Home < Stanford Blockchain < Stanford Blockchain Conference 2019 < Thundercore

Thundercore

Transcript By: Bryan Bishop

Tags: Consensus

Category: Conference

Thundercore consensus

Elaine Shi

https://twitter.com/kanzure/status/1090773938479624192

Introduction

Synchronous with a chance of partition tolerance. Thank you for inviting me. I am going to be talking about some new updates. It’s joint work with my collaborators. The problem is state-machine replication, sometimes called blockchain or consensus. These terms are going to be the same thing in this talk. The nodes are trying to agree on a linearly updated transaction log.

State-machine replication

We care about consistency and liveness. Consistency is honest nodes agreeing on a log, and liveness is that whenever I make a transaction then it should get into the transaction logs fairly quickly. I don’t want to wait forever for my coffee.

What makes the problem challenging is that the nodes can be compromised and the corrupted nodes can behave arbitrarily. Even under adversarial conditions, we want to satisfy the important security properties as well.

This talk

In this talk, I am going to describe thunderella consensus protocol which is a fast consensus protocol that works in a decentralized environment and it’s the protocol that our engineers made. I talked about thunderella last year actually.

I want to focus on the more exciting part which is that Thunderella is a scalable blockchain. When our engineers are implementing it, they find something that seems like a flaw- a confirmed transaction can sometimes become undone. This was puzzling in the beginning because we had painstakingly written 74 pages of proofs that this protocol was correct.

Upon examination, we found it’s not a problem with the proofs. They are perfectly fine. It’s a flaw in the undelrying synchronous model. This is a model where we have studied consensus for the past 30 years. So in the second part of the talk, I want to rethink what should be the right model for studying these practical consensus protocols and I’ll talk about how to fix the problem too.

Thunderella

The scenario we are considering here is the following. The guy in the center will be called the proposer or leader. We have a committee of voters that elect the leader maybe through state distribution. We look at one snapshot. And someone is corrupted. In this protocol, a block is proposed. For the rest of the talk, I am going to call this notarization. Honest nodes must vote uniquely at each epoch or sequence number. If you are honest, you are going to vote on the first proposal and only that proposal you don’t vote for anything else. With this invariant in mind, I can give you a very simple consistency proof and the proof goes as follows. This assumes that less than half of the nodes are corrupt. The venn diagram must intersect at an honest node. There are only n nodes in total. The intersection of these two sets must be large. Now, if the number of correct nodes is less than the half, then we are sure that in this intersection there lives an honest guy in the intersection.. he is going to vote uniquely at every sequence number and the only reasonable explanation is that the blue is the same as the orange. So extremely simple group.

When we have honest majority, when the majority of voters are honest, we can achieve consistency which doesn’t rely on the leader being honest. I only require that honest nodes vote uniquely and I don’t rely on anything about the leader.

Now if you want liveness and want the protocol to make progress, you need a stronger assumption. In particular, you want to assume that 3/4ths of the nodes are honest and online, and also we need the leader to be honest and online. So now this benevolent dictator of liveness might not be so benevolent after all, the world will stop.

So how do we make the protocol really decentralized? We want both consistency and liveness under the weaker condition of honest majority only. The way tha tthunderella achieves this property is by introducing another slow chain, like bitcoin or ethereum but it could be a proof-of-stake slow chain. The slow chain is going to provide consistency and slow liveness for the honest majority. If this voting protocol ever fails, you fall back to the slow chain and you can figure out how to switch leader or how to continue the protocol and you can discuss how to fix this on the blockchain. The nice thing about this paradigm is that almost all the time, the protocol will be operating on the fast chain, so we can do it in a couple of round trips.

So what I haven’t explained is how we do the fallback. I’ll quickly explain that because it matters to the rest of the talk later. To explain the fallback, I have to describe two things- one is how to detect fast-path failure, and the other is how to do the fallback. Fallback is important for failure detection we use a heartbeat mechanism. Post periodic heartbeats to the slowchain by a signed hash of the current fast path lock by the committee. When it has enough votes, it becomes notarized and it becomes a heartbeat that gets periodically posted to the slow chain. The heartbeat is a keep alive mechanism, and it gives you a periodic checkpoint of the fast-path lock.

If however, at some point you notice that a large number of blocks have gone by without a single heartbeat, it’s a security parameter, then something must be wrong. Maybe the fast path failed, the leader crashed, or the leader is trying to cnesor transactions, or the committee is unhappy and has stopped signing. And then we want to fallback. We have consistency at every sequence number, but at the moment, people might not agree on the log before fallback completes. So when he decides to fallback, someone slow might see 3 transactions whereas I might see 6 transactions. So that’s why we need to use slowchain to reach agreement on the fast-path lock before we do a fallback.

If on the fast path you have confirmed a transaction, then you should post it to the green set of blocks. By posting these blocks to the slowchain, you can make sure that they get picked up. You don’t have to post anything if it has already been checked by a heartbeat.

This approach, the nice thing we can achieve is that we can achieve consistency and liveness under the stronger set of conditions. So most of the time the protocol is going to operate in this orange regime.

Recap: Thunderella

Thunderella usually has a single round of voting when things are good. When things are bad, we fallback to the slowchain and then rebootstrap the fast path from that.

Also in our paper

This talk might make it sound simpler than it really is. There’s leader/committee reconfiguration and many other topics in the paper.

What is the flaw in Thunderella?

