USENIX '05 Paper
[USENIX '05 Technical Program]
Attrition Defenses for a Peer-to-Peer Digital Preservation System
Stanford University, CA
Intel Research, Berkeley, CA
HP Labs, Palo Alto, CA
David S. H. Rosenthal
Stanford University Libraries, CA
Harvard University, Cambridge, MA
Denial of Service (DoS) attacks are among the most difficult for distributed systems to resist. Distinguishing legitimate requests for service from the attacker's requests can be tricky, and devoting substantial effort to doing so can easily be self-defeating. The term DoS was introduced by Needham  with a broad meaning but over time it has come to mean high-bit-rate network-level flooding attacks  that rapidly degrade the usefulness of the victim system. In addition to DoS, we use the term attrition to include also moderate- or low-bit-rate application-level attacks that impair the victim system.
The mechanisms described in this paper are aimed at equipping the LOCKSS (LOCKSS is a trademark of Stanford University -- it stands for "Lots Of Copies Keep Stuff Safe") peer-to-peer (P2P) digital preservation system to resist attrition attacks. The system is in use at about 80 libraries worldwide; publishers of about 2000 titles have endorsed its use. Cooperation among peers reduces the cost and increases the reliability of preservation, eliminates the need for backup, and greatly reduces other operator interventions.
A loyal (non-malign) peer participates in the LOCKSS system for two reasons: to achieve regular reassurance that its content agrees with the consensus of the peers holding copies of the same content, and if it does not, to obtain the needed repair. The goal of an attrition adversary is to prevent loyal peers from successfully determining the consensus of their peers or from obtaining requested repairs for so long that undetected storage problems such as natural ``bit rot'' or human error corrupt their content. Other types of resource waste may be inconvenient but have no lasting effect on this system.
In prior work  we defended LOCKSS peers against attacks seeking to corrupt their content. That system, however, remained vulnerable to application-level attrition; about 50 malign peers could abuse the protocol to prevent a network of 1000 peers from auditing and repairing their content.
We have developed a set of defenses, some adapted from other systems, whose combination in a P2P context provides novel and effective protection against attrition. These defenses include admission control, desynchronization, and redundancy. Admission control, effected via rate limitation, first-hand reputation, and effort balancing, ensures that legitimate requests can be serviced even during malicious request floods. Desynchronization ensures that progress continues even if some suppliers of a needed service are currently too busy. Redundancy ensures that the attacker cannot incapacitate the system by targeting only few peers at a time. Our defenses may not all be immediately applicable to all P2P applications, but we believe that many systems may benefit from a subset of these defenses, and that our analysis of the effectiveness of these defenses is more broadly useful.
This paper presents a new design of the LOCKSS protocol that makes four contributions. First, we demonstrate via simulation how our new design ensures that application-level attrition, no matter how powerful the attacker, is less effective than simple network flooding. We do this while retaining our previous resistance against other adversaries. Second, we show that even network-level flooding attacks that continuously prevent all communication among a majority of the peers must last for months to affect the system significantly. Such attacks are orders of magnitude more powerful than those observed in practice . Third, since resource management lies at the crux of attrition attacks and their defenses, we extend our prior evaluation  to deal with numerous concurrently preserved archival units of content competing with each other for resources. Finally, resource over-provisioning is essential in defending against attrition attacks. We show that with a practical amount of over-provisioning we can defend the LOCKSS system from an arbitrarily powerful attrition adversary.
In the rest of this paper, we first describe our application. We continue by outlining how we would like this application to behave under different levels of attrition attack. We give an overview of the LOCKSS protocol, describing how it incorporates each of our attrition defenses. We then explain the results of a systematic exploration of simulated attacks against the resulting design, showing that it successfully defends against attrition attacks at all layers, from the network level up through the application protocol. Finally, we describe how the new LOCKSS protocol compares to our previous work, as well as other related work.
In this section, we provide an overview of the digital preservation problem for academic publishing. We then present and justify the set of design goals required of any solution to this problem, setting the stage for the LOCKSS approach in subsequent sections.
Academic publishing has migrated to the Web , placing society's scientific and cultural heritage at a variety of risks such as confused provenance, accidental editing by the publisher, storage corruption, failed backups, government or corporate censorship, and vandalism. The LOCKSS system was designed  to provide librarians with the tools they need to preserve their community's access to journals and other Web materials.
Any solution must meet six stringent requirements. First, since under U.S. law copyrighted Web content can only be preserved with the owner's permission , the solution must accommodate the publishers' interests. Requiring publishers, for example, to offer perpetual no-fee access or digital signatures on content makes them reluctant to give that permission. Second, a solution must be extremely cheap in terms of hardware, operating cost, and human expertise. Few libraries could afford  a solution involving handling and securely storing off-line media, but most can afford the few cheap off-the-shelf PCs that provide sufficient storage for tens of thousands of journal-years. Third, the existence of cheap, reliable storage cannot be assumed; affordable storage is unreliable [22,38]. Fourth, a solution must have a long time horizon. Auditing content against stored digital signatures, for example, assumes not only that the cryptosystem will remain unbroken, but also that the secrecy, integrity, and availability of the keys are guaranteed for decades. Fifth, a solution must anticipate adversaries capable of powerful attacks sustained over long periods; it must withstand these attacks, or at least degrade slowly and gracefully while providing unambiguous warnings . Sixth, a solution must not require a central locus of control or administration, if it is to withstand concentrated technical or legal attacks.
Two different architectures have been proposed for preserving Web journals. The centralized architecture of a ``trusted third party'' archive requires publishers to grant a third party permission, under certain circumstances, to republish their content. Obtaining this permission involves formidable legal and business obstacles . In contrast, the distributed architecture of the LOCKSS system consists of many individual archives at subscribing (second party) libraries. Readers only access their local library's copy, whose subscription already provides them access to the publisher's copy. Most publishers see this as less of a risk to their business, and are willing to add this permission to the subscription agreement. It is thus important to note that our goal is not to minimize the number of replicas consistent with content safety. Instead, we strive to minimize the per-replica cost of maintaining a large number of replicas. We trade extra replicas for fewer lawyers, an easy decision given their relative costs.
The LOCKSS design is extremely conservative, making few assumptions about the infrastructure. Although we believe this is appropriate for a digital preservation system, less conservative assumptions are certainly possible. Increasing risk can increase the amount of content that can be preserved with given computational power. Limited amounts of reliable, write-once memory would allow audits against local hashes, a reliable public key infrastructure might allow publishers to sign their content and peers to audit against the signatures, and so on. Conservatively, the assumptions underlying such optimizations could be violated without warning at any time; the write-once memory might be corrupted or mishandled, or a private key might leak. Thus, these optimizations still require a distributed audit mechanism as a fallback. The more a peer operator can do to avoid local failures the better the system works, but our conservative design principles lead us to focus on mechanisms that minimize dependence on these efforts.
