# Bounding Communication Cost in Dynamic Load Balancing of Distributed Hash Tables

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2 The Peer-to-Peer networks which we are considering are based on consistent hashing [1] with ring topology like Chord [2], Tapestry [3], Pastry [4], and a topology inspired by de Bruijn graph [5, 6]. In short, Distributed Hash Tables work as follows. Each peer chooses a number between 0 and 1 and takes responsibility for the interval from the chosen place to the next node. When a peer wants to insert or find an item x in the network, it computes a global hash function on the item s name which is a value h(x) between 0 and 1. It can then contact the node responsible for h(x) using the network structure that has been built among the participating nodes. There are several approaches to build an efficient topology, however the exact structure of the network is not relevant for us. It is only important that each node has direct links to its successor and predecessor on the ring and that there is a routine that lets any node contact the node responsible for any given point in the network in time D. A crucial parameter of a network defined in this way is its smoothness which is the ratio of the length of the longest interval to the length of the shortest interval. The smoothness is a parameter, which informs about three aspects of the load balance. Storage load of a peer: The longer its interval is, the more data has to be stored in the peer. On the other hand, if there are n nodes and Ω(n log n) items distributed (pseudo-)randomly on the ring, then, with high probability, the items are distributed evenly among the peers provided that the smoothness is constant. Degree of a node: A longer interval has a higher probability of being contacted by many short intervals which increases its in-degree. Congestion and dilation: Having constant smoothness is necessary to keep these routing parameters small, for example in [5,6]. Even if we choose the points for the nodes fully randomly, the smoothness is as high as Ω(n logn) with constant probability, 3 whereas we would like it to be constant (n denotes the current number of nodes). 1.1 Related work Load balancing has been a crucial issue in the field of Peer-to-Peer networks since the design of the first network topologies like Chord [2]. In [7] the concept of virtual servers was introduced and it was proposed that each real peer works as log n virtual servers, thus greatly decreasing the probability that some peer will get a large part of the ring. Some extensions of this method were proposed in [8] and [9], where more schemes based on virtual servers were introduced and experimentally evaluated. Unfortunately, such an approach increases the degree of each peer by a factor of log n, because each peer has to keep all the links of all its virtual servers. 3 With constant probability (w.c.p.) means with probability at least 1 ǫ for an arbitrary constant ǫ > 0.

3 The paradigm of many random choices [10] was used by Byers et al [11] and by Naor and Wieder [5,6]. When a peer joins, it contacts log n random places in the network and chooses to cut the longest of all the found intervals. This yields constant smoothness with high probability. 4 A similar approach was proposed in [12]. It extensively uses the structure of the hypercube to decrease the number of random choices to one and the communication to only one node and its neighbors. It also achieves constant smoothness with high probability. The two approaches above have a certain drawback. They both assume that peers join the network sequentially. What is more important, they do not provide analysis for the problem of balancing the intervals afresh when peers leave the network. Karger and Ruhl [13,14] propose a scheme, in which each node chooses Θ(log n) places in the network and takes responsibility for only one of them. This can change, if some nodes leave or join, but each node migrates only among the Θ(log n) places it chose, and after each operation Θ(log log n) nodes have to migrate on expectation. The advantage of our algorithm is that it performs a migration only if there really is a too long interval in the network. Both their and our algorithms use only tiny messages for checking the network state, and in both approaches the number of messages in half-life 5 can be bounded by Θ(log n) per peer. Their scheme is claimed to be resistant to attacks thanks to the fact that each node can only join in logarithmically bounded number of places on the ring. However, in [15] it is stated that such a scheme cannot be secure and that more sophisticated algorithms are needed to provide provable security. The reasoning for this is that with IPv6 protocol the adversary has access to thousands of IP numbers, and she can join the system with the ones falling into an interval that she has chosen. She does not have to join the system with each possible IP to check if this IP is useful, because the hash functions are public and she can compute them offline. Manku [16] presented a scheme based on a virtual binary tree that achieves constant smoothness with low communication cost for peers joining or leaving the network. It is also shown that the smoothness can be diminished to as low as (1 + ǫ) with communication cost per operation increased to O(1/ǫ). All the nodes form a binary tree, where some of them (called active) are responsible for perfect balancing of subtrees rooted at them. Our scheme treats all peers evenly and is substantially simpler. In [17] the authors present a scheme that extends [16] and generalizes [12]. Nodes are grouped into clusters of size h. Each node checks k clusters when it joins the network (In [12], h = log n and k = 1). The authors prove that it is sufficient that h k = Ω(log n) to guarantee that the resulting smoothness 4 With high probability (w.h.p.) means with probability at least 1 O ( 1 n l ) for an arbitrary constant l. 5 For now, we define half-life of the network to be the time it takes for half of the peers in the system to arrive or depart.

