WO2019126311A1 - Fast and partition-resilient blockchains - Google Patents

Abstract

In a transaction system in which transactions are organized in blocks, a new block Br of valid transactions is constructed, relative to a sequence of prior blocks B⁰,..., B^r-1, by having an entity determine a quantity Q from the prior blocks, having the entity use a secret key in order to compute a string S uniquely associated to Q and the entity, having the entity compute from S a quantity T that is one of: S itself, a function of S, and/or a hash value of S, having the entity determine whether T possesses a given property, and, if T possesses the given property, having the entity digitally sign a hash value H of B ^r and make available S, B ^r and a digitally signed version of H, wherein, B ^r may be proposed in different steps of the round r and may be re-proposed multiple times during round r, and an entity may verify a hash value H of a block B independent of whether the entity has received B or not.

Description

Fast and Partition- Resilient Blockchains

CROSS REFERENCE TO RELATED APPLICATIONS

This application claims priority to U.S. Prov. App. No. 62/607,558, filed December 19, 2017, and entitled FASTER BYZANTINE AGREEMENT IN PROPAGATION NETWORKS WITH > 2/3 HONEST MAJORITY, and to U.S. Prov. App. No. 62/632,944, filed February 20, 2018, and entitled ALGORAND, and to U.S. Prov. App. No. 62/643,331, filed March 15, 2018, and entitled INCENTIVES AND TRANSACTION FEES IN ALGORAND, and to U.S. Prov. App. No. 62/777,410, filed December 10, 2018, and entitled VIRTUAL BLOCKCHAIN PROTOCOLS FOR FAIR ELECTRONIC EXCHANGE, and to U.S. Prov. App. No. 62/778,482, filed December 12, 2018, and entitled VIRTUAL BLOCKCHAIN PROTOCOLS FOR FAIR ELECTRONIC EXCHANGE, which are all incorporated by reference herein.

TECHNICAL FIELD

This application relates to the field of electronic transactions and more particularly to the field of distributed public ledgers, securing the contents of sequence of transaction blocks, and the verification of electronic payments.

BACKGROUND OF THE INVENTION

A blockchain consists of an augmentable sequence of blocks: B , l¾, · · · , wherein each block consists of a number of transactions, the hash of the previous block, and other data e.g., the index of the block, time information, etc. Useful properties of a blockchain are that (PI) there is a unique block corresponding to each index 1, 2, . . . , (P2) every user in the system eventually learns the content of every block, (P3) no one can alter the content or the order of the blocks, and (P4) any valid transactio will eventually enter a block in the chain.

Users can digitally sign messages, and thus each user possesses at least one public key and a corresponding secret key. In a blockchain, in general, one knows the public keys but not necessarily the user who owns it. Accordingly, we may identify a public key with its owner.

Several blockchain systems require a block to be certified by the digital signatures of sufficiently many users in the system. In some systems such certifying users belong to a fixed set of users. In some other systems they belong to a dynamically changing set. The latter is preferrable, because an adversary would have a harder time to corrupt a dynamically changing set, particularly if the set is not only dynamic, but unpredictable as well.

A particularly effective way of selecting a set of users in a verifiable but unpredictable way is the cryptographic sortition technique is described in published PCT patent application PCT/US2017/031037, which is incorporated by reference herein. Here, a user i belongs to a set of users empowered to act in some step s during the production of block number r based on the result of a computation that i performs via a secret key of his, using inputs s and r, and possibly other inputs and other data (e.g., the fact that the user has joined the system at least k blocks before block r, for some given integer k ) . For instance, Ls computation may involve Ls digital signature, of such inputs, hashing and checking whether the hash is less than a given threshold t. (Indeed, like any other string, a hashed value can be interpreted in some standard way as a number.) In this case, s/’⁵ = s/⁵ is defined to be the credential of i for step s about block r. Such credential proves to anyone that i is indeed entitled to produce a (preferably signed) message rn^ , his voting message for step s in round r, in the process aimed at producing block r. In fact, /’s digital signatures can be checked by anyone, and anyone can hash a given value, and then check whether the result is indeed less than (or equal to) a given number.

A blockchain works by propagating messages (e.g., blocks, transactions, voting messages, digital signatures, etc) . Typically, but not exclusively, messages are propagated by gossiping them in a peer-to-peer fashion, or via relays. Several blockchain systems require the propagation network to guarantee the delivery of messages propagated by every honest user to other honest users within a bounded delay. Some further require the users to have (almost) aligned system clocks, so that the users propagate messages in a synchronized way e.g., the users enter step 2 in the generation of block 100 at time 11:20:00am EST, the voting messages for this step are delivered by time 11:20:05am EST, and the users enter step 3 of block 100 then. A less demanding and thus preferable requirement, as imposed by Algorand, is that the users’ clocks have (almost) the same speed, but the actual times shown on the clocks can be arbitrarily far from each other. A user starts his own step s in the generation of block r based on the messages he has received from the propagation network, and ends it based on messages received and how long his own clock has advanced since he started this step.

When the propagation network satisfies this requirement, Algorand ensures that the adversary cannot prevent the blockchain from functioning properly (including achieving properties PI P4). However, this relies on the adversary not attacking the propagation network itself. Such attacks include any effort the adversary may take in order to violate the bounded delay of message delivery for a sufficiently large amount of users e.g., by partitioning the users into two groups of equal size and controlling the message delivery channels between them, so that a message propagated by a user from group 1 may have indefinite delay before reaching any user in group 2.

It is thus desirable to weaken this requirement and provide blockchains and electronic money systems that do not suffer from the inefficiencies and insecurities of known decentralized approaches.

SUMMARY OF THE INVENTION.

According to the system described herein, an entity manages a transaction system in which transactions are organized in a sequence of blocks that are certified by digital signatures of a sufficient number of verifiers by the entity proposing a hash of a block B’ that includes new valid transactions relative to a sequence of certified blocks B°, . . . , B ^r_1 if no rth block B^r has been certified and by the entity proposing a hash of the block B^r if the rth block B^r has been verified by a sufficient number of other entities. A block may be certified by the entity only in response to confirming transactions for the block and confirming that the block was constructed and propagated by an entity entitled to construct and propagate the block. The entity may propose a hash value by digitally signing the hash value to provide a digitally-signed version of the hash value and the entity may propagate the digitally-signed version of the hash value to a network that includes other entities. If no rth block B^r has been certified, the entity may also digitally sign and propagate the block B' . The entity may determine a quantity Q from the prior blocks and may use a secret key in order to compute a string S uniquely associated with Q and computes from S a quantity T that is S itself, a function of S, and/or a hash value of S and the entity may determine whether to propose a hash value by determining whether T possesses a given property. S may be a signature of Q under a secret key of the entity, T may be a hash of S and T may possess the given property if T is less than a given threshold. The entity may be part of a network of entities and a particular one of the entities may construct and propagates the block B^r . The rth block B^r may be determined to be certified by the entity if the entity receives an indication that at least a predetermined number of the entities individually certify a hash value corresponding to the rth block B^r . In response to the entity receiving the indication that a predetermined number of the entities individually certified the rth block B^r , the entity may increment r to begin adding additional blocks to the sequence of blocks. The particular one of the entities may be individually chosen by a predetermined number of the entities to be a leader. The rth block B^r may be determined to be certifiable by the entity if the entity receives an indication that at least a predetermined number of the entities individually verify receiving an indication that the particular one of the entities has provided a hash value corresponding to the rth block B^r to each of the predetermined number of the entities.

According further to the system described herein, an entity manages a transaction system in which transactions are organized in a sequence of certified blocks by the entity receiving a hash value of a block B^r from an other entity that generated the block based on new valid transactions relative to a sequence of certified blocks B°, . . . , B^r_1 ,the entity certifying the block B^r in response to a sufficient number of other entities having indicated receipt of the hash value of the block B^r from the other entity and the hash value being valid for the block B^r , the entity generating a new block B’ based on new valid transactions relative to a sequence of certified blocks B°, . . . , B ^{r_ 1} in response to an insufficient number of the other entities indicating receipt of the hash value of the block B^r from the other entity, B’ being different from B^r , and the entity incrementing r to begin adding additional blocks to the sequence of blocks in response to the entity receiving the indication that a predetermined number of the entities individually certified the rth block B^r or a predetermined number of the entities individually certified the new block B’. The blocks may be certified by digital signatures. New blocks may be proposed by different ones of the entities until receiving the indication that a predetermined number of the entities individually certified a previously proposed block. The entity may provides an indication that a new block should be generated in response to the hash value not being valid for the block B^r . The entity may generate a new block B’ based on new valid transactions relative to a sequence of certified blocks B°, . . . , B ^r_1 in response to a sufficient number of the other entities providing an indication that a new block should be generated. The entity may provide an indication that the hash value of the block B^r should be propagated in response to a sufficient number of the other entities having indicated receipt of the hash value of the block B^r from the other entity and the hash value being valid for the block B^r .

According further to the system described herein, an entity verifies a proposed hash value of a new block B^r of transactions relative to a given a sequence of blocks, B°, . . . , B^r_1 , without access to the new block B^r in a transaction system in which transactions are organized in blocks and blocks are certified by a set of digital signatures by having the entity determine a quantity Q from the prior blocks, having the entity compute a digital signature S of Q, having the entity compute from S a quantity T that is S itself, a function of S, and/or hash value of S, having the entity determine whether T possesses a given property, and, if T possesses the given property, having the entity verify the proposed hash value of the new block B^r independent of confirming whether the proposed hash value corresponds to the new block B^r . The entity may propagate the proposed hash value of the new block B^r prior to receiving the new block B^r .

According furtheer to the system described herein, in a transaction system in which transactions are organized in blocks, a new block B^r of valid transactions is constructed, relative to a sequence of prior blocks B°, B¹, . . . , B^{r_ 1}, by having an entity determine a quantity Q from the prior blocks, having the entity use a secret key in order to compute a string S uniquely associated to Q, having the entity compute from S a quantity T that is S itself, a function of S, and/or a hash value of S, having the entity determine whether T possesses a given property and, if T possesses the given property, having the the entity compute a hash value H of B^r , digitally sign H and make available to others S, B^r and a digitally signed version of H. The secret key may be a secret signing key corresponding to a public key of the entity and S' is a digital signature of Q by the entity. T may be a binary expansion of a number and satisfies the given property if T is less than a given number p. S may be made available by making S deducible from B^r . Each user may have a balance in the transaction system and p may vary for each user according to the balance of each user.

According further to the system described herein, selecting a subset of users in a blockchain system to verify a data string m relative to a sequence of prior blocks B°, . . . , B^{1" 1} , includes causing at least some of the users to determine a quantity Q from the prior blocks, causing at least some of the users to compute a digital signature S of Q and other information, causing at least some of the users to determine a hash value of the digital signature, causing at least some of the users to compare the hash value to a pre-determined threshold, and causing the subset of the users to digitally sign m together with other information and make available to others S and a digitally signed version of m in response to the hash value being below a pre-determined threshold for each of the subset of users. The digital signature may be credentialed if the hash value is below a pre-determined threshold. Each user may have a balance in the transaction system and the pre-determined threshold may vary for each user according to the balance of each user. The pre-determined threshold for each user may be chosen to cause the subset of the users to contain a minimum number of the users. The data string m may be a hash value of a new block B^r . The data string m may be verified by at least a given number of credentialed signatures of m.

According further to the system described herein, selecting a subset of users in a blockchain system to certify a new block B^r relative to a sequence of prior blocks B°, . . . , B^{r 1} , includes causing at least some of the users to determine a quantity Q from the prior blocks, causing at least some of the users to compute a digital signature S of Q and other information, causing at least some of the users to determine a hash value of the digital signature, causing at least some of the users to compare the hash value to a pre-determined threshold, causing the subset of the users to determine B^r is valid relative to B°, . . . , B ^r_1 in response to the hash value being below a pre-determined threshold for each of the subset of users and causing the subset of the users to digitally sign a hash value H of B^r together with other information and make available to others S and a digitally signed version of H. A particular one of the users may digitally sign the new block B^r only if the particular one of the users verifies information provided in the new block B^r . Each user may have a balance in the transaction system and the pre-determined threshold may vary for each user according to the balance of each user. The pre-determined threshold for each user may be chosen to cause the subset of the users to contain a minimum number of the users. The block B^r may be certified by at least a given number of credentialed signatures of H from users who have determined B^r is valid relative to B°, . . . , B^{r 1}.

According further to the system described herein, computer software, provided in a non- transitory computer-readable medium, includes executable code that implements any of the steps described herein.

The present invention dispenses with the requirement on the propagation network’s message delivery delay for the security of certifying new blocks. A new block is first prepared (e.g., proposed, propagated, and/or agreed upon by at least some users) and then it is certified. A user who has receive a newly constructed block, a hash value of a new block, and/or credentialed signatures within a desired time interval proceeds to verify and/or certify the new block. However, we wish to certify a new block even when messages propagated in the network may be indefinitely delayed. The certification of a block B guarantees that certain valuable properties apply to the block. A typical main property is to enable a user, even a user who has not participated to or observed the preparation of a block B, to ascertain that B has been added to the blockchain, or even that B is the rth block in the blockchain. Another valuable property (often referred to as finalization) guarantees that B will not disappear from the blockchain, due to a soft fork, even in the presence of a partition of the communication network on which the blockchain protocol is executed. A partition of the network may cause the users to be separated into multiple groups, with messages propagated from one group not reaching users in other groups. A partition may be resolved after an indefinite amount of time, after which the network again guarantees message delivery after bounded delays.

Assume that a block B has been prepared, in any fashion and in any number of steps. Realizing that a block has been properly prepared requires time and effort, and the verification of various pieces of evidence. A certificate of B consists of a given number of users’ digital signatures with valid credentials. Such a certificate of B vouches that the users who have produced such signatures have participated to or observed the preparation of B. At least, it vouches that, if one of the digital signatures of the certificate has been produced by an honest user, then that user has checked that B has been properly prepared.

In the system described herein, while a user is collecting evidences for one block B being properly prepared as the rth block in the blockchain relative to prior blocks B°, . . . , R^r_1 , the user may construct and propose a new block B' as the rth block in the blockchain relative to B°, . . . , B ^{r_ 1} if he has evidence that a certificate of B has not been generated in the system. A user proposes B' by determining a quantity Q from prior blocks, computing a string S uniquely associated to Q using a secret key, computing from S a quantity T that is S itself, a function of S, and/or a hash value of S, determining whether T possesses a given property and, if T possesses the given property, computing a hash value H' of B¹ , digitally signing H' and making available to others S, B ' and a digitally signed version of H' . The secret key may be a secret signing key corresponding to a public key of the entity and S' is a digital signature of Q by the entity. T may be a binary expansion of a number and satisfies the given property if T is less than a given number p. S may be made available by making S deducible from B' . Each user may have a balance in the transaction system and p may vary for each user according to the balance of each user.

In the system described herein, while a user is collecting evidences for one block B being properly prepared as the rth block in the blockchain relative to prior blocks B°, . . . , B^{r 1} , the user may re-propose B as the rth block in the blockchain if he has evidence that a certificate of B may have been generated in the system but have not been made available to him. A user may re-propose B without having received B itself but rather having received a given number of users’ digital signatures with valid credentials verifying a hash value of B. A user may re- propose B by determining a quantity Q from prior blocks, computing a string S uniquely associated to Q using a secret key, computing from S a quantity T that is S itself, a function of S, and/or a hash value of S, determining whether T possesses a given property and, if T possesses the given property, digitally signing a hash value H of B and making available to others S and a digitally signed version of H. The secret key may be a secret signing key corresponding to a public key of the entity and S' is a digital signature of Q by the entity. T may be a binary expansion of a number and satisfies the given property if T is less than a given number p. S may be made available by making S deducible from B' . Each user may have a balance in the transaction system and p may vary for each user according to the balance of each user.

Proposing new blocks and re-proposing existing blocks may happen indefinite amount of times during the generation of a rth block of the blockchain and may be carried out by different users. A block B may have more than one certificates generated from different steps. However, during the generation of a rth block of the blockchain, one and only one block will have a certificate and thus be considered by the users as the rth block of the blockchain. The efficiency of the system described herein derives from the following facts. First, a user i may verify and/or re-propose a hash value of a block B before the user may receive B itself. Second, a new block B' may be proposed as the rth block before all users may have collected evidence that a previously proposed block B as the rth block does not have a certificate generated in the system. Indeed, B' may be proposed as soon as one user has collected such evidence. Third, a previously proposed block B as the rth block of the blockchain may be re-proposed before all users may have collected evidence that a certificate for B may have been generated in the system but not made available to all. Indeed, B may be re-proposed as soon as one user has collected such an evidence.

An evidence may consists of a set of credentialed signatures properly formed to verify a data string m. Evidences for different purposes may consist of different numbers of signatures. The security of the system described herein derives from proper choices of pre- determined thresholds that users compare the hash values of their signatures to when verifying different data strings, and from proper choices of numbers of signatures sufficient to form evidences for different purposes. For instance, let p be the maximum percentage of malicious users in the system. Typically, malicious users are in a minority e.g., p < 1/3. Then the pre-determined threshold t and the sufficient number n of signatures forming a certificate for a block may be chosen so that, with sufficiently high probability, (a) for any possible block value B, there are n or more credentialed signatures of honest users to form a certificate for B and (b) in any certificate of B, more than 2/3 of credentialed signatures belongs to honest users.

The system described herein is agnostic about whether ephemeral keys are used in the blockchain: when users propose new blocks, re-propose existing blocks, or verify data strings, the users may use long-term secret keys to generate credentialed signatures, where the keys may be used repeatedly during the life time of the system, or the users may use ephemeral secret keys where one key is used only once, or the users may use combinations of long-term keys and ephemeral keys.

As part of the system described herein, a user i may not only have a credentialed signature for participating in the generation of a block, but also a credentialed signature with a weight (essentially a credentialed signature associated with a number of votes). Indeed, the weight of Fs credential signature for participating in the generation of a block may depend on how much money i has in the system. Indeed, rather than having a single pre-determined threshold t for all users for participating in block generation, each user i may have his own threshold ¾ that is higher the higher Fs amount of money is. And a user i may have different thresholds for participating in block generation in different ways e.g., proposing a new block, re- proposing a block, or verifying a data string m of a specific format. For simplicity, but without limitation intended, we shall continue to describe our system treating a user i with a weight-n credentialed signature as n users, each having a (weight-1) credentialed signature.