With the application of digital preservation for academic publishing in mind, we tackle the ``abstract'' problem of auditing and repairing replicas of distinct archival units or AUs (a year's run of an on-line journal, in our target application) preserved by a population of peers (libraries) in the face of attrition attacks. For each AU it preserves, a peer starts out with its own, correct replica (obtained from the publisher's Web site), which it can only use to satisfy local read requests (from local patrons) and to assist other peers with replica repairs. In the rest of this paper we refer to AUs, peers, and replicas, rather than journals and libraries.
In this section we present the adversary we model, our security goals and the framework for our defenses.
Our conservative design philosophy leads us to assume a powerful adversary with several important abilities. Pipe stoppage is his ability to prevent communication with victim peers for extended periods by flooding links with garbage packets or using more sophisticated techniques . Total information awareness allows him to control and monitor all of his resources instantaneously. He has unconstrained identities in that he can purchase or spoof unlimited network identities. Insider information provides him complete knowledge of victims' system parameters and resource commitments. Masquerading means that loyal peers cannot detect him, as long as he follows the protocol. Finally, he has unlimited computational resources, though he is polynomially bounded in his computations (i.e., he cannot invert cryptographic functions).
The adversary employs these capabilities in effortless and effortful attacks. An effortless attack requires no measurable computational effort from the attacker and includes traditional DoS attacks such as pipe stoppage. An effortful attack requires the attacker to invest in the system with computational effort.
The overall goals of the LOCKSS system are that, with high probability, the consensus of peers reflects the correct AU, and readers access good data. In contrast, an attrition adversary's goal is to decrease significantly the probability of these events by preventing peers from auditing their replicas for a long time, long enough for undetected storage problems such as ``bit rot'' to occur.
Severe but narrowly focused pipe stoppage attacks in the wild last for days or weeks . Our goal is to ensure that, in the very least, the LOCKSS system withstands or degrades gracefully with even broader such attacks sustained over months. Beyond pipe stoppage, attackers must use protocol messages to some extent. We seek to ensure the following three conditions. First, a peer manages its resources so as to prevent exhaustion no matter how much effort is exerted by however many identities requesting service. Second, when deciding which requests to service, a peer gives preference to requests from those likely to behave properly (i.e., ``ostensibly legitimate''). And third, at every stage of a protocol exchange, an ostensibly legitimate attacker expends commensurate effort to that which he imposes upon the defenders.
We seek to curb the adversary's success by modeling a peer's processing of inbound messages as a series of filters, each costing a certain amount to apply. A message rejected by a filter has no further effect on the peer, allowing us to estimate the cost of eliminating whole classes of messages from further consideration. Each filter increases the effort a victim needs to defend itself, but limits the effectiveness of some adversary capability. The series of filters as a whole is sound if the cost of applying a filter to the input stream passed through its preceding filter is low enough to permit the system to make progress. The filters include a volume filter, a reciprocity filter, and a series of effort filters.
The volume filter models a peer's network connection. It represents the physical limits on the rate of inbound messages that an adversary can force upon the peer. It is an unavoidable filter; no adversary can push data through a victim's network card at a rate greater than the card's throughput. Soundness requires the volume filter to restrict the volume of messages enough that processing costs at the next filter downstream are low. This condition can be enforced either through traffic shaping or via the low-tech choice of configuring peers with low-bandwidth network cards.
The reciprocity filter takes inbound messages at the maximum rate exported by the volume filter and further limits them by rejecting those sent from peers who appear to be misbehaving. A peer's reciprocity filter favors those of its peers who engage it with requests at the same average rate as it engages them. The filter further penalizes those peers it has not known for long enough to evaluate their behavior. In this sense, the reciprocity filter implements a self-clocking invariant, by which inbound traffic exiting the filter mirrors in volume traffic originated at the peer. Thus on average the number of requests passed to the next filter matches the number of requests inflicted by the peer upon others.
The effort filters focus on the balance of effort expended by the peer and a correspondent peer while the two are cooperating on an individual content audit request. These filters ensure that the computational effort imposed upon a potential victim peer by its ostensibly legitimate correspondent is matched by commensurate effort borne by that correspondent. For example, an attacker can only trick its victim peer into cryptographically hashing large amounts of data by first performing the same hash itself (or other effort equivalent to the same hash). As a result, these filters enforce the invariant that ostensible legitimacy costs the attacker as much as it allows the attacker to inflict on its victim. Furthermore, the effort filters ensure that a peer can detect at a low cost that an attacker has abandoned ostensible legitimacy.
In summary, these filters take an input stream of protocol messages and reduce it to levels consistent with legitimate traffic in terms of volume (volume filter), then in number of individual messages per source (reciprocity filter), and then in effort induced per message (effort filters). Malicious interactions that pass all filters can ultimately affect the victim peer adversely, but are ensured to impose no more than manageable additional burden on the victim peer and are guaranteed to cost the attacker as much burden in the process. The former guarantee is essential for the correct operation of good peers in all cases, whereas the latter is only meaningful when the adversary is resource-constrained.
We show in Section 7.4 that the most effective strategy for effortful attacks is to emulate legitimacy, and that even this has minimal effect on the utility of the system. Effortless attacks, such as traditional distributed DoS (DDoS) attacks, are more effective but must be maintained for a long time against most of the peer population to degrade the system significantly (Section 7.2).
The LOCKSS audit process is a sequence of ``opinion polls'' conducted by every peer on each of its AU replicas. At intervals, typically every 3 months, a peer (the poller) picks a random sample of peers that it knows to be preserving an AU, and invites those peers as voters into a poll. Each voter individually hashes a poller-supplied nonce and its replica of the AU to produce a fresh vote, which the poller tallies. If the poller is outvoted in a landslide (e.g., it disagrees with 80% of the votes), it assumes its replica is corrupt and repairs it from a disagreeing voter. The roles of poller and voter are distinct, but every peer plays both.
The general structure of a poll follows the timeline of Figure 1. A poll consists of two phases: vote solicitation and evaluation. In the vote solicitation phase the poller requests and obtains votes from as many voters in its sample of the population as possible. Then the poller begins the evaluation phase, during which it compares these votes to its own replica, one hashed content block at a time, and tallies them. If the hashes disagree the poller may request repair blocks from its voters and reevaluate the block. If in the eventual tally, after any repairs, the poller agrees with the landslide majority, it sends a receipt to each of its voters and immediately starts a new poll. Peers interleave progress on their own polls with voting in other peers' polls, spreading each poll over a period chosen so that polls on a given AU occur at a rate much higher than that of undetected storage errors.
The outcome of a poll is determined by the votes of the inner circle peers, chosen at the start of the poll by the poller from its reference list for the AU. The reference list contains mostly peers that have agreed with the poller in recent polls on the AU, and a few peers from its static friends list, maintained by the poller's operator.