4 is constant, with high probability. Similarly to the other two approaches, this algorithm does not consider nodes leaving the network. 1.2 Our results In [18] we showed a self-stabilizing distributed algorithm that rebalances the system in the static case, that is when no nodes join or leave it. Such strong assumption was needed in the analysis, but the algorithm itself was designed to work in a dynamic environment. In this paper we extend and modify the scheme so, that it can be analyzed as an ever-running process. We show formally that the communication in a half-life is bounded to log n messages per node. Unbounded communication was the main drawback of our previous scheme. Consider a scenario in which the sequence of join and leave requests is generated by an oblivious adversary and the algorithm responsible for keeping the system connected is optimal with respect to communication cost. We show that our algorithm keeps the system smooth with only slight delay with additional communication cost by a logarithmic factor larger than the cost of keeping the system connected. Consider any phase of length equal to half-life, in which the sequence of join and leave requests is generated by an oblivious adversary. We show that our algorithm smooths the imbalance present in the beginning of the phase. As a byproduct we present a scheme, which gives a constant approximation of the number of nodes in the network n. The scheme needs the knowledge of an upper bound of log n. Finally, we present the results of the experimental evaluation that we performed. It shows that the constants used in the theoretical analysis can be notably reduced in practice. 2 The algorithm We assume that each node knows an upper bound on log n, the logarithm of the current number n of nodes in the system. In order to avoid complex terms in the analysis, we assume that is a sharp bound (up to constant factors), that is = Θ(log n). This bound should be global for the whole network. Most of the existing systems already assume such knowledge. It will be used to estimate the value of n, in order to enable each node to decide whether its current load is below or above the average. An assumption already mentioned in the introduction is that each node can communicate directly with its successor and predecessor on the ring. Let D denote the dilation of the network, i.e. the maximum number of hops needed to route from any node to any other node. D is O(log n) in Chord [2] and O(log n/ log log n) in de Bruijn graph [5, 6]. In Subsection 2.1 we introduce a notion of weight of an interval. It is a measure supplementary to the interval s length. We show the relationships between these two measures and how to use them to efficiently approximate the number

5 of participants of the network. In Subsection 2.2 we state the algorithm and in Subsection 2.3 we show that it achieves constant smoothness and its communication is bounded. In Section 3 we present the experimental evaluation of our algorithm. 2.1 Estimating the current number of nodes The goal of this subsection is to provide a scheme, which, for an arbitrary node i responsible for an interval I i of sufficient length, returns an estimate n i of the total number of nodes, so that each n i is within a constant factor of n, with high probability. Let be the upper bound on log n mentioned above. Each node keeps connections to δ = α random positions on the ring, for sufficiently large global constant α. We assume each of these connections is visible for the node responsible for the interval containing the chosen position. These connections are called markers and define weights of intervals and nodes in the following way. Definition 1. For a node i and its interval I i, the weight of the interval I i, denoted w(i i ), is the number of markers which fall into I i. The weight of a node is the weight of the interval it is responsible for. It is clear that E[w(I i )] = δ l(i i ) n, so n i = w(ii) δ l(i i) is a good estimate for n. In the following lemma we state precisely the relation between the length and the weight of an interval and prove, using Chernoff bounds, that the quality of the estimate is sufficient. Lemma 1. There exist constants w min < 1 2, l min and l max, such that for arbitrary interval I i and timestep t, if α is large enough, the following holds with high probability. If w(i i ) w min δ, then l min w(ii) δ n l(i i) l max w(ii) δ n. Proof. Consider an interval I of length l(i) lmin n and let I be any interval containing I such that l(i ) = lmin n. Let X = w(i ) be the random variable denoting how many of the δ n markers hit I ; we have E[X] = l min δ = Θ(δ) = Ω(log n). Using Chernoff bound, we obtain X = Θ(δ), with high probability. Since I is contained in I, its weight is always smaller than the weight of I, so w(i) = O(δ), with high probability. Thus, if an interval I has weight) w(i) = Ω, which w min δ = Ω(δ) then, with high probability, l(i) lmin n proves the first inequality. ( w(i) δ n For the proof of the second inequality, we assume that l(i) lmin n. The random variable X = w(i) has expectation E[X] = l(i) δ n = Ω(log n), so with high probability, w(i) = Ω(l(I) δ n). This finishes the proof. It is hard for nodes to compare their length to the average length (which is 1/n) due to the lack of knowledge about the exact number of participants n. On the other hand it is easy for them to compare their weight to the average weight, since the latter is a global constant equal to δ.

