDMA schedule to map fairly into each slot, thesource node to each destination node. For example, Node 0 is additional constraint is imposed that , whererequesting four minislots to Node 1, one minislot to Node 2, is any positive integer.and five minislots to Node 3. With , if Node 0 receives The effect of the P/R algorithm is that each node is guaran-allocations for all 10 requests, it will obviously be able to teed at least of its transmission capacity to everyuse six at most, however in the absence of any information other node. A node may be able to obtain up to the entire avail-regarding the requests made by other nodes, the R/A protocol able capacity, depending on the requests made by other nodes.considers all requests independently. The preferential source The access to this capacity is shared in a statistically fair waynodes for the data minislots at each of the destination nodes among those requesting it. Under heavy uniform loading the net-are shown in Fig. 10(b), in which the number in each box work limits to the optimal fixed TDMA schedule and does notrepresents the node number to which the corresponding data suffer the instabilities of the cyclic algorithm. It can be shown,
SPENCER AND SUMMERFIELD: MAC PROTOCOL FOR WAVELENGTH-ROUTED PONS 1667 N MFig. 10. Example of the P/R allocation scheme for a 4-node ( = 4) network with six data minislots ( = 6). (a) Requests made to the nodes. (b) Preferentialsources for the data minislots at each node. (c) Preferential allocations made. (d) Possible random allocations made. (e) Summary of the allocations received byeach node. (f) Actual allocations used.although space does not permit its inclusion in this paper, that Simulation of the P/R allocation algorithm resulted in theall three allocation algorithms produce the same number of allo- delay distributions shown in Fig. 11(a)–(d). The guaranteed ca-cations, the difference between them being in the data minislots pacity of two minislots per slot and the buffer size of 40 min-chosen. Furthermore, allocation clash resolution by the P/R al- islots per node per destination ensure that the worst case delaygorithm can be shown to result in a mean performance that is is limited to just above 20 slots. The protocol delay is alwaysalways better than the random algorithm’s, and that under cer- present, but is masked by pipelining at high loading.tain loads (including uniform) the P/R algorithm achieves the Fig. 12 shows the average delay and packet loss under themaximum possible minislot usage. P/R allocation algorithm as a function of the mean network
1668 JOURNAL OF LIGHTWAVE TECHNOLOGY, VOL. 18, NO. 12, DECEMBER 2000 (a) (b) (c) (d)Fig. 11. Distribution of packet delays for the P/R allocation algorithm in a three-node network under symmetric loading, for average network loads of (a) 50%,(b) 65%, (c) 80%, and (d) 125%.Fig. 12. Average delay and packet loss as a function of offered load for the P/R allocation algorithm in a three-node network under symmetric loading.loading. The maximum average delay under overload conditions considered. Therefore, in the remainder of the paper only theis 360 s (18 slots), comparable to the desirable behavior of the P/R algorithm is considered.cyclic scheme. The knee of the delay characteristic occurs at amean network loading level of 80%. At high loading there is full IV. PERFORMANCE SIMULATIONSutilization of the available network capacity under the I-TDMAschedule. A. Scaling Performance of the Protocol The results in this section illustrate the clear superiority of By choosing a value for the slot period in conjunctionthe P/R allocation algorithm. It is stable under heavy loading with the number of nodes a tradeoff can be made betweenand overload conditions, and exhibits the lowest delay, delay the minimum delay at low loads, and the cost of the per-slotvariance and highest average throughput of the three algorithms overheads. As discussed previously, in practice, is a free
SPENCER AND SUMMERFIELD: MAC PROTOCOL FOR WAVELENGTH-ROUTED PONS 1669Fig. 13. Average packet delay for the WRAP protocol under symmetric loading, as a function of total offered network load, for a channel capacity of B = 1Gb/s per wavelength and a fixed slot duration of 15 s. Results are shown for three-node (squares), five-node (diamonds), seven-node (triangles), and nine-node(circles) networks.Fig. 14. Total throughput for the WRAP protocol under symmetric loading, as a function of total offered network load, for a channel capacity of B = 1 Gb/s perwavelength and a fixed slot duration of 15 s. Results are shown for three-node (squares), five-node (diamonds), seven-node (triangles), and nine-node (circles)networks.parameter in star networks but may be constrained by the the overall network capacity grows as but the number offiber length in ring networks. Simulations were carried out source/destination