A poll is considered successful if its result is based on a minimum number of inner circle votes, the quorum, which is typically 10, but may change according to the application's needs for fault tolerance. To ensure that a poll is likely to succeed, a poller invites into its poll a larger inner circle than the quorum (typically, twice as large). If at first try an inner circle peer fails to respond to an invitation, or refuses it, the poller contacts a different inner circle voter, retrying the reluctant peer later in the same vote solicitation phase.
An individual vote solicitation consists of four messages (see Figure 1): Poll, PollAck, PollProof, and Vote. For the duration of a poll, a poller establishes an encrypted TLS session with each voter individually, via an anonymous Diffie-Hellman key exchange. Every protocol message is conveyed over this TLS session, either keeping the same TCP connection from message to message, or resuming the TLS session over a new one.
The Poll message invites a voter to participate in a poll on an AU. The invited peer responds with a PollAck message, indicating either a refusal to participate in the poll at the time, or an acceptance of the invitation, if it can compute a vote within a predetermined time allowance. The voter commits and reserves local resources to that effect. The PollProof message supplies the voter with a random nonce to be used during vote construction. To compute its vote, the voter uses a cryptographic hash function to hash the nonce supplied by the poller, followed by its replica of the AU, block by block. The vote consists of the running hashes produced at each block boundary. Finally, the voter sends its vote back to the poller in a Vote message.
These messages also contain proofs of computational effort, such as those introduced by Dwork et al. , sufficient to ensure that, at every protocol stage, the requester of a service has more invested in the exchange than the supplier of the service (see Section 5.1).
The poller uses the vote solicitation phase of a poll not only to obtain votes for the current poll, but also to discover new peers for its reference list from which it can solicit inner circle votes in future polls.
Discovery is effected via nominations included in Vote messages. A voter picks a random subset of its current reference list, which it includes in the Vote message. The poller accumulates these nominations. When it concludes its inner circle solicitations, it chooses a random sample of these nominations as its outer circle. It proceeds to solicit regular votes from these outer circle peers in a manner identical to that used for inner circle peers.
The purpose of the votes obtained from outer circle voters is to show the ``good behavior'' of newly discovered peers. Those who perform correctly, by supplying votes that agree with the prevailing outcome of the poll, are added into the poller's reference list at the conclusion of the poll; the outcome of the poll is computed only from inner-circle votes.
For each hash computed by the poller for an AU block, there are three possibilities: first, the landslide majority of inner-circle votes (e.g., 80%) agree with the poller; in this case, the poller considers the audit successful up to this block and proceeds with the next block. Second, the landslide majority of inner-circle votes disagree with the poller; in this case, the poller regards its own replica of the AU as damaged, obtains a repair from one of the disagreeing voters (via the RepairRequest and Repair messages), and reevaluates the block hoping to find itself in the landslide majority, as above. Third, if there is no landslide majority of agreeing or disagreeing votes, the poller deems the poll inconclusive, raising an alarm that requires attention from a human operator.
Throughout the evaluation phase, the poller may also decide to obtain a repair from a random voter, even if one is not required (i.e., even if the corresponding block met with a landslide agreement). The purpose of such frivolous repairs is to prevent targeted free-riding via the refusal of repairs; voters are expected to supply a small number of repairs once they commit to participate in a poll, and are penalized otherwise (Section 5.1).
If the poller hashes all AU blocks without raising an alarm, it concludes the poll by sending an evaluation receipt to each voter (with an EvaluationReceipt message), containing cryptographic proof that it has evaluated received votes. The poller then updates its reference list by removing all voters whose votes determined the poll outcome and by inserting all agreeing outer-circle voters and some peers from the friends list (for details see ). The poller then restarts a poll on the same AU, scheduling it to conclude one interpoll interval into the future.
Here we outline the attrition defenses of the LOCKSS protocol: admission control, desynchronization, and redundancy. These defenses raise system costs for both loyal peers and attackers, but favor ostensible legitimacy. Given a constant amount of over-provisioning, loyal peers continue to operate at the necessary rate regardless of the attacker's power. Many systems over-provision resources to protect performance from known worst-case behavior (e.g., the Unix file system ).
In prior work  we applied some of these defenses (such as redundancy and some aspects of admission control, including rate limitation and effort balancing) to combat powerful attacks aiming to modify content without detection or to discredit the intrusion detection system with false alarms. In this work, we combine these previous defenses with new ones to defend against attrition attacks as well.
The purpose of the admission control defense is to ensure that a peer can control the rate at which it considers poll invitations from others, favoring invitations from those who operate at roughly the same rate as itself and penalizing others. We implement admission control using three mechanisms: rate limitation, first-hand reputation, and effort balancing.
Rate Limitation: Without limits on the rate at which they attempt to service requests, peers can be overwhelmed by floods of ostensibly valid requests. Rate Limitation suggests that peers should initiate and satisfy requests no faster than necessary rather than as fast as possible. Because readers access only their local LOCKSS peer, the audit and repair protocol is not subject to end-users' unpredictable request patterns. The protocol can proceed at its own pace, providing an interesting test case for rate limitation.
We identify three possible attacks based on deviation from the necessary rate of polling. A poll rate adversary seeks to trick victims into either decreasing (e.g., through back-off behavior) or increasing (e.g., through recovery from a failed poll) their rate of calling polls. A poll flood adversary seeks, under a multitude of identities, to invite victims into as many frivolous polls as possible to crowd out the legitimate poll requests and thereby reduce the ability of loyal peers to audit and repair their content. A vote flood adversary seeks to supply as many bogus votes as possible to exhaust loyal pollers' resources in useless but expensive proofs of invalidity.
Peers defend against all these adversaries by setting their rate limits autonomously, not varying them in response to other peers' actions. Responding to adversity (inquorate polls or perceived contention) by calling polls more frequently could aggravate the problem; backing off to a lower rate of polls would achieve the adversary's aim of slowing the detection and repair of damage. Kuzmanovic et al.  describe a similar attack in the context of TCP retransmission timers. Because peers do not react, the poll rate adversary has no opportunity to attack. The price of this fixed rate of operation is that, absent manual intervention, a peer may take several interpoll intervals to recover from a catastrophic storage failure.
The poll flood adversary tries to get victims to over-commit their resources or at least to commit excessively to the adversary. To prevent over-commitment, peers maintain a task schedule of their promises to perform effort, both to generate votes for others and to call their own polls. If the effort of computing the vote solicited by an incoming Poll message cannot be accommodated in the schedule, the invitation is refused. Furthermore, peers limit the rate at which they even consider poll invitations (i.e., establishing a secure session, checking their schedule, etc.). A peer sets this rate limit for considering poll invitations according to the rate of poll invitations it sends out to others; this is essentially a self-clocking mechanism. We explain how this rate limit is enforced in the first-hand reputation description below. We evaluate our defenses against poll flood strategies in Section 7.3.