6 2.2 The balancing algorithm Our algorithm reorganizes the positions of the nodes on the ring, so that the resulting intervals I i have weights between w min δ and 32 δ, where w min is the constant defined in Lemma 1. By Lemma 1, this implies that in the end the smoothness is constant, with high probability. We call the intervals of weight at most 16 δ light and intervals of weight at least 32 δ heavy. Intervals of intermediate weight are called medium. The algorithm minimizes the weight of the heaviest interval, but we also have to take care that no interval is too light. This is needed, since smoothness demands not only the longest interval to be sufficiently short but also the shortest interval to be sufficiently long. Therefore, before we begin the routine we force all the nodes responsible for intervals with weight smaller than w min δ to leave the network; we call such nodes very light. Doing this, we assure that the weight of the lightest interval in the network is bounded from below. We have to explain why this does not destroy the structure of the network. That is, that the algorithm does not worsen the balance of the system instead of improving it and that the removed nodes can still communicate with others. First of all, it is possible that we remove a huge fraction of the nodes. It is even possible that a very long interval appears even though the network was very near to a balanced state before. This is not a problem, since the algorithm rebalances the system. Besides, if this algorithm is used also for new nodes at the moment of joining, this initialization is not necessary. We do not completely remove the nodes with too light intervals from the network. The number of nodes n, and thus also the weight δ of a perfect interval is unaffected, and the removed nodes act as though they were simple light nodes. Each of the very light nodes knows its virtual position on the ring and contacts the network through the node responsible for that position. We assume that the node is not heavy since it could immediately split using the light node. The routine works differently for different nodes, depending on the node s interval weight. The heavy nodes only wait for help proposals in order to split their intervals. The medium nodes do nothing as they are too light to split their intervals and too heavy to help. In order to be able to bound the communication cost we distinguish two types of light nodes. A light node can be passive and do nothing until it is stimulated, i.e. until it loses or receives a marker, which is noticed by a loss or gain of a connection. Then the node becomes active and stays active until it successfully sends a constant (A) number of help proposals. An active node sends help proposals to random places until it finds a heavy node or it becomes passive. In order not to create any heavy nodes, an active node sends a proposal only if its predecessor is light and has been locked. Locking means that the predecessor will not send any proposals, and thus guarantees that the predecessor will not migrate. The routines for two types of light nodes are depicted in Figure 1. The parameters A and F, responsible for the maximum activity time and the forwarding distance, are sufficiently large constants, which will be defined later.

7 The very light nodes act in the following way. The light node responsible for such nodes informs them of its markers and they are stimulated at the same time as it is. They behave in the same way as the light nodes except that they do not have to lock anything. passive light (permanently in staying state) if (any change in markers) { change activity level to A become active } active light if (locked) wait until unlocked by successor if (predecessor is light) { do perform one of the following choosing equiprobably at random let the successor lock me and wait until unlocked try to lock the predecessor until (predecessor locked) p := random(0..1) P := the node responsible for p if (P is heavy and migrating to p does not produce a very light interval) migrate to p else contact consecutively the node P and its F successors on the ring if (a contacted node R is heavy) migrate to the middle of the interval of R unlock the predecessor } if (any change in markers) change activity level to A decrease activity level if (activity level = 0) become passive Fig.1. The algorithm for two types of light nodes 2.3 Performance and communication cost In [18] we showed an algorithm that can be treated as the algorithm we are currently analyzing with parameter A =, and we proved that it is self-stabilizing, that is after a constant number of rounds the smoothness becomes constant, provided that no new nodes join the system and no nodes leave it during this time.