Below, after quickly recalling the traditional Algorand system, we provide an example of the preferred embodiment, without any limitation intended, based on Algorand.

BRIEF DESCRIPTION OF THE DRAWINGS .

Embodiments of the system described herein are explained in more details in accordance with the figures of the drawings, which are briefly described as follows.

FIG. 1 is a schematic representation of a network and computing stations according to an embodiment of the system described herein.

FIG. 2 is a schematic and conceptual summary of the first step of Algorand system, where a new block of transactions is proposed.

FIG. 3 is a schematic and conceptual summary of the agreement and certification of a new block in the Algorand system. FIG. 4 is a schematic diagram illustrating a Merkle tree and an authenticating path for a value contained in one of its nodes.

FIG. 5 is a schematic diagram illustrating the 8 Merkle trees corresponding to the first 8 blocks constructed in a blocktree.

FIG. 6 is a schematic representation of a partitioned network and computing stations according to an embodiment of the system described herein.

DETAILED DESCRIPTION OF VARIOUS EMBODIMENTS .

The system described herein provides a mechanisms for distributing transaction verification and propagation so that no entity is solely responsible for performing calculations to verify and/or propagate transaction information. Instead, each of the participating entities shares in the calculations that are performed to propagate transaction in a verifiable and reliable manner.

Referring to FIG. 1, a diagram 20 shows a plurality of computing workstations 22a-22c connected to a data network 24, such as the Internet. The workstations 22a-22c communicate with each other via the network 24 to provide distributed transaction propagation and verification, as described in more detail elsewhere herein. The system may accommodate any number of workstations capable of providing the functionality described herein, provided that the workstations 22a-22c are capable of communicating with each other. Each of the workstations 22a-22c may independently perform processing to propagate transactions to all of the other workstations in the system and to verify transactions, as described in more detail elsewhere herein.

FIG. 2 diagrammatically and conceptually summarizes the first step of a round r in the Algorand system, where each of a few selected users proposes his own candidate for the rth block. Specifically, the step begins with the users in the system, a, . . . , z, individually undergo the secret cryptographic sortition process, which decides which users are selected to propose a block, and where each selected user secretly computes a credential proving that he is entitled to produce a block. Only users b, d, and h are selected to propose a block, and their respectively computed credentials are s/’¹ , s

and s/’¹. Each selected user i assembles his own proposed block, R/, ephemerally signs it (i.e., digitally signs it with an ephemeral key, as explained later on) , and propagates to the network together with his own credential. The leader of the round is the selected user whose credential has the smallest hash. The figure indicates the leader to be user d. Thus his proposed block, B_d ^r , is the one to be given as input to the Binary agreement protocol.

FIG. 3 diagrammatically and conceptually summarizes Algorand’s process for reaching agreement and certifying a proposed block as the official rth block, B^r . Since the first step of Algorand consists of proposing a new block, this process starts with the second step. This step actually coincides with the first step of Algorand’s preferred Byzantine agreement protocol, BA*. Each step of this protocol is executed by a different “committee” of players, randomly selected by secret cryptographic sortition (not shown in this figure). Accordingly, the users selected to perform each step may be totally different. The number of Steps of BA* may vary. Fig. 3 depicts an execution of BA* involving 7 steps: from Algorand’s 2 through Algorand’s step 8. The users selected to perform step 2 are a, e, and q. Each user i 6 {a, e, q} propagates to the network his credential, s , that proves that i is indeed entitled to send a message in step 2 of round r of Algorand, and his message proper of this step, m/⁵, ephemerally signed. Steps 3-7 are not shown. In the last step 8, the figure shows that the corresponding selected users, b, /, and x, having reached agreement on B^r as the official block of round r, propagate their own ephemeral signatures of block B^r (together, these signatures certify B^r ) and their own credentials, proving that they are entitled to act in Step 8.

FIG. 4 schematically illustrates a Merkle tree and one of its autenticating path. Specifically, Fig. 4. A illustrates a full Merkle tree of depth 3. Each node x, where x is denoted by a binary string of length < 3, stores a value v_x. If x has length < 2, then v_x = H{v_x o, ¾i) . For the Merkle tree of Fig. 4. a, Fig. 4.B illustrates the authenticating path of the value t¾io·

FIG. 5 schematically illustrates the 8 Merkle trees, corresponding to the first 8 blocks constructed in a blocktree, constructed within a full binary tree of depth 3. In Fig. 5.i, nodes marked by an integer belong to Merkle tree Ϊ). Contents of nodes marked by i (respectively, by ί ) are temporary (respectively, permanent) .

The description herein focuses on transactions that are payments and on describing the system herein as a money platform. Those skilled in the art will realize that the system described herein can handle all kinds of transactions as well.

The system described herein has a very flexible design and can be implemented in various, but related, ways. We illustrate its flexibility by detailing two possible embodiments of its general design. From them, those skilled in the art can appreciate how to derive all kinds of other implementations as well.

To facilitate understanding the invention, and allow to internal cross reference of its various parts, we organize its presentation in numbered and titled sections. The first sections are common to both of the detailed embodiments.

1 Introduction

Money is becoming increasingly virtual. It has been estimated that about 80% of United States dollars today only exist as ledger entries. Other financial instruments are following suit.

In an ideal world, in which we could count on a universally trusted central entity, immune to all possible cyber attacks, money and other financial transactions could be solely electronic. Unfortunately, we do not live in such a world. Accordingly, decentralized cryptocurrencies, such as Bitcoin, and“smart contract” systems, such as Ethereum, have been proposed. At the heart of these systems is a shared ledger that reliably records a sequence of transactions, as varied as payments and contracts, in a tamperproof way. The technology of choice to guarantee such tamperproofness is the blockchain. Blockchains are behind applications such as cryptocurrencies, financial applications, and the Internet of Things. Several techniques to manage blockchain- based ledgers have been proposed: proof of work , proof of stake, practical Byzantine fault- tolerance, or some combination.

Currently, however, ledgers can be inefficient to manage. For example, Bitcoin’s proof-of- work approach requires a vast amount of computation, is wasteful and scales poorly. In addition, it de facto concentrates power in very few hands.

We therefore wish to put forward a new method to implement a public ledger that offers the convenience and efficiency of a centralized system run by a trusted and inviolable authority, without the inefficiencies and weaknesses of current decentralized implementations. We call our approach Algorand, because we use algorithmic randomness to select, based on the ledger constructed so far, a set of verifiers who are in charge of constructing the next block of valid transactions. Naturally, we ensure that such selections are provably immune from manipulations and unpredictable until the last minute, but also that they ultimately are universally clear.

Algorand’s approach is quite democratic, in the sense that neither in principle nor de facto it creates different classes of users (as“miners” and“ordinary users” in Bitcoin) . In Algorand “all power resides with the set of all users” .

One notable property of Algorand is that its transaction history may fork only with very small probability (e.g., one in a trillion, that is, or even 11U¹⁸) . Algorand can also address some legal and political concerns. The Algorand approach applies to blockchains and, more generally, to any method of generating a tamperproof sequence of blocks. We actually put forward a new method alternative to, and more efficient than, blockchains that may be of independent interest.

1.1 Bitcoin’s Assumption and Technical Problems

Bitcoin is a very ingenious system and has inspired a great amount of subsequent research. Yet, it is also problematic. Let us summarize its underlying assumption and technical problems which are actually shared by essentially all cryptocurrencies that, like Bitcoin, are based on proof- of-work.

For this summary, it suffices to recall that, in Bitcoin, a user may own multiple public keys of a digital signature scheme, that money is associated with public keys, and that a payment is a digital signature that transfers some amount of money from one public key to another. Essentially, Bitcoin organizes all processed payments in a chain of blocks, Hi, H2, each consisting of multiple payments, such that, all payments of B , taken in any order, followed by those of B 2, in any order, etc., constitute a sequence of valid payments. Each block is generated, on average, every 10 minutes.

This sequence of blocks is a chain , because it is structured so as to ensure that any change, even in a single block, percolates into all subsequent blocks, making it easier to spot any alteration of the payment history. (As we shall see, this is achieved by including in each block a cryptographic hash of the previous one.) Such block structure is referred to as a blockchain.

Assumption: Honest Majority of Computational Power Bitcoin assumes that no malicious entity (nor a coalition of coordinated malicious entities) controls the majority of the computational power devoted to block generation. Such an entity, in fact, would be able to modify the blockchain, and thus re-write the payment history, as it pleases. In particular, it could make a payment p, obtain the benefits paid for, and then“erase” any trace of p.

Technical Problem 1 : Computational Waste Bitcoin’s proof-of-work approach to block generation requires an extraordinary amount of computation. Currently, with just a few hundred thousands public keys in the system, the top 500 most powerful supercomputers can only muster a mere 12.8% percent of the total computational power required from the Bitcoin players. This amount of computation would greatly increase, should significantly more users join the system.

Technical Problem 2: Concentration of Power Today, due to the exorbitant amount of computation required, a user, trying to generate a new block using an ordinary desktop (let alone a cell phone), expects to lose money. Indeed, for computing a new block with an ordinary computer, the expected cost of the necessary electricity to power the computation exceeds the expected reward. Only using pools of specially built computers (that do nothing other than “mine new blocks”) , one might expect to make a profit by generating new blocks. Accordingly, today there are, de facto, two disjoint classes of users: ordinary users, who only make payments, and specialized mining pools, that only search for new blocks.

It should therefore not be a surprise that, as of recently, the total computing power for block generation lies within just five pools. In such conditions, the assumption that a majority of the computational power is honest becomes less credible.

Technical Problem 3: Ambiguity In Bitcoin, the blockchain is not necessarily unique. Indeed its latest portion often forks : the blockchain may be say B , . . . , ¾, B_k' _+l , B_k' ₊₂, according to one user, and B , . . . , ¾, B_k+1 , B_k+2, B_k+3 according another user. Only after several blocks have been added to the chain, can one be reasonably sure that the first k + 3 blocks will be the same for all users. Thus, one cannot rely right away on the payments contained in the last block of the chain. It is more prudent to wait and see whether the block becomes sufficiently deep in the blockchain and thus sufficiently stable.

Separately, law-enforcement and monetary-policy concerns have also been raised about Bitcoin.¹

1.2 Algorand, in a Nutshell

Setting Algorand works in a very tough setting. Briefly,

(a) Permissionless and Permissioned Environments. Algorand works efficiently and securely even in a totally permissionless environment, where arbitrarily many users are allowed to join the system at any time, without any vetting or permission of any kind. Of course, Algorand works even better in a permissioned environment.

(b) Very Adversarial Environments. Algorand withstands a very powerful Adversary, who can

(1) instantaneously corrupt any user he wants, at any time he wants, provided that, in a permissionless environment, 2/3 of the money in the system belongs to honest user. (In a permissioned environment, irrespective of money, it suffices that 2/3 of the users are honest.)

(2) totally control and perfectly coordinate all corrupted users; and

(3) schedule the delivery of all messages, provided that each message m sent by a honest user reaches reaches all (or sufficiently many of) the honest users within a time X_m, which solely depends on the size of m.

Main Properties Despite the presence of our powerful adversary, in Algorand

• The amount of computation required is minimal. Essentially, no matter how many users are present in the system, each of fifteen hundred users must perform at most a few seconds of computation.

• A new block is generated quickly and will de facto never leave the blockchain. That is, Algorand’s blockchain may fork only with negligible probability (i.e., less than one in a trillion or 10 ¹⁸). Thus, users can relay on the payments contained in a new block as soon as the block appears.

• All power resides with the users themselves. Algorand is a truy distributed system. In particular, there are no exogenous entities (as the “miners” in Bitcoin), who can control which transactions are recognized.

¹The (pseudo) anonymity offered by Bitcoin payments may be misused for money laundering and/or the financing of criminal individuals or terrorist organizations. Traditional banknotes or gold bars, that in principle offer perfect anonymity, should pose the same challenge, but the physicality of these currencies substantially slows down money transfers, so as to permit some degree of monitoring by law-enforcement agencies.

The ability to “print money” is one of the very basic powers of a nation state. In principle, therefore, the massive adoption of an independently floating currency may curtail this power. Currently, however, Bitcoin is far from being a threat to governmental monetary policies, and, due to its scalability problems, may never be. Algorand’s Techniques.

1. A NEW AND FAST BYZANTINE AGREEMENT PROTOCOL . Algorand generates a new block via an inventive cryptographic, message-passing , binary Byzantine agreement (BA) protocol, BA*. Protocol BA* not only satisfies some additional properties (that we shall soon discuss) , but is also very fast. Roughly said, its binary-input version consists of a 3-step loop, in which a player i sends a single message rr¾ to all other players. Executed in a complete and synchronous network, with more than 2/3 of the players being honest, with probability > 1/3, after each loop the protocol ends in agreement. (We stress that protocol BA* satisfies the original definition of Byzantine agreement, without any weakenings.)

Algorand leverages this binary BA protocol to reach agreement, in our different commu nication model, on each new block. The agreed upon block is then certified, via a prescribed number of digital signature of the proper verifiers, and propagated through the network.

2. SECRET CRYPTOGRAPHIC SORTITION. Although very fast, protocol BA* would benefit from further speed when played by millions of users. Accordingly, Algorand chooses the players of BA* to be a much smaller subset of the set of all users. To avoid a different kind of concentration-of-power problem, each new block B^r will be constructed and agreed upon, via a new execution of BA*, by a separate set of selected verifiers, SV^r. In principle, selecting such a set might be as hard as selecting B^r directly. We traverse this potential problem by a novel approach that we term secret cryptographic sortition. Sortition is the practice of selecting officials at random from a large set of eligible individuals. (Sortition was practiced across centuries: for instance, by the republics of Athens, Florence, and Venice. In modern judicial systems, random selection is often used to choose juries. Random sampling has also been advocated for elections.) In a decentralized system, of course, choosing the random coins necessary to randomly select the members of each verifier set SV^r is problematic. We thus resort to cryptography in order to select each verifier set, from the population of all users, in a way that is guaranteed to be automatic (i.e., requiring no message exchange) and random. In a similar fashion we select a user, the leader, in charge of proposing the new block B^r , and the verifier set SV^r , in charge to reach agreement on the block proposed by the leader. The inventive system leverages some information, Q ^r_1 , that is deducible from the content of the previous block and is non-manipulatable even in the presence of a very strong adversary.

3. THE QUANTITY (SEED) Q^r . We use the the last block

in the blockchain in order to automatically determine the next verifier set and leader in charge of constructing the new block B'^’ . The challenge with this approach is that, by just choosing a slightly different payment in the previous round, our powerful Adversary gains a tremendous control over the next leader. Even if he only controlled only 1/1000 of the players/money in the system, he could ensure that all leaders are malicious. (See the Intuition Section 4.1.) This challenge is central to all proof- of-stake approaches, and, to the best of our knowledge, it has not, up to now, been satisfactorily solved.

To meet this challenge, we purposely construct, and continually update, a separate and carefully defined quantity, Q^r , which provably is, not only unpredictable, but also not influentiable, by our powerful Adversary. We may refer to Q^r as the rth seed, as it is from Q^r that Algorand selects, via secret cryptographic sortition, all the users that will play a special role in the generation of the rth block. The seed Q^r will be deducible from the block B ^{r 1}.

4. SECRET CREDENTIALS. Randomly and unambiguously using the current last block, B^{1 1} , in order to choose the verifier set and the leader in charge of constructing the new block, B^r , is not enough. Since B ^r_1 must be known before generating B^r , the last non- influentiable quantity Q ^r_1 deducible from B ^r_1 must be known too. Accordingly, so are the verifiers and the leader in charge to compute the block B^r . Thus, our powerful Adversary might immediately corrupt all of them , before they engage in any discussion about B^r , so as to get full control over the block they certify.

To prevent this problem, leaders (and actually verifiers too) secretly learn of their role, but can compute a proper credential , capable of proving to everyone that indeed have that role. When a user privately realizes that he is the leader for the next block, first he secretly assembles his own proposed new block, and then disseminates it (so that can be certified) together with his own credential. This way, though the Adversary will immediately realize who the leader of the next block is, and although he can corrupt him right away, it will be too late for the Adversary to influence the choice of a new block. Indeed, he cannot “call back” the leader’s message no more than a powerful government can put back into the bottle a message virally spread by WikiLeaks.

As we shall see, we cannot guarantee leader uniqueness, nor that everyone is sure who the leader is, including the leader himself! But, in Algorand, unambiguous progress will be guaranteed.

5. PLAYER REPLACEABILITY. After he proposes a new block, the leader might as well“die” (or be corrupted by the Adversary) , because his job is done. But, for the verifiers in SV^r , things are less simple. Indeed, being in charge of certifying the new block B^r with sufficiently many signatures, they must first run Byzantine agreement on the block proposed by the leader. The problem is that, no matter how efficient it is, BA* requires multiple steps and the honesty of > 2/3 of its players. This is a problem, because, for efficiency reasons, the player set of BA* consists the small set SV^r randomly selected among the set of all users. Thus, our powerful Adversary, although unable to corrupt 1/3 of all the users, can certainly corrupt all members of SV^r !

Fortunately we’ll prove that protocol BA* , executed by propagating messages in a peer-to- peer fashion, is player-replaceable. This novel requirement means that the protocol correctly and efficiently reaches consensus even if each of its step is executed by a totally new, and randomly and independently selected, set of players. Thus, with millions of users, each small set of players associated to a step of BA* most probably has empty intersection with the next set.

In addition, the sets of players of different steps of BA* will probably have totally different cardinalities. Furthermore, the members of each set do not know who the next set of players will be, and do not secretly pass any internal state.

The replaceable-player property is actually crucial to defeat the dynamic and very powerful Adversary we envisage. We believe that replaceable-player protocols will prove crucial in lots of contexts and applications. In particular, they will be crucial to execute securely small sub protocols embedded in a larger universe of players with a dynamic adversary, who, being able to corrupt even a small fraction of the total players, has no difficulty in corrupting all the players in the smaller sub-protocol.

