NSDI 2009: Day 2

Evaluation/Correctness

SPLAY: Distributed Systems Evaluation Made Simple (or How to Turn Ideas into Live Systems in a Breeze)

  • Large-scale distributed applications are difficult and costly to develop, deploy, test and tune. Moving from simulations to real deployments must bridge a “simplicity gap”.
  • Often focus on developers’ technical skills rather than the algorithms.
  • Motivated to use real testbeds (PlanetLab, Emulab, ModelNet, idle desktop machines, etc.) for development but these are hard to deploy on, so we need tools to simplify and accelerate development, resource management, deployment and control.
  • Simplify development: could use a high-level language or an abstract network substrate. Simplify deployment: resource selection, deployment and monitoring. Control application: centralised logging, replay churn/network conditions dynamics.
  • SPLAY is intended for easy prototyping, multiple-testbed deployment, distributed systems teaching (hands-on teaching), and using idle resources for distributed systems research.
  • SPLAY architecture: daemons written in C on the testbed machines (testbed agnostic (abstract away the type of testbed and OS) and lightweight). Results are retrieved by remote logging. Daemon instantiates distributed system participant (e.g. BitTorrent client or Chord node).
  • SPLAY language based on Lua (made for interaction with C, bytecode-based with garbage collection). Language close to pseudocode, so that programmers can focus on the algorithm. Programs written at the RPC level. Applications can be run locally for debugging and testing, and have a comprehensive set of libraries (compatible with wrapped C libraries).
  • SPLAY controller may be distributed, and performs deployment and control support. Can for example emulate churn or network packet loss.
  • Resource isolation is a requirement for non-dedicated testbeds (e.g. where you have multiple users or you want to sandbox testbed code). Need to limit the resources used and enforce protection. All access to the system is directed through libraries.
  • SPLAY can do Chord in 59 LOC (+17 for fault tolerance and 26 to use a leafset). Pastry takes 265 LOC. Also did Scribe, SplitStream, WebCache, BitTorrent (420 LOC, much of them for protocol parsing), Cyclon, Epidemic and Trees.
  • Interface for resource selection (web or command-line), using criteria on load, resource availability and location (e.g. from PlanetLab metadata).
  • Chord implementation is validated to provide the correct results (route length, rather than performance) which demonstrates the validity of using a real deployment instead of simulation.
  • SPLAY allows real traces of PlanetLab to be replayed on a testbed (for reproducible experiments). Can aslo use synthetic descriptions in a DSL (say when fractions of nodes join and leave).
  • Evaluated synthetic churn effect on routing performance in a Pastry DHT. Also looked at trace-driven churn (trace from OverNet file-sharing network), with variable speedup factors.
  • Evaluated runtime overhead. Can run up to 1263 nodes of Pastry on a machine with 2GB of RAM before swapping is triggered.
  • Also evaluated robustness for long-running applications (web cache application).
  • Q: how does SPLAY help you to validate simulation results? This is not the goal of SPLAY: instead it helps you to run a real application rather than simulate it. Could you reuse the Lua code? Yes, Lua is a good fit for this.
  • Q: if you get a non-deterministic bug on PlanetLab, it can be hard to reproduce, but does SPLAY help you with this? Don’t solve the distributed snapshot problem, but if the problem is due to churn, SPLAY will help you to reproduce it.
  • Q: does the framework allow querying network conditions (for developing network-aware distributed algorithms)? Don’t have direct access from the present Lua libraries to low-level network conditions, but anything that can be expressed in C can be used, so this could be added.

Modeling and Emulation of Internet Paths

  • How do you evaluate a distributed system? (e.g. a DHT, CDN, P2P….) Set up a topology on which you are going to test the system, then run it on Emulab, which will emulate the network between them.
  • Emulation is repeatable, real-world environment and controllable (dedicated nodes and network parameters).
  • Hitherto all about link emulation: not the paths between nodes.
  • Goal is path emulation: two nodes in the real world (opposite coasts of the US) with a path between them that you want to emulate.
  • Could do link-by-link emulation: need to know topology of the path, the capacities along that path, the queue sizes, and the cross-traffic along the path. Most researchers don’t have that information.
  • Instead abstract the path and do end-to-end emulation. Focus on high-level network characteristics (RTT and available bandwidth).
  • Obvious solution: given a link emulator, turn it into a path emulator by mapping the actions of the link emulator onto high-level characteristics. Link emulators have queues but it’s hard to determine the queue size from the path (set it to a default and we’ll adjust it later). Then parameterise a traffic shaper. And finally induce a fixed delay on every packet (based on measured real-world RTT, to estimate non-queuing delay). Must do this in both directions, which may be independent.
  • Used this on a real link with iperf running bandwidth measurements in both directions. 8.0% error on the forward path and 7.2% error on the reverse path. Room for improvement, but not too bad: good news! Bad news is on RTT: obvious emulation gives a fixed RTT which is an order of magnitude higher than on the real path. Also, for asymmetric paths (6.4Mbps/2.6Mbps): 50.6% error on the forward path (smaller than real path) and 8.5% error on the reverse path.
  • Why are these errors arising? TCP increases congestion window until it sees a loss, but there are no losses until the queue fills up. The large delays we were seeing were queuing delays. More than 200ms of delay was due to queuing delay.
  • Maximum tolerable queuing delay is ((maximum window size / target bandwidth) - base RTT). Total queue size must be <= product of capacity and maximum tolerable delay. Can also calculate a lower bound for the queue size. But this gives a lower bound greater than the upper bound! Solution is to distinguish capacity from available bandwidth. (ABW is the rate that my packets drain from the queue.) Lower bound is independent of capacity, but upper bound grows as capacity grows, so this gives a region in which we can select a viable queue size.
  • How do we separate the emulation of available bandwidth and capacity? Set queue based on constraints rather than a reasonable default. The traffic shaper should emulate the capacity (high fixed bandwidth, not available bandwidth). Delay is the same. Now introduce constant bitrate cross traffic into the queue and sink it after the shaper. Just enough such that only the available bandwidth is available inside the traffic shaper. Again, do this in both directions.
  • Use CBR traffic instead of TCP cross-traffic because TCP cross-traffic backs off. If the cross-traffic backs off, you would see a higher bandwidth than is realistic.
  • But how can CBR traffic be reactive? Approximate aggregate available bandwidth as a function of the number of foreground flows. Change CBR traffic based on flow count.
  • Evaluated the error in emulation (from obvious solution to new emulation): forward 50.6% down to 4.1%; reverse 8.5% down to 5.0%. Delay is less noisy than the real environment, but is in the correct order of magnitude.
  • Tested system with BitTorrent running on it (12 clients, 1 seed). Isolated capacity and queue size changes. Makes a difference when compared to obvious solution.
  • Another principle is modelling shared bottlenecks (details in the paper).
  • Q: RTT was shown as time-series, but should you show the PDF/CDF of RTTs? Do they also follow the real world? Very unlikely to follow what is seen in the real world, constraining in a simple way. Focussing on higher-level aspects that would matter at endpoints.
  • Q: did you compare results to real-world on PlanetLab? Hard to characterise ground-truth on PlanetLab. Used widely, but hard to distinguish between host conditions (overloaded host) between network conditions: this impacts things like BitTorrent applications. See FlexLab paper. Weren’t confident in the ability to do that experiment.
  • Q: there is great value in having repeatable experiments, but a problem with Emulab is parameterisation (too much to choose), and it would be good to come up with a “standard scenario” (”500 node March 2009 settings”)? Difficult to do but have taken steps in that direction. FlexMon did wide-area bandwidth and delay measurements on PlanetLab. Unfortunately, this is hard to scale up.
  • Q: do you have a sense of whether the byte-based queue size is realistic? Talked in terms of packets to make it simpler, but implemented in terms of bytes.