The vote flood adversary is hamstrung by the fact that votes can be supplied only in response to an invitation by the putative victim poller, and pollers solicit votes at a fixed rate. Unsolicited votes are ignored.
First-hand reputation: A peer locally maintains and uses first-hand reputation (i.e., history) for other peers. For each AU it preserves, each peer maintains a known-peers list containing an entry for every peer that has encountered in the past, tracking 's exchange of votes with . The entry holds a reputation grade for , which takes one of three values: debt, even, or credit. A debt grade means that has supplied with fewer votes than has supplied . A credit grade means has supplied with fewer votes than has supplied . An even grade means that and are even in their recent exchanges of votes. Entries in the known-peers list ``decay'' with time toward the debt grade.
In a protocol interaction, the poller and a voter each modify the grade assigned to the other depending on their respective behaviors. If the voter supplies a valid vote and valid repairs for any blocks the poller requests, then the poller increases the grade it assigns to the voter (from debt to even, from even to credit, or from credit to credit) and the voter correspondingly decreases the grade it assigns to the poller. If either the poller or the voter misbehave (e.g., the voter commits to supplying a vote but does not, or the poller does not send a valid evaluation receipt), then the other peer decreases its grade to debt. This is similar to the reciprocative strategy of Feldman et al. , in that it penalizes peers who do not reciprocate. This reputation system thus reduces free-riding, as it is not possible for a peer to maintain an even or credit grade without providing valid votes.
Peers randomly drop some poll invitations arriving from previously unknown peers and from known pollers with a debt grade. To discourage identity whitewashing the drop probability imposed on unknown pollers is higher than that imposed on known indebted pollers. Invitations from known pollers with an even or credit grade are not dropped.
Invitations from unknown or indebted pollers are subject to a rigid rate limit; after it admits one such invitation for consideration, a voter enters a refractory period. Like the known-peers list, refractory periods are maintained on a per AU basis. During a refractory period, a voter automatically rejects all invitations from unknown or indebted pollers. Consequently, during every refractory period, a voter admits at most one invitation from unknown or indebted peers, plus at most one invitation from each of its fellow peers with a credit or even grade.
Since credit and even grades decay with time, the total ``liability'' of a peer in the number of invitations it can admit per refractory period is limited to a small constant number. The duration of the refractory period is thus inversely proportional to the rate limit imposed by the peer on the per AU poll invitations it considers.
If a victim peer's clock could be sped up over several poll intervals then the refractory period could be shortened, increasing the effectiveness of poll flood attacks. The victim would call polls at a faster rate, indebting the victim to its peers and making its invitations less likely to be accepted. However, halving the refractory period from 24 to 12 hours has little effect (see Section 7.4). Doubling the rate of issuing invitations does not affect other peers significantly since the invitations are not accepted. Further, an attack via the Network Time Protocol  that doubles a victim's clock rate for months on end would be easy to detect.
Continuous triggering of the refractory period can stop a victim voter from accepting invitations from unknown peers who are loyal; this can limit the choices of voters a poller has to peers that know the poller already. To reduce this impediment to diversity, we institute the concept of peer introductions. A peer may introduce to others those peers it considers loyal; peers introduced this way bypass random drops and refractory periods. Introductions are bundled along with nominations during the regular discovery process (Section 4.2). Specifically, a poller randomly partitions the peer identities in a Vote message into outer circle nominations and introductions. A poll invitation from an introduced peer is treated as if coming from a known peer with an even grade. This unobstructed admission consumes the introduction such that at most one introduction is honored per (validly voting) introducer, and unused introductions do not accumulate. Specifically, when consuming the introduction of peer by peer for AU , all other introductions of other introducees by peer for AU are ``forgotten,'' as are all introductions of peer for by other introducers. Furthermore, introductions by peers who have entered and left the reference list are also removed, and the maximum number of outstanding introductions is capped.
Effort Balancing: If a peer expends more effort to react to a protocol message than did the sender of that message to generate and transmit it, then an attrition attack need consist only of a flow of ostensibly valid protocol messages, enough to exhaust the victim peer's resources.
Real-world attackers may be very powerful but their resources are finite; markets have arisen to allocate pools of compromised machines to competing uses . Raising the computational cost of attacking one target system both absolutely and relative to others will reduce the frequency of attacks. Our simulations are conservative; the unconstrained adversary has ample power for any attack. But our design is more realistic. It adapts the ideas of pricing via processing  to discourage attacks from resource-constrained adversaries by effort balancing our protocol. We inflate the cost of a request by requiring it to include a proof of computational effort sufficient to ensure that the total cost of generating the request exceeds that imposed on the receiver both for verifying the effort proof and for satisfying the request. We favor Memory-Bound Functions (MBF)  rather than CPU-bound schemes such as ``client puzzles''  for this purpose, because the spread in memory system performance is smaller than that of CPU performance .
Applying an effort filter at each step of a multi-step protocol defends against three attack patterns: first, desertion strategies in which the attacker stops taking part some way through the protocol, having spent less effort in the process than the effort inflicted upon his victim; second, reservation strategies that cause the victim to commit resources the attacker does not use, making those resources unavailable to other, useful tasks; and, third, wasteful strategies in which service is obtained but the result is not ``consumed'' by the requester as expected by the protocol, in an attempt to minimize the attacker's total expended effort.
Pollers could mount a desertion attack by cheaply soliciting an expensive vote. To discourage this, the poller must include provable effort in its vote solicitation messages (Poll and PollProof) that in total exceeds, by at least an amount described in the next paragraph, the effort required by the voter to verify that effort and to produce the requested vote. Producing a vote amounts to fetching an AU replica from disk, hashing it, and shipping back to the poller one hash per block in the Vote message.
Voters could mount a desertion attack by cheaply generating a bogus vote in response to an expensive solicitation, returning garbage instead of block hashes to waste not merely the poller's solicitation effort but also its effort to verify the hashes. Because the poller evaluates the vote one block at a time, it costs the effort of hashing one block to detect that the vote disagrees with its own AU replica, which may mean either that the vote is bogus, or that the poller's and voter's replicas of the AU differ in that block. The voter must therefore include in the Vote message provable effort sufficient to cover the cost of hashing a single block and of verifying this effort. The extra effort in the solicitation messages referred to above is required to cover the generation of this provable effort.
Pollers could mount a reservation attack by sending a valid Poll message to cause a voter to reserve time for computing a vote in anticipation of a PollProof message the poller never sends. When a voter accepts a poller's invitation, it reserves a block of time in the future to compute the vote. When it is time to begin voting, the voter sets a timeout and waits for the poller to send a PollProof if it has not done so already. If the timeout expires, the voter can reschedule the remainder of the block of time as it pleases. The attack exploits the voter's inability to reallocate the timeout period to another operation by asking for a vote and then never sending a PollProof. To discourage this, pollers must include sufficient introductory effort in Poll messages to match the opportunity cost the voter experienced while waiting for the timeout.