An Additional Property /Technique: Lazy Honesty A honest user follows his prescribed instructions, which include being online and run the protocol. Since, Algorand has only modest computation and communication requirement, being online and running the protocol “in the background” is not a major sacrifice. Of course, a few “absences” among honest players, as those due to sudden loss of connectivity or the need of rebooting, are automatically tolerated (because we can always consider such few players to be temporarily malicious) . Let us point out, however, that Algorand can be simply adapted so as to work in a new model, in which honest users to be offline most of the time. Our new model can be informally introduced as follows.

Lazy Honesty. Roughly speaking, a user i is lazy-but-honest if (1) he follows all his prescribed instructions, when he is asked to participate to the protocol, and (2) he is asked to participate to the protocol only rarely, and with a suitable advance notice.

With such a relaxed notion of honesty, we may be even more confident that honest people will be at hand when we need them, and Algorand guarantee that, when this is the case,

The system operates securely even if, at a given point in time, the majority of the participating players are malicious.

2 Preliminaries

2.1 Cryptographic Primitives

Ideal Hashing. We shall rely on an efficiently computable cryptographic hash function, H, that maps arbitrarily long strings to binary strings of fixed length. Following a long tradition, we model H as a random oracle , essentially a function mapping each possible string s to a randomly and independently selected (and then fixed) binary string, H{s) , of the chosen length.

In our described embodiments, H has 256-bit long outputs. Indeed, such length is short enough to make the system efficient and long enough to make the system secure. For instance, we want H to be collision-resilient. That is, it should be hard to find two different strings x and y such that H{x) = H{y) . When H is a random oracle with 256-bit long outputs, finding any such pair of strings is indeed difficult. (Trying at random, and relying on the birthday paradox, would require 2²⁵⁶/² = 2¹²⁸ trials.)

Digital Signing. Digital signatures allow users to to authenticate information to each other without sharing any sharing any secret keys. A digital signature scheme consists of three fast algorithms: a probabilistic key generator G, a signing algorithm S, and a verification algorithm

V.

Given a security parameter k, a sufficiently high integer, a user i uses G to produce a pair of fc-bit keys (i.e., strings) : a“public” key pki and a matching“secret” signing key ski. Crucially, a public key does not “betray” its corresponding secret key. That is, even given knowledge of pki, no one other than i is able to compute ski in l^ess than astronomical time.

User i uses ski to digitally sign messages. For each possible message (binary string) m, i first hashes m and then runs algorithm S on inputs H(m ) and ski ^s° ^as to produce the fc-bit string

sig_pki {m) = S(H(m) , ski) ?

The binary string sig_pk {m) is referred to as F s digital signature of m (relative to pki) , and can be more simply denoted by sigfim ) , when the public key pki is clear from context.

Everyone knowing pki can use it to verify the digital signatures produced by i. Specifically, on inputs (a) the public key ki of a player i, (b) a message m, and (c) a string s, that is, F s alleged digital signature of the message m, the verification algorithm V outputs either YES or NO.

The properties we require from a digital signature scheme are:

1. Legitimate signatures are always verified: If s = sigfim) , then V(pki, m, s) = YES; and

²Since H is collision-resilient it is practically impossible that, by signing m one“accidentally signs” a different message m! . 2. Digital signatures are hard to forge: Without knowledge of ski the time to find a string s such that V(pki, m, s) = YES, for a message m never signed by i, is astronomically long.

(Following strong security requirements, this is true even if one can obtain the signature of any other message.)

Accordingly, to prevent anyone else from signing messages on his behalf, a player i must keep his signing key ski secret (hence the term “secret key”) , and to enable anyone to verify the messages he does sign, i has an interest in publicizing his key pki (hence the term“public key”) .

Signatures with Message Retrievability In general, a message m is not retrievable from its signature sigfim). In order to virtually deal with digital signatures that satisfy the conceptually convenient “message retrievability” property (i.e., to guarantee that the signer and the message are easily computable from a signature, we define

SIG_pk^m) = (i, m, sig_pk^m)) and SIGi{m) = {i, m, sigfim)), if pki is clear.

Unique Digital Signing. We also consider digital signature schemes (G, S, V) satisfying the following additional property.

3. Uniqueness. It is hard to find strings pk' , m, s, and s' such that

s s' and V{pk', m, s) = V{pk', m, s') = 1.

(Note that the uniqueness property holds also for strings pk' that are not legitimately generated public keys. In particular, however, the uniqueness property implies that, if one used the specified key generator G to compute a public key pk together with a matching secret key sk, and thus knew sk, it would be essentially impossible also for him to find two different digital signatures of a same message relative to pk.)

Remarks

• FROM UNIQUE SIGNATURES TO VERIFIABLE RANDOM FUNCTIONS . Relative to a digital signature scheme with the uniqueness property, the mapping m— H{sigi{m )) associates to each possible string m, a unique, randomly selected, 256-bit string, and the correctness of this mapping can be proved given the signature sigfim) .

That is, ideal hashing and digital signature scheme satisfying the uniqueness property essentially provide an elementary implementation of a verifiable random function (VRF) .

A VRF is a special kind of digital signature. We may write VRFfim) to indicate such a special signature of i of a message m. In addition to satisfy the uniqueness property, verifiable random functions produce outputs that are guaranteed to be sufficiently random. That is, VRFfim) is essentially random, and unpredictable until it is produced. By contrast, SIGfim) need not be sufficiently random. For instance, user i may choose his public key so that SIGfim) always is a fc-bit string that is (lexicographically) small (i.e., whose first few bits could always be Os) . Note, however, that, since H is an ideal hash function, H{SIGi{m )) will always be a random 256-bit string. In our preferred embodiments we make extensive use of hashing digital signatures satisfying the uniqueness property precisely to be able to associate to each message m and each user i a unique random number. Should one implement Algorand with VRFs, one can replace H{SIGi{m )) with V RFfim) . In particular, a user i need not first to compute SIGfim ) , then H{SIGi{m )) (in order, say to compare H{SIGi{m )) with a number p) . He might directly compute V RFfim) . In sum, it should be understood that H{SIGi{m )) can be interpreted as VRFi(m) , or as a sufficiently random number, easily computed by player i, but unpredictable to anyone else, unambiguously associated to i and m.

• THREE DIFFERENT NEEDS FOR DIGITAL SIGNATURES . In Algorand, a user i relies on digital signatures for

(1) Authenticating i’s own payments. In this application, keys can be “long-term” (i.e., used to sign many messages over a long period of time) and come from a ordinary signature scheme.

(2) Generating credentials proving that i is entitled to act at some step s of a round r. Here, keys can be long-term, but must come from a scheme satisfying the uniqueness property.

(3) Authenticating the message i sends in each step in which he acts. Here, keys must be ephemeral (i.e., destroyed after their first use), but can come from an ordinary signature scheme.

• A SMALL-COST SIMPLIFICATION. For simplicity, we envision each user i to have a single long- term key. Accordingly, such a key must come from a signature scheme with the uniqueness property. Such simplicity has a small computational cost. Typically, in fact, unique digital signatures are slightly more expensive to produce and verify than ordinary signatures.

2.2 The Idealized Public Ledger

Algorand tries to mimic the following payment system, based on an idealized public ledger.

1. The Initial Status. Money is associated with individual public keys (privately generated and owned by users). Letting pk , . . . , pk_j be the initial public keys and a , . . . , a_j their respective initial amounts of money units, then the initial status is

-¾ = (r&i, ai), . . . , (/ , ¾·) ,

which is assumed to be common knowledge in the system.

2. Payments. Let pk be a public key currently having a > 0 money units, pk! another public key, and a! a non-negative number no greater than a. Then, a (valid) payment p is a digital signature, relative to pk, specifying the transfer of a' monetary units from pk to pk' , together with some additional information. In symbols, p = SIG_pk(pk, pk^,, a^,, I, H(T ) ) ,

where / represents any additional information deemed useful but not sensitive (e.g., time information and a payment identifier) , and T any additional information deemed sensitive (e.g., the reason for the payment, possibly the identities of the owners of pk and the pk' , and so on) .

We refer to pk (or its owner) as the payer , to each pk' (or its owner) as a payee , and to a' as the amount of the payment p.

Free Joining Via Payments. Note that users may join the system whenever they want by generating their own public/secret key pairs. Accordingly, the public key pk' that appears in the payment p above may be a newly generated public key that had never“owned” any money before. 3. The Magic Ledger. In the Idealized System, all payments are valid and appear in a tamper proof list L of sets of payments“posted on the sky” for everyone to see:

L = PAY¹, PAY², . . . ,

Each block PAY^r+1 consists of the set of all payments made since the appearance of block PAY^r . In the ideal system, a new block appears after a fixed (or finite) amount of time.

Discussion.

• More General Payments and Unspent Transaction Output. More generally, if a public key pk owns an amount a, then a valid payment p of pk may transfer the amounts a , a'₂, . . ., respectively to the keys pk^ pk^, . . ., so long as

a) < a.

In Bitcoin and similar systems, the money owned by a public key pk is segregated into separate amounts, and a payment p made by pk must transfer such a segregated amount a in its entirety. If pk wishes to transfer only a fraction a' < a of a to another key, then it must also transfer the balance, the unspent transaction output, to another key, possibly pk itself.

Algorand also works with keys having segregated amounts. However, in order to focus on the novel aspects of Algorand, it is conceptually simpler to stick to our simpler forms of payments and keys having a single amount associated to them.

• Current Status. The Idealized Scheme does not directly provide information about the current status of the system (i.e., about how many money units each public key has) . This information is deducible from the Magic Ledger.

In the ideal system, an active user continually stores and updates the latest status information, or he would otherwise have to reconstruct it, either from scratch, or from the last time he computed it. (Yet, we later on show how to augment Algorand so as to enable its users to reconstruct the current status in an efficient manner.)

• Security and “Privacy” . Digital signatures guarantee that no one can forge a payment of another user. In a payment p, the public keys and the amount are not hidden, but the sensitive information T is. Indeed, only H(I) appears in p, and since H is an ideal hash function, H(I) is a random 256-bit value, and thus there is no way to figure out what I was better than by simply guessing it. Yet, to prove what I was (e.g., to prove the reason for the payment) the payer may just reveal I. The correctness of the revealed I can be verified by computing H(I) and comparing the resulting value with the last item of p. In fact, since H is collision resilient, it is hard to find a second value T' such that H(I) = H Z').

2.3 Basic Notions and Notations

Keys, Users, and Owners Unless otherwise specified, each public key (“key” for short) is long-term and relative to a digital signature scheme with the uniqueness property. A public key i joins the system when another public key j already in the system makes a payment to i.

For color, we personify keys. We refer to a key i as a“he”, say that i is honest, that i sends and receives messages, etc. User is a synonym for key. When we want to distinguish a key from the person to whom it belongs, we respectively use the term“digital key” and“owner”.

Permissionless and Permissioned Systems. A system is permissionless, if a digital key is free to join at any time and an owner can own multiple digital keys; and its permissioned, otherwise. Unique Representation Each object in Algorand has a unique representation. In particular, each set {(x, y, z, . . .) : x £ X, y £ Y, z £ Z, . . .} is ordered in a pre-specified manner: e.g., first lexicographically in x, then in y, etc.

Same-Speed Clocks There is no global clock: rather, each user has his own clock. User clocks need not be synchronized in any way. We assume, however, that they all have the same speed.

For instance, when it is 12pm according to the clock of a user i, it may be 2:30pm according to the clock of another user j, but when it will be 12:01 according to f’s clock, it will 2:31 according to j’s clock. That is, “one minute is the same (sufficiently, essentially the same) for every user” .

Rounds Algorand is organized in logical units, r = 0, 1, . . ., called rounds.

We consistently use superscripts to indicate rounds. To indicate that a non- numerical quantity Q (e.g., a string, a public key, a set, a digital signature, etc.) refers to a round r, we simply write Q^r . Only when Q is a genuine number (as opposed to a binary string interpretable as a number), do we write Q^^r so that the symbol r could not be interpreted as the exponent of Q.

At (the start of a) round r > 0, the set of all public keys is PK^r , and the system status is

(r)

where a) is the amount of money available to the public key i. Note that PK^r is deducible from S^r , and that S^r may also specify other components for each public key i.

For round 0, PK° is the set of initial public keys, and S'⁰ is the initial status. Both PK° and S'⁰ are assumed to be common knowledge in the system. For simplicity, at the start of round r, so are PK \ . . . , PK^r and S¹ , . . . , S^r.

In a round r, the system status transitions from S^r to S^r+1 : symbolically,

Round r: S^r ® S^r+1.

Payments In Algorand, the users continually make payments (and disseminate them in the way described in subsection 2.7) . A payment p of a user i 6 PK^r has the same format and semantics as in the Ideal System. Namely,

p = SIG_t{i, i', a, I, H{l )) .

Payment p is individually valid at a round r (is a round-r payment , for short) if (1) its amount a is less than or equal to ^\ and (2) it does not appear in any official payset PAY^r' for r ' < r. (As explained below, the second condition means that p has not already become effective.

A set of round-r payments of i is collectively valid if the sum of their amounts is at most

Paysets A round-r payset V is a set of round-r payments such that, for each user i, the payments of i in V (possibly none) are collectively valid. The set of all round-r paysets is PAY(r) . A round-r payset V is maximal if no superset of is a round-r payset.

We actually suggest that a payment p also specifies a round p, p = SIGi{p, i, i a, I, H (I)) , and cannot be valid at any round outside [p, p + k] , for some fixed non-negative integer fc.^{3 3}This simplifies checking whether p has become“effective” (i.e. , it simplifies determining whether some payset PAY^r contains p. When k = 0, if p = SIGi (r, i, i' , a, I , H(X)) , and p ^ PAY^r , then i must re-submit p. Official Paysets For every round r, Algor and publicly selects (in a manner described later on) a single (possibly empty) payset, PAY^r , the round’s official pay set. (Essentially, PAY^r represents the round-r payments that have “actually” happened.)

As in the Ideal System (and Bitcoin) , (1) the only way for a new user j to enter the system is to be the recipient of a payment belonging to the official payset PAY^r of a given round r; and (2) PAY^r determines the status of the next round, S'^r+1, from that of the current round, S^r . Symbolically,

PAY^r : S^r ® S^r+1.

Specifically,

1. the set of public keys of round r + 1, PK^r+1 , consists of the union of PK^r and the set of all payee keys that appear, for the first time, in the payments of PAY^r ; and

2. the amount of money

that a user i owns in round r + 1 is the sum of <¾(r) i.e., the amount of money i owned in the previous round (0 if i ^ PK^r) and the sum of amounts paid to i according to the payments of PAY^r .

In sum, as in the Ideal System, each status S^r+l is deducible from the previous payment history:

PAY°, . . . , PAY^r.

2.4 Blocks and Proven Blocks

In Algorand_Q, the block B^r corresponding to a round r specifies: r itself; the set of payments of round r, PAY^r ; a quantity

to be explained, and the hash of the previous block,

Thus, starting from some fixed block B°, we have a traditional blockchain:

B¹ = (l, PAY¹ , SIG_efiQ ) , H(B⁰)) , B² = (2, PAY² , SIGffiQ¹) , ^(B¹)) , . . .

In Algorand, the authenticity of a block is actually vouched by a separate piece of information, a“block certificate” CERT^r, which turns B^r into a proven block, B^r . The Magic Ledger, therefore, is implemented by the sequence of the proven blocks,

Discussion As we shall see, CERT^r consists of a set of digital signatures for H{B^r ) , those of a majority of the members of SV^r , together with a proof that each of those members indeed belongs to SV^r . We could, of course, include the certificates CERT^r in the blocks themselves, but find it conceptually cleaner to keep it separate.)

In Bitcoin each block must satisfy a special property, that is, must “contain a solution of a crypto puzzle”, which makes block generation computationally intensive and forks both inevitable and not rare. By contrast, Algorand’s blockchain has two main advantages: it is generated with minimal computation, and it will not fork with overwhelmingly high probability. Each block B * is safely final as soon as it enters the blockchain.

2.5 Acceptable Failure Probability

To analyze the security of Algorand we specify the probability, F, with which we are willing to accept that something goes wrong (e.g., that a verifier set SV^r does not have an honest majority). As in the case of the output length of the cryptographic hash function H, also F is a parameter. But, as in that case, we find it useful to set I⁷ to a concrete value, so as to get a more intuitive grasp of the fact that it is indeed possible, in Algorand, to enjoy simultaneously sufficient security and sufficient efficiency. To emphasize that F is parameter that can be set as desired, in the first and second embodiments we respectively set

F = 11G¹² and F = HG¹⁸ .

Discussion Note that 11D¹² is actually less than one in a trillion, and we believe that such a choice of F is adequate in our application. Let us emphasize that 11D¹² is not the probability with which the Adversary can forge the payments of an honest user. All payments are digitally signed, and thus, if the proper digital signatures are used, the probability of forging a payment is far lower than 10 ¹², and is, in fact, essentially 0. The bad event that we are willing to tolerate with probability F is that Algorand’s blockchain forks. Notice that, with our setting of F and one-minute long rounds, a fork is expected to occur in Algorand’s blockchain as infrequently as (roughly) once in 1.9 million years. By contrast, in Bitcoin, a forks occurs quite often.

A more demanding person may set i⁷ to a lower value. To this end, in our second embodiment we consider setting F to 11D¹⁸. Note that, assuming that a block is generated every second , 10¹⁸ is the estimated number of seconds taken by the Universe so far: from the Big Bang to present time. Thus, with F = 10 ¹⁸ , if a block is generated in a second, one should expect for the age of the Universe to see a fork.

2.6 The Adversarial Model

Algorand is designed to be secure in a very adversarial model. Let us explain.

Honest and Malicious Users A user is honest if he follows all his protocol instructions, and is perfectly capable of sending and receiving messages. A user is malicious (i.e., Byzantine , in the parlance of distributed computing) if he can deviate arbitrarily from his prescribed instructions.

The Adversary The Adversary is an efficient (technically polynomial-time) algorithm, personified for color, who can immediately make malicious any user he wants, at any time he wants (subject only to an upperbound to the number of the users he can corrupt) .

The Adversary totally controls and perfectly coordinates all malicious users. He takes all actions on their behalf, including receiving and sending all their messages, and can let them deviate from their prescribed instructions in arbitrary ways. Or he can simply isolate a corrupted user sending and receiving messages. Let us clarify that no one else automatically learns that a user i is malicious, although Ls maliciousness may transpire by the actions the Adversary has him take.