MODIST: Transparent Model Checking of Unmodified Distributed Systems

  • Distributed systems are hard to get right (complicated protocols and code to implement them). With no centralised view of the entire system, large range of failures to tolerate and increasingly large scale.
  • Normally do some kind of randomised testing, but this is low coverage and non-determininstic (very hard to reproduce).
  • MoDist is a model checker for distributed systems. It’s comprehensive. It runs in-situ (unmodified, real implementations). And it’s deterministic (allows replay).
  • Applied to Berkeley DB, Paxos-MPS (in production service for Microsoft’s data centres), and PacificA. Found 35 bugs of which 31 have been found by developers.
  • Look at BDB replication. Based on Paxos (single primary, multiple secondaries): primary can read and write, secondary can only read. On primary failure, secondaries will elect new primary. On duplicate primary, degrade both and re-elect. But there was a bug in the leader election protocol: could lead to a secondary node receiving an request-for-updates message, which is unexpected and causes the secondary to crash.
  • MoDist took about an hour to make the bug show up, and outputs a trace which can be replayed.
  • Goal is to explore all states and actions of the system. Model checking makes rare actions (failures, crashes) appear as often as common ones, which drives you into corner cases.
  • Look at real processes, communicating using messages, which may also be multi-threaded. Set of normal actions (send/recv message or run thread) and rare actions (message delay, link failure, machine crash) which may be injected.
  • Ideal is to explore all actions. But this leads to a combinatorial explosion. Built-in checks for crashes, deadlocks and infinite loops. Can also have user-written checks (local and global assertions). MoDist amplifies the checks that you give to it.
  • Avoid redundancy by exploring one interleaving of independent actions.
  • Challenges are exposing actions, checking timeout code, simulating realistic failures and scheduling actions (avoiding deadlocks and maintaining extensibility).
  • To check a system, must know and control the actions of the system. Previous work required users either to write an application in a special language, or port it into a fake environment. MODEST uses Explode (OSDI 2006) to interlace control needed into the checked system. Fake environment perturbs the system and can introduce false positives.
  • MODIST instead inserts an interpretation frontend that intercepts RPC API calls and sends these to a MODIST backend. Backend interposes on all RPCs and schedules all intercepted API calls. This is transparent and simple (does not perturb the system and cause false positives). Frontend is stateless: all state in the backend. Only frontend is OS-dependent; backend is OS-independent.
  • Frontend intercepts 82 API functions (e.g. networking, thread synchronisation). Most wrappers are simple, either returning a failure or calling the real API function. Each wrapper has an average of 67 lines of code.
  • Timeout checking is complicated by the heavy use of implicit timers (e.g. a comparison with gettime()). Could intercept gettime() but what should be returned? Don’t know what the comparison is (because it isn’t an API call), and previous work to address this was manual.
  • Instead, do a static, symbolic analysis. Observe that time values are used in simple ways (db_timespec, mostly with +, - and sometimes * or /). Also that timeouts are checked in the vicinity of gettime() calls, 12 out of 13 are within a few lines. Static intra-procedural symbolic analysis can discover implicit timers (much simpler than approaches like KLEE).
  • Checked systems of size up to 172.1KLOC, and found 35 bugs. 10 bugs in protocols, 25 in implementations. All bugs were previously unknown.
  • Q: how does it compare to DPOR (designed for multi-core systems)? Work with general distributed systems as well as multithreaded systems. Different kinds of failures in distributed system, so not clear how this can be mapped to a multithreaded system.