Pollers could mount a wasteful attack by soliciting expensive votes and then discarding them unevaluated. To discourage this we require the poller, after evaluating a vote, to supply the voter with an unforgeable evaluation receipt proving that it evaluated the vote. Voters generate votes and pollers evaluate them using very similar processes: generating or validating effort proofs and hashing blocks of the local AU replica. Conveniently, generating a proof of effort using our chosen MBF mechanism also generates 160 bits of unforgeable byproduct. The voter remembers the byproduct; the poller uses it as the evaluation receipt to send to the voter. If the receipt matches the voter's remembered byproduct the voter knows the poller performed the necessary effort, regardless of whether the poller was loyal or malicious.
Section 7.4 shows how effort balancing fares against all three types of attacks mounted by pollers. We omit the evaluation of these attacks by voters, since they are rendered ineffective by the rate limits described above.
The Filters Revisited: Figure 2 illustrates how the defenses of rate limitation, first-hand reputation, and effort balancing, enforced as serial filters over incoming traffic (see Section 3.3), can protect LOCKSS peers from attrition attackers. Among the peers with an initially good standing ( through ), and maintain a steady balance of requested votes throughout the time interval to . Note that asks for two votes in close succession; this is an instance of a peer expending its ``credit.'' In contrast, requests many more votes in close succession than justified by its grade and is downgraded to the debt grade by the reciprocity filter, eventually becoming subject to the refractory period. behaves with ostensible legitimacy with regards to the rate of invitations it sends, but misbehaves by deserting (e.g., by not supplying correct effort proofs) and, as a result, is downgraded to the debt grade by the effort filter. 's subsequent invitations are subject to the refractory period. is initially unknown and therefore subject to the refractory period, but behaves ostensibly legitimately and is upgraded to even or credit grade, freeing itself from the refractory period. Peers , , , and request many more votes than reasonable and occasionally send simultaneous traffic spikes which exceed link capacity; they are thinned out by the volume filter along with other peers' traffic. These peers, as well as misbehaving peers and , share the same refractory period and therefore only one invitation from them can be accepted per refractory period.
The desynchronization defense avoids the kind of inadvertent synchronization observed in many distributed systems, typically by randomization. Examples include TCP sender windows at bottleneck routers, clients waiting for a busy server, and periodic routing messages . Peer-to-peer systems in which a peer requesting service must find many others simultaneously available to supply that service (e.g., in a read-one-write-many fault-tolerant system ) may encounter this problem. If they do, even absent an attack, moderate levels of peer busyness can prevent the system from delivering services. In this situation, a poll flood attacker may only need to increase peer busyness slightly to have a large effect.
Simulations of poll flood attacks on an earlier version of the protocol  showed this effect. Loyal pollers suffered because they needed to find a quorum of voters who could simultaneously vote on an AU. They had to be chosen at random to make directed subversion hard for the adversary. They also needed to have free resources at the specified time, in the face of resource contention from other peers competing for voters on the same or other AUs. Malign peers had no such constraints, and could invite victims one by one into futile polls.
Peers avoid this problem by soliciting votes individually rather than synchronously, extending the period during which a quorum of votes can be collected before they are all evaluated. A poll is thus a sequence of two-party interactions rather than a single multi-party interaction.
If the survival of, or access to, an AU relied only on a few replicas, an attrition attack could focus on those replicas, cutting off the communication between them needed for audit and repair. Each LOCKSS peer preserving an AU maintains its own replica and serves it only to its local clients. This massive redundancy helps resist attacks in two ways. First, it ensures that a successful attrition attack must target most of the replicas, typically a large number of peers. Second, it forces the attrition attack to suppress the communication or activity of the targeted peers continuously for a long period. Unless the attack does both, the targeted peers recover by auditing and repairing themselves from the untargeted peers, as shown in Section 7.2. This is because massive redundancy allows peers at each poll to choose a sample of their reference list that is bigger than the quorum and continue to solicit votes from them at random times for the entire duration of a poll (typically 3 months) until the voters accept. Further, the margin between the rate at which peers call polls and the rate at which they suffer undetected damage provides redundancy in time. A single failed poll has little effect on the safety of its caller's replica.
In this section we give details about the simulation environment and the metrics we use to evaluate the system's effectiveness in meeting its goals.
We measure the effectiveness of our defenses against the attrition adversary using four metrics:
Access failure probability: To measure the success of an attrition adversary at increasing the probability that a reader obtains a damaged AU replica, we compute the access failure probability as the fraction of all replicas in the system that are damaged, averaged over all time points in the experiment.
Delay ratio: To measure the degradation an attrition adversary achieves, we compute the delay ratio as the mean time between successful polls at loyal peers with the system under attack divided by the same measurement without the attack.
Coefficient of friction: To measure the cost of an attack to loyal peers, we measure the coefficient of friction, defined as the average effort expended by loyal peers per successful poll during an attack divided by their average per-poll effort absent an attack.
Cost ratio: To compare the cost of an effortful attack to the adversary and to the defenders, we compute the cost ratio, which is the ratio of the total effort expended by the attackers during an attack to that of the defenders.
We run our experiments using Narses , a discrete-event simulator that provides facilities for modeling computationally expensive operations, such as computing MBF efforts and hashing documents. Narses allows experimenters to pick from a range of network models that trade off speed for accuracy. A simplistic network model that accounts for network delays but not congestion, except for the side-effects of a pipe stoppage adversary's artificial congestion, suffices for our current focus on application-level effects. Peers' link bandwidths are uniformly distributed among three choices: 1.5, 10, and 100 Mbps, and latencies are uniformly distributed between 1 and 30 ms.
Nodes in the system are divided into two categories: loyal peers and the adversary's minions. Loyal peers are uncompromised peers that execute the protocol correctly. Adversary minions are nodes that collaborate to execute the adversary's attack strategy.
We conservatively simulate the adversary as a cluster of nodes with as many IP addresses and as much compute power as he needs. Each adversary minion has complete and instantaneous knowledge of all adversary state and has a magically incorruptible copy of all AUs. Other assumptions about our adversary that are less relevant to attrition can be found in .
To distill the cost of an attack from other efforts the adversary might have to shoulder (e.g., to masquerade as a loyal peer), in these experiments he is completely outside of the network of loyal peers. Loyal peers never ask his minions to vote in polls and he only asks loyal peers to vote in his polls. This differs from LOCKSS adversaries we have studied before .
