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Flp Impossibility Argumentative Essay

The PODC Influential Paper Award was created to acknowledge an outstanding paper on the principles of distributed computing, whose significance and impact on the theory and/or practice of distributed computing has been evident for at least a decade. The following paper received this award at PODC 2001.

Michael J. Fischer, Nancy A. Lynch and Michael S. Paterson, “Impossibility of Distributed Consensus with One Faulty Process,” Journal of the ACM, April 1985, 32(2):374-382.

Members of the selection committee were Richard Ladner, Yoram Moses, Michael Saks, and Nir Shavit (chair). The following descriptions of the winning paper’s contributions were written by Jennifer Welch and Vassos Hadzilacos (presented here with minor editing by Nir Shavit). Jennifer Welch writes the following:

The result of this paper (commonly known as FLP) is that, surprisingly, it is impossible for a set of processors in an asynchronous distributed system to agree on a binary value, even if only a single processor is subject to an unannounced crash. Although the result was motivated by the problem of committing transactions in distributed database systems, the proof is sufficiently general that it directly implies the impossibility of a number of related problems, including consensus.

This result has had a monumental impact in distributed computing, both theory and practice. Systems designers were motivated to clarify their claims concerning under what circumstances the systems work.

On the theory side, people have attempted to get around the impossibility result by changing the system assumptions or the problem statement. Work on changing the system assumptions includes the study of partially synchronous models and of various kinds of failure detectors. Modified problem statements include randomized algorithms, approximate agreement, k-set agreement, and condition-based approaches.

The proof technique used in FLP, valency arguments, has been used and adapted to show many other impossibility and lower bound results in distributed computing. These include impossibility results for consensus, k-set consensus, and renaming in various models, and lower bounds on contention and on the number of rounds for synchronous consensus.

The FLP result forms the basis of work on the wait-free hierarchy, in which data types are classified and compared according to the maximum number of processes for which they can solve wait-free consensus. The calculation of consensus numbers relies on valency arguments.

Finally, work on applying ideas from topology to fault-tolerant distributed computing were inspired by the posing of the k-set consensus problem, which in turn was inspired by the FLP result.

Vassos Hadzilacos writes the following:

This paper proved the most fundamental result in fault-tolerant distributed computing. The result asserts that a particularly simple and natural problem, consensus, is not solvable in the very attractive asynchronous model of computation, even under strong assumptions about the number and nature of failures that may occur – specifically, even if communication is perfectly reliable, if at most one process can fail, and if it can do so only by crashing.

This result has been extremely influential both because of its implications and because of the proof technique it pioneered. Its importance stems from the fact that consensus lies at the heart of many practical problems, including atomic commit, leader election, atomic broadcast, and the maintenance of consistent replicated data. The unsolvability result has motivated researchers to explore models that retain as many of the asynchronous model’s attractive features as possible, while making consensus (and related practical problems) solvable. These explorations include randomization, partially synchronous models, and unreliable failure detectors. The proof technique introduced in the nominated paper, now known as “the bivalence argument”, is notable for its originality, its elegance, and its wide use to obtain other impossibility results in distributed computing. The paper is written masterfully, and although it contains an unusually deep result, it is accessible even to advanced undergraduate computer science students.

The following is a brief historical perspective of the paper written by the authors themselves:

Nancy Lynch and Mike Fischer began working on the distributed consensus problem in early 1980 and proved a lower bound on the number of rounds needed to reach agreement in a synchronous distributed setting.

During 1981-82, Lynch worked with Danny Dolev and others on an approximate version of the consensus problem. For this problem, it turned out that the algorithms for the synchronous setting extended easily to the asynchronous setting, simply by adding enough extra processes to accommodate the differences in process views due to asynchrony. This observation led her to wonder whether synchronous algorithms for exact agreement could similarly be adapted to the asynchronous setting. She made no progress on such an adaptation, however.

Independently, Fischer was introduced to the problem of asynchronous consensus by Butler Lampson during a visit to Xerox PARC in early summer 1982. Lampson suggested an algorithm to solve the problem, and Fischer was anxious to understand it and to prove it correct. Although he made some progress on the former, he failed utterly on the latter.

Later that summer, Fischer, Lynch, and Mike Paterson tried to resolve this problem during a visit by Fischer to MIT, a subsequent visit by Paterson to Yale, and some phone calls. We worked simultaneously on trying to produce an algorithm and trying to prove that no such algorithm exists. Only gradually did we come to realize that the problem is fundamentally unsolvable.

The intuition behind the impossibility proof is pretty simple: Initially, either decision, 0 or 1, is possible. Assume a correct algorithm. At some point in time the system as a whole must commit to one value or the other. That commitment must result from some action of a single process. Suppose that process fails. Then there is no way for the other processes to know the commitment value; hence, they will sometimes make the wrong decision. Contradiction!