CrystalBall: Predicting and Preventing Inconsistencies in Deployed Distributed Systems

  • Errors remain in distributed systems even after the system is deployed (DoS, data loss, loss of connectivity), and manifest themselves as loss of safety properties.
  • Paxos is a fault tolerant protocol for achieving consensus, and is integrated in many deployed distributed systems.
  • Want to use increased computing power and bandwidth to make it easier to find errors in distributed systems. Use spare power to do state exploration that can predict what sequences of user actions might lead to lead to violations of user-specified consistency properties.
  • Compared against classic model checking, which is limited in the depth that can be achieved. Compared against replay-based/live predicate checking (state of the art distributed debugging). Deep online debugging periodically starts state space exploration to find inconsistencies. CrystalBall also does execution steering to avoid states which would lead to inconsistency.
  • High-level overview: service is implemented as a state machine with some amount of local state and several handlers. Runtime manages the timers and messages to/from the network. CrystalBall controller takes information from the state machine (neighbour info and periodic local checkpoints). The user puts safety properties into the controller. Controller invokes consequence prediction algorithm, looking for violations (passed back to controller). Controller can then put event filters (for steering) into the runtime.
  • State space exploration using MaceMC, run enabled handlers in a set of nodes (live code) and check safety properties in each state. Want to reduce state space coverage to predict future inconsistencies quickly. Want to explore all interleavings of network messages (important for race conditions), but not all interleavings of local node actions.
  • Standard hashing does not re-explore system-wide states. Instead use hashing to remove previously explored local actions of nodes. Network messages can still be interleaved.
  • Aim is to increase resilience of deployed systems, with a focus on generic runtime mechanisms. Want to prevent inconsistencies without introducing new ones (do no harm). But this is a hard problem.
  • Execution steering uses “sound filters” which have behaviour equivalent to events that could happen in normal running (e.g. TCP connection broken, UDP packet loss).
  • Model checker might not have enough time to detect inconsistencies. Deal with this using an “immediate safety check”.
  • Evaluated live using 6–100 participants on 25 machines. Implemented in Mace, and used ModelNet to emulated wide-area Internet behaviour. Looked at RandTree, Chord and Bullet’ systems. Found 7 inconsistencies that weren’t found by MaceMC or a few years of debugging.
  • Looked at execution steering for Paxos. Injected 2 bugs and induced live scenarios that violate the Paxos safety property. Ran each bug 100 times, with a random [0, 20] seconds between injections. Avoided violation in 95% of cases.
  • Also evaluated the performance impact. (Download times using Bullet’.) Less than a 5% slowdown: due to the bandwidth spent on shipping checkpoints.
  • Q: could you comment on the CPU overhead of running these state-space explorations? One core is dedicated to this and maxed out. The state machine is single-threaded.
  • Q: as machines get more powerful, you would expect them to handle more load… how much is the amount that should be spent on model checking? Well, we haven’t looked into that, but could parallelise state-space exploration and use multiple cores.
  • Q: how big is the state space in bytes? For 8 levels (depth reached at runtime), need only 600KB, which fits easily in L2 cache.
  • Q: is there any way of quantifying what “relevant” means for the relevant states that you aim to explore? The states that we are exploring are obviously quite relevant, better than random walk.
  • Q: Paxos example is artificial because you have 3 replicas and 2 faults? There were separate failures and a network partition? This failure still shouldn’t have happened and CrystalBall detected it.
  • Q: ditto. Paxos should have used stable storage for the proposed values? This was a bug in the implementation.

Wide-Area Services and Replication

Tolerating Latency in Replicated State Machines Through Client Speculation

  • Want to make fault tolerant systems faster in a distributed environment.
  • Simple service configuration with client and server, exchanging messages. Server has some state. For FT, replicate the service as a replicated state machine. Replicas need to agree on the order in which they execute the client requests. Need to reach consensus.
  • Problem is that RSMs have high latency: must block until enough non-faulty replies are received. Also have to deal with geographic distribution (to avoid correlated failures).
  • Idea is to use speculative execution in the RSM. Speculate before consensus reached. Without faults, any reply will have the consensus value.
  • Client can take a checkpoint of its state and execute speculatively while waiting on consensus. When consensus is reached, commit the state and continue executing. Obviously, if we speculate based on a faulty reply, we can rollback to the speculation point.
  • First reply latency sets the critical path (need speed and accuracy here), not the completion latency. Can therefore achieve high throughput, support good batching, have stability under contention and maintain a smaller number of replicas.
  • What if you make a request while executing speculatively? (Buy a Corvette after speculating that we’ve won the lottery.) Could hold this because it is external output, which gives bad performance (just as bad as before). Could use a distributed commit/rollback protocol but this makes state tracking complex.
  • Instead explicitly encode dependencies as predicates. New message is “buy if win=yes”. Replicas need to log past replies to test these predicates. Local decision at the replicas matches the client.
  • Applied speculative execution to PBFT (PBFT-Client Speculation, PBFT-CS). Move the execution stage to the beginning (reply from the primary). If the primary is non-faulty, it is in a natural place to give a high-quality speculation.
  • Added tentative execution to PBFT-CS, a read-only optimisation, a failure threshold and proved correctness.
  • Evaluated with benchmarks: a shared counter and NFS (running an Apache httpd build). Three topologies (primary-local, primary-remote and uniform). For primary-local on shared counter, the runtime is roughly equal to the non-replicated case (as the network delay is varied). Performs better than PBFT and Zyzzyva (which see run time increase linearly with network delay). Also evaluated with failures (still performs well with 1% failure). Also looked at throughput. If latency is the limiting factor and you’re not operating at peak capacity, then speculative execution is a win. However, if the server is fully loaded (bound by throughput), client speculation will not improve things (and may introduce some overhead).
  • Q: have you thought about adaptively using client speculation under low load and Zyzzyva under high load to improve overall throughput? Either client or the primary could switch off speculation when load is high.
  • Q: how do you deal with RSMs getting into divergent states because of speculation (i.e. by having external operations)? Could fall back to blocking or distributed commit/rollback.