9 5. After a half-life the phase ends, and the adversary is switched off. The algorithm is allowed to communicate and migrate in an extension of Θ(log n) timesteps. In order to precisely express the performance of the algorithm, we consider the following virtual process after the half-life has ended. For each node which left during the considered half-life, we re-insert it into the network provided that such an insertion does not create a very light interval. We perform the reinsertions sequentially in any order, so that there are no ambiguities with parallel checks. If we prove that after performing such process the system is balanced, we prove that our algorithm copes with the imbalance present in the network at the beginning of the half-life. The imbalance created by the adversarial deletions of nodes will be repaired in the next half-life. In the rest of this subsection we prove the following theorem characterizing the communication cost of the algorithm in a half-life and stating how well the algorithm rebalances the system. Theorem 2. In the game described above, the following properties hold with high probability: 1. The communication cost of the algorithm in half-life is at most O(log n) times larger than the communication cost of maintaining the system. 2. If we re-insert the nodes removed in the current half-life as described above, the smoothness is constant at the end of the half-life. For further analysis we need the following claims about the behavior of the network in a half-life. Lemma 2. Denote the number of nodes in the network in the beginning of a half-life by n 0, the number of nodes present at time t by n t, and the number of markers at time t as m t, for t 0. Then during a half-life 1. the number of nodes in the network at time t is bounded by n0 2 n t 2 n 0, 2. the number of markers in the network at time t is bounded by n0 δ 2 m t 2 n 0 δ. 3. the number of insert / delete operations in a half-life ranges from n0 5 n0 2 to 2, Proof. Let the nodes present at time t 0 be called old and the nodes inserted during half-life new. Before the half-life ends the adversary can remove at most n 0 2 old nodes and insert at most 2 n 0 new nodes. Thus, the total number of nodes ranges from n0 2 to 2 n 0 and the number of markers ranges from n0 δ 2 to 2 n 0 δ. It is straightforward that removing n 0 /2 nodes is the fastest way to end the half-life, whereas inserting n 0 1 new nodes and removing them and n 0 /2 old ones is the best way for the half-life to last long. Thus, the last statement follows. We are now ready to prove the communication-competitiveness of the algorithm.

10 Proof (of Property 1 of Theorem 2). For the exact analysis we modify our algorithm slightly. Instead of δ markers, each node inserts δ/2 insertion-markers and δ/2 deletion-markers into the system. If a node u stores a marker of a node v then u is activated in two cases: if v joins the network and the marker is an insertion-marker or if v leaves the network and the marker is a deletion-marker. Since the adversary is oblivious, we may think that the deletion-markers are generated at the moment when a node leaves the network. Thus, when a node joins or leaves the network, it contacts Θ(δ) places chosen uniformly at random at the moment in which the operation happens. Since the total number of joins and leaves in a half-life is Θ(n), the total number of contacted places in a half-life is Θ(n δ), all of which are independent. Since each light node has length Θ(1/n) during the whole half-life we can use Chernoff bounds to show that each light node is activated Θ(δ) times, with high probability. Each activation is responsible for a constant number of messages, so the total number of messages sent by a light node is O(δ) = O(log n) in a half-life, with high probability. Each of these messages is sent to a distance of O(D+logn) = O(log n) and the total communication cost is O(log 2 n) messages per node. By Theorem 1, this proves the first property of Theorem 2. It remains to prove that such communication is still sufficient to keep the system balanced. First, we show that the number of proposals in a half-life is sufficiently high. Lemma 3. There exist n0 16 light nodes alive during the whole half-life which send the maximum number A of proposals and still are not accepted anywhere. With high probability, a fraction 1 3 of the messages they want to send is sent in the half-life or in the extension. Proof. There are at least n 0 /2 nodes which are alive during the whole half-life - we call them immortals. Since an interval is light when its weight is below 16 δ and, by Lemma 2, the total weight in the network is at most 2 n 0 δ, at most n 0 /8 nodes are not light, and thus at every timestep at least 3 8 n 0 immortals are light. Consider the distribution of the immortals in the beginning of the half-life and ignore other nodes for a while. We couple the immortals in the following way. Any immortal is chosen as the first element of the first pair and its successor is the second element of the first pair. Consecutive immortals (according to the order on the ring) build consecutive pairs. We want to bound the number of pairs with two light nodes. There are n 0 /4 pairs, at most n 0 /8 of which can have at least one middle or heavy node, thus in at least n 0 /8 pairs, an immortal will be willing to migrate when it is activated. Notice that presence of other nodes can only decrease the weight and cannot discourage any immortals from sending proposals. As a node only wishes to share its load with another node, if its current weight exceeds 32 δ, n 0 /16 immortals suffice to store all load. Thus, only n 0 /16 immortals have a chance to migrate and at least n 0 /16 send the maximum number A of messages.