This simple argument turned out to be surprisingly difficult to make precise. The subtlety of the argument became apparent to us when we tried to explain the result to others: many people expressed skepticism and disbelief. It took a lot of time, and careful polishing of the proof on our part, before the correctness of the result became generally accepted.

Little did we imagine how influential the paper would turn out to be! We assumed that the main value of our impossibility result was to close off unproductive lines of research on trying to find fault-tolerant consensus algorithms. But much to our surprise, it opened up entirely new lines of research. There has been analysis of exactly what assumptions about the distributed system model are needed for the impossibility proof. Many related distributed problems to which the proof also applies have been found, together with seemingly similar problems which do have solutions. Eventually a long line of research developed in which primitives were classified based on their ability to implement wait-free fault-tolerant consensus.

We thank the Awards Committee for the recognition and the prize. We especially thank the PODC community for the subsequent research which allowed our paper to be deemed influential.

This page maintained by Gil Neiger.

Last modified: January 23, 2003

In this lecture we'll see the,

um, FLP proof of the impossibility of consensus

in asynchronous distributed systems.

So consensus is impossible to solve

in the asynchronous distributed system.

This is a result that was proved,

uh, in a now-famous result by Fischer, Lynch and Patterson

in 1983, also known as the FLP result,

using the, uh, first letters of the last names

of the three coauthors of that paper.

Uh, before then, many distributed systems designers

had been claiming 100% reliability,

uh, for their systems,

and this result stopped them dead in their tracks.

A lot of claims of reliability vanished overnight,

and one of the, uh, long-term side effects of this

was the multiple 9s of reliability,

um, offerings that vendors now publicize

for their products and services.

So once again, just to remind you, um,

the asynchronous distributed system has

no bounds on message delays and p-or processing delays

or, uh, clock drift rates.

These might be arbitrarily long or arbitrarily short.

The consensus problem requires each process, uh, p,

uh, to decide on the same value.

So each process p has a state

which consists of the program counter, the registers,

the stack, the local variables, the heap,

anything else that you would consider to be a part of the,

uh, core, uh, dump of the process.

It also has initial, um, an input register xp,

which is initially either 0 or 1.

Different processes might have different,

uh, input register values,

that is, those processes', uh, um, proposal to the group,

and then each process also has an output register yp

which is initially undecided but needs to be set to 0 or 1.

The only constraint is once you set the output register

you cannot change it.

And remember that each,

uh, process has its own output register,

um, and you want to make sure, uh, that the consensus,

uh, uh, protocol at the end, um,

uh, has all the non-faulty processes

set their output variables to be all-0s or-or all-1s.

So you want an all-0s decision

among the non-faulty processes

or an all-1s decision among the non-faulty processes.

Once again, this problem just by itself, just with these, uh,

two, uh, constraints is enough, uh, to solve consensus

because you can just say,

"Well, everyone set their output variables to be 0

all the time, and that solves it,"

but that is not interesting or, uh, useful at all.

So we have the non-triviality clause that says that

at least one initial system state leads to each

of the above outcomes,

meaning that at least one initial system state

leads to an all-0s outcome,

and at least one initial system state

leads to an all-1s outcome.

So let's try to set up the proof.

Uh, for the impossibility proof,

uh, we'll consider a more restrictive system model

and an easier problem.

Well, essentially this is okay because if you can show that

an easier problem is also impossible to solve,

then obviously the consensus problem,

the harder consensus problem,

is easier to solve, is impossible to solve.

Uh, if you also sh-if you show

that in a more restrictive system model,

uh, the consensus problem is, uh, impossible to solve,

then obviously in the less restrictive system model,

um, it's impossible to solve, as well.

So, uh, what do I mean by restrictive system model?

Uh, instead of considering the entire network,

we'll consider the network to be a simple global message buffer.

When a process p sends a message to a process p', message m,

the message m just gets deposited

in the global message buffer.

Subsequently, p' may, uh, call the global message buffer

with, uh, receive,

and this may return either the message m that is waiting for it

or it may return null,

and it may continue returning null, uh, for a while,

or maybe even forever,

because the message m might be delayed for arbitrarily long.

Okay, so we have abstracted out our network to be just this,

uh, global message buffer

that is sitting underneath all the processes.

Uh, we also define, uh, the state of a process,

which we have seen before.

It consists of its, uh, program counter, heap,

registers, stack and everything else,

along with the input and output variables,

but we also define a state for the entire system.

We call this global state a configuration.

Uh, now, uh, the configuration or global state

consists of a collection of states,

one state for each process,

alongside the state of the global buffer itself.

So the state of all the processes

along with the state of the network is,

uh, the global state, or the configuration.

Now we also define an event.