We evaluate the preservation of a collection of AUs distributed among a population of loyal peers. For simplicity in this stage of our exploration, we assume that each AU is 0.5 GBytes (a large AU in practice). Each peer maintains 50 to 600 AUs. All peers have replicas of all AUs; we do not yet simulate the diversity of local collections we expect will evolve over time. These simplifications allow us to focus our attention on the common performance of our attrition resistance machinery, ignoring for the time being how that performance varies when AUs vary in size and popularity. Note that our 600 simulated AUs total about 10% of the size of the annual AU intake of a large journal collection such as that of Stanford University Libraries. Adding the equivalent of 10 of today's low-cost PCs per year and consolidating them as old PCs are rendered obsolete is an affordable deployment scenario for such a library. We set all costs of primitive operations (hashing, encryption, L1 cache and RAM accesses, etc.) to match the capabilities of a low-cost PC.
All simulations have a constant loyal peer population of 100 nodes and run for 2 simulated years, with 3 runs per data point. Each peer runs a poll on each of its AUs on average every 3 months. Each poll uses a quorum of 10 peers and considers landslide agreement as having a maximum of 3 disagreeing votes. These parameters were empirically determined from previous iterations of the deployed beta protocol. We set the fixed drop probability to be 0.90 for unknown peers and 0.80 for indebted peers.
We set the fixed drop probability for indebted peers and the cost of verifying an introductory effort so that the cumulative introductory effort expended by an effortful attack on dropped invitations is more than the voter's effort to consider the adversary's eventually admitted invitation. Since an adversary has to try with indebted identities on average 5 times to be admitted (thanks to the admission probability), we set the introductory effort to be 20% of the total effort required of a poller; by the time the adversary has gotten his poll invitation admitted, even if he defects for the rest of the poll, he has already expended on average 100% of the effort he would have, had he behaved well in the first place.
Memory limits in the Java Virtual Machine prevent Narses from simulating more than about 50 AUs/peer in a single run. We simulate 600-AU collections by layering 50 AUs/peer runs, adding the tasks caused by one layer's 50 AUs to the task schedule for each peer accumulated during the preceding layers. In effect, layer is a simulation of 50 AUs on peers already running a realistic workload of AUs. The effect is to over-estimate the peer's busyness for AUs in higher layers and under-estimate it for AUs in lower layers; AUs in a layer compete for the resources left over by lower layers, but AUs in lower layers are unaffected by the resources used in higher layers. We have validated this technique against unlayered simulations in smaller collections, as well as against simulations in which inflated per-AU preservation costs cause similar levels of peer load; we found negligible differences.
We are currently exploring the parameter space but use the following heuristics to help determine parameter values. The refractory period of one day allows for 90 invitations from unknown or indebted peers to be accepted per 90-day interpoll interval; in contrast, a peer requires an average of 30 votes per poll and, because of self-clocking, should be able to accept at least an average of 30 poll invitations per interpoll interval. Consequently, the one-day refractory period allows up to a total of 120 invitations per poll period, four times the rate of poll invitations that should be expected in the absence of attacks.
The probability of access failure summarizes the success of an attrition attack. We start by establishing a baseline rate of access failures absent an attack. We then assess the effectiveness against this baseline of the effortless attacks we consider: network-level flooding attacks on the volume filter in Section 7.2, and Sybil attacks on the reciprocity filter in Section 7.3. Finally, in Section 7.4 we assess against this baseline each of the effortful attacks corresponding to each effort filter.
In each case we show the effect of increasing scales of attack on the access failure probability, and relevant supporting graphs including the delay ratio, the coefficient of friction, and for effortful attacks the cost ratio.
Our mechanisms for defending against an attrition adversary raise the effort required per loyal peer. To achieve a bound on access failure probabilities, one must be willing to over-provision the system to accommodate the extra effort. Over-provisioning the system by a constant factor defends it against application-level attrition attacks of unlimited power (Sections 7.3 and 7.4).
The LOCKSS polling process is intended to detect and recover from storage damage that is not detected locally, from causes such as ``bit rot,'' human error and attack. Our simulated peers suffer such damage at rates of one block in 1 to 5 disk years (50 AUs per disk). This is an aggressively inflated rate of undetected damage, given that, for instance, it is 125-400% the rate of detected failures in Talagala's study of IDE drives in a large disk farm . Experience with the IDE drives in deployed LOCKSS peers covers about 10 times as many disk years but with less reliable data collection; it suggests much lower detected failure rates.
Figure 3 plots access failure probability versus the interpoll interval. It shows that as the interpoll interval increases relative to the mean interval between storage failures, access failure probability increases because damage takes longer to detect and repair. The access failure probability is similar for a 50-AU collection all the way up to a 600-AU collection (we omit intermediate collection sizes for clarity).
For comparison purposes in the rest of the experiments, the baseline access failure probability of for a 50-AU collection and of for a 600-AU collection correspond to our interpoll interval of 3 months and a storage damage rate of one block per 5 disk years. With these parameters, a machine preserving 600 AUs has an average load of 9%, and a machine preserving 50 AUs has a 0.7% average load.
The ``pipe stoppage'' adversary models packet flooding and more sophisticated attacks . This adversary suppresses all communication between some proportion of the total peer population (its coverage) and other LOCKSS peers. During a pipe-stoppage attack, local readers may still access content. The adversary subjects a victim to a period of pipe stoppage lasting between 1 and 180 days. Each attack is followed by a 30-day recuperation period, during which communication is restored to the victim; this pattern is repeated for the entire experiment. To lower the probability that a recuperating peer can contact another peer, the adversary schedules his attacks such that there is little overlap in peers' recuperation periods. We performed experiments with an adversary that schedules his attacks so that all victims' recuperation periods completely overlap, but found that the low-overlap adversary caused more damage, so we present results from the low-overlap adversary.
Figure 4 plots the access failure probability versus the attack duration for varying coverage values (10 to 100%). As expected, the access failure probability increases as the coverage of the attack increases, though the attack covering 70% of the peer population is almost as effective as the 100% attack. In the extreme, the 180 day attack over 100% of the 600-AU collection raises the access failure probability to ; this is within tolerable limits for services open to the Internet.
For attacks between 20% and 60% coverage, the access failure probability peaks at an attack duration of 60 days and decreases for larger durations. The 180 day attack is less damaging for these coverage values because, while the adversary focuses on a smaller number of peers for a longer time, the rest of the peers continue polling. The 30 to 60 day attacks cycle across more victims and interrupt more polls, wasting peers' time and tarnishing their reputations, while 1 to 10 day attacks are too short to interrupt many polls. As the attack coverage grows from 70%, the 180 day attack disables such a significant portion of the network that the peers free of attack have great difficulty finding available peers and the access failure probability increases beyond the 60 day attack.
Figures 5 and 6 plot the delay ratio and coefficient of friction, respectively, versus attack duration. We find that attacks must last longer than 30 days to raise the delay ratio by an order of magnitude. Similarly, the coefficient of friction during repeated attacks that last less than a few days each is negligibly greater than 1. For very long attacks that completely shut down the victim's Internet, the coefficient can reach 6700, making pipe stoppage the most cost-effective strategy for the attrition adversary.