This powerful adversary however,

• Does not have unbounded computational power and cannot successfully forge the digital signature of an honest user, except with negligible probability; and

• Cannot interfere in any way with the messages exchanges among honest users.

Furthermore, his ability to attack honest users is bounded by one of the following assumption.

Honesty Majority of Money We consider a continuum of Honest Majority of Money (HMM) assumptions: namely, for each non-negative integer k and real h > 1/2,

HHM_k > h: the honest users in every round r owned a fraction greater than h of all money in the system at round r— k. Discussion. Assuming that all malicious users perfectly coordinate their actions (as if controlled by a single entity, the Adversary) is a rather pessimistic hypothesis. Perfect coordination among too many individuals is difficult to achieve. Perhaps coordination only occurs within separate groups of malicious players. But, since one cannot be sure about the level of coordination malicious users may enjoy, we’d better be safe than sorry.

Assuming that the Adversary can secretly, dynamically, and immediately corrupt users is also pessimistic. After all, realistically, taking full control of a user’s operations should take some time.

The assumption HMM^ > h implies, for instance, that, if a round (on average) is implemented in one minute, then, the majority of the money at a given round will remain in honest hands for at least two hours, if k = 120, and at least one week, if k = 10, 000.

Note that the HMM assumptions and the previous Honest Majority of Computing Power assumptions are related in the sense that, since computing power can be bought with money, if malicious users own most of the money, then they can obtain most of the computing power.

2.7 The Communication Model

We envisage message propagation i.e., “peer-to-peer gossip”⁴ to be the only means of communication, and assume that every propagated message reaches almost all honest users in a timely fashion. We essentially assume that each message m propagated by honest user reaches, within a given amount of time that depends on the length of m, all honest users. (It actually suffices that m reaches a sufficiently high percentage of the honest users.)

3 The BA Protocol BA * in a Traditional Setting

As already emphasized, Byzantine agreement is a key ingredient of Algorand. Indeed, it is through the use of such a BA protocol that Algorand is unaffected by forks. However, to be secure against our powerful Adversary, Algorand must rely on a BA protocol that satisfies the new player-replaceability constraint. In addition, for Algorand to be efficient, such a BA protocol must be very efficient.

BA protocols were first defined for an idealized communication model, synchronous complete networks (SC networks). Such a model allows for a simpler design and analysis of BA protocols. Accordingly, in this section, we introduce a new BA protocol, BA* , for SC networks and ignoring the issue of player replaceability altogether. The protocol BA* is a contribution of separate value. Indeed, it is the most efficient cryptographic BA protocol for SC networks known so far.

To use it within our Algorand protocol, we modify BA* a bit, so as to account for our different communication model and context.

We start by recalling the model in which BA* operates and the notion of a Byzantine agreement.

3.1 Synchronous Complete Networks and Matching Adver saries

In a SC network, there is a common clock, ticking at each integral times r = 1, 2, . . .

⁴Essentially, as in Bitcoin, when a user propagates a message m, every active user i receiving m for the first time, randomly and independently selects a suitably small number of active users, his “neighbors”, to whom he forwards m, possibly until he receives an acknowledgement from them. The propagation of m terminates when no user receives m for the first time. At each even time click r, each player i instantaneously and simultaneously sends a single message rn^ (possibly the empty message) to each player j, including himself. Each rn^ is correctly received at time click r + 1 by player j, together with the identity of the sender i.

Again, in a communication protocol, a player is honest if he follows all his prescribed instructions, and malicious otherwise. All malicious players are totally controlled and perfectly coordinated by the Adversary, who, in particular, immediately receives all messages addressed to malicious players, and chooses the messages they send.

The Adversary can immediately make malicious any honest user he wants at any odd time click he wants, subject only to a possible upperbound t to the number of malicious players. That is, the Adversary“cannot interfere with the messages already sent by an honest user i”, which will be delivered as usual.

The Adversary also has the additional ability to see instantaneously, at each even round, the messages that the currently honest players send, and instantaneously use this information to choose the messages the malicious players send at the same time tick.

3.2 The Notion of a Byzantine Agreement

The notion of Byzantine agreement might have been first introduced for the binary case, that is, when every initial value consists of a bit. However, it was quickly extended to arbitrary initial values. By a BA protocol, we mean an arbitrary-value one.

Definition 3.1. In a synchronous network, let V be a n-player protocol, whose player set is common knowledge among the players, t a positive integer such that n > 2t + 1. We say that V is an arbitrary- value (respectively, binary; (n, t)-Byzantine agreement protocol with soundness s £ (0, 1) if, for every set of values V not containing the special symbol T (respectively, for V = {0, 1};, in an execution in which at most t of the players are malicious and in which every player i starts with an initial value 6 V, every honest player j halts with probability 1, outputting a value outi £ V U {T} so as to satisfy, with probability at least s, the following two conditions:

1. Agreement: There exists out £ V U {T} such that outi = out for all honest players i.

2. Consistency:

We refer to out as P’s output, and to each outi ^as player Is output.

3.3 The BA Notation

In our BA protocols, a player is required to count how many players sent him a given message in a given step. Accordingly, for each possible value v that might be sent,

(or just _ί(n) when s is clear) is the number of players j from which i has received v in step s.

Recalling that a player i receives exactly one message from each player j, if the number of players is n, then, for all i and s, _u Fί(ⁿ) ^{= n}-

3.4 The New Binary BA Protocol BBA*

In this section we present a new binary BA protocol, BBA* , which relies on the honesty of more than two thirds of the players and is very fast: no matter what the malicious players might do, each execution of its main loop not only is trivially executed, but brings the players into agreement with probability 1/3.

In BBA *, each player has his own public key of a digital signature scheme satisfying the unique-signature property. Since this protocol is intended to be run on synchronous complete network, there is no need for a player i to sign each of his messages.

Digital signatures are used to generate a sufficiently common random bit in Step 3. (In Algorand, digital signatures are used to authenticate all other messages as well.)

The protocol requires a minimal set-up: a common random string r, independent of the players’ keys. (In Algorand, r is actually replaced by the quantity Q^r .)

Protocol BBA * is a 3-step loop, where the players repeatedly exchange Boolean values, and different players may exit this loop at different times. A player i exits this loop by propagating, at some step, either a special value 0* or a special value 1*, thereby instructing all players to “pretend” they respectively receive 0 and 1 from i in all future steps. (Alternatively said: assume that the last message received by a player j from another player i was a bit b. Then, in any step in which he does not receive any message from i, j acts as if i sent him the bit b.)

The protocol uses a counter 7, representing how many times its 3-step loop has been executed. At the start of BBA*, 7 = 0. (One may think of 7 as a global counter, but it is actually increased by each individual player every time that the loop is executed.)

There are n > 3t + l, where t is the maximum possible number of malicious players. A binary string x is identified with the integer whose binary representation (with possible leadings 0s) is x; and lsb(x) denotes the least significant bit of x.

PROTOCOL BBA*

(COMMUNICATION) STEP 1. [Coin-Fixed-To-0 Step] Each player i sends bi .

(COMMUNICATION) STEP 2. [Coin-Fixed-To-1 Step] Each player i sends bi .

2.1 If ff_i iX) > 2t + 1, then i sets bi = 1, sends 1*, outputs outi = 1, and HALTS.

2.2 If #?(0) > 2t + 1, then i set bi = 0.

2.3 Else, i sets bi = 1.

(COMMUNICATION) STEP 3. [Coin- Genuinely-Flipped Step] Each player i sends bi and 5/G_j(r, 7) .

Theorem 3.1. Whenever n > 3i + 1, BBA * is a binary ( n, t) -BA protocol with soundness 1.

A proof of Theorem 3.1 can be found in https://people.csail.mit.edu/silvio/ SelectedScien- tificPapers/DistributedComputation/BYZANTINEAGREEMENTMADETRIVIAL. 15 pdf. 3.5 Graded Consensus and the Protocol GC

Let us recall, for arbitrary values, a notion of consensus much weaker than Byzantine agreement.

Definition 3.2. Let V be a protocol in which the set of all players is common knowledge, and each player i privately knows an arbitrary initial value

We say that V is an (n, t)-graded consensus protocol if, in every execution with n players, at most t of which are malicious, every honest player i halts outputting a value-grade pair (_¾, _<¾), where g^ G {0, 1, 2}, so as to satisfy the following three conditions:

1. For all honest players i and j, \g_%— Qj \ < 1.

The following two-step protocol GC is a graded consensus protocol in the literature. To match the steps of protocol Algorand^ of section 4.1, we respectively name 2 and 3 the steps of GC. (Indeed, the first step of Algorand^ is concerned with something else: namely, proposing a new block.)

PROTOCOL GC

STEP 3. Each player i sends to all players the string x if and only if i£j{x) > 2/ + 1.

OUTPUT DETERMINATION. Each player i outputs the pair (_¾, _<¾) computed as follows:

Since protocol GC is a protocol in the literature, it is known that the following theorem holds.

Theorem 3.2. If n > 3t + 1, then GC is a ( n, t)-graded broadcast protocol.

3.6 The Protocol BA ^k

We now describe the arbitrary-value BA protocol BA* via the binary BA protocol BBA * and the graded-consensus protocol GC. Below, the initial value of each player i is v^.

PROTOCOL BA*

STEPS 1 AND 2. Each player i executes GC, on input v^, so as to compute a pair (i¾, _<¾) .

STEP 3 , . . . Each player i executes BBA * with initial input 0, if gi = 2, and 1 otherwise so as to compute the bit outi.

OUTPUT DETERMINATION.

Theorem 3.3. Whenever n > 3t + 1, BA* is a ( n, t)-BA protocol with soundness 1. Proof. We first prove Consistency, and then Agreement.

PROOF OF CONSISTENCY. Assume that, for some value v G V, v = v. Then, by property 3 of graded consensus, after the execution of GC, all honest players output (v, 2) . Accordingly, 0 is the initial bit of all honest players in the end of the execution of BBA *. Thus, by the Agreement property of binary Byzantine agreement, at the end of the execution of BA* , outi = 0 for all honest players. This implies that the output of each honest player i in BA*

= v. □

PROOF OF AGREEMENT. Since BBA* is a binary BA protocol, either

(A) outi = 1 for all honest player i, or

(B) outi = 0 for all honest player i.

In case A, all honest players output _L in BA* , and thus Agreement holds. Consider now case B. In this case, in the execution of BBA*, the initial bit of at least one honest player i is 0. (Indeed, if initial bit of all honest players were 1, then, by the Consistency property of BBA* , we would have out_j = 1 for all honest j.) Accordingly, after the execution of GC, i outputs the pair (v, 2) for some value v. Thus, by property 1 of graded consensus, g_j > 0 for all honest players j. Accordingly, by property 2 of graded consensus, V_j = v for all honest players j. This implies that, at the end of BA* , every honest player j outputs v. Thus, Agreement holds also in case B. □

Since both Consistency and Agreement hold, BA* is an arbitrary-value BA protocol. ■

Protocol BA* works also in gossiping networks, and in fact satisfies the player replaceability property that is crucial for Algorand to be secure in the envisaged very adversarial model.

The Player Replaceability of BBA * and BA* Let us now provide some intuition of why the protocols BA* and BBA * can be adapted to be executed in a network where communication is via peer-to-peer gossiping, satisfy player replaceability. For concreteness, assume that the network has 10M users and that each step x of BBA * (or BA*) is executed by a committee of 10,000 players, who have been randomly selected via secret cryptographic sortition, and thus have credentials proving of being entitled to send messages in step x. Assume that each message sent in a given step specifies the step number, is digitally signed by its sender, and includes the credential proving that its sender is entitled to speak in that step.

First of all, if the percentage h of honest players is sufficiently larger than 2/3 (e.g., 75%) , then, with overwhelming probability, the committee selected at each step has the required 2/3 honest majority.

In addition, the fact that the 10,000-strong randomly selected committee changes at each step does not impede the correct working of either BBA * or BA* . Indeed, in either protocol, a player i in step s only reacts to the multiplicity with which, in Step s— 1, he has received a given message m. Since we are in a gossiping network, all messages sent in Step s— 1 will (immediately, for the purpose of this intuition) reach all users, including those selected to play in step s. Furthermore because all messages sent in step s— 1 specify the step number and include the credential that the sender was indeed authorized to speak in step s— 1. Accordingly, whether he happened to have been selected also in step s— 1 or not, a user i selected to play in step s is perfectly capable of correctly counting the multiplicity with which he has received a correct step s— 1 message. It does not at all matter whether he has been playing all steps so far or not. All users are in“in the same boat” and thus can be replaced easily by other users. 4 Two Embodiments of Algorand

As discussed, at a very high level, a round of Algorand ideally proceeds as follows. First, a randomly selected user, the leader, proposes and circulates a new block. (This process includes initially selecting a few potential leaders and then ensuring that, at least a good fraction of the time, a single common leader emerges.) Second, a randomly selected committee of users is selected, and reaches Byzantine agreement on the block proposed by the leader. (This process includes that each step of the BA protocol is run by a separately selected committee.) The agreed upon block is then digitally signed by a given threshold (¾) of committee members. These digital signatures are propagated so that everyone is assured of which is the new block. (This includes circulating the credential of the signers, and authenticating just the hash of the new block, ensuring that everyone is guaranteed to learn the block, once its hash is made clear.)

In the next two sections, we present two embodiments of the basic Algorand design, Algorand ( and Algorand'₂ , that respectively work under a proper majority-of-honest-users assumption. In Section 7 we show how to adopts these embodiments to work under a honest- majority-of-money assumption.

Algorand ( only envisages that > 2/3 of the committee members are honest. In addition, in Algorand ( , the number of steps for reaching Byzantine agreement is capped at a suitably high number, so that agreement is guaranteed to be reached with overwhelming probability within a fixed number of steps (but potentially requiring longer time than the steps of Algorand ₂) . In the remote case in which agreement is not yet reached by the last step, the committee agrees on the empty block, which is always valid.

Algorand₂ envisages that the number of honest members in a committee is always greater than or equal to a fixed threshold tn (which guarantees that, with overwhelming probability, at least 2/3 of the committee members are honest). In addition, Algorand '₂ allows Byzantine agreement to be reached in an arbitrary number of steps (but potentially in a shorter time than Algorand ().

Those skilled in the art will realize that many variants of these basic embodiments can be derived. In particular, it is easy, given Algorand₂, to modify Algorand ( so as to enable to reach Byzantine agreement in an arbitrary number of steps.

Both embodiments share the following common core, notations, notions, and parameters.

4.1 A Common Core

Objectives Ideally, for each round r, Algorand should satisfy the following properties:

1. Perfect Correctness. All honest users agree on the same block B^r .

2. Completeness 1. With probability 1, the block B^r has been chosen by a honest user.

(Indeed a malicious user may always choose a block whose payset contains the payments of just his“friends”.)

Of course, guaranteeing perfect correctness alone is trivial: everyone always chooses the official payset PAY^r to be empty. But in this case, the system would have completeness 0. Unfortunately, guaranteeing both perfect correctness and completeness 1 is not easy in the presence of malicious users. Algorand thus adopts a more realistic objective. Informally, letting h denote the percentage of users who are honest, h > 2/3, the goal of Algorand is

Guaranteeing, with overwhelming probability, perfect correctness and completeness close to h.

Privileging correctness over completeness seems a reasonable choice: payments not processed in one round can be processed in the next, but one should avoid forks, if possible. Led Byzantine Agreement Disregarding excessive time and communication for a moment, perfect Correctness could be guaranteed as follows. At the start of round r, each user i proposes his own candidate block B^'· . Then, all users reach Byzantine agreement on just one of the candidate blocks. As per our introduction, the BA protocol employed requires a 2/3 honest majority and is player replaceable. Each of its step can be executed by a small and randomly selected set of verifiers, who do not share any inner variables.

Unfortunately, this approach does not quite work. This is so, because the candidate blocks proposed by the honest users are most likely totally different from each other. Indeed, each honest user sees different payments. Thus, although the sets of payments seen by different honest users may have a lot of overlap, it is unlikely that all honest users will construct a propose the same block. Accordingly, the consistency agreement of the BA protocol is never binding, only the agreement one is, and thus agreement may always been reached on _L rather than on a good block.

Algorand ' avoids this problem as follows. First, a leader for round r, i^r , is selected. Then, i^r propagates his own candidate block, Bfi^' . Finally, the users reach agreement on the block they actually receive from i^r . Because, whenever i^r is honest, Perfect Correctness and Completeness 1 both hold, Algorand ' ensures that i^r is honest with probability close to h.

Leader Selection In Algorand’s, the rth block is of the form

As already mentioned in the introduction, the quantity Q ^r_1 is carefully constructed so as to be essentially non-manipulatable by our very powerful Adversary. (Later on in this section, we shall provide some intuition about why this is the case.) At the start of a round r, all users know the blockchain so far, B°, . . . , B^{1 1} , from which they deduce the set of users of every prior round: that is, PK¹ , . . . , PK^{r l}. A potential leader of round r is a user i such that

.H {SIG_t {r, l, Q^{r 1})) < p .

Let us explain. Note that, since the quantity Q ^{r_ 1} is deducible from block B ^{r_ 1}, because of the message retrievability property of the underlying digital signature scheme. Furthermore, the underlying signature scheme satisfies the uniqueness property. Thus, SIGi (r, l, Q^{r 1}^ is a binary string uniquely associated to i and r. Accordingly, since H is a random oracle, H {SIG_t (r , 1, !) ) is a random 256-bit long string uniquely associated to i and r. The symbol in front of H ( SIGi (r, 1 , Q^{r 1})) is the decimal (in our case, binary ) point, so that ri = .H ( SIGi (r, 1, Q^{r _ 1}) ) is the binary expansion of a random 256-bit number between 0 and 1 uniquely associated to i and r. Thus the probability that r* is less than or equal to p is essentially p.

The probability p is chosen so that, with overwhelming (i.e., 1— F) probability, at least one potential verifier is honest. (If fact, p is chosen to be the smallest such probability.)

Note that, since i is the only one capable of computing his own signatures, he alone can determine whether he is a potential verifier of round 1. However, by revealing his own credential, sI^' = SIGi (r, 1, Q^{r _1}) , i can prove to anyone to be a potential verifier of round r.