This is slightly different from the Lamport events

you've seen before.

An event consists of, uh,

three steps which are executed atomically, or in one shot.

Uh, the event starts with the receipt of a message

by a process, say a process p.

Then the message is processed by the process p.

Uh, this may change the recipient's state.

Process p's state might change as a result of this,

and p might also resi-decide to send out some messages,

as a result of this receipt,

and those messages that then result,

are then deposited in the global message buffer.

Uh, so all these three, uh, steps put together,

uh, determine an event.

So an event essentially consists

of a process p receiving a message,

uh, processing it and then depositing the resulting,

uh, messages into the global message buffer and then,

uh, that's an event.

Next we'll define a schedule.

A schedule is simply a linear sequence of events,

so one event followed by another event followed by another event,

that's a schedule, okay?

So here on the left is an example of a schedule.

You start with a configuration or a global state;

we label that as c.

An event e' is applied to it.

This means that process p' receives m',

uh, processes it and deposits any resulting messages

into the global message buffer.

That changes the state of process p'.

It also changes the state

of the global message buffer potentially,

and that means the configuration itself has changed,

and it has changed to something else which we label as c'.

A different event, e'',

will change the configuration again similarly to, uh,

another configuration c''.

Now, these two events, e' followed by e'',

is a schedule, because it's a linear sequence of events,

and we call that a schedule s.

When the schedule s is applied on the initial configuration c,

this c here is the same as this c here,

it results in the configuration c''.

Again, this c'' is the same as the c, this c''.

So the left side of the slide

is equivalent to the right side of the slide.

'Kay, so the schedule is, essentially,

a compact way of representing a sequence of events

rather than mentioning each event separately.

So, here is our first, uh, lemma, or our first result.

It says that disjoint schedules are commutative.

What does this mean?

If I have a schedule s1,

consisting of some sequence of events,

and another schedule s2,

consisting of another sequence of events,

if the sets of receiving processes in s1,

remember that each schedule consists of a set of messages

being received as a set of processes,

if I consider all these processes in s1,

which received messages,

and all the processes in s2, that receive messages,

if these two sets are disjoint,

meaning there are no processes in common,

uh, receiving messages in both s1 and s2,

then these schedules are commutative.

In other words, their order can be flipped.

So if you apply s1 first on a configuration c

and then s2 to reach, to reach a state c'',

then in a different scenario,

if you apply s2 first on, uh, c, and then s1,

you would reach the same final configuration,

c'', okay?

So, uh, w-why is this true?

Well, um, this is true because

these are disjoint sets of receiving processes,

and applying s1 or s2 in any order

would result in the same final outcome.

In fact, interweaving s1 and s2

would also result in the same outcome,

which would be the configuration c''.

So earlier, uh, we saw consensus problem here.

We tried to prove the impossibility

about an easier consensus problem where some process,

not just all, but some process, eventually sets its yp variable,

its upper variable, to be 0 or a 1.

'Kay, and also we'll assume that only one process crashes,

but we are free to choose which process, uh, crashes.

Uh, we define configurations to have valences.

Uh, configuration C may have

a set of decision values V reachable from it,

and since we are only considering 0 or 1 decisions,

um, there might be either 2 or 1 decisions reachable from it.

If both the decisions, both,

and all-0s and an all-1s outcome are reachable from it,

then we say that the size of the valency is 2,

and we say that the configuration C is bivalent.

If only one decision is reachable from it,

either a 0, an all-0s decision,

or an all-1s decision, uh, not both, uh,

then the configuration C is said to be, uh,

0-valent or 1-valent, respectively.

Bivalent essentially means

that the configuration C is unpredictable.

That is, it doesn't know which value it's going to reach,

and essentially what we're going to show is that, um,

a system can always be kept in the bivalent state.

That it-that is, it can always be prevented

from ever making a decision

and ever being sure about its decision.

So the FLP proof shows two things.

First, it shows that there is some initial configuration,

the global state, that is bivalent by itself,

and second it shows that there is some sequence,

there is always some sequence of events that happen,

uh, in a system that start from a bivalent configuration

that keeps the system state or the configuration also bivalent.

So, essentially, there is always some, uh,

things that could happen in the system, um, that, uh,

keeps the system moving

from one bivalent configuration to another,

and so prevents the system from ever reaching a decision,

uh, with respect to consensus.

So let's show the first, um, part of this proof,

that's the second lemma.

Uh, i-uh, here we show that

some initial configuration is bivalent.

Well let's assume, uh, that this is not true;

let's prove it by contradiction.

Let's assume that all initial configurations are either

0-valent or 1-valent, okay?

Now, if there are N processes in the system,

there are 2^N positive initial configurations.

Well, why is this?

Well each of the processes can propose either 0 or 1

for its input variable,

and so you have two possibilities for each process,

and so this means that there are 2^N

possible initial configurations.