Cimbiosys: A Platform for Content-based Partial Replication

  • Scenario of photo sharing: Alice stores her complete photo collection on her home PC, tags and labels all photos. The metadata is used to route photos to devices and services (public -> Flickr, 5* -> digital photo frame, family -> Mum). She might update these photos (do red-eye reduction) and these should propagate everywhere. She might upload photos from elsewhere (an internet cafĂ©).
  • Two key replication requirements. Not all devices are interested in all data (content-based selection, device-specific filter). Not all devices talk to all other devices (need flexible, arbitrary communication patterns). Also eventual consistencty, incremental updates, conflict detection and resolution, bandwidth efficiency, etc.
  • Item store is a database of XML objects stored on each device. Application only accesses items in the item store on that device, but will then be synchronised by Cimbiosys (through Sync and Comm components).
  • Filtered synchronisation protocol. Target device sends list of its item store and its filter to the source. The source sends each item that is unknown to the device and matches the filter. Then the target adds these items to its store and updates it knowledge.
  • Main contribution is a protocol that has eventual filter consistency (device’s item store = filter(whole collection)) (correctness property), and eventual knowledge singularity (device’s knowledge = single version vector) (performance property). Extensive model checking has confirmed that these properties hold.
  • Concentrate on eventual filter consistency here.
  • Problem of partial sync: how can you sync from laptop (family photos) through Facebook (public photos only) to the home PC (all photos)? Facebook just gets public family photos.
  • Needed to invent “Item-Set Knowledge”. Knowledge is the set of versions known to device: either stored, obsolete or out-of-filter. The knowledge fragment is a set of items (or * for all items) and a version vector. A device knows about all versions in the version vector for all items in the set of items.
  • By acquiring knowledge using the sync protocol, you receive items at most once.
  • What happens when an item is updated such that it moves out of a filter? (Remove family keyword from a member of the family who’s been disowned.) Need to add a new type of message to the sync protocol: a move-out notification. Send one of these if the version is unknown to the target device and the item does not match the target’s filter and (optionally) if the item is stored at the target. Don’t want to send too many notifications for items that were never stored (not correctness but performance issue).
  • Also need to handle move-out notification chains. Send if the the source’s filter is as broad as the target’s, the source is as knowledgeable as the target and the source does not store an item stored at the target device.
  • Another problem is filter changes. Could just create a new empty collection. But that would be user-unfriendly. Want to discard things outside the new filter, retain the intersection of the old and new filters and acquire things not inside the old filter. So use a knowledge retraction protocol.
  • Evaluated based on a C# implementation (confirmed in Mace), by looking at the average number of inconsistent items per replica.
  • Q: what are the semantics of filter? Does it have any compositional properties? Only requirement on the filter is that it can make a binary decision based on an item and its metadata. How do you make the filters consistent? No notion of consistency for a filter.
  • Q: to what degree do failures or data corruption fit into this protocol? Don’t do anything about Byzantine failures; assume that each device has persistent storage.
  • Q: what is the trust model here, and how do you deal with not-fully-trusted cloud services? We have an access control policy that says what devices can perform what actions on items in the collection.
  • Q: does the originator of the item specify the policy? The owner of the collection specifies the policy. Also use signatures to prove the origin of an item.
  • Q: how is Cimbiosys different from Practi? Practi lets you build protocols as policies. But you would still have to specify policy on top of it (not built into the system).

RPC Chains: Efficient Client-Server Communication in Geodistributed Systems

  • RPC chains are a new distributed computing primitive. Web services run things over a large number of machines, e.g. webmail. Client talks to a frontend server, but that might talk to an authentication server, a storage server, an advert server, etc. The webmail app is a composition of more-basic services. RPC is used to build this composition. (Could be Sun RPC, RMI, SOAP, etc.)
  • If RPC is synchronous, it is neither efficient nor natural. As the number of services involved grows, scalability becomes a challenge. As does heterogeneity. And geodiversity. RPC is rigid and inefficient, so we want to give developers more control, and better tools for doing this.
  • A more natural path would be for the request to flow through each of the backend services.
  • Related work in continuations and function shipping (also mobile agents, active networks, distributed workflows, and MapReduce/Dryad).
  • Assume applications built as a composition of services, where services export service functions via RPC. Services have a fixed location but may be shared. Also assume a single administrative domain (security is less of an issue, orthogonal).
  • RPC has a simple interface which hides the complexity of the distributed environment. Goal is to preserve simplicity while providing finer-grain control.
  • Embed chaining logic as part of an RPC call. Chaining function is portable code that connects services.
  • Implemented as a C# prototype. Chaining functions are implemented as C# static methods, which are stored in a central repository and cached by hosts. Example applications are a storage service and a webmail service.
  • Used an NFS server with an RPC interface. Modified client applications to do a chained copy (third party transfer). NFS server was unmodified. Evaluated it with a client in Redmond and two servers in Mountain View. Chain copy gave a 5x improvement in performance. Chain copy gets a peak throughput of 10.4MB/s against 4.5MB/s for RPC copy.
  • Chaining function enables chains of arbitrary length, dynamic chains and parallel chains. Can also optimise by composing sub-chains recursively (e.g. if a storage server must invoke a backup).
  • Now maintain a stack of chaining functions and execution state. Push the current chaining function and state onto the stack when a subchain starts, and pop it when it ends. Also evaluated chain copy with composition turned on (12–20% savings).
  • Chain splitting and merging. Multiple chains can execute concurrently and services may invoke multiple subchains in parallel. This requires support for chain splitting and merging. Need a way of specifying this. Need a dedicated merge host and merge function.
  • Also issues of debugging and profiling, exceptions, broken chains, isolation (limit damage caused by a bug in the chaining function) and dealing with legacy RPC servers.
  • Q: what happens to worst case performance (presuming you have to set higher timeouts)? Nodes along the path report back to the source to monitor liveness (of the request). This is handled as a broken chain.
  • Q: is the process of adding chains automated/automatable? Not automated at present, but in future work we’ll try to do this. Instead provide an abstraction close enough to RPC that it shouldn’t be difficult.