pair routes grows as . These obser-for , 5, 7, and 9 and , 25, 35, and 45 s vations are confirmed in the throughput results in Fig. 14, whichwith symmetric loading to provide insight into this tradeoff. show clearly that the fractional utilization of the raw bandwidthSymmetric loading was used as it fully utilizes the network is reduced for higher numbers of nodes. In this sense, the uni-capacity under high loads, and is therefore a case of particular form symmetric loading example represents a worst-case for theinterest in which the results are relatively easy to interpret. protocol.The number of minislots was chosen to be The relative effect of the increasing overheads with increasingand the total buffer capacity was fixed at 1 Mbyte per node, number of nodes can be mitigated by making correspondingdivided equally amongst all possible destinations. All network changes to the slot period , at the expense of some addi-and protocol overheads were again included, using the same tional delay under light loading. The effect of changing both theparameters discussed previously. slot length and the number of nodes, while keeping them propor- In the first set of simulations the number of nodes was in- tional was simulated. The delay as a function of the total offeredcreased while keeping constant at 15 s. The average delay load to the network is shown in Fig. 15, and the correspondingas a function of the total offered load to the network is shown throughput is shown in Fig. 16. The overheads remain at aboutin Fig. 13, and the corresponding total throughput is shown in the same fraction of the network capacity and as a result theFig. 14. The worsening delay performance, with the knee of the knees of the delay curves occur at roughly the same fractionalcurve occurring at a relatively small fraction of the total net- loading in all cases. Similarly, the fractional utilization underwork capacity as the number of nodes increases, is due mainly overload conditions is approximately the same in all cases. Theto the effect of increasing overheads. The increase in the max- slight increase in delay under light loading is difficult to see inimum delay under overload is due to the reduced capacity be- Fig. 15, as it is small compared to the queuing delays whichtween each source/destination pair, resulting from the fact that occur under moderate to heavy loads.
1670 JOURNAL OF LIGHTWAVE TECHNOLOGY, VOL. 18, NO. 12, DECEMBER 2000Fig. 15. Average packet delay for the WRAP protocol under symmetric loading, as a function of total offered network load, for a channel capacity of B = 1Gb/s per wavelength and a slot duration which is increased in proportion to the number of nodes. Results are shown for three-nodes and t = 15 s (squares),five-nodes and t = 25 s (diamonds), seven-nodes and t = 35 s (triangles), and nine-nodes and t = 45 s (circles).Fig. 16. Total throughput for the WRAP protocol under symmetric loading, as a function of total offered network load, for a channel capacity of B = 1 Gb/s perwavelength and a slot duration which is increased in proportion to the number of nodes. Results are shown for three-nodes and t = 15 s (squares), five-nodesand t = 25 s (diamonds), seven-nodes and t = 35 s (triangles), and nine-nodes and t = 45 s (circles). We conclude that the WRAP protocol scales reasonably well topology. In addition, FatMAC has been specifically optimizedwith increasing numbers of nodes, so long as the slot period for multiprocessor interconnection, and as such tends to tradeis similarly scaled. This scaling property is the result of the off bandwidth efficiency in order to obtain lower minimumfact that each node brings with it additional capacity, in the delay. The FatMAC protocol is a time-slotted protocol withform of an added wavelength. The additional delay caused by two phases. In the first, control slots containing the destinationthe longer slot period under lightly loaded conditions must be of the packet at the head of the queue are broadcast by eachconsidered. If the geographical size of the network increases as node. In the second phase the packet from the head of eachnodes are added then the minimum delay remains proportional node’s queue is sent, with the order of transmissions to ato the inherent propagation delays. Since propagation delays are destination determined by the number of control slots with thatunavoidable, the penalty due to the protocol delays in such a destination’s address in the previous phase. This two phasecase is minimal. This additional delay for light loads is small procedure is then repeated. A simulation comparison betweenin comparison with queuing delays and it is deterministic, and WRAP, I-TDMA, and FatMAC has been carried out in order tothus unlikely to adversely affect multimedia traffic. The effect assess the relative performance of our protocol.of this delay on applications