Now we create a lattice here,

this is, of course, a virtual lattice,

uh, where the nodes in the lattice

are the initial configurations,

so there are 2^N, uh, nodes in this lattice.

This lattice is essentially a hypercube, uh, with dimension N.

uh, we, uh, link two, uh, configurations together,

we join them by an edge, uh,

if they're d-if they differ in the initial xp,

the initial input variable value for exactly one process, okay?

Uh, this means that, uh,

you know, suppose I have, uh, 2 processes, P1 and P2,

uh, then I'm going to have a lattice that has, uh, four, um,

uh, nodes in it, four initial configurations

where the initial values for, uh, the, um, uh, for the, uh,

uh, uh, ini-for the input variables are 0,0,1,0,1,1,

and um, uh, 0,1.

And in this, uh, the 0,0 node is going to be linked

to the 1,0 node because they,

uh, differ in the input variable values for P1 only,

exactly 1 process.

Also, the 1,1, uh, node is going to be linked to the, um,

1,0 node because, uh, these 2 configurations differ

in the input variable values for P2.

And so, essentially, the hypercube in this 2 process case

looks like a square.

The hypercube for the 3 process case

looks like a cube,

and so on and so forth, okay?

Now, essentially, um, this, uh, here we are saying

that each configuration is either 0-valent or 1-valent,

there are no bivalent configurations.

So we tag each configuration with a 0 or a 1

according to its, uh, valency, either 0-valent or 1-valent.

And because there is at least one 0-valent state,

at least one configuration stacked with a 0,

and at least one 1-valent state,

at least one configuration tagged with a 1,

and because everything is connected, uh,

in just one hypercube,

it has to be the case that at least one 0-valent configuration

is linked directly to a 1-valent configuration, okay?

This, you can, uh, imagine the hypercube and, uh,

you will see that this is true.

So this means that these two configurations differ

in the input variables for exactly one process,

say that process is p,

and let's say we consider around

where this process p has crashed;

that is, it is silent throughout.

Both the initial configurations are indistinguishable,

because the only thing that differed

between these configurations is the state of p,

but p has crashed, so as far as the system running is concerned,

p has no effect on it,

but this means that both these initial configurations

are in fact the same.

One of them will, in fact, result in an all-0's outcome,

the 0-valent configuration,

and the other one will result in a one-in an all-1's outcome

because it's a 1-valent configuration.

So this initial configuration,

either one of these two configurations

where p has crashed is, in fact, a bivalent configuration

because it may result in an all-0's decision

or it may result in an all-1's decision.

'Kay, so we have shown

that when you have one process that could crash,

and we can choose which process is the one that crashes,

you can have at least one bivalent configuration

that is the initial configuration in the system.

Okay, so that's the first part of the proof.

Next we need to show that,

um, starting from a bivalent configuration,

there is always another bivalent configuration that is reachable.

Notice that this proof doesn't say

that you can never reach consensus ever.

It says that there is always some way in which

you can be prevented from reaching consensus.

Let the red configuration be a bivalent configuration,

and let, uh, the event e,

which consists of the process p receiving a message m,

that is the global message buffer in the red configuration,

the sum event that is applicable to the initial configuration,

so m is in the global message buffer in the red configuration.

Now let's put our hand on e and prevent e from being applied.

This means that you might be able to still apply

some other events on the red configuration,

and there are some configurations

that you might be able to reach,

starting from this red configuration.

We call that set to be C, okay?

Those are the blue configurations

that are shown in the triangle.

These are the configurations

that are reached without applying the special event e.

why are we not applying e?

You'll see in a moment, there is a reason for it.

Now, if you take any one of these

blue or the red configurations in the-in the triangle,

and you apply the single event e to it,

the special event e to it,

you will reach another dark blue event.

Let that set of events, the dark blue set of events,

be called as D, 'kay?

Once again, D, any event in,

any configuration D is reached by applying the special event e

on any one of the configurations in the triangle.

Now, this is the summary of what we have discussed.

You have the initial bivalent configuration, the red one.

You don't apply the event e to it, and you reach,

and all the possible states, uh,

that are reachable are in the triangle.

You take one of the configurations in the triangle

and you apply the event e to it,

you'll reach a state that is,

or a configuration that is in the set D.

Okay, so we claim that

the set D contains a bivalent configuration, okay,

and, again, the proof here is by contradiction.

If you can show that the state D contains

a bivalent configuration,

then you can show that

there is a sequence that consists of at least one event

that starts from a bivalent configuration, the red one,

that also leads to another bivalent configuration.

Let's assume the contradiction.

Suppose that D only has 0 and 1-valent contradiction, uh,

configurations and no, uh, bivalent ones.

Okay, So there are these states in D,

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