11 Assume that in order to send a message, a node does not have to lock its predecessor. Then all A messages sent by a node will be sent in time O(log n). Suppose that the part of the algorithm responsible for trying to lock a neighbor is performed immediately. Then at every timestep if a node is not sending a message, it means that either it cannot lock its predecessor or that it has been locked by its successor. In either case it implies that one of its neighbors is sending a message. This means that at least 1 3 of the messages will be delivered in the time, in which all messages would be delivered if locking was immediate. We assign additional Θ(log n) timesteps to each active node. Each time a node fails to lock a predecessor or fails to let the successor lock itself, because the other node tries to synchronize with its other neighbor, we pay for such waste of time from these additional timesteps. If within these Θ(log n) timesteps a node succeeds A times, it manages to send all its messages. At each such trial the node succeeds with probability at least 1 4, as at least one of its neighbors tries to contact it at the same time and the probability that at a timestep they try to contact each other is exactly 1 4. With high probability, there is at least one success in Θ(log n) timesteps, and thus also A successes in Θ(log n) trials. Also, all nodes use Θ(log n) additional timesteps to succeed A times, w.h.p. Below we prove the second part of Theorem 2. As writing all constants exactly significantly decreases the readability of the proof, we simplify it by assuming that nodes operate on lengths instead of weights and that an interval is considered very short, if its length is below 1 n. This assumption is justified by Lemma 1. Proof (of Property 2 of Theorem 2). We know that in the half-life (or in the extension guaranteed by the definition of the model) Θ(n) proposals from immortals are generated in the system, and that we can make the constant in the Θ notation arbitrarily large just by increasing the A constant in the algorithm. Let y [0, 1) be an arbitrary real number on the ring. Consider the interval I y of length log n n starting in y. Let x and z be equal to y log n n and y + log n n respectively. I x and I z are intervals of the same length as the length of I y and are the predecessor and successor of I y on the ring. Using Chernoff bound we can show that with high probability the number of proposals from immortals that fall into I x is Θ(log n). We say that a proposal reaches z, if it contacts the node responsible for z. The parameter F is chosen so, that even if all intervals between x and z have the lowest possible length, a proposal travelling from x and forwarded through all the nodes can reach z. At this point of the analysis we allow the adversary to be adaptive. She can choose, which of the mortal nodes die and when. When she removes a node from the network, a heavy interval can appear, which should be split by an immortal. Our analysis is based on the fact that when an immortal splits a heavy interval, it does it at least until the end of the half-life. The adversary may remove some nodes from I y after all immortals have been forwarded through I y and the algorithm repairs such imbalance in the next half-life. Below we show that the following event happens at least 2 log n times: a proposal that hit I x reaches z at the moment when there is a node in I z. If such