using the network is an important In the first set of simulations, all protocols were simulatedtopic, but is beyond the scope of the present paper. for the symmetrically-loaded 3-node case, and for loads ranging from 10 to 150%. In order to make the comparisons as fair asB. Comparison with Other Applicable Protocols possible, the I-TDMA and FatMAC slot lengths were set equal to the WRAP minislot length. To accommodate FatMAC, the Two previously-proposed MAC protocols which could be network topology is assumed to be a star, the optical transmit-applied to WR-PONs of the type considered here are I-TDMA ters are assumed to have a WDM broadcast facility, and propa- and FatMAC . The former protocol is applicable to gation delays in the network are assumed to be negligible, i.e.,both stars and rings, while the latter is only practical in a star less than the transmission time of the smallest data or protocol
SPENCER AND SUMMERFIELD: MAC PROTOCOL FOR WAVELENGTH-ROUTED PONS 1671 (a) (b)Fig. 17. (a) Comparison of the average packet delay for the WRAP (squares), I-TDMA (diamonds), and FatMAC (triangles) protocols for three-node networksunder symmetric loading, as a function of the offered load. (b) Corresponding comparison of packet loss.unit. While both WRAP and I-TDMA can accommodate prop- TABLE II ASYMMETRIC TRAFFIC LOAD DISTRIBUTIONS FOR SIMULATIONS OF ANagation delays within their slot structures, FatMAC cannot tol- EIGHT-NODE NETWORKerate propagation delays which exceed its small control slot du-ration. The delay and loss of these protocols as a function of thefractional offered load to the network are shown in Fig. 17(a)and (b) respectively. As expected, FatMAC gives low delaysunder light load, but has poor performance as the loading is in-creased. I-TDMA has similar but better performance to WRAPfor the symmetric loading, as it does not incur as much protocoldelay and has lower protocol overheads. While I-TDMA is op-timal for symmetric loading with small traffic arrival variations,it is not flexible, in that the maximum available capacity for anysource/destination route is limited to of that obtain-able under WRAP. To illustrate this, an eight-node network with traffic loadsas shown in Table II was simulated. Equivalent slot/minislotlengths were used for the different protocols, and buffersize per source/destination route chosen to be 40 times thecapacity of a slot/minislot. Loading from 10 to 150% ofthe tabulated load was applied, this factor being called the this throughput for an example node, Node 4, into the traffic“Traffic Matrix Scaling Factor” in the figures. The nodes sent to the four destinations making up the total (i.e., Nodeshave been numbered 0 through 7. For each protocol we have 0, 3, 5, and 6); and 3) the corresponding average delaysplotted: 1) the throughput per node, measured as the total experienced by traffic sent from Node 4 to each of the fouroutput to all other nodes in the network; 2) a breakdown of destinations. Fig. 18 shows the throughput per node for a)
1672 JOURNAL OF LIGHTWAVE TECHNOLOGY, VOL. 18, NO. 12, DECEMBER 2000 (a) (b) (c)Fig. 18. Total throughput achieved by each node in an eight-node network under the asymmetric traffic pattern shown in Table II as a function of the factor bywhich the traffic matrix is scaled. Simulation results are shown for (a) I-TDMA, (b) FatMAC, and (c) WRAP.I-TDMA, b) FatMAC, and c) WRAP; Fig. 19(a)–(c) shows tination capacity is sought. Only Node 0 is able to obtain allthe breakdown for traffic sent from node 4; and Fig. 20(a)–(c) of the desired throughput as the loading is increased, as indi-shows the corresponding delays. cated by the linear growth of throughput as a function of loading It can be seen from Fig. 18(a), Fig. 19(a), and Fig. 20(a) that in Fig. 18(a). This is because the total output from Node 0 isthe fixed bandwidth allocations of I-TDMA lead to low limiting uniformly distributed to all seven remote nodes, a circumstancethroughput and high delays if more than the available per-des- which is perfectly accommodated by the I-TDMA schedule. The
SPENCER AND SUMMERFIELD: MAC PROTOCOL FOR WAVELENGTH-ROUTED PONS 1673 (a) (b) (c)Fig. 19. Breakdown of the throughput achieved by Node 4 to each of the four destination nodes, Node 0, Node 3, Node 5, and Node 6, corresponding to the totalthroughput shown in Fig. 18. Simulation results are shown for (a) I-TDMA, (b) FatMAC, and (c) WRAP.throughput of all other nodes is severely limited due to exhaus- exhausts the capacity of the slots corresponding to those desti-tion of the capacity of the per-destination slots in the transmis- nations in the schedule. The onset of this capacity exhaustion ission schedule. This is demonstrated by the flattening out of the accompanied by the large increases in queuing delays observedthroughput curves in Fig. 18(a). For the particular case of Node in Fig. 20(a).4, shown in Fig. 19(a) and 20(a), the small proportion of traffic Fig. 18(b) shows that FatMAC obtains better overalldestined to Nodes 0 and 3 is able to be transmitted. However the throughput than I-TDMA. Fig. 20 also shows that it has themajority of traffic, which is destined for Nodes 5 and 6, rapidly lowest delays of all three protocols under light loading and