The leader i^r is defined to be the potential leader whose hashed credential is smaller that the hashed credential of all other potential leader j: that is,

< H{a^r-^,s) .

Note that, since a malicious i^r may not reveal his credential, the correct leader of round r may never be known, and that, barring improbable ties, i^r is indeed the only leader of round r.

Let us finally bring up a last but important detail: a user i can be a potential leader (and thus the leader) of a round r only if he belonged to the system for at least k rounds. This guarantees the non-manipulatability of Q^r and all future Q-quantities. In fact, one of the potential leaders will actually determine Q^r .

Verifier Selection Each step s > 1 of round r is executed by a small set of verifiers, SV^r,s. Again, each verifier i 6 SV^r,s is randomly selected among the users already in the system k rounds before r, and again via the special quantity Q ^r_1. Specifically, i 6 PK^{r k} is a verifier in SV^r’^s, if

Once more, only i knows whether he belongs to SV^ but, if this is the case, he could prove it by exhibiting his credential s^tU = H(SIGi (r, s, Q^{7"- 1}) ) . A verifier i 6 SV^r,s sends a message, m/⁵, in step s of round r, and this message includes his credential s/’⁵, so as to enable the verifiers f the nest step to recognize that rrfiY is a legitimate step-s message.

The probability p^r is chosen so as to ensure that, in SV^r,s, letting figood be the number of honest users and fibad the number of malicious users, with overwhelming probability the following two conditions hold.

For embodiment Algorandfi

(1) figood > 2 · fibad and

(2) figood + 4 · fibad < 2 n, where n is the expected cardinality of SV^r,s.

For embodiment Algorandfi.

(1) figood > t and

(2) figood + 2 fibad < 2 t , where tn is a specified threshold.

These conditions imply that, with sufficiently high probability, (a) in the last step of the BA protocol, there will be at least given number of honest players to digitally sign the new block B^r , (b) only one block per round may have the necessary number of signatures, and (c) the used BA protocol has (at each step) the required 2/3 honest majority.

Clarifying Block Generation If the round-r leader i^r is honest, then the corresponding block is of the form

where the payset PAY^r is maximal (recall that all paysets are, by definition, collectively valid.) Else (i.e., if i^r is malicious) , B^r has one of the following two possible forms:

In the first form, PAY^r is a (non-necessarily maximal) payset and it may be PAY^r = 0; and i is a potential leader of round r. (However, i may not be the leader i^r . This may indeed happen if if i^r keeps secret his credential and does not reveal himself.)

The second form arises when, in the round-r execution of the BA protocol, all honest players output the default value, which is the empty block B_e ^t in our application. (By definition, the possible outputs of a BA protocol include a default value, generically denoted by _L. See section 3.2.)

Note that, although the paysets are empty in both cases, B^r = (r, 0, SIGi

) and B_e ^t are syntactically different blocks and arise in two different situations: respectively,“all went smoothly enough in the execution of the BA protocol” , and“something went wrong in the BA protocol, and the default value was output” .

Let us now intuitively describe how the generation of block B^r proceeds in round r of Algorand' . In the first step, each eligible player, that is, each player i 6 PK^{r k}, checks whether he is a potential leader. If this is the case, then i is asked, using of all the payments he has seen so far, and the current blockchain, B°, . . . , B^{r 1}, to secretly prepare a maximal payment set, PAY⁷ , and secretly assembles his candidate block, B⁷ = ( r, PAY , SIGi (Q^{7 1}) , H ( B That, is, not only does he include in B^" _· , as its second component, the just prepared payset, but also, as its third component, his own signature of Q^{r_ 1}, the third component of the last block, B^{r l} . Finally, he propagates his round-r-step-1 message, mf’¹, which includes (a) his candidate block Bt, (b) his proper signature of his candidate block (i.e., his signature of the hash of B- , and (c) his own credential o⁷’¹, proving that he is indeed a potential verifier of round r.

(Note that, until an honest i produces his message m⁷’¹ , the Adversary has no clue that i is a potential verifier. Should he wish to corrupt honest potential leaders, the Adversary might as well corrupt random honest players. However, once he sees , since it contains Fs credential, the Adversary knows and could corrupt i, but cannot prevent m⁷’¹ , which is virally propagated, from reaching all users in the system.)

In the second step, each selected verifier j 6 SV ^7,2 tries to identify the leader of the round. Specifically, j takes the step-1 credentials, o⁷’¹ , . . . ,

contained in the proper step-1 message m⁷’¹ he has received; hashes all of them, that is, computes cR)¹ _? finds the

credential, s , whose hash is lexicographically minimum; and considers P to be the leader of round r.

Recall that each considered credential is a digital signature of Q ^{r_ 1} , that SIGi (r, 1, Q^{7 1}) is uniquely determined by i and Q^{7 1} , that H is random oracle, and thus that each H(SIGi ( r , 1, Q^{7 1}) is a random 256-bit long string unique to each potential leader i of round r.

From this we can conclude that, if the 256-bit string Q ^r_1 were itself randomly and independently selected , than so would be the hashed credentials of all potential leaders of round r. In fact, all potential leaders are well defined, and so are their credentials (whether actually computed or not) . Further, the set of potential leaders of round r is a random subset of the users of round r— k, and an honest potential leader i always properly constructs and propagates his message m(, which contains Fs credential. Thus, since the percentage of honest users is h, no matter what the malicious potential leaders might do (e.g., reveal or conceal their own credentials) , the minimum hashed potential- leader credential belongs to a honest user, who is necessarily identified by everyone to be the leader i⁷ of the round r. Accordingly, if the 256-bit string Q ^{r_ 1} were itself randomly and independently selected, with probability exactly h (a) the leader i⁷ is honest and (b) i_j = i⁷ for all honest step-2 verifiers j.

In reality, the hashed credential are, yes, randomly selected, but depend on Q^{r_ 1} , which is not randomly and independently selected. A careful analysis, however, guarantees that Q ^{r 1} is sufficiently non-manipulatable to guarantee that the leader of a round is honest with probability h' sufficiently close to h: namely, h' > h²(l + h— h²) . For instance, if h = 80%, then h' > .7424.

Having identified the leader of the round (which they correctly do when the leader i^r is honest) , the task of the step-2 verifiers is to start executing BA* using as initial values what they believe to be the block of the leader. Actually, in order to minimize the amount of communication required, a verifier j 6 SV^r,2 does not use, as his input value v'_· to the Byzantine protocol, the block B_j that he has actually received from i_j (the user j believes to be the leader) , but the the leader, but the hash of that block, that is, v'_· = H{Bi) . Thus, upon termination of the BA protocol, the verifiers of the last step do not compute the desired round-r block B^r , but compute (authenticate and propagate) H{B^r) . Accordingly, since H{B^r) is digitally signed by sufficiently many verifiers of the last step of the BA protocol, the users in the system will realize that H{B^r) is the hash of the new block. However, they must also retrieve (or wait for, since the execution is quite asynchronous) the block B^r itself, which the protocol ensures that is indeed available, no matter what the Adversary might do.

Asynchrony and Timing Algorand^ and Algorand^ have a significant degree of asynchrony. This is so because the Adversary has large latitude in scheduling the delivery of the messages being propagated. In addition, whether the total number of steps in a round is capped or not, there is the variance contribute by the number of steps actually taken.

As soon as he learns the certificates of B°, . . . , IT^-1 , a user i computes Q ^r_1 and starts working on round r, checking whether he is a potential leader, or a verifier in some step s of round r.

Assuming that i must act at step s, in light of the discussed asynchrony, i relies on various strategies to ensure that he has sufficient information before he acts.

For instance, he might wait to receive at least a given number of messages from the verifiers of the previous step (as in Algorand ^) , or wait for a sufficient time to ensure that he receives the messages of sufficiently many verifiers of the previous step (as in Algorand^) .

The Seed Q^r and the Look-Back Parameter k Recall that, ideally, the quantities Q^r should random and independent, although it will suffice for them to be sufficiently non- manipulatable by the Adversary.

At a first glance, we could choose Q ^r_1 to coincide with H (PAy^r_1) . An elementary analysis reveals, however, that malicious users may take advantage of this selection mechanism.⁵ Some additional effort shows that myriads of other alternatives, based on traditional block quantities are easily exploitable by the Adversary to ensure that malicious leaders are very frequent. We instead specifically and inductively define our brand new quantity Q^r so as to be able to prove that it is non-manipulatable by the Adversary. Namely,

Q^r =

if B^r is not the empty block, and Q^r = H(Q^{r 1} , r) otherwise.

The intuition of why this construction of Q^r works is as follows. Assume for a moment that Q ^r_1 is truly randomly and independently selected. Then, will so be Q^r ? When £^r is honest the answer is (roughly speaking) yes. This is so because

H(SIG_r (·), r ) : {0, l}²⁵⁶ ® {0, l}²⁵⁶

is a random function. When £^r is malicious, however, Q^r is no longer univocally defined from Q ^r_1 and £^r . There are at least two separate values for Q^r . One continues to be Q^r = H(SIGe_r (Q^{r 1}) , r) , and the other is H(Q^{r 1} , r) . Let us first argue that, while the second

⁵We are at the start of round r 1. Thus, Q^{r 2} = PAY^{r 2} is publicly known, and the Adversary privately knows who are the potential leaders he controls. Assume that the Adversary controls 10% of the users, and that, with very high probability, a malicious user w is the potential leader of round r 1. That is, assume that

H (SIG (r 2, 1, Q^{r 2})) is so small that it is highly improbable an honest potential leader will actually be the leader of round r 1. (Recall that, since we choose potential leaders via a secret cryptographic sortition mechanism, the Adversary does not know who the honest potential leaders are.) The Adversary, therefore, is in the enviable position of choosing the payset PAY' he wants, and have it become the official payset of round r 1. However, he can do more. He can also ensure that, with high probability, (*) one of his malicious users will be the leader also of round r, so that he can freely select what PAY^r will be. (And so on. At least for a long while, that is, as long as these high-probability events really occur.) To guarantee (*), the Adversary acts as follows. Let PAY' be the payset the Adversary prefers for round r 1. Then, he computes H(PAY') and checks whether, for some already malicious player z, SIG_z (r, l, H(PAY')) is particularly small, that is, small enough that with very high probability z will be the leader of round r. If this is the case, then he instructs w to choose his candidate block to be = (r 1, PAY', H(B^{r 2}). Else, he has two other malicious users x and y to keep on generating a new payment r' , from one to the other, until, for some malicious user z (or even for some fixed user z) p[ ( SIG (PAY' U {p})) is particularly small too. This experiment will stop quite quickly. And when it does the Adversary asks w to propose the candidate block

= (r 1, PAY' U {p}, H(B^{r 2})). choice is somewhat arbitrary, a second choice is absolutely mandatory. The reason for this is that a malicious i^r can always cause totally different candidate blocks to be received by the honest verifiers of the second step.⁶ Once this is the case, it is easy to ensure that the block ultimately agreed upon via the BA protocol of round r will be the default one, and thus will not contain anyone’s digital signature of Q ^r_1. But the system must continue, and for this, it needs a leader for round r. If this leader is automatically and openly selected, then the Adversary will trivially corrupt him. If it is selected by the previous Q ^{r_ 1} via the same process, than £^r will again be the leader in round r + 1. We specifically propose to use the same secret cryptographic sortition mechanism, but applied to a new Q-quantity: namely, H(Q^{r 1}, r) . By having this quantity to be the output of H guarantees that the output is random, and by including r as the second input of H, while all other uses of H have either a single input or at least three inputs, “guarantees” that such a Q^r is independently selected. Again, our specific choice of alternative Q^r does not matter, what matter is that i^r has two choice for Q^r , and thus he can double his chances to have another malicious user as the next leader.

The options for Q^r may even be more numerous for the Adversary who controls a malicious i^r . For instance, let x, y, and z be three malicious potential leaders of round r such that

and particulary small. That is, so small that there is a good chance that

is smaller of the hashed credential of every honest potential leader. Then, by asking x to hide his credential, the Adversary has a good chance of having y become the leader of round r— 1. This implies that he has another option for Q^r: namely, H ( SIG_y (Q^{r 1}) , r) . Similarly, the Adversary may ask both x and y of withholding their credentials, so as to have z become the leader of round r— 1 and gaining another option for Q^r: namely, H ( SIG_Z (Q^{r 1}) , r) .

Of course, however, each of these and other options has a non-zero chance to fail, because the Adversary cannot predict the hash of the digital signatures of the honest potential users.

A careful, Mar kov-chain- like analysis shows that, no matter what options the Adversary chooses to make at round r— 1, as long as he cannot inject new users in the system, he cannot decrease the probability of an honest user to be the leader of round r + 40 much below h. This is the reason for which we demand that the potential leaders of round r are users already existing in round r— k. It is a way to ensure that, at round r— k, the Adversary cannot alter by much the probability that an honest user become the leader of round r. In fact, no matter what users he may add to the system in rounds r— k through r, they are ineligible to become potential leaders (and a fortiori the leader) of round r. Thus the look-back parameter k ultimately is a security parameter. (Although, as we shall see in section ??, it can also be a kind of “convenience parameter” as well.)

Ephemeral Keys Although the execution of our protocol cannot generate a fork, except with negligible probability, the Adversary could generate a fork, at the rth block, after the legitimate block r has been generated.

Roughly, once B^r has been generated, the Adversary has learned who the verifiers of each step of round r are. Thus, he could therefore corrupt all of them and oblige them to certify a new block B^r . Since this fake block might be propagated only after the legitimate one, users that have been paying attention would not be fooled.⁷ Nonetheless, B^r would be syntactically correct and we want to prevent from being manufactured.

⁶For instance, to keep it simple (but extreme),“when the time of the second step is about to expire”, £^r could directly email a different candidate block Bi to each user i. This way, whoever the step-2 verifiers might be, they will have received totally different blocks.

⁷Consider corrupting the news anchor of a major TV network, and producing and broadcasting today a We do so by means of a new rule. Essentially, the members of the verifier set SV^r,s of a step s of round r use ephemeral public keys pk^ to digitally sign their messages. These keys are single-use-only and their corresponding secret keys sk 'Y are destroyed once used. This way, if a verifier is corrupted later on, the Adversary cannot force him to sign anything else he did not originally sign.

Naturally, we must ensure that it is impossible for the Adversary to compute a new key p^rY and convince an honest user that it is the right ephemeral key of verifier i 6 SV^r,s to use in step s.

4.2 Common Summary of Notations, Notions, and Parameters

Notations

• r > 0: the current round number.

• s > 1: the current step number in round r.

• B^r: the block generated in round r.

• PK^r: the set of public keys by the end of round r— 1 and at the beginning of round r.

• S^r: the system status by the end of round r— 1 and at the beginning of round r.⁸

• PAY^r : the payset contained in B^r .

• i^r: round-r leader. i^r chooses the payset PAY^r of round r (and determines the next Q^r) .

• Q^r : the seed of round r, a quantity (i.e., binary string) that is generated at the end of round r and is used to choose verifiers for round r + 1. Q^r is independent of the paysets in the blocks and cannot be manipulated by i^r .

• SV^r,s: the set of verifiers chosen for step s of round r.

• SV^r: the set of verifiers chosen for round r, SV^r = U_s>i SV^r,s.

• M SV^r,s and HSV^r,s: respectively, the set of malicious verifiers and the set of honest verifiers in SV^r’^s. MSV^r’^s U HSV^r’^s = SV^r’^s and MSV^r’^s P HSV^r’^s = 0.

• n i e Έ⁺ and n 6 Z⁺: respectively, the expected numbers of potential leaders in each SV^1"’¹ , and the expected numbers of verifiers in each SV^r,s, for s > 1.

Notice that n « n, since we need at least one honest honest member in SV^7"’¹, but at least a majority of honest members in each SV^r,s for s > 1.

• ft £ (0, 1) : a constant greater than 2/3. h is the honesty ratio in the system. That is, the fraction of honest users or honest money, depending on the assumption used, in each PK^r is at least h.

• H: a cryptographic hash function, modelled as a random oracle.

• _L: A special string of the same length as the output of H.

• F 6 (0, 1): the parameter specifying the allowed error probability. A probability < F is considered“negligible” , and a probability > 1— W is considered“overwhelming” .

• p_h 6 (0, 1): the probability that the leader of a round r, i^r , is honest. Ideally p_h = h. With the existence of the Adversary, the value of

will be determined in the analysis.

newsreel showing secretary Clinton winning the last presidential election. Most of us would recognize it as a hoax. But someone getting out of a coma might be fooled.

⁸In a system that is not synchronous, the notion of “the end of round r— 1” and “the beginning of round r” need to be carefully defined. Mathematically, PK^r and S^r are computed from the initial status S° and the blocks B¹ , . . . , B^r-¹. • k £ Z⁺ : the look-back parameter. That is, round r— k is where the verifiers for round r are chosen from namely, SV^r C PK^{r k 9}

• pi £ (0, 1) : for the first step of round r, a user in round r— k is chosen to be in SV⁷'¹ with probability pi = _|p^__{fc |} .

• p £ (0, 1): for each step s > 1 of round r, a user in round r— k is chosen to be in SV^r,s with probability

• CERT^r : the certificate for B^r . It is a set of tn signatures of H{B^r) from proper verifiers in round r.

proven block.

A user i knows B^r if he possesses (and successfully verifies) both parts of the proven block. Note that the CERT^r seen by different users may be different.

• T T : the (local) time at which a user i knows B^r . In the Algorand protocol each user has his own clock. Different users’ clocks need not be synchronized, but must have the same speed. Only for the purpose of the analysis, we consider a reference clock and measure the players’ related times with respect to it.

• a'C and 0^r£^s·. respectively the (local) time a user i starts and ends his execution of Step s of round r.

• A and l: essentially, the upper-bounds to, respectively, the time needed to execute Step 1 and the time needed for any other step of the Algorand protocol.

Parameter A upper-bounds the time to propagate a single 1MB block.

Parameter l upperbounds the time to propagate one small message per verifier in a Step s > 1.

We assume that A < 4L.

Notions

• Verifier selection.