Botnets

Studying Spamming Botnets Using Botlab

  • Botnets are a big problem (spam, DDoS, phishing often use botnets as the underlying infrastructure). They are a network of compromised or infected machines. However, they are hard to study and hence not well understood. Malware authors go to great lengths of obfuscation.
  • Automate botnet analysis using a black-box approach: execute the binary and study its behaviour. Also want to automate the finding and execution of bot binaries. Needs to be scalable and safe.
  • Peek into botnets using spam: a high-volume activity of botnets. Can collect information about botnets from the spam. Synthesising multiple sources of data in real-time gives accurate, high-fidelity information.
  • Characterise botnets by size, behaviour, types of spam sent, etc. New and better data gives better answers about this and improves our understanding about the botnet ecosystem.
  • Botlab is a system for monitoring botnets. It obtains bot binaries, perform an initial analysis, execute safely, and output data.
  • Malware is traditionaly collected through honeypots. Collected 2000 binaries over a 2-month period from honeypots. Spam botnets are in a new generation that propagates through social engineering. Need to make honeypots perform active crawling to accumulate this. Also use a constant spam feed from UW’s mail servers (2.5 million emails per day of which 90% spam; 200000 email addresses; over 100000 unique URLs each day; 1% point to malicious binaries).
  • Want to discard duplicate binaries, but a simple hash is insufficient (because of obfuscation, repacking, etc.). Use network fingerprinting (DNS lookups, IPs, ports connected to) to build a behavioural profile instead. Also used to find binaries which perform VMM detection (if fingerprint is different between VMM and bare-metal, then you have something that tries to detect VMMs).
  • Want to run bots safely. Botlab should not cause harm. But it must be effective: you can’t just prevent traffic from leaving the system. Botlab drops traffic to privileged ports and known vulnerabilities; limits connection and data rates (to avoid DDoS); and redirects SMTP traffic to a fake mail server.
  • Manual adjustments (e.g. when a bot verifies over SMTP to its C&C server) needed to ensure that the bot functions.
  • Botlab tries to send 6 million emails per day (to hotmail, gmail, yahoo) from about 10 bot instances. Local view of spam producers but global view of spam produced.
  • Combining the spam sources gives a different perspective. Spam is received from almost every bot in the world: this gives a local view of the spam produced but a global view of the spam producers.
  • Want to combine these two sources of data. Observe that spam subjects are carefully chosen: there is no overlap in subjects sent by different botnets (489 different subjects/day/botnet).
  • Who is sending all the spam? 79% from just 6 botnets. 35% from the largest botnet (Srizbi).
  • What are some characteristics of the prominent botnets? Most contact only a small number of C&C servers. Send spam at 20–2000 messages/minute. Mailing lists are 100 million — 1 billion with a maximum overlap of 30% across two botnets. Active size is 16000–130000 nodes.
  • Are spam campaigns tied to specific botnets? Spam campaign defined as the contents of the webpage to which the spam URL points.
  • How does web hosting for spam relate to botnets? Many-to-many relationship between campaigns and web hosts. Does spam from a single botnet point to a single set of web servers? No, which suggests that hosting spam campaigns is a 3rd party service not tied to botnets. 80% of spam points to just 57 web server IPs.
  • Could use Botlab as a real-time database of bot-generated spam (including web links), which could be used for safer web browsing and better spam filtering.
  • Q: do botnets apply a sophisticated mechanism to back-off under user activity? Yes, some clever botnets try not to inconvenience you at all (Srizbi would look for mouse activity), but we weren’t actively using machines, so they worked at full throughput.
  • Q: what would the real sending rate be based on the presence of user activity? How would this affect the rate? Have not looked at this as it would depend on user behaviour.
  • Q: since number of botnets is small, could a botnet have several different masters? Based on the C&C servers, which imply the control structure of the botnet.
  • Q: who is providing the web hosting (as surely they would have a huge bandwidth provision)? We have information about this, but they serve static pages of constant size, and the bandwidth requirement is based on the small fraction who actually read the pages. Many hosted in South Korea.
  • Q: do bots contact their C&C server for reactivation? Not really, more for new data etc., except when they crash and require reactivation.

Not-a-Bot: Improving Service Availability in the Face of Botnet Attacks

  • [Firefox crashed and lost my notes on this! Fortunately, other people are also blogging!]
  • Basic idea was to attest to real human interaction, using a small trusted attester. Attester is separated from the untrusted OS and applications using virtualisation (Xen disaggregation in this case). TPM sealing used to protect a key that is used for attestation, and stored in memory in the attester. Human interaction can trigger an attestation which may be used to sign e.g. emails or other interactions (clicks on adverts, votes, etc.). This gives a small window of opportunity in which a botnet could send fraudulent clicks, but this is pretty short, and the evaluation shows that it cuts down on the amount of spam that may be sent or fradulent clicks that may be registered.
  • Q: if you can outsource CAPTCHA, couldn’t you outsource TPMs? No real benefit from this.
  • Q: a viable commercial OS needs to let you install third-party device drivers, so even if you assume trustworthy drivers, wouldn’t some drivers still need to be able to generate keystrokes? (Also, what about remote access to a machine?) Virtualisation is a convenient means to bootstrap this process. Could equally use trusted path techniques from Intel and AMD.