As attack durations grow to 30 days and beyond, the adversary succeeds in decreasing the total number of successful polls. For example, attacks against 100% of the population with a 30 day duration reduce the number of successful polls to the number absent attack. However, the average machine load during recuperation remains within 2 to 3 times the baseline load -- a result of designing the protocol to limit increases in resource consumption while under attack. Fewer successful polls and nearly constant resource consumption for increasing attack durations drives up the average cost of a successful poll, and with it the coefficient of friction.
The reciprocity adversary attacks our admission control defenses aiming to reduce the likelihood of a victim admitting a loyal poll request by triggering that victim's refractory period as often as possible. This adversary sends cheap garbage invitations to varying fractions of the peer population for varying periods of time separated by a fixed recuperation period of 30 days. The adversary sends invitations using poller addresses unknown to the victims. These, when eventually admitted, cause those victims to enter their refractory periods and drop all subsequent invitations from unknown and indebted peers.
Figure 7 shows that these attacks have little effect. The access failure probability is raised to when the duration of the attack reaches the entire duration of our simulations (2 years) for full population coverage and a 600-AU collection. At that attack intensity, loyal peers no longer admit poll invitations from unknown or indebted loyal peers, unless supported by an introduction. This causes discovery to operate more slowly; loyal peers waste their resources on introductory effort proofs that are summarily rejected by peers in their refractory period. This wasted effort, when sustained over years, raises the coefficient of friction by 33%, (much less than the friction caused by pipe stoppage), and raises average machine load from 9% to 11%. The delay ratio is largely unaffected by this adversary. Consequently, the first effect of this adversary, increasing load in loyal peers, is tolerable given a practical level of over-provisioning.
We switch our attention to the other effect of this adversary, namely, the suppression of invitations from unknown or indebted peers, which introductions are intended to mitigate. We have repeated the experiments with 600 AUs, in which the adversary attacks 100% of the peer population, with introductions disabled. Without introductions, the shorter attacks cause a higher coefficient of friction, much closer to pipe stoppage attacks, whereas longer attacks are largely unaffected. For comparison, suppressing introductions for attack durations of 10 days raises the coefficient of friction from 1.03 to 1.16, vs. 1.51 for pipe stoppage; in contrast, suppressed introductions for attack durations of six months raises the coefficient of friction from 1.34 to 1.36, vs. 6700 for pipe stoppage. The absence of introductions does not make this attack markedly worse in terms of load increase.
The major consequence of unknown and indebted invitation suppression without introductions is that victims call polls almost exclusively composed of voters from their friends list, who are more likely to accept a poll invitation from a fellow friend. This reliance increases as the attack lasts longer. It is undesirable because it allows an adversary to predict closely the membership of a poll (mostly the poller's friends), promoting focused poll disruptions. The main function of introductions is thus to ensure the unpredictability of poll memberships.
Note that techniques such as blacklisting, commonly used to defeat denial-of-service attacks in the context of email spam, or server selection  by which pollers only invite voters they believe will accept, could significantly reduce the friction caused by the admission control attack. However, we have yet to explore whether these defenses are compatible with our goal of protecting against subversion attacks that operate by biasing the opinion poll sample toward corrupted peers .
To attack filters downstream of the reciprocity filter, the adversary must get through as fast as possible. We consider an attack by a ``brute force'' adversary who continuously sends enough poll invitations with valid introductory efforts to get past the random drops; such invitations cannot arrive from credit or even identities at the steady attack state, because they are more frequent than what is considered legitimate. Since unknown peers suffer more random drops than peers in debt, the adversary launches attacks from indebted addresses. We conservatively initialize all adversary addresses with a debt grade at all loyal peers. We also give the adversary an oracle that allows him to inspect all the loyal peers' schedules. This avoids his wasting introductory efforts due to scheduling conflicts at loyal peers.
Once through the reciprocity filter, the adversary can defect at any stage of the protocol exchange: after providing the introductory effort in the Poll message (INTRO) by never following up with a PollProof, after providing the remaining effort in the PollProof message (REMAINING) by never following up with an EvaluationReceipt, or not defecting at all (NONE).
Table 1 shows that the brute force adversary's most cost-effective strategy (i.e., with the lowest cost ratio metric) is to participate fully in the protocol; doing so he can raise loyal peers' preservation cost (i.e., their coefficient of friction) to a factor of ( for the large collection, which equates to an average machine load of 21%). To defend against this increase in cost, LOCKSS peers must over-provision their resources by a reasonable amount. The baseline probability of access failure rises to at a cost almost identical to that incurred by the defenders (a cost ratio of ). Fortunately, this continuous attack even from a brute force adversary unconcerned by his own effort expenditure is unable to increase the access failure probability of the victims greatly; the rate limits prevent him from bringing his advantage in resources to bear. Similar behavior in earlier work  prevents a different unconstrained adversary from stealthily modifying content.
We measured the effectiveness of the refractory period in rate limiting poll flood attacks against the brute force adversary that does not defect, since this strategy has the best cost/benefit ratio among the brute force strategies. Figure 8 shows the coefficient of friction during a brute force attack on 50 AUs where the refractory period varies from 1 to 96 hours. With a shorter refractory period, poll invitations from the attacker are accepted by the victims at a greater rate, driving up the coefficient of friction. With the refractory period at one hour, the average machine load at the victim peers is 21%. If only 50 AUs consume 21% of a peer's processing time, an average peer cannot support 600 AUs while under attack. With the refractory period of 24 hours, the peers' average load supporting 50 AUs is only 2%.
On the other hand, the graph shows that lengthening the refractory period beyond 24 hours would not greatly reduce the coefficient of friction. Furthermore, increasing the refractory period decreases the probability of a peer accepting legitimate poll invitations from unknown or indebted peers, since voters accept fewer of these invitations per unit time. A very long refractory period stifles the discovery process of pollers finding new voters and causes increased reliance on a poller's friends list. Similar behavior occurs when introductions are removed from the protocol (see Section 7.3).
Thus a shorter refractory period increases the probability of voters accepting invitations from legitimate, unknown pollers, but it also increases damage during a poll flood attack. Our choice of 24 hours limits the harm an attacker can do while accepting enough legitimate poll invitations from unknown or indebted peers for the discovery process to function.
In the analysis above, we conservatively assume that the brute force adversary uses attacking identities in the debt grade of their victims. Space constraints lead us to omit experiments with an adversary whose minions may be in either even or credit grade. This adversary polls a victim only after he has supplied that victim with a vote, then defects in any of the ways described above. He then recovers his grade at the victim by supplying an appropriate number of valid votes in succession. Each vote he supplies is used to introduce new minions that thereby bypass the victim's admission control before defecting. This attack requires the victim to invite minions into polls and is sufficiently rate-limited to be less effective than brute force. It is further limited by the decay of first-hand reputation toward the debt grade. We leave the details for an extended version of this paper.