12 an event happens, then either there are no heavy nodes in I y, or the immortal is inserted somewhere between y and the node in I z. At most 2 log n immortals can be inserted in the interval I y I z, so after 2 log n such events there are no heavy nodes in I y. First, we show that Θ(log n) of the proposals reach z in the process of forwarding (not necessarily when there exists a node in I z ). This is true, since in I x and I y there is place for at most 2 log n of them. Notice that we cannot be sure which proposals (in order of coming) reach z: it may happen that the first one gets through, right after that many nodes die in the network, and the next immortals split the heavy intervals that appear in I x and I y. We introduce a notion of a steady barrier as the first node behind z on the ring which will survive until the end of the half-life. At the beginning of the process we can bound the distance between y and the barrier by 1. When during our process an immortal reaches z and there is no node in I z then the immortal is inserted between y and the current steady barrier and becomes the new steady barrier. In the worst case, the distance between y and the steady barrier halves each time when it happens, so after fewer than log n such events there is a steady barrier in I z. Thus, log n immortals can be used to assure that there is a steady barrier in z and further 2 log n to assure that I y is balanced. n log n If we divide the ring into log n disjoint intervals each of length n, then the probability that the algorithm fails for at least one of them is at most n times larger than the probability of failure for a single interval. Thus, our algorithm succeeds on the whole ring with high probability. 3 Experimental evaluation In this paper we focused on theoretical proofs of correctness and efficiency of our algorithm. Additionally, we performed a series of experiments on a single machine, in order to show that constants emerging from our analysis are reasonable. We did not experimentally compare our approach with others, as such comparisons should probably be done in a large distributed environment. We did not model the dynamics of the system, either, as our approach copes with adversarial dynamics and in this matter we consider proofs of correctness to be much better than any experiments. In all tests the underlying space was the ring of natural numbers {0,...,2 64 1} with arithmetic modulo There were two groups of experiments. The first group was to find the best parameters for the algorithm in the static case. In the tests in this group n nodes were inserted in the very beginning and all lived until the end of the test. The number of nodes n was 2 5, 2 11, 2 14 or 2 17 ; they were chosen to be powers of two only for easier implementation. We present the results for the highest number of nodes. The first parameter we tested was the number of markers that each node should insert into the network. The number of markers δ ranged from 8 to 128 and the other parameters were set as follows. Each packet was forwarded through 128 nodes to assure quick convergence. An interval could send a proposal, if

13 the sum of the weight of itself and its successor was at most δ. An interval was considered very light if its weight was below δ/2. We measured the final smoothness in the system and noticed that choosing a value of 40 already yields smoothness below 20 and values larger than 64 do not improve the system much. The final smoothness for different δ parameters is depicted in Figure The resulting smoothness Number of markers per node Fig.2. Smoothness according to approximation of log n With the number of markers fixed to 64 we performed tests which had to estimate the best definition of a very light interval. The forward parameter F was 128 again. The final smoothness for different weights of a very light interval is depicted in Figure 3. We consider the constant 56 to be the best The resulting smoothness Very Light Interval parameter Fig.3. Smoothness according to the weight of the very light interval The last set of tests for the static scenario was developed to optimize the distance to which a packet should be forwarded. This time the output of the analysis was not smoothness but the time it takes for the scheme to stabilize. From Figure 4 it can be seen that already very short forwarding gives good results. We chose the forward parameter to be 16. To test the behavior of our algorithm in a dynamic scenario, we took the model from [20]. For each time step the number of nodes joining at this time