BotGraph: Large Scale Spamming Botnet Detection

  • Issue is web-account abuse attack. Zombie hosts can perform automated signup using CAPTCHA solvers. These accounts can then be used to send out spam.
  • Want to detect abused accounts based on Hotmail logs. Input is user activity traces (signup, login, email-sending), and goal is to stop aggressive account signup to limit outgoing spam.
  • At present, the attack is stealthy and large-scale. We need low false positive and false negative rates.
  • Designed a graph-based approach to detect attacks. Identifies a user-user graph to capture bot-account correlations. Identified 26 million bot accounts with a low false positive rate. Implemented using Dryad/DryadLINQ on a 240-machine cluster.
  • Graph of signup history for a particular IP. Notice a spike in the signup count. Could use exponential weighted moving average algorithm to predict future. Where there is a large prediction error, you have an anomalous window, during which we suppose malicious accounts are being created.
  • Can detect stealthy accounts using graphs. Observe that bot accounts work collaboratively. Normal users share IP addresses in a single AS with DHCP assignment. Bot users are likely to share different IPs across ASs. Bot users form a giant connected component while normal users do not. Can use random graph theory to detect this. So detect giant connected components from the user-user graph, then use a hierarchical algorithm to identify correct groupings. Then prune normal-user groups (e.g. due to cell phone users, Facebook applications, etc.).
  • Increase edge weight threshold until the connected component breaks up.
  • Implemented in parallel on DryadLINQ. EWMA-based signup abuse detection can partition data by IP and achieve real-time detection. The user-user graph construction uses two algorithms and some optimisations and can process 200–300GB of data in 1.5 hours on 240 machines. Connected component extraction using divide-and-conquer in just 7 minutes.
  • Graph construction by selecting ID group by IP (map phase), then generating potential edges (reduce phase). Then select IP group by ID pair (map) and calculate edge weight (reduce). Problem is that the number of weight-1 edges is two orders of magnitude larger than other weights. Their computation/communication is unnecessary.
  • Second algorithm does selective filtering, which saves transferring weight-1 edges between nodes. Also used optimisations for compression and broadcast (and the Join functionality).
  • Evaluated detection on two datasets (Jun 2007 and Jan 2008). Three types of data: signup, login and sendmail logs. Bot IPs went from 82k to 241k between datasets (user accounts from 4.83 million to 16.41 million). The anomaly window shrank from 1.45 to 1.01 days.
  • Validated using a manual check (sample groups sent to Hotmail team; almost no false positives). Also through comparison with known-spammers (detected 86% of complained-about accounts, and 54% of the detected accounts are new findings). False positive rate is very low.
  • How can you evade BotGraph? Be stealthy in your signups (sign up less). Also fix a binding to IP or AS, which lowers your utilisation rate. Accounts bound to a single host will easily be grouped. Or send few emails (like a normal user). All these approaches limit spam throughput.
  • Q: what is the relationship between this and Sybil-detection using a random walk? Not sure if random walk can be used to detect spam. Why do bots have to communicate with each other? Graph is just a way to cluster the bots.
  • Q: false positive rate of 0.5% on tens of millions of accounts seems like a concern? Absolute value is pretty large. Quite conservative; real false positive rate may be lower. Could also be a starting point from which more sophisticated and costly approaches could be used.

Network Management

Unraveling the Complexity of Network Management

  • Enterprise networks are complicated: topologies, diverse devices, tweaked network configuration and diverse goals.
  • Example of a configuration change: adding a new department with hosts spread across three buildings. Need to reconfigure routers, and an error can lead to outages or loopholes.
  • Complexity leads to misconfiguration. There is no metric that captures this, and it’s difficult to reason about the difficulty of future changes, or for selecting between possible changes.
  • Defined a set of complexity metrics, based on an empirical study of complexity in 7 networks. Metrics were validated by questionnaire sent to network operators (in public and private enterprises). Questionnaire had tasks to quantify complexity, either network-specific or common to all operators. The metrics focus on layer 3.
  • Complexity is unrelated to size or line count. Largest mean file size was a “simple” configuration. So was highest number of routers.
  • Implementation complexity (referential dependence and different roles for routers) and inherent complexity (uniformity). Inherent complexity provides a lower bound for implementation complexity.
  • Referential dependency metric. Look at referential graph in the stanza of the configuration file. (e.g. Router stanza contains line that references an interface stanza.) Could have intra- and inter-file links. Inter-file links correspond to global network symbols, e.g. the subnet and VLANs. Operators try to reduce dependency chains in their configurations so as to have few moving parts (dependencies). Metric should capture the difficult of setting up layer 3 functionality and the extent of dependencies.
  • Metric is the number of referential links, normalised by the number of devices. Greater number of links implies higher complexity.
  • Another metric is the number of routing instances (i.e. a partition of routing protocols into largest atomic domains of control).
  • Largest network (83 routers) has only 8 average ref links per router, which is low (simple).
  • Gave operators a task of adding a new subnet at a randomly chosen router. Metric was monotonically increasing but not absolute.
  • Inherent complexity: policies determine a network’s design and configuration complexity. Where policies are uniform, this is easy to configure, but special cases make it hard to configure. Challenge was to mine implemented policies and quantify similarities and consistency.
  • Policies were captured with reachability sets (i.e. the set of packets allowed between 2 routers). These imply a connectivity matrix between routers, which is affected by data/control plane mechanisms. Get a uniformity metric, which is the entropy of reachability sets. Simple policies show entropy values close to ideal. Simple policies have filtering at higher levels. Also discovered a bug in one configuration because it had close-to-ideal entropy when it should not have.
  • Some networks are simple, but most are complex. Most networks studied have inherently simple policies (so it was more implementation complexity).
  • In one network, get a high referential link count due to dangling references (to interfaces). Another was complex because the network was in the middle of a restructuring.
  • Future work to look at ISP networks, and consider absolute versus relative complexity.
  • Q: did operators introduce the right kind of complexity? Yes, they knew what they were doing. Does metric help them or did they know they were in trouble? They knew they were in trouble.
  • Q: is there a reason for normalising by the number of devices? This helps compare between two networks of different sizes. However, it’s a first attempt, and this may be refined.
  • Q: have you thought about the complexity of provisioning against the runtime complexity? Something might be easy to provision but could go wrong horribly if it fails? Not looked at that yet.