For each round r and step

Each user i 6 PK^{r k} privately computes his signature using his long-term key and decides whether i £ SV^r,s or not. If i £ SV^r,s, then SIGi(r, s,

is i’s (r, s)-credential, compactly denoted by [^,!> .

For the first step of round r, SV^r'¹ and s)^"’¹ are similarly defined, with p replaced by p . The verifiers in SV ^r’¹ are potential leaders.

• Leader selection.

User i £ SV^r'¹ is the leader of round r, denoted by £^r ,

for all potential leaders j £ SV^r,1. Whenever the hashes of two players’ credentials are compared, in the unlikely event of ties, the protocol always breaks ties lexicographically according to the (long-term public keys of the) potential leaders.

By definition, the hash value of player £^{r ,}s credential is also the smallest among all users in PK^{r k} . Note that a potential leader cannot privately decide whether he is the leader or not, without seeing the other potential leaders’ credentials.

Since the hash values are uniform at random, when SV^r,1 is non-empty, £^r always exists and is honest with probability at least h. The parameter n_\ is large enough so as to ensure that each SV^r,1 is non-empty with overwhelming probability.

⁹Strictly speaking, “r k” should be“max{0, r fc}”. • Block structure.

A non-empty block is of the form B^r = (r, PAY' ,

and an empty block is of the form

Note that a non-empty block may still contain an empty payset PAY^r, if no payment occurs in this round or if the leader is malicious. However, a non-empty block implies that the identity of £^r , his credential s// and SIGp-

have all been timely revealed. The protocol guarantees that, if the leader is honest, then the block will be non-empty with overwhelming probability.

• Seed Q^r .

If B^r is non-empty, then Q^r = H(SIGe (Q^{r 1}) , r) , otherwise Q^r = H(Q^{r 1} , r).

Parameters

• Relationships among various parameters.

The verifiers and potential leaders of round r are selected from the users in PK^{r k}, where k is chosen so that the Adversary cannot predict Q ^r_1 back at round r— k— 1 with probability better than F : otherwise, he will be able to introduce malicious users for round r— k, all of which will be potential leaders/ verifiers in round r, succeeding in having a malicious leader or a malicious majority in SV^r,s for some steps s desired by him.

For Step 1 of each round r, n is chosen so that with overwhelming probability, SV 0.

• Example choices of important parameters.

The outputs of H are 256-bit long.

h = 80%, m = 35.

A = 1 minute and L = 15 seconds.

• Initialization of the protocol.

The protocol starts at time 0 with r = 0. Since there does not exist“H^-1” or“CERT ¹” , syntactically

is a public parameter with its third component specifying Q^-1 , and all users know

time 0.

5 Algorand^

In this section, we construct a version of Algorand ' working under the following assumption.

HONEST MAJORITY OF USERS ASSUMPTION: More than 2/3 of the users in each PK^r are honest.

In Section 7, we show how to replace the above assumption with the desired Honest Majority of Money assumption.

5.1 Additional Notations and Parameters

Notations

• m £ Z⁺ : the maximum number of steps in the binary BA protocol, a multiple of 3. • L^r < m/ 3: a random variable representing the number of Bernoulli trials needed to see a 1, when each trial is 1 with probability and there are at most mj 3 trials. If all trials fail then L^r = m/3. U will be used to upper-bound the time needed to generate block B^r .

• i_// = + 1: the number of signatures needed in the ending conditions of the protocol.

• CERT^r: the certificate for B'^’ . It is a set of tn signatures of H{B^r) from proper verifiers in round r.

Parameters

• Relationships among various parameters.

For each step s > 1 of round r, n is chosen so that, with overwhelming probability,

| HSV^r'^s \ > 2\MSV^r,s \ and \HSV^r'^s\ + 4| MSV^r'^s\ < 2 n.

The closer to 1 the value of h is, the smaller n needs to be. In particular, we use (variants of) Chernoff bounds to ensure the desired conditions hold with overwhelming probability.

m is chosen such that L^r < mj 3 with overwhelming probability.

• Example choices of important parameters.

F = 11G¹².

1500, k = 40 and m = 180.

5.2 Implementing Ephemeral Keys in Algorand

As already mentioned, we wish that a verifier i 6 SV^r,s digitally signs his message m^r/^s of step s in round r, relative to an ephemeral public key pk^r/^s , using an ephemeral secrete key sk^r/^s that he promptly destroys after using. We thus need an efficient method to ensure that every user can verify that pk^r/^s is indeed the key to use to verify i’s signature of m^r/^s . We do so by a (to the best of our knowledge) new use of identity-based signature schemes.

At a high level, in such a scheme, a central authority A generates a public master key, PMK, and a corresponding secret master key, SMK. Given the identity, U, of a player U, A computes, via SMK, a secret signature key sk_j relative to the public key U, and privately gives sk _j to U. (Indeed, in an identity-based digital signature scheme, the public key of a user U is U itself!) This way, if A destroys SMK after computing the secret keys of the users he wants to enable to produce digital signatures, and does not keep any computed secret key, then U is the only one who can digitally sign messages relative to the public key U. Thus, anyone who knows“C/’s name”, automatically knows C/’s public key, and thus can verify C/’s signatures (possibly using also the public master key PMK) .

In our application, the authority A is user i, and the set of all possible users U coincides with the round-step pair (r, s ) in say S = {/} x {r' , . . . , r' + 10⁶} x {1, . . . , m + 3}, where r' is a given round, and m + 3 the upperbound to the number of steps that may occur within a round. This way, pk^rf^s = (i, r, s), so that everyone seeing i’ s signature SIG^^r _r,s (m^r _i ^,s) can, with overwhelming probability, immediately verify it for the first million rounds r following r' .

In other words, i first generates PMK and SMK. Then, he publicizes that PMK is C’s master public key for any round r 6 \r' , r' + 10⁶] , and uses SMK to privately produce and store the secret key sk/’^s for each triple ( i, r , s ) 6 S. This done, he destroys SMK. If he determines that he is not part of SV^r,s, then i may leave sk^r/^s alone (as the protocol does not require that he aunthenticates any message in Step s of round r). Else, i first uses sk/’^s to digitally sign his message m^r/^s , and then destroys sk^r/^s . Note that i can publicize his first public master key when he first enters the system. That is, the same payment p that brings i into the system (at a round r' or at a round close to r') , may also specify, at Fs request, that Fs public master key for any round r 6 [r', r' + 10⁶] is PMK e.g., by including a pair of the form ( PMK , [r', r' + 10⁶] ).

Also note that, since m + 3 is the maximum number of steps in a round, assuming that a round takes a minute, the stash of ephemeral keys so produced will last i for almost two years. At the same time, these ephemeral secret keys will not take i too long to produce. Using an elliptic-curve based system with 32B keys, each secret key is computed in a few microseconds. Thus, if m + 3 = 180, then all 180M secret keys can be computed in less than one hour.

When the current round is getting close to r' + 10°, to handle the next million rounds, i generates a new (RMK' , SMK') pair, and informs what his next stash of ephemeral keys is by for example having SIGi{PMK' , [r' + 10° + l, r' + 2 · 10° + 1]) enter a new block, either as a separate“transaction” or as some additional information that is part of a payment. By so doing, i informs everyone that he/she should use PMK¹ to verify Z’s ephemeral signatures in the next million rounds. And so on.

(Note that, following this basic approach, other ways for implementing ephemeral keys without using identity-based signatures are certainly possible. For instance, via Merkle trees.¹⁰) Other ways for implementing ephemeral keys are certainly possible e.g., via Merkle trees.

5.3 Matching the Steps of Algorand’_x with those of BA*

As we said, a round in Algorand ^ has at most m + 3 steps.

STEP 1. In this step, each potential leader i computes and propagates his candidate block B , together with his own credential, s^’¹.

Recall that this credential explicitly identifies i. This is so, because a ¹ = SIGi(r, 1, Q^{r _1}) .

Potential verifier i also propagates, as part of his message, his proper digital signature of Not dealing with a payment or a credential, this signature of i is relative to his ephemeral public key

that is, he propagates sig_pkr,i (H(B [)).

Given our conventions, rather than propagating B^ and sig _kr,i (H(B )) , he could have propagated S I G _kr,i {H (B )) . However, in our analysis we need to have explicit access to sig_p x {H{Bl)) ^P‘

STEPS 2. In this step, each verifier i sets /'( to be the potential leader whose hashed credential is the smallest, and B- to be the block proposed by /'(. Since, for the sake of efficiency, we wish to agree on H(B^r), rather than directly on B^r , i propagates the message he would have propagated in the first step of BA* with initial value = H(B^' _· ). That is, he propagates after ephemerally signing it, of course. (Namely, after signing it relative to the right ephemeral public key, which in this case is pk^².) Of course too, i also transmits his own credential.

Since the first step of BA* consists of the first step of the graded consensus protocol GC, Step 2 of Algorand ' corresponds to the first step of GC.

¹⁰In this method, i generates a public-secret key pair (pfc”’^s, sfc(^,s) for each round-step pair (r, s) in— say— {V, . . . , r' + 10⁶} x { 1, . . . , m + 3}. Then he orders these public keys in a canonical way, stores the _jith public key in the _jith leaf of a Merkle tree, and computes the root value ¾, which he publicizes. When he wants to sign a message relative to key pfc(’^s, i not only provides the actual signature, but also the authenticating path for pk^’ relative to ¾. Notice that this authenticating path also proves that pfc”’^s is stored in the _jith leaf. Form this idea, the rest of the details can be easily filled. STEPS 3. In this step, each verifier i £ SV^r,‘2 executes the second step of BA* . That is, he sends the same message he would have sent in the second step of GC. Again, Fs message is ephemerally signed and accompanied by Fs credential. (From now on, we shall omit saying that a verifier ephemerally signs his message and also propagates his credential.)

STEP 4. In this step, every verifier i 6 SV^{r, i} computes the output of GC, ( Vi ) , and ephemerally signs and sends the same message he would have sent in the third step of BA* , that is, in the first step of BBA* , with initial bit 0 if g = 2, and 1 otherwise.

STEP S = 5, . . . , m + 2. Such a step, if ever reached, corresponds to step s— 1 of BA* , and thus to step s— 3 of BBA*.

Since our propagation model is sufficiently asynchronous, we must account for the possibility that, in the middle of such a step s, a verifier i 6 SV^r,s is reached by information proving him that block B^r has already been chosen. In this case, i stops his own execution of round r of Algorand' , and starts executing his round-(r + 1) instructions.

Accordingly, the instructions of a verifier i 6 SV^r,s, in addition to the instructions corresponding to Step s— 3 of BBA*, include checking whether the execution of BBA* has halted in a prior Step s'. Since BBA * can only halt is a Coin-Fixed-to-0 Step or in a Coin-Fixed-to-1 step, the instructions distinguish whether

A (Ending Condition 0): s'— 2 º 0 mod 3, or

B (Ending Condition 1): s'— 2 º 1 mod 3.

In fact, in case A, the block B^r is non-empty, and thus additional instructions are necessary to ensure that i properly reconstructs B^r , together with its proper certificate CERT^r . In case B, the block B^r is empty, and thus i is instructed to set B^r = B_e ^t = (r, 0, Q^{r 1} , H(B^{r 1})), and to compute CERT^r.

If, during his execution of step s, i does not see any evidence that the block B^r has already been generated, then he sends the same message he would have sent in step s— 3 of BBA* .

STEP m + 3. If, during step m + 3, i £ SV^r,m+s sees that the block B^r was already generated in a prior step s', then he proceeds just as explained above.

Else, rather then sending the same message he would have sent in step m of BBA*, i is instructed, based on the information in his possession, to compute B^r and its corresponding certificate CERT^r .

Recall, in fact, that we upperbound by m + 3 the total number of steps of a round.

5.4 The Actual Protocol

Recall that, in each step s of a round r, a verifier i £ SV^r,s uses his long-term public-secret key pair to produce his credential, s(’⁵ = SIGi{r, s, Q^{1 1}), as well as SIGi (z ^-1) in case s = 1. Verifier i uses his ephemeral secret key sk^ to sign his (r, s)-message m ^s . For simplicity, when r and s are clear, we write esigi{x) rather than sig_pkr,s (x^ to denote Fs proper ephemeral signature of a value x in step s of round r, and write ESIGi(x) instead of SIG_pkr,s x ) to denote (i, x, esigi(x)) .

Step 1: Block Proposal

Instructions for every user i 6 PK^{r k}: User i starts his own Step 1 of round r as soon as he knows B ^{r 1}. • User i computes Q^{r 1} from the third component of B^{r 1} and checks whether i 6 SV ^7,1 or not.

• If i then i stops his own execution of Step 1 right away.

• If i 6 SV^1"’¹ , that is, if i is a potential leader, then he collects the round-r payments that have been propagated to him so far and computes a maximal payset PAY? from them.

Next, he computes his“candidate block”

= (r, PAY[ , SIGi(Q^{r 1}) , H(B^{r 1})) . Finally, he computes the message

, destroys his ephemeral secret key s q’¹ , and then propagates r,l Remark. In practice, to shorten the global execution of Step 1, it is important that the (r, l)-messages are selectively propagated. That is, for every user i in the system, for the first (r, l)-message that he ever receives and successfully verifies,¹¹ player i propagates it as usual. For all the other (r, l)-messages that player i receives and successfully verifies, he propagates it only if the hash value of the credential it contains is the smallest among the hash values of the credentials contained in all (r, l)-messages he has received and successfully verified so far. Furthermore, as suggested by Georgios Vlachos, it is useful that each potential leader i also propagates his credential u ¹ separately: those small messages travel faster than blocks, ensure timely propagation of the m^’¹’ s where the contained credentials have small hash values, while make those with large hash values disappear quickly.

¹¹ That is, all the signatures are correct and both the block and its hash are valid although i does not check whether the included payset is maximal for its proposer or not. Step 2: The First Step of the Graded Consensus Protocol GC

Instructions for every user i G PK^{r k}: User i starts his own Step 2 of round r as soon as he knows B^{1 1}.

• User i computes Q ^r_1 from the third component of B^{r 1} and checks whether i E SV ^r’² or not.

• SV^r,‘2 then i stops his own execution of Step 2 right away.

• If i 6 SV^r'² , then after waiting an amount of time t_{2 —} A + L, i acts as follows.

1. He finds the user i such that H(s_i ^¹) < H^s^¹) for all credentials a⁷ _·'¹ that are part of the successfully verified (r, l)-messages he has received so far.¹²

2. If he has received from i a valid message m/’¹ =

esig_£(H(B^)), s/’¹),¹³ then i sets v = H{B_i ^r)·, otherwise i sets v = _L.

3. i computes the message m ^7,2 = (ESIGi(v^/ _i) , a^7,‘2) ,¹⁴ destroys his ephemeral secret key sfc/’², and then propagates m/’².

Step 3: The Second Step of GC

Instructions for every user i 6 PK^{r k}: User i starts his own Step 3 of round r as soon as he knows B ^r_1.

• User i computes Q ^r_1 from the third component of B ^{r_ 1} and checks whether i 6 SV^r,s or not.

• SV^r,s, then i stops his own execution of Step 3 right away.

• If i 6 SV^r,s, then after waiting an amount of time /¾— t₂ + 2l = 3l + L, i acts as follows.

1. If there exists a value n' _L such that, among all the valid messages m ^r _· ² he has received, more than 2/3 of them are of the form ESIG_j v'), a^r£²), without any contradiction,¹⁵ then he computes the message m/’³ = ( ESIGi(v'), a /’³) . Otherwise, he computes m/’³ (ESIGi(-L), s/’³).

2. i destroys his ephemeral secret key sk^ , and then propagates m .

Step 4: Output of GC and The First Step of BBA *

Instructions for every user i 6 PK^{r k} : User i starts his own Step 4 of round r as soon as he knows B^{r 1}.

¹²Essentially, user i privately decides that the leader of round r is user £.

¹³Again, player t’s signatures and the hashes are all successfully verified, and PAY[ in SJ is a valid payset for round r— although i does not check whether PAY[ is maximal for l or not.

¹⁴The message m/ signals that player i considers

to be the hash of the next block, or considers the next block to be empty.

¹⁵That is, he has not received two valid messages containing ESIG_j (v') and a different ESIG_j (v") respectively, from a player j. Here and from here on, except in the Ending Conditions defined later, whenever an honest player wants messages of a given form, messages contradicting each other are never counted or considered valid. • User i computes Q^{r 1} from the third component of B^{r 1} and checks whether i 6 SV^r,4 or not.

• If then i his stops his own execution of Step 4 right away.

• If i 6 SV^7"’⁴, then after waiting an amount of time i₄ =

+ 2l = 5l + L, i acts as follows.

1. He computes and the output of GC, as follows.

(a) If there exists a value v' _L such that, among all the valid messages m^! _· ³ he has received, more than 2/3 of them are of the form ( ESIG_j (v'), a ⁷’³), then he sets

(b) Otherwise, if there exists a value n' _L such that, among all the valid messages rry ³ he has received, more than 1/3 of them are of the form ( ESIG_j(v') , a f’³) , then he sets

(c) Else, he sets

2. He computes bi , the input of BBA *, as follows:

bi = 0 if gi = 2, and ¾ = 1 otherwise.

3. He computes the message

, destroys his ephemeral secret key sfc/’⁴, and then propagates m/’⁴.

Step s, 5 < s < m + 2, s— 2 º 0 mod 3: A Coin-Fixed- To-0 Step of BBA *

Instructions for every user i 6 PK^{r k}: User i starts his own Step s of round r as soon as he knows B^{1 1}.

• User i computes Q ^r_1 from the third component of B^{1 1} and checks whether i 6 SV^r,s.

• If i SV^r,s, then i stops his own execution of Step s right away.

• If i 6 SV^r,s then he acts as follows.

- He waits until an amount of time t_s = t_{s- i} + 2l = (2s— 3)A + L has passed.

- Ending Condition 0: If, during such waiting and at any point of time, there exists a string r 1 and a step s' such that

(a) 5 < s' < s, s'— 2 º 0 mod 3 that is, Step s' is a Coin-Fixed- To-0 step,

(b) i has received at least t = ^ + 1 valid messages m^r-'^{s 1} = (ESIG_j 0) , ESIG_j (v) , a_j ^{r S 4}) , ¹⁷ and

(c) i has received a valid message m¹ _·

esig_j{H{B^r _· )) , a_j ¹) with v = H(B_j ), then, i stops his own execution of Step s (and in fact of round r) right away without propagating anything; sets B^r = B^r _· and sets his own CERT^r to be the set of messages m^r- ^{s 1} of sub-step (b).¹⁸

1⁶It can be proved that the u' in case (b), if exists, must be unique.