14 Steps to converge The Forwad parameter Fig.4. Convergence according to the forward parameter step was chosen according to the Poisson distribution with mean λ, and for each node its time to stay in the system was chosen according to the exponential distribution with mean µ. For parameters (λ, µ) equal to (10, 100) and (100, 100) the smoothness never exceeded 14 and was in 97% time steps at most 9. For comparison, if we insert the nodes uniformly and independently at random the smoothness is approximately 23 for 2 14 and 127 for 2 17 nodes. References 1. Karger, D.R., Lehman, E., Leighton, F.T., Levine, M.S., Lewin, D., Panigrahy, R.: Consistent Hashing and Random Trees: Distributed Caching Protocols for Relieving Hot Spots on the World Wide Web. In: Proc. of the 29th ACM Symp. on Theory of Computing (STOC). (1997) Stoica, I., Morris, R., Karger, D.R., Kaashoek, M.F., Balakrishnan, H.: Chord: A Scalable Peer-to-Peer Lookup Service for Internet Applications. In: Proc. of the ACM SIGCOMM. (2001) Hildrum, K., Kubiatowicz, J.D., Rao, S., Zhao, B.Y.: Distributed Object Location in a Dynamic Network. In: Proc. of the 14th ACM Symp. on Parallel Algorithms and Architectures (SPAA). (2002) Rowstron, A., Druschel, P.: Pastry: Scalable, Decentralized Object Location, and Routing for Large-Scale Peer-to-Peer Systems. Lecture Notes in Computer Science (2001) Naor, M., Wieder, U.: A Simple Fault Tolerant Distributed Hash Table. In: 2nd International Workshop on Peer-to-Peer Systems (IPTPS). (2003) 6. Naor, M., Wieder, U.: Novel Architectures for P2P Applications: the Continuous- Discrete Approach. In: Proc. of the 15th ACM Symp. on Parallel Algorithms and Architectures (SPAA). (2003) Dabek, F., Kaashoek, M.F., Karger, D., Morris, R., Stoica, I.: Wide-Area Cooperative Storage with CFS. In: Proceedings of the 18th ACM Symposium on Operating Systems Principles (SOSP). (2001) 8. Rao, A., Lakshminarayanan, K., Surana, S., Karp, R.M., Stoica, I.: Load Balancing in Structured P2P Systems. In: 2nd International Workshop on Peer-to-Peer Systems (IPTPS). (2003)

15 9. Godfrey, B., Lakshminarayanan, K., Surana, S., Karp, R.M., Stoica, I.: Load Balancing in Dynamic Structured P2P Systems. In: 23rd Conference of the IEEE Communications Society (INFOCOM). (2004) 10. Michael Mitzenmacher, Andra W. Richa, R.K.S.: The Power of Two Random Choices: A Survey of Techniques and Results. In: Handbook of Randomized Computing. P. Pardalos, S.Rajasekaran, J.Rolim, and Eds. Kluwer, editors. (2000) 11. Byers, J.W., Considine, J., Mitzenmacher, M.: Simple Load Balancing for Distributed Hash Tables. In: 2nd International Workshop on Peer-to-Peer Systems (IPTPS). (2003) 12. Adler, M., Halperin, E., Karp, R.M., Vazirani, V.: A stochastic process on the hypercube with applications to peer-to-peer networks. In: Proc. of the 35th ACM Symp. on Theory of Computing (STOC). (2003) Karger, D.R., Ruhl, M.: Simple Efficient Load Balancing Algorithms for Peer-to- Peer Systems. In: 3rd International Workshop on Peer-to-Peer Systems (IPTPS). (2004) 14. Karger, D.R., Ruhl, M.: Simple Efficient Load Balancing Algorithms for Peer-to- Peer Systems. In: Proc. of the 16th ACM Symp. on Parallelism in Algorithms and Architectures (SPAA). (2004) Awerbuch, B., Scheideler, C.: Group Spreading: A protocol for provably secure distributed name service. In: Proc. of the 31st Int. Colloquium on Automata, Languages, and Programming (ICALP). (2004) Manku, G.S.: Balanced Binary Trees for ID Management And Load Balance in Distributed Hash Tables. In: Proc. of the 23rd annual ACM symposium on Principles of Distributed Computing (PODC). (2004) Kenthapadi, K., Manku, G.S.: Decentralized Algorithms Using Both Local and Random Probes for P2P Load Balancing. In: Proc. of the 17th ACM Symp. on Parallel Algorithms and Architectures (SPAA). (2005) 18. Bienkowski, M., Korzeniowski, M., Meyer auf der Heide, F.: Dynamic Load Balancing in Distributed Hash Tables. In: 4th International Workshop on Peer-to-Peer Systems (IPTPS). (2005) 19. Liben-Nowell, D., Balakrishnan, H., Karger, D.R.: Analysis of the Evolution of Peer-to-Peer Systems. In: Proc. of the 21st annual ACM symposium on Principles of Distributed Computing(PODC). (2002) Pandurangan, G., Raghavan, P., Upfal, E.: Building Low-Diameter Peer-to-Peer Networks. In: Proc. of the 42nd IEEE symposium on Foundations of Computer Science (FOCS). (2001)

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