NetPrints: Diagnosing Home Network Misconfigurations Using Shared Knowledge

  • Typical home network has multiple devices connected to a (wireless) cable modem, which connects to the internet. You might be running diverse applications and have diverse firewall and security requirements. Very heterogeneous. Worse, there is no network administrator, so how do you manage it?
  • Looked at examples of problems faced in home networking. Some caused by home router misconfiguration, some by end-host misconfiguration and some by remote-host misconfiguration (that may nevertheless be solved locally, e.g. by changing MTU).
  • Users take on average 2 hours to solve these problems. NetPrints is network problem fingerprinting. It automates problem diagnosis using “shared knowledge”. NetPrints subscribers occasionally submit network configuration information to the NetPrints service. On receiving a problem notice, the NetPrints service can suggest a solution.
  • In context with rule-based techniques: these are too application specific (require too many rules). Also local configuration issue resolvers (like Autobash, etc.). NetPrints combines these approaches. It has to deal with unstructured, heterogeneous environments, and solve problems due to the interaction of multiple configurations.
  • Two basic assumptions: connectivity is available (application-layer problems only; knowledge base could be shipped offline, however), and we don’t look at performance (only “good” and “bad” states).
  • Example of user tries a VPN connection from home, and it fails. Enters application name into NetPrints, configurations are scraped, and this information is shipped off to the NetPrints server. NetPrints suggests a fix, and the client applies it directly.
  • Three diagnosis strategies: snapshot-based (collect snapshots from different users), change-based (collect the configuration changes that a user makes) and symptom-based (collect signatures of problems from network traffic).
  • NetPrints has two operating modes: normal (collecting information) and diagnose mode (when people complain).
  • The configuration scraper scrapes from the router (using UPnP to get basic information, and the web-based interface (HTTP Request Hijacking)), the end-host (interface-specific parameters, patches, software versions and firewall rules) and the remote system (composition of local and remote configurations).
  • Server knowledgebase stores per-application decision trees with the popular C4.5 decision tree learning algorithm.
  • Evaluation methodology: testbed with 7 different wireless routers, clients running the VPN client sent configuration information to the NetPrints service (6000 configuration parameters per snapshot), and then service learned these using C4.5.
  • Configuration tree is not too large for the applications we have seen.
  • Configuration mutation can be done automatically by walking the decision tree. However, don’t want useless advice (e.g. change your router manufacturer) when a soft parameter could be changed. Therefore track the frequency of configuration changes and use this to inform what the cost of a change would be.
  • Another technique that uses network traffic signatures and change trees to diagnose other problems.
  • Evaluated in three different scenarios. First: VPN client in home network talks to VPN server outside. Second: want connect from outside to an FTP server inside. Third: file sharing within the home network.
  • Found some intuitive inferences (need pptp_pass=1 for VPN to work), and some surprising inferences (to do with the stateful firewall being off).
  • NetPrints uses labelled data to learn its knowledgebase. a 13–17% mislabelling causes only a 1% error in diagnosis.
  • Q: what if you have a problem that none of the existing decision trees have seen? Did you consider merging decision trees? If you don’t have enough data, you should be able to use the knowledgebase of a similar application to solve the problem. Currently looking at calculating the similarity of applications.
  • Q: what if there are user-specific constraints that the decision tree would suggest you break? No user-specific weights in the current application. Could send back multiple suggestions (with weights) and combine them with local user policy.
  • Q: do you have a sense of whether the trees would still apply if you looked at other problems (beyond connectivity management)? Could have, e.g., an application that fails only with certain inputs, but not with others. Haven’t faced that yet, but intuition is that the trees would still be pretty small.
  • Q: what happens when the user fails to report something that could be crucial? System is limited to configuration that we can capture. Doesn’t deal with transient outages. Could look at work from CHI?

Green Networked Systems

Somniloquy: Augmenting Network Interfaces to Reduce PC Energy Usage

  • Power and energy efficiency are key drivers today. How can I make my laptop last longer? How can I lower the power consumption of my desktop machine (environmental impact/cost)?
  • IT equipment consumes significant power, but shutdown opportunities are rarely used. 67% of office PCs are left on after work hours (sleep and hibernate modes used in less than 4% of these PCs). Home PCs left on for 34% of the time; 50% of the time not being used. Also a problem at CSE@UCSD.
  • People leave machines on to maintain state, occasional remote access (SSH/VNC; administrator access) and active applications running all night (long downloads, maintaining IM presence).
  • Occasional access and active applications cannot be handled by sleep modes (maintaining state obviously can).
  • Hosts are today either active/awake or inactive/asleep. Power consumption is 100x in awake compared to asleep. The network assumes that hosts are always connected. What we really want is something that provides the functionality of an awake host while only consuming the power of sleep mode.
  • Want to do this with availability across the entire protocol stack without making changes to the infrastructure or user behaviour. Can we achieve this with a low-power secondary processor (low-power CPU, DRAM and flash memory storage)? Want power consumption when secondary processor is active to be ~1W. Could we put this on the NIC?
  • Add a separate power domain to the NIC, powered on when host is asleep, with secondary processor device, and same MAC/IP address as the PC. Can wake up host when needed, but also handle some applications using application stubs on the secondary processor.
  • Stateless applications can be supported using “filters”; compare to Wake-on-LAN (which is either impractical (too many wakeups) or it affects usability (need infrastructure to send special packets to do the waking)). Somniloquy specifies filters at any layer of the network stack.
  • Stateful applications can be supported using “stubs”: need application specific code on the secondary processor, but the secondary processor is limited in resources. Stub is simplified version of the application; stub code is generated manually. Investigating how to do this automatically. Done for BitTorrent, web downloads and IM.
  • Built using the gumstix platform. PXA270 processor with full TCP/IP stack. USB connection to PC for sleep detection/wakeup trigger; power while asleep and IP networking for data. Wired and wireless prototypes. *-1NIC (augmented NIC) and *-2NIC (uses PC-internal interface while awake and simplifies legacy support).
  • Built a physical prototype. Two USB interfaces (one for power and USB networking; other for control).
  • Evaluated to maintain network reachability, and look at stateless and stateful applications.
  • Reachability: 4–5 second outage while desktop goes to sleep or comes back from sleep. This is because sleep transition is not optimised for latency.
  • Look at setup latency for stateless applications: incoming TCP SYN causes wakeup. Additional Somniloquy latency is 3–10s. As a proportion of an interactive session, this is probably OK.
  • Gumstix power consumption is between 290mW (WiFi) and 1W (Ethernet).
  • Lowest power consumption without using Somniloquy (in Dell Optiplex 745) is 93.1W (down from 102.1W in normal idle state). In S3 suspend state, this is just 1.2W. Total power consumption for Somniloquy is 5W. Assuming a 45-hour work week, could save 620kWh per year, which is $56, or some amount of carbon dioxide.
  • Extends laptop battery life from 6 hours to 60 hours. (Power drops from 11W to 1W.)
  • Used desktop PC trace data from next talk, using 24 desktop PCs (ON, sleep, idle, OFF durations). Identified energy savings.
  • Energy savings using stateful applications: web download stub. 200MB flash; download when Desktop PC is asleep. Wake up PC to upload data whenever needed. Uses 92% less energy than using the host PC for the download.
  • Other details in the paper: how to build application stubs, and a model of energy savings validated with real measurements.
  • Q: how is this different from the small screen attached to the laptop? Windows Sideshow. This is not an active device; only shows you information from before the laptop went to sleep. This could be augmented with Somniloquy.
  • Q: why do you have the design constraint of only making changes to the client and not to the network? For individual desktops and laptops, Somniloquy is the way to go. For cost, it might be better to have a dedicated machine, but this adds security issues and other overheads.
  • Q: could you not look at zombie state resume? Unfortunate thing is that currently you have an either/or state, and everything is either on or off.
  • Q: what is the energy cost associated with sleeping (to disk) and resuming (from disk)? Could you end up waking and sleeping more (due to e.g. the wakeup filters for stateless applications), and hence using more applications? Results account for the energy cost in suspend in resume. As long as it’s not once every one or two minutes, we’ll be alright.
  • Q: how much in the way of latency savings could be achieved by integrating the device into the motherboard? Important design point is that you don’t require everything to be on to power the device.