¹⁷Such a message from player j is counted even if player i has also received a message from j signing for 1. Similar things for Ending Condition 1. As shown in the analysis, this is done to ensure that all honest users know B^r within time l from each other.

¹⁸User i now knows B^r and his own round r finishes. He still helps propagating messages as a generic user, but does not initiate any propagation as a (r, s)-verifier. In particular, he has helped propagating all messages in his CERT^r , which is enough for our protocol. Note that he should also set b = 0 for the binary BA protocol, but b is not needed in this case anyway. Similar things for all future instructions. Ending Condition 1: If, during such waiting and at any point of time, there exists a step s' such that

(a’) 6 < s' < s, s'— 2 º 1 mod 3 that is, Step s' is a Coin-Fixed- To-1 step, and (b’) i has received at least t valid messages m^r-'^{s 1} = (ESIG_j (l) , ESIG_j(v_j) , r,s'— l _\

s3 ) ^>

then, i stops his own execution of Step s (and in fact of round r) right away without propagating anything; sets B^r = H/; and sets his own CERT^r to be the set of messages m^!- ^{s 1} of sub-step (b’) .

Otherwise, at the end of the wait, user i does the following.

He sets to be the majority vote of the V_j in the second components of all the valid m^!- ^{s 1}’s he has received.

He computes ¾ as follows.

If more than 2/3 of all the valid m^!- ^{s 1}’s he has received are of the form

Else, he sets

= 0.

He computes the message

= ( ESIGi{bi ) , ESIGi(vi), s/’⁵) , destroys his ephemeral secret key sk^ , and then propagates

Step s, 6 < s < m + 2, s— 2 º 1 mod 3: A Coin-Fixed- To-1 Step of BBA *

Instructions for every user i 6 PK^{r k}: User i starts his own Step s of round r as soon as he knows B^{1 1}.

• User i computes Q ^r_1 from the third component of B ^r_1 and checks whether i 6 SV^r,s or not.

• If / / SV^r,s, then i stops his own execution of Step s right away.

• If i 6 SV^r,s then he does the follows.

He waits until an amount of time t_s = (2 s— 3)A + L has passed.

Ending Condition 0: The same instructions as Coin-Fixed- To-0 steps.

Ending Condition 1: The same instructions as Coin-Fixed- To-0 steps.

Otherwise, at the end of the wait, user i does the following.

He sets to be the majority vote of the V_j in the second components of all the valid m^!- ^{s 1}’s he has received.

He computes ¾ as follows.

If more than 2/3 of all the valid m^!- ^{s 1}’s he has received are of the form

¹⁹In this case, it does not matter what the ¾’s are. Else, he sets ¾ = 1.

He computes the message m/’⁵ = (ESIGi(bi) , ESIGi(vi) , a^^,s) , destroys his ephemeral secret key sk^ , and then propagates m/’⁵.

Step s, 7 < s < m + 2, s— 2 º 2 mod 3: A Coin- Genuinely-Flipped Step of BBA *

Instructions for every user i 6 PK^{r k}: User i starts his own Step s of round r as soon as he knows B ^r_1.

• User i computes Q ^r_1 from the third component of B ^r_1 and checks whether i 6 SV^r,s or not.

• If / / SV^r,s , then i stops his own execution of Step s right away.

• If i 6 SV^r,s then he does the follows.

He waits until an amount of time t_s = (2 s— 3)l + L has passed.

Ending Condition 0: The same instructions as Coin-Fixed- To-0 steps.

Ending Condition 1: The same instructions as Coin-Fixed- To-0 steps.

Otherwise, at the end of the wait, user i does the following.

He sets to be the majority vote of the V_j in the second components of all the valid m^!- ^{s 1}’s he has received.

He computes ¾ as follows.

If more than 2/3 of all the valid m^!- ^{s 1}’s he has received are of the form

Else, let SV ^{,s 1} be the set of (r, s— 1)- verifiers from whom he has received a valid message m^!- ^{s 1}. He sets ¾ = lsb

.

He computes the message m/’⁵ = (ESIGi(bi) , ESIGi(vi) , a^^,s) , destroys his ephemeral secret key

and then propagates m'C .

Step m + 3: The Last Step of BBA * ²⁰

Instructions for every user i 6 PK^{r k}: User i starts his own Step m + 3 of round r as soon as he knows B^{r 1}.

• User i computes Q ^{r_ 1} from the third component of B ^r_1 and checks whether i 6 SV^r,m+s or not.

• If i / 5V^r,m+3, then i stops his own execution of Step m + 3 right away.

then he does the follows.

²⁰With overwhelming probability BBA * has ended before this step, and we specify this step for completeness. - He waits until an amount of time t_m+3 = t_{m+ 2} + 2l = (2 m + 3)l + L has passed.

- Ending Condition 0: The same instructions as Coin-Fixed- To-0 steps.

- Ending Condition 1: The same instructions as Coin-Fixed- To-0 steps.

- Otherwise, at the end of the wait, user i does the following.

He sets outi = 1 and B^r = B .

He computes the message m’^m+3 = ESIG/outf) , ESIGi(H(B^r)), s/^hί+Ά) , destroys his ephemeral secret key sfc’^m+3, and then propagates m’^m+3 to certify B^r .²¹

Reconstruction of the Round-r Block by Non- Verifiers

Instructions for every user i in the system: User i starts his own round r as soon as he knows B^{r_ 1}, and waits for block information as follows.

- If, during such waiting and at any point of time, there exists a string v and a step s' such that

(a) 5 < s' < m + 3 with s'— 2 º 0 mod 3,

(b) i has received at least t valid messages mf^{r s 1} = (ESIG_j( 0), ESIG_j (v) , af^{r s 3}) , and

(c) i has received a valid message m¹ _· ¹ = (B/ esig_j {H{B^r _· )),

with v = H(B‘ _· ) , then, i stops his own execution of round r right away; sets B^r = B ; and sets his own CERT^r to be the set of messages mf^{r s 1} of sub-step (b).

- If, during such waiting and at any point of time, there exists a step s' such that

(a’) 6 < s' < m + 3 with s'— 2 º 1 mod 3, and

(b’) i has received at least t valid messages mf^{r s 1}

then, i stops his own execution of round r right away; sets B^r = Bz ; and sets his own CERT^r to be the set of messages m^!- ^{s 1} of sub-step (b’) .

- If, during such waiting and at any point of time, i has received at least tn valid messages m^r,^m+³ _ ^ stops his own execution of round r

right away, sets B^r = B , and sets his own CERT^r to be the set of messages mb^m+3 for 1 and H(B ) .

6 Algorand'₂

In this section, we construct a version of Algorand' working under the following assumption.

HONEST MAJORITY OF USERS ASSUMPTION: More than 2/3 of the users in each PK^r are honest.

In Section 7, we show how to replace the above assumption with the desired Honest Majority of Money assumption.

²¹ A certificate from Step m + 3 does not have to include ESIGi(outi). We include it for uniformity only: the certificates now have a uniform format no matter in which step they are generated. 6.1 Additional Notations and Parameters for Algorand'₂

Notations

• m £ Z⁺: a pragmatic upper-bound to the number of steps that, with overwhelming probability, will actually taken in one round. (As we shall see, parameter m controls how many ephemeral keys a user prepares in advance for each round.)

• 17 \ a random variable representing the number of Bernoulli trials needed to see a 1 , when each trial is 1 with probability

L^r will be used to upper-bound the time needed to generate block B^r .

• i_ff : a lower-bound for the number of honest verifiers in a step s > 1 of round r, such that with overwhelming probability (given n and p ) , there are > i_ff honest verifiers in SV^r,s.

Parameters

• Relationships among various parameters.

For each step s > 1 of round r, n is chosen so that, with overwhelming probability,

\HSV^r'^s\ > ί_H and \HSV^r'^s\ + 2\MSV^r'^s\ < 2 t_H.

Note that the two inequalities above together imply \HSV^r,s\ > 2\M SV^r,s\ : that is, there is a 2/3 honest majority among selected verifiers.

The closer to 1 the value of h is, the smaller n needs to be. In particular, we use (variants of) Chernoff bounds to ensure the desired conditions hold with overwhelming probability.

• Specific choices of important parameters.

F = 11G¹⁸.

4000, i_ff ss 0.69n, k = 70.

6.2 Implementing Ephemeral Keys in Algorand'₂

Recall that a verifier i 6 SV^r,s digitally signs his message m^r7 of step s in round r, relative to an ephemeral public key p q’^s, using an ephemeral secrete key sJtf^{r s} that he promptly destroys after using. When the number of possible steps that a round may take is capped by a given integer m, we have already seen how to practically handle ephemeral keys. For example, as we have explained in Algorand (where m = m + 3), to handle all his possible ephemeral keys, from a round r' to a round r' + 10⁶, i generates a pair ( PMK , SMK ) , where PMK public master key of an identity based signature scheme, and SMK its corresponding secret master key. User i publicizes PMK and uses SMK to generate the secret key of each possible ephemeral public key (and destroys SMK after having done so) . The set of Fs ephemeral public keys for the relevant rounds is S = {/} x fir' , . . . , r' + 10⁶} x {1, . . . , m}. (As discussed, as the round r' + 10⁶ approaches, i“refreshes” his pair {PMK, SMK) .)

In practice, if m is large enough, a round of Algorand'₂ will not take more than m steps. In principle, however, there is the remote possibility that, for some round r the number of steps actually taken will exceed m. When this happens, i would be unable to sign his message m^r7 for any step s > m, because he has prepared in advance only m secret keys for round r. Moreover, he could not prepare and publicize a new stash of ephemeral keys, as discussed before. In fact, to do so, he would need to insert a new public master key PMK ' in a new block. But, should round r take more and more steps, no new blocks would be generated.

However, solutions exist. For instance, i may use the last ephemeral key of round r, pkfi¹¹ , as follows. He generates another stash of key-pairs for round r e.g., by (1) generating another master key pair {PMK , SMK) ·, (2) using this pair to generate another, say, 10⁶ ephemeral O

keys sr . ,m+I ⁶

, sk corresponding to steps m + 1, ..., m + 10⁶ of round r; (3) using sk⁷^ to digitally sign PMK (and any (r, m) -message if i £ SV ^7,m) , relative to and (4) erasing

SMK and sk ⁷'^m . Should i become a verifier in a step m + s with s 6 {1, . . . , 10⁶}, then i digitally signs his (r, m + s)-message pi^7,m+:3 relative to his new key r^^,m+3 = (i, r, m + s). Of course, to verify this signature of i, others need to be certain that this public key corresponds to Z’s new public master key PMK. Thus, in addition to this signature, i transmits his digital signature of PMK relative to pk⁷M .

Of course, this approach can be repeated, as many times as necessary, should round r continue for more and more steps! The last ephemeral secret key is used to authenticate a new master public key, and thus another stash of ephemeral keys for round r. And so on.

6.3 The Actual Protocol Algorand'₂

Recall again that, in each step s of a round r, a verifier i 6 SV^7,S uses his long-term public- secret key pair to produce his credential, a^7,s = SIGi (r, S Q^{7 1} ) , as well as SIGi (Q^{7 1}) in case s = 1. Verifier i uses his ephemeral key pair, (pk^7,s , sk^7,s) , to sign any other message m that may be required. For simplicity, we write esigi (m) , rather than sig_pkr^ (m) , to denote Fs proper ephemeral signature of m in this step, and write ESIGi{m) instead of SIG_pkr,s (m ) = ( i, m, esigi (m )).

Step 1: Block Proposal

Instructions for every user i 6 PK^{r k}: User i starts his own Step 1 of round r as soon as he has CERT^{7 1} , which allows i to unambiguously compute H^B^{7- 1}) and Q ^r_1.

• User i uses Q^{r 1} to check whether i 6 SV ^7,1 or not. If i ^ SV⁷’¹ , he does nothing for Step

1.

• If i 6 AV^7"’¹ , that is, if i is a potential leader, then he does the following.

(a) If i has seen B°, . . . , B^{7 1} himself (any B^ = B{ can be easily derived from its hash value in CERT ¹ and is thus assumed “seen”) , then he collects the round-r payments that have been propagated to him so far and computes a maximal payset PAY from them.

(b) If i hasn’t seen all B°, . . . , B ^{r 1} yet, then he sets PAY J = 0.

(c) Next, i computes his“candidate block” B ⁷ = (r, PAY^, SIGi(Q^{7 1}) , H(B^{7 1})) .

(c) Finally, i computes the message

, destroys his ephemeral secret key sk⁷’¹, and then propagates two messages,

) , separately but simultaneously.²²

Selective Propagation

To shorten the global execution of Step 1 and the whole round, it is important that the (r, 1)- messages are selectively propagated. That is, for every user j in the system,

²²When i is the leader, SIG (Q^{r 2}) allows others to compute Q^r H(SIG (Q^{r 1} ) , r) . • For the first (r, l)-message that he ever receives and successfully verifies,²³ whether it contains a block or is just a credential and a signature of Q^{r 1} , player j propagates it as usual.

• For all the other (r, l)-messages that player j receives and successfully verifies, he propagates it only if the hash value of the credential it contains is the smallest among the hash values of the credentials contained in all (r, 1) -messages he has received and successfully verified so far.

• However, if j receives two different messages of the form m ^7,1 from the same player i,²⁴ he discards the second one no matter what the hash value of Fs credential is.

Note that, under selective propagation it is useful that each potential leader i propagates his credential s⁷’¹ separately from m⁷’¹ :²⁵ those small messages travel faster than blocks, ensure timely propagation of the m⁷’¹’ s where the contained credentials have small hash values, while make those with large hash values disappear quickly.

Step 2: The First Step of the Graded Consensus Protocol GC

Instructions for every user i £ PK^{r k}: User i starts his own Step 2 of round r as soon as he has CERT^{1" 1}.

• User i waits a maximum amount of time t_2— A + L. While waiting, i acts as follows.

1. After waiting for time 2L, he finds the user i such that H(s ^¹) < H^s⁷'¹) for all credentials a⁷ _·'¹ that are part of the successfully verified (r, l)-messages he has received so far.²⁶

2. If he has received a block B^{1" 1} , which matches the hash value H^B^{1 1}) con tained in CERT' ¹ ,²⁷ and if he has received from l a valid message

=

{B_i ^r, esigi{H{B_i ^r)) , a_i ^r,1) ^2?’ then i stops waiting and sets

, £) .

3. Otherwise, when time t₂ runs out, i sets v = _L.

4. When the value of v has been set, i computes Q ^r_1 from CERT ^{r_ 1} and checks whether i £ SV^r’² or not.

5. If i £ SV^r'² , i computes the message m ^7,2 = (ESIGi^) , s⁷’²) ,²⁹ destroys his ephemeral secret key sk ^7,2 , and then propagates m⁷’². Otherwise, i stops without propagating anything.

²³That is, all the signatures are correct and, if it is of the form ml’¹, both the block and its hash are valid — although j does not check whether the included payset is maximal for i or not.

²⁴Which means i is malicious.

²⁵We thank Georgios Vlachos for suggesting this.

²⁶Essentially, user i privately decides that the leader of round r is user £.

²⁷Of course, if CERT ^{r 1} indicates that B”^{- 1} = Bz ¹, then i has already“received” B”^{- 1} the moment he has CERT^{r 1}.

²⁸Again, player Z’s signatures and the hashes are all successfully verified, and PAY in B_t ^r is a valid payset for round r— although i does not check whether PAY[ is maximal for l or not. If B J contains an empty payset, then there is actually no need for i to see S^{r_ 1} before verifying whether SJ is valid or not.

²⁹The message m”’² signals that player i considers the first component of to be the hash of the next block, or considers the next block to be empty. Step 3: The Second Step of GC

Instructions for every user i G PK^{r k}: User i starts his own Step 3 of round r as soon as he has CERT^{1" 1}.

• User i waits a maximum amount of time

— t₂ + 2l = 3l + L. While waiting, i acts as follows.

r 2

1. If there exists a value v such that he has received at least tu valid messages m - of the form ( ESIG_j{v ), s^’²) , without any contradiction,³⁰ then he stops waiting and sets v' = v.

2. Otherwise, when time t₃ runs out, he sets v' = _L.

3. When the value of v' has been set, i computes Q ^{r_ 1} from CERT ^{r_ 1} and checks whether i G SV^{r ,3} or not.

4. If i G SV^r,s, then i computes the message ’³ = (ESIGi(v') ,

, destroys his ephemeral secret key sfc’³, and then propagates m’³. Otherwise, i stops without propagating anything.

Step 4: Output of GC and The First Step of BBA *

Instructions for every user i G PK^{r k}: User i starts his own Step 4 of round r as soon as he finishes his own Step 3.

• User i waits a maximum amount of time 2l.³¹ While waiting, i acts as follows.

1. He computes

and <¾, the output of GC, as follows.

(a) If there exists a value n' _L such that he has received at least tn valid messages

, then he stops waiting and sets

= v' and g _j = 2.

(b) If he has received at least t valid messages

= (ESIG_j (A-), s^’³), then he stops waiting and sets

= _L and <¾ = 0.³²

(c) Otherwise, when time 2l runs out, if there exists a value v¹ _L such that he has received at least \^ \ valid messages m^!A = ( ESIG_j{v') , a -’³) , then he sets Vi = v' and gi = l.³³

(d) Else, when time 2l runs out, he sets

= _L and <¾ = 0.

2. When the values and gi have been set, i computes ¾, the input of BBA* , as follows: bi = 0 if gi = 2, and ¾ = 1 otherwise.

3. i computes Q ^{r_ 1} from CERT ^{r_ 1} and checks whether i G SV^{r, i} or not.

4. If i G SV^7"’⁴, he computes the message m_·’⁴ = (ESIGi(bi) , ESIGi(vi), s[’⁴), destroys his ephemeral secret key sfcb⁴, and propagates m_·’⁴. Otherwise, i stops without propagating anything.