It’s a flaw in the synchronous model. I would describe to you a scenario where this can take place. One thing that is interesting is that in this scenario, everyone is benign and nobody has malicious intent. But if a few nodes crash in a specific pattern, then a confirmed transaction can be undone. The leader makes a proposal, everyone votes, and the leader collects a bunch of notes and notarization and sends the notarization to Coinbase. Only Coinbase has received the notarization. At this very instant, maybe the leader crashes. So he is out of the picture. Coinbase believes the transaction is confirmed. The leader went away, and everyone is going to try to fallback, and they will post the notarized transactions they have seen since the last checkpoint to the slowchain. After some time, the rest of the network comes together and have fixed the problem they have re-bootstrapped the fast path.

Paradox

On the one hand we have this mathematical proof, and on the other hand, it seems like I have described a pretty serious flaw. As I already explained, the problem is not with the proof. It’s actually with the underlying synchronous model. It assumes a synchronous network and we don’t have that.

Synchronous model

What is the synchronous model? We assume that when honest node sends message, the messages will be delivered in a bounded amount of time. The message delay is at most 1 round. In case of the temporary outage in that last scenario, no matter how short your outage is, the model will treat you as faulty. The protocol was not required to provide any security guarantees for any faulty nodes because who cares about a corrupted malicious node.

Partial synchrony

You might be saying, Elaine, yes, we know this and it’s been known for 30 years. Yes, so that’s why we are also looking at asynchronous models. There, the message delay can be arbitrarily long for partial synchrony regime. If you have a protocol that is provably secure in the partial synchrony, the faulty node that just came back online, will be okay.

If you want partition tolerance, you have to suffer from a well known law. Any partial synchronous protocol cannot tolerate more than 1/3rd corruption. In the synchronous model, you can tolerate <1/2 corruptions.

Can we achieve the best of both worlds?

Given the classical insights we have, the answer should be no. But perhaps it’s not completely hopeless because who says that synchrony has to be binary? Why does it have to be either synchronous or not synchronous? And why does partition tolerance have to be an attribute? So we’re going to carefully quantify how synchronous our network is, and how partition tolerant our protocol is.

Imagine we had a set of nodes and they were honest with good network connectivity and they can send messages to each other in a single round. But maybe there’s some other nodes with unstable connectivity and they go online/offline frequently, and then there’s a set of corrupt nodes which we don’t care about achieving anything with. But the faulty ones and the honest good connectivity ones should be able to achieve something together. We don’t want to penalize offline nodes, which can’t have liveness, but they should be able to have consistency and liveness when they come back online.

Lazy coauthor model

So can we do this? Well, you need some assumptions. You have to assume that the green node set is larger than 1/2. For best-possible partition tolerance, it could be that– every node may be offline at some point. No node can guarantee that they will always be online. Not even Google. The real model we work with is what I would like to call the lazy coauthor model. Some of the authors are online on Monday and others are online on other days of the week. We want to be able to write a paper in this manner and we want to be able to for instance make the next Crypto deadline. It can’t be the case that at the end of the day we discover everyone has a different theorem in mind. We want everyone to be proving the same thereom. For the online nodes, the message can be received by the people who are online in the future. Whenever you are offline, your messages can get erased and there’s no guarantee of delivery sent by an offline node. So how can we write the paper such that we meet the deadline and prove consistency?

Solution

The fix has two pieces. We need to fix the fast path and the slow chain. Fixing the slow chain is a separate story on its own so I won’t cover it in this talk. But I’ll talk about fixing the fast path and fallback. The fast path fix is very simple. I can describe it in one slide. Earlier, assume that honest nodes vote only after the previous coin is notarized. There is a notary sequence. Let’s be more patient, and let’s wait for the next block to be notarized too. We’re always going to chop up the last notarized block at the end. The reason for this is that if I see a notarization on a block, the only thing I am sure of is that man ypeople have voted on that block. Imagine honest nodes will only vote for the next block when the parent is notarized. So if I see notarization on the next block too, it means many people have seen a notarization for the parent. This gives me more confidence because even if I crash at this point, other people have seen the notarization of the parent chain and some of them will be online and be able to post it to the slowchain. You can formalize this into a proof.

Additional results

What I have talked about today is just the tip of the iceberg. There’s a few papers talking about the models and practical construction. We study not just consensus but also computation. We want best possible partition tolerant protocol with optimistic responsiveness, and we want privacy. We have a practical variant with liveness only during “periods of synchrony”, under standard cryptographic assumptions.

Conclusion

I would like to conclude by making a couple of philosophical observations. As it turns out, this class of protocols that has best possible partition tolerance, in the white class in this picture, is a strict subset of synchronous honest majority protocols. The way to think about this is that our model is a refinement of classical synchronicity. We can tease out which of those protocols have the drawback we don’t care about in practice. If we look at a bunch of existing classical synchronous protocols, perhaps non-surprisingly, none of them belong the white model set. It’s not like any honest classic honest majority protocol would satisfy this property. So what this means is that if you’re a blockchain company that wants to deploy a synchronous model, and you don’t have a security proof then you shouldn’t be able to eat or sleep. If it belongs to the yellow set that’s a little better but not good enough.

Throughout the talk, I’ve purposefully hidden this black set from you. There’s a classical corrupt majority model like Dolev-Strong but unfortnuately we show that if you want to tolerate corrupt majority then it can’t be best-possible partition tolerant. If you reflect on what this means, then one way to interpret this is that the classical synchronous model is kind of like a mismatch for what we care about in practice. If you work with classical synchrony, then we would be misled to think that tolerating more corruptions is always better than tolerating fewer. But from the example you saw today, I hope I have convinced that you should always choose from the white set rather than from the black set.

That’s the end of my talk. We will be launching mainnet very soon. So that’s the little rocket there, going to the moon.