In this section we first describe the most significant ways in which the new LOCKSS protocol differs from our previous efforts. We then list work that describes the nature and types of denial of service attacks, as well as related work that applies defenses similar to ours.
The protocol described here is derived from earlier work  in which we covered the background of the LOCKSS system. That protocol used redundancy, rate limitation, effort balancing, bimodal behavior (polls must be won or lost by a landslide) and friend bias (soliciting some percentage of votes from peers on the friends list) to prevent powerful adversaries from modifying the content without detection, or discrediting the system with false alarms. In this work, we target the protocol's vulnerability to attrition attacks by reinforcing our previous defenses with admission control, desynchronization, and redundancy.
Another major difference between our prior work and the protocol described in this paper is our treatment of repair. In the previous protocol voting and repair were separated into two phases. When pollers determined repair was necessary, they requested a complete copy of the document from the publisher, if still available, or from a peer for whom they had previously supplied votes. This had at least three problems. First, pollers requested repairs only when needed, signaling the vulnerability of those pollers' content to an adversary. Second, the repair mechanism was only exercised when content recovery was needed. Mechanisms exercised only during emergencies are unlikely to work . Finally, this left the system more vulnerable to free-riding, since a peer could supply votes but later defect when the poller requested a costly repair. We address all three problems through restructuring the repair mechanism (as described in Section 4.3) to integrate block repairs, including ``frivolous'' repairs, into the actual evaluation of votes.
A third significant difference in the protocol supports our desynchronization defense. In the previous protocol, loyal pollers needed to find a quorum of voters who could simultaneously vote on an AU. Instead, the poller now solicits and obtains votes one at a time, across the duration of a poll, and only evaluates the outcome of a poll once it has accumulated all requisite votes.
Our attrition adversary draws on a wide range of work in detecting , measuring , and combating [2,27,41,42] network-level DDoS attacks capable of stopping traffic to and from our peers. This work observes that current attacks are not simultaneously of high intensity, long duration, and high coverage (many peers) .
Redundancy is a key to survival during some DoS attacks, because pipe stoppage appears to other peers as a failed peer. Many systems use redundancy to mask storage failure . Byzantine Fault Tolerance  is related to the LOCKSS opinion polling mechanism in its goal of managing replicas in the face of attack. It provides stronger guarantees but requires that no more than one third of the replicas are faulty or misbehaving. In a distributed system, such as the LOCKSS system, that is spread across the Internet, we cannot assume an upper bound on the number of misbehaving peers. We therefore aim for system performance to degrade gracefully with increasing numbers of misbehaving peers, rather than fail suddenly when a critical threshold is reached. Routing along multiple redundant paths in Distributed Hash Tables (DHTs) has been suggested as a way of increasing the probability that a message arrives at its intended recipient despite nodes dropping messages due to malice  or pipe stoppage .
Rate limits are effective in slowing the spread of viruses [43,48]. They have also been suggested for limiting the rate at which peers can join a DHT [6,47] as a defense against attempts to control part of the hash space. Our work suggests that DHTs will need to rate limit not only joins but also stores to defend against attrition attacks. Another study  suggests that the increased latency this causes will not affect users' behavior.
Effort balancing is used as a defense against spam, which may be considered an application-level DoS attack and has received the bulk of the attention in this area. Our effort balancing defense draws on pricing via processing concepts . We measure cost by memory cycles [1,14]; others use CPU cycles [4,15] or even Turing tests . Crosby et al.  show that worst-case behavior of application algorithms can be exploited in application-level DoS attacks; our use of nonces and the bounded verification time of MBF avoid this risk. In the LOCKSS system we avoid strong peer identities and infrastructure changes, and therefore rule out many techniques for excluding malign peers such as Secure Overlay Services .
Related to first-hand reputation is the use of game-theoretic analysis of peer behavior by Feldman et al.  to show that a reciprocative strategy in admission control policy can motivate cooperation among selfish peers.
Admission control has been used to improve the usability of overloaded services. For example, Cherkasova et al.  propose admission control strategies that help protect long-running Web service sessions (i.e., related sequences of requests) from abrupt termination. Preserving the responsiveness of Web services in the face of demand spikes is critical, whereas LOCKSS peers need only manage their resources to make progress at the necessary rate in the long term. They can treat demand spikes as hostile behavior. In a P2P context, Daswani et al.  use admission control (with rate limiting) to mitigate the effects of a query flood attack against superpeers in unstructured file-sharing peer-to-peer networks.
Golle and Mironov  provide compliance enforcement in the context of distributed computation using a receipt technique similar to ours. Random auditing using challenges and hashing has been proposed [9,47] as a means of enforcing trading requirements in some distributed storage systems.
In DHTs waves of synchronized routing updates caused by joins or departures result in instability during periods of high churn. Bamboo's  desynchronization defense using lazy updates is effective.
We have three immediate goals for future work. First, we observe that although the protocol is symmetric, the attrition adversary's use of it is asymmetric. It may be that adaptive behavior of the loyal peers can exploit this asymmetry. For example, loyal peers could modulate the probability of acceptance of a poll request according to their recent busyness. The effect would be to raise the marginal effort required to increase the loyal peer's busyness as the attack effort increases. Second, we need to understand how our defenses against attrition work in a more dynamic environment, where new loyal peers continually join the system over time. Third, we need to consider combined adversary strategies; an adversary could weaken the system with an attrition attack in preparation for some other type of attack.
The defenses of this paper equip the LOCKSS system to resist attrition well. First, application-level attrition attacks, even from adversaries with no resource constraints and sustained for two years, can be defeated with reasonable over-provisioning. Such over-provisioning is natural in our application, but further work may significantly reduce the required amount. Second, the strategy that provides an unconstrained adversary with the greatest impact on the system is to behave as a large number of new loyal peers. Third, network-level attacks do not affect the system significantly unless they are (a) intense enough to stop all communication between peers, (b) widespread enough to target all of the peers, and (c) sustained over months.
Digital preservation is an unusual application, in that the goal is to prevent things from happening. The LOCKSS system resists failures and attacks from powerful adversaries without normal defenses such as long-term secrets and central administration. The techniques that we have developed may be primarily applicable to preservation, but we hope that our conservative design will assist others in building systems that better meet society's need for more reliable and defensible systems.
Both the LOCKSS project and the Narses simulator are hosted at SourceForge, and both carry BSD-style Open Source licenses. Implementation of this protocol in the production LOCKSS system is in progress.
This work is supported by the National Science Foundation (Grant No. 9907296) and by the Andrew W. Mellon Foundation. Any opinions, findings, and conclusions or recommendations expressed here are those of the authors and do not necessarily reflect the views of these funding agencies.
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