1674 JOURNAL OF LIGHTWAVE TECHNOLOGY, VOL. 18, NO. 12, DECEMBER 2000 (a) (b) (c)Fig. 20. Breakdown of the average delay experienced by packets sent from Node 4 to each of the four destination nodes, Node 0, Node 3, Node 5, and Node 6,corresponding to the throughput per destination shown in Fig. 19. Simulation results are shown for (a) I-TDMA, (b) FatMAC, and (c) WRAP.overload conditions. However, Fig. 19(b) demonstrates that to Nodes 0, 3, and 6 could be improved, at the expense of aalthough Node 4 obtains higher overall throughput under reduction in throughput to Node 5. An additional problemFatMAC than I-TDMA, this is the result of higher throughput with FatMAC is that it suffers from head-of-line blocking as itto Node 5, with traffic to Nodes 0, 3, and 6 fairing somewhat only uses one queue per node. The loading levels at which theworse. This is the result of a fairness problem in FatMAC throughput to each destination begins to limit is more clearlywhich we have not addressed, although a possible solution observed from the locations of the knees in the delay curves ofis discussed in . If this problem were solved, throughput Fig. 20(b).
SPENCER AND SUMMERFIELD: MAC PROTOCOL FOR WAVELENGTH-ROUTED PONS 1675 Fig. 18(c) shows that the WRAP protocol achieves the highest the lowest delay and delay variance, and the highest averagemaximum throughput for all nodes (note that the vertical scale is throughput of the three algorithms considered.twice that of the other two graphs in the Figure). Furthermore, it The scaling behavior of the protocol for increasing numbershares the available capacity most efficiently amongst the nodes, of nodes was investigated by simulation, and found to be ac-as evidenced by the fact that the throughput increases linearly ceptable provided that the slot length is increased in proportionwith offered load for all nodes to a scaling factor of over 50%. the number of nodes. This maintains the same relative protocolWhen the throughput does level out, it is due to the fact that overheads but does so at the expense of an increase in the in-the total offered traffic from all sources to a particular desti- trinsic protocol delay.nation exceeds the total capacity available to that destination. Simulations have shown that WRAP provides both fairFor example, the total traffic offered to the network destined for and flexible access to the network. It allows bandwidth to beNode 6 is the highest, so nodes transmitting to that node (i.e., allocated according to demand, while providing a guaranteedNodes 4 and 7) are the first to be affected. Although Node 0 is minimum bandwidth allocation between each source-destina-also transmitting to Node 6, it is not affected in this example, as tion pair. Comparisons were performed with two alternativeits demand never exceeds the guaranteed minimum provided by protocols, I-TDMA and FatMAC. For symmetric networkthe preferential I-TDMA allocations. This is further illustrated loading I-TDMA performs optimally, and has no additionalin Fig. 19(c) which shows that the throughput to Node 6 is the protocol overheads. Under these conditions, WRAP limitsfirst to level out, at a scaling factor of about 60%. As the total to the I-TDMA behavior under heavy loads, but has highernetwork load is increased, throughput to Node 6 from Node 4 overheads, and larger intrinsic protocol delays. FatMAC hasactually declines, due to increased traffic demands from other the lowest intrinsic delay, however this is obtained at thesource nodes, however it can never fall below the guaranteed expense of network utilization. Under asymmetric trafficminimum. Once again, the knees in the delay curves shown in conditions WRAP provides the best throughput performanceFig. 20(c) are the clearest indicator of the loading level at which in comparison to FatMAC and I-TDMA, offering the I-TDMAthe throughput reaches a maximum. capacity as a guaranteed minimum, and maintaining low delays Based on the simulation results in this section, we conclude for comparatively higher offered loads. We thus conclude thatthat I-TDMA is only appropriate in an environment where the of the three protocols considered here, WRAP is the best-suitedtotal traffic load on the network is distributed evenly amongst to general-purpose data communications applications such asall source-destination pairs. This is unsurprising, as it utilizes local, campus and metropolitan area networks.a fixed TDMA schedule. 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