Skilled in the Art of Being Idle: Reducing Energy Waste in Networked Systems

  • Idle network systems draw significant power even when idle. Either go into a low-power sleep state (S3) and lose network presence; or remain powered on and waste power.
  • Several proposals have looked at this: Wake-on-(W)LAN, special wakeup packets, Network Connection Proxy. These systems have so far seen little use. But there has been little evaluation of the potential for energy savings, and little exploration of the design space for a network proxy.
  • Is the problem worth solving? What is the design space? What protocols should be handled and how?
  • This work is trace-driven evaluation of the broader benefits and design tradeoffs. Focus on the types of energy savings that can be obtained with simple techniques.
  • Collected traces from 250 Intel host computers (90% laptops, 10% desktops), in the office and at home. Over 4 weeks in spring 2007. Traces contain packet traces, flow information, logs of keyboard and mouse activity, power state, etc.
  • Look at desktop machines and their power states. On average, they are idle for >50% of the time. Therefore desktops waste >60% of their energy while idle. Given there are 170 million desktop PCs in the US, this translates into 60TWh/year wasted (about $6billion).
  • Do we need proxying or should we just Wake-on-LAN? Depends on the time that it takes to wake up and go back to sleep. PCs receive 1–4 packets per second while idle. Optimistic assumption that transition time (sleep-to-wake-to-sleep) is 10s. Office environment does not give any savings.
  • Differentiate between relevant and irrelevant packets, and only wake for relevant? Or respond with some proxy device? Or do complex processing on the proxy? Option between transparent (default wake) or non-transparent (default ignore) solution.
  • Deconstruct by protocol. Some protocols are responsible for “poor sleep”. Enumerate them and decide how to handle them (ignore/respond/wake).
  • Define half-time sleep measures a protocol’s role in preventing sleep. How much can we sleep when waking for protocol P. Compute sleep for discrete transition times. ts_50 (half-time sleep) is the largest transition time for which sleep > 50%. The lower ts_50(P) is for protocol P, the worse an offender P is.
  • Majority of traffic is broadcast or multicast. Mostly useless network chatter. ARP can be handled by simple responses (just know IP address of a particular machine). IPX and NetBIOS packets are probably uninteresting to the machine. Can also ignor HSRP and PIM. IGMP and SSDP can be handled by simple responses.
  • Look now at unicast. Look at TCP and UDP ports. Also TCP keep-alives and TCP SYNs. Some can be handled with easy approaches. But some (SMB/CIFS and DCE/RPC) require special handling because they are used by many applications.
  • Transparent proxies might be good for home, but not office computers. Cannot handle unicast well unless it’s complex. Non-transparent proxies (even simple ones) are very efficient.
  • Architecture is a list of (trigger, action) rules. Agnostic to where the proxy runs (NIC, server running on same LAN, router, firewall). Example implementation as a standalone machine on the same LAN, implemented in Click. Used a simple (non-transparent) set of rules. No explicit state transfer between sleeping machine and the proxy; just learn state by sniffing traffic. So no modifications needed at end systems.
  • Q: where is the line between idleness that can be exploited for future better performance and idleness when we should power down? Have scheduled wakeups when we do things in batches. Even when exploiting idleness, we don’t do it at 100% utilisation. Instead, have periodic spells of high utilisation scheduled.
  • Q: [Missed.] If the machine was asleep to start with, some traffic wouldn’t arise. This solution couldn’t really help.
  • Q: do you think that the problem can be solved completely by proxy, or would it be better to involve applications and protocol development? Not sure what the best answer would be, but it would definitely be better to have protocols that work well with power saving. Have seen examples of bad applications that wake the computer up every couple of minutes.
  • Q: [Missed.] If we had the same periodicity of exchanges, we could just be awake for that, but we are not sure how to design protocols to achieve that.

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