³⁰That is, he has not received two valid messages containing ESIG_j (v) and a different ESIG_j iv) respectively, from a player j. Here and from here on, except in the Ending Conditions defined later, whenever an honest player wants messages of a given form, messages contradicting each other are never counted or considered valid.

³¹Thus, the maximum total amount of time since i starts his Step 1 of round r could be t = tz + 2l = 5l + L. 3²Whether Step (b) is in the protocol or not does not affect its correctness. However, the presence of Step (b) allows Step 4 to end in less than 2l time if sufficiently many Step-3 verifiers have“signed _L.”

³³It can be proved that the v' in this case, if exists, must be unique. Step s, 5 < s < m + 2, s— 2 º 0 mod 3: A Coin-Fixed- To-0 Step of BBA*

Instructions for every user i 6 PK^{r k}: User i starts his own Step s of round r as soon as he finishes his own Step s— 1.

• User i waits a maximum amount of time 2l.³⁴ While waiting, i acts as follows.

Ending Condition 0: If at any point there exists a string ¾ 1 and a step s' such that

(a) 5 < s' < s, s'— 2 º 0 mod 3 that is, Step s' is a Coin-Fixed- To-0 step,

(b) i has received at least t valid messages

(c) i has received a valid message

) with j being the second component of v,

then, i stops waiting and ends his own execution of Step s (and in fact of round r) right away without propagating anything as a (r, s)-verifier; sets H{B^r) to be the first component of v and sets his own CERT^r to be the set of messages m^!- ^{s 1} of step (b) together

Ending Condition 1: If at any point there exists a step s' such that

(a’) 6 < s' < s, s'— 2 º 1 mod 3 that is, Step s' is a Coin-Fixed- To- 1 step, and (b’) i has received at least t valid messages m^r- ^{s 1} = (ESIG_j (l) , ESIG_j(v_j) , r,s' l 37

j b

then, i stops waiting and ends his own execution of Step s (and in fact of round r) right away without propagating anything as a (r, s)-verifier; sets B^r = B_t ^r and sets his own CERT^r to be the set of messages rn - ^{s 1} of sub-step (b’) .

If at any point he has received at least tn valid m^!- ^{s 1}’s of the form

(ESIG_j(l) , ESIG_j(v_j) , a^rG ⁴) , then he stops waiting and sets ¾ = 1.

If at any point he has received at least tn valid m^!- ^{s 1}’s of the form

( ESIG_j{tS) , ESIG_j{v_j) , a^!C ⁴) , but they do not agree on the same v, then he stops waiting and sets ¾ = 0.

Otherwise, when time 2l runs out, i sets ¾ = 0.

When the value 6* has been set, i computes Q ^r_1 from CERT ^r_1 and checks whether i G SV^r'³.

If i E SV^r,s, i computes the message

being the value he has computed in Step 4, destroys his ephemeral secret key sk^ , and then propagates m'C . Otherwise, i stops without propagating anything.

³⁴Thus, the maximum total amount of time since i starts his Step 1 of round r could be t_s = t_s_i + 2l = (2s— 3)l + L.

³⁵ Such a message from player j is counted even if player i has also received a message from j signing for 1. Similar things for Ending Condition 1. As shown in the analysis, this is to ensure that all honest users know CERT^r within time l from each other.

³⁶User i now knows H(B^r ) and his own round r finishes. He just needs to wait until the actually block B^r is propagated to him, which may take some additional time. He still helps propagating messages as a generic user, but does not initiate any propagation as a (r, sj-verifier. In particular, he has helped propagating all messages in his CERT^r , which is enough for our protocol. Note that he should also set bi = 0 for the binary BA protocol, but bi is not needed in this case anyway. Similar things for all future instructions.

³⁷In this case, it does not matter what the ¾’s are. Step s, 6 < s < m + 2, s— 2 º 1 mod 3: A Coin-Fixed- To-1 Step of BBA *

Instructions for every user i E PK^{r k}: User i starts his own Step s of round r as soon as he finishes his own Step s— 1.

• User i waits a maximum amount of time 2L. While waiting, i acts as follows.

Ending Condition 0: The same instructions as in a Coin-Fixed- To-0 step.

Ending Condition 1: The same instructions as in a Coin-Fixed- To-0 step.

If at any point he has received at least t valid m!- ^{s 1}’s of the form (ESIG_j( 0), ESIG_j (v_j), a^r-^{,s 1} then he stops waiting and sets ¾ = 0.³⁸

Otherwise, when time 2l runs out, i sets ¾ = 1.

When the value ¾ has been set, i computes Q ^{r_ 1} from CERT ^{r_ 1} and checks whether i G SV^r,s.

If i £ SV^r,s, i computes the message m’³ = ( ESIGiipi ) , ESIGi(vi) , s)^"’⁵) with being the value he has computed in Step 4, destroys his ephemeral secret key sk^ , and then propagates m!^^s . Otherwise, i stops without propagating anything.

Step s, 7 < s < m + 2, s— 2 º 2 mod 3: A Coin- Genuinely-Flipped Step of BBA *

Instructions for every user i E PK^{r k}: User i starts his own Step s of round r as soon as he finishes his own step s— 1.

• User i waits a maximum amount of time 2l. While waiting, i acts as follows.

Ending Condition 0: The same instructions as in a Coin-Fixed- To-0 step.

Ending Condition 1: The same instructions as in a Coin-Fixed- To-0 step.

If at any point he has received at least t_jj valid m^!- ^{s 1}’s of the form {ESIG_j{ 0), ESIG_j (v_j), a^r-^{,s 1} then he stops waiting and sets ¾ = 0.

If at any point he has received at least tn valid m^!- ^{s 1}’s of the form {ESIG_j{ 1), ESIG_j (v_j), a^r-^{,s 1} then he stops waiting and sets ¾ = 1.

Otherwise, when time 2l runs out, letting SV ^{,s 1} be the set of (r, s— l)-verifiers from whom he has received a valid message rn - ^{s 1} , i sets b-i = lsb

.

When the value

has been set, i computes Q ^{r_ 1} from CERT ^{r_ 1} and checks whether i E SV^r’^s.

If i E SV^r,s, i computes the message m’^s = ( ESIGi{bi ) , ESIGi(vi) , o^^,s) with v being the value he has computed in Step 4, destroys his ephemeral secret key sk^ , and then propagates m -’^s. Otherwise, i stops without propagating anything.

Remark. In principle, as considered in subsection 6.2, the protocol may take arbitrarily many steps in some round. Should this happens, as discussed, a user i E SV^r,s with s > m has exhausted his stash of pre-generated ephemeral keys and has to authenticate his (r, s)-message nJP by a“cascade” of ephemeral keys. Thus Fs message becomes a bit longer and transmitting these longer messages will take a bit more time. Accordingly, after so many steps of a given round, the value of the parameter l will automatically increase slightly. (But it reverts to the original l once a new block is produced and a new round starts.)

³⁸Note that receiving tu valid (r, s l)-messages signing for 1 would mean Ending Condition 1. Reconstruction of the Round-r Block by Non- Verifiers

Instructions for every user i in the system: User i starts his own round r as soon as he has CERT^{1" 1}.

• i follows the instructions of each step of the protocol, participates the propagation of all messages, but does not initiate any propagation in a step if he is not a verifier in it.

• i ends his own round r by entering either Ending Condition 0 or Ending Condition 1 in some step, with the corresponding CERT^r .

• From there on, he starts his round r + 1 while waiting to receive the actual block B^r (unless he has already received it) , whose hash H{B^r) has been pinned down by CERT^r . Again, if CERT^r indicates that B^r = Bz, the i knows B^r the moment he has CERT^r .

7 Protocol Algorand' with Honest Majority of Money

We now, finally, show how to replace the Honest Majority of Users assumption with the much more meaningful Honest Majority of Money assumption. The basic idea is (in a proof-of-stake flavor) “to select a user i 6 PK^{r k} to belong to SV^r,s with a weight (i.e., decision power) proportional to the amount of money owned by h”³⁹

By our HMM assumption, we can choose whether that amount should be owned at round r— A: or at (the start of) round r. Assuming that we do not mind continual participation, we opt for the latter choice. (To remove continual participation, we would have opted for the former choice. Better said, for the amount of money owned at round r— k— 2, 000.)

There are many ways to implement this idea. The simplest way would be to have each key hold at most 1 unit of money and then select at random n users i from PK^{r k} such that

(r) ₁

ai = ! ·

The Next Simplest Implementation

The next simplest implementation may be to demand that each public key owns a maximum amount of money M, for some fixed M. The value M is small enough compared with the total amount of money in the system, such that the probability a key belongs to the verifier set of more than one step in say k rounds is negligible. Then, a key i 6 PK^{r k}, owning an

(r)

amount of money a) ¹ in round r, is chosen to belong to SV^r,s if

(rj

.H (SIG_i (r, s, Q^{r 1}) ) £ p · ^ .

And all proceeds as before.

A More Complex Implementation

The last implementation“forced a rich participant in the system to own many keys” .

An alternative implementation, described below, generalizes the notion of status and consider each user i to consist of K + 1 copies (i, v) , each of which is independently selected to be a verifier, and will own his own ephemeral key (pk^rC , sk^rT) in a step s of a round r. The

(r)

value K depends on the amount of money a) ¹ owned by i in round r.

Let us now see how such a system works in greater detail.

³⁹We should say pK^r-^k-^{2 000} _{so as} to replace continual participation. For simplicity, since one may wish to require continual participation anyway, we use PK^{r k} as before, so as to carry one less parameter. Number of Copies Let n be the targeted expected cardinality of each verifier set, and let

(r)

a be the amount of money owned by a user i at round r. Let A^r be the total amount of money owned by the users in PK^{r k} at round r, that is,

If i is an user in PK^{r k} , then V s copies are (i, 1) , . . . , (/, K + 1), where

EXAMPLE. Let n = 1, 000, A^r = 10⁹,

= 3.7 millions. Then,

Verifiers and Credentials Let i be a user in PK^{r k} with K + 1 copies.

For each v = 1 , . . . , K, copy ( i, v ) belongs to SV^r,s automatically. That is, Fs credential is ^s _ί '_{n —} SIGi{{i, n) , r, s, Q^{r _ 1}) , but the corresponding condition becomes

< 1, which is always true.

For copy (i, K + 1) , for each Step s of round r, i checks whether

If so, copy {i, K + 1) belongs to SV^r,s. To prove it, i propagates the credential

(r)

EXAMPLE. AS in the previous example, let n = IK, a = 3.7 M, A^r = IB, and i has 4 copies: (i, 1) , . . . , (i, 4) . Then, the first 3 copies belong to SV^r,s automatically. For the 4th one, conceptually, Algorand ' independently rolls a biased coin, whose probability of Heads is 0.7. Copy (7, 4) is selected if and only if the coin toss is Heads.

(Of course, this biased coin flip is implemented by hashing, signing, and comparing as we have done all along in this application so as to enable i to prove his result.)

Business as Usual Having explained how verifiers are selected and how their credentials are computed at each step of a round r, the execution of a round is similar to that already explained.

8 Algorand with the Assumption of No Network Partition

Essentially, in Algorand, blocks are gnerated in rounds. In a round r,

(1) A properly credentialed leader proposes a new block and then (2) Properly credentialed users run, over several steps, a proper Byzantine agreement (BA) protocol on the block proposed.

The preferred BA protocol is BA* . The block proposal step can be considered step 1, so that the steps of BA* are 2, 3,...

Only a proper user i, randomly selected among the users in the system, is entitled to send a message rrfiA in step s of round r. Algorand is very fast and secure because such a user i checks whether he is entitled to speak. If this is the case, user i actually obtains a proof, a credential. If it is his turn to speak in step s of round r, i propagates in the network both his credential, s’⁵, and his digitally signed message

The credential proves to other users that they should take in consideration the message

A necessary condition for user i to be entitled to speak in step s of round r is that he was already in the system a few rounds ago. Specifically, k rounds before round r, where A: is a parameter termed the‘look-back’ parameter. That is, to be eligible to speak in round r, i must belong to PK^{r k}, the set of all public keys/users already in the system at round r— k. (Users can be identified with their public keys.) This condition is easy to verify in the sense that it is derivable from the blockchain.

The other condition is that

tf(SJG_i(r, s, Q^{r 1})) < p

where p is a given probability that controls the expected number of verifiers in SV^r,s, that is, the set of users entitled to speak in step s of round r. If this condition is satisfied, then i’s credential is defined to be

Of course, only i can figure out whether he belongs to SV^r,s. All other users, who lack knowledge of i’s secret signing key, have no idea about it. However, if i £ SV^r,s, then i can demonstrate that this is the case to anyone by propagating his credential given the blockchain so far. Recall in fact that (1) Q ^{r_ 1} is easily computable from the previous block, B^r_1 , although essentially unpredictable sufficiently many blocks before, and (2) anyone can verify i’s digital signatures (relative to his long-term key in the system) .

Recall that, in the versions of Algorand so far, new blocks are proposed only once in a round r, that is, in step 1. The BA protocol has the users reach consensus on one of them (or the empty block) , and does not further propose new blocks or re- propose blocks that have already been proposed for round r. When the network is not partitioned and the upper-bounds for the time to propagate messages are met, the users reach consensus efficiently and securely.

9 Algorand Resilient Against Network Partition

Let us describe a new embodiment of Algorand, Algorand2, that dispenses the assumption of no network partition. We present the new protocol under the Honest Majority of Users (HMU) assumption. Using the same approach as in Section 7, the HMU assumption can be replaced with the Honest Majority of Money (HMM) assumption.

9.1 Communication Model

When the network is not partitioned, small messages take time l to propagate to all honest users, and blocks take time A to propagate to all honest users, as before. When the network is partitioned into more than one groups of users, the adversary determines whether a message m propagated by a user from one group will be delivered to users in other groups, who in other groups will receive m, and when they will receive m. A network partition may be resolved at an indefinite time in the future and messages propagated during the partition are delivered to all users after the partition is resolved. For simplicity, but without limitation intended, we describe the new embodiment assuming messages propagated during a network partition are delivered to all users immediately after the partition is resolved. For example, if a network partition lasts from time t to time t₂, and let M be the set of messages propagated during the partition, then all users receive messages in M by time t₂. Those skilled in the art will realize that the system described herein can handle other situations where messages in M take certain amount of time to reach all users or are re-propagated by users who have received them.

9.2 Choices of Parameters

The expected committee size n and the threshold i_ff are chosen according to the following conditions. Let PK be the set of users, HPK and MPK respectively the set of honest and malicious users. Let HPK be an arbitrary subset of HPK with half the size. When each user i 6 PK is selected independently and randomly with probability p = let HSV i and

MSV respectively be the set of selected ones from HPK and MPK. Then with overwhelming probability,

\HSV_! \ + \MSV\ < i_ff.

Moreover, let HSV be the set of selected ones from HPK. Then with overwhelming probability,

\HSV\ > t_H.

Note that the above two conditions imply \HSV\ > 2\MSV\ .

For example, when h = 80% and PK is large enough, we may choose n = 3, 500 and t_H = 2, 625.

9.3 General Structure

Rounds. The protocol generates one block every round. A round consists of periods 1, 2, ... and a period consists of steps 1, 2, .... At any moment in time, each user i is working on exactly one round-period pair. In particular, we use r.p to refer to period p of round r.

In step 1 of period 1, users propose new blocks. In step 1 of following periods, users propose new blocks or re-propose blocks that have been proposed in earlier periods.

Committees. Each step s of period r.p has a committee chosen by cryptographic self selection, denoted by SV^r,p,s. We use the same look-back parameter as in Section 4.1, denoted by k. For example, k = 70. A user i is eligible to be selected in round-r committees if i £ PK^{r k}. The committee for Step 1 of each period has expected size n (e.g., 35) and all other committees have expected size n. Committee members for Step 1 are referred to as potential leaders.

Note that for simplicity, but without limitation intended, we describe the new embodiment herein with the same expected committee size n for all steps other than step 1 of each period. Those skilled in the art will realize that different committees may have different sizes and can appreciate how to derive all kinds of other implementations as well.

Keys. All credentials for cryptographic self-selection are signed with users’ long-term keys for a digital signature scheme with unique signatures , so are the random seeds Q^r specified in the blocks. All other messages are signed using ephemeral keys of corresponding steps. In general we will use SIGi{m) to denote user i’s signature for message m, without specifying the keys. Note that for simplicity, but without limitation intended, we describe the new embodiment herein with ephemeral keys. Those skilled in the art can appreciate how to derive other implementations with message-credentialed blockchains, including using techniques introduced in existing versions of Algorand.

Definition 9.1. Credential: User i’s credential a^rf^p,s for a round r, period p and step s is SIGi(Q^{r 1}, r , p, s).

A committee member for a step always propagates his corresponding credential together with his message for that step, and we will not explicitly mention the propagation of credentials.

Definition 9.2. Leader: The leader £_{r p} for period r.p is

.

When a user i identifies his own leader for period r.p, /y_r.p, i sets to be the user

Definition 9.3. Valid block: We call a block proposed during round r valid if and only if all its transactions are valid with respect to blocks B°, B¹, · · · , B ^{r_ 1} and the seed Q^r specified by it follows the rules of the protocol.

Voting Messages. The committee members generate three types of voting messages.

Definition 9.4. Cert- vote: User i’s cert- vote for a value v for period r.p is the signature SIGi(v,“cert” , r.p) .

We say a user i cert-votes a value v for period r.p when he propagates SIGfiv ,“cert” , r.p).

Definition 9.5. Soft-vote: User i’s soft- vote for a value v for period r.p is the signature SIGfiv,“soft”, r.p).

We say a user i soft- votes a value v for period p when he propagates SIGfiv,“soft” , r.p) .

Definition 9.6. Next-vote: User i’s next- vote for a value v for period r.p and step s is the signature SIGfiv,“next” , r.p. s) .

We say a user i next-votes a value v when he propagates SIGfiv ,“next” , r.p. s) .⁴⁰

The values v that will be voted upon are either values in the range of the hash function H or a special symbol _L of the same length but outside the range of H .⁴¹

Timers. Each user i keeps a timer clocki, which he resets to 0 every time he starts a new period. As long as i remains in the same period, clocki keeps counting. The users’ individual timers do not need to be synchronized or almost synchronized. The only requirement is that they have the same speed.

⁴⁰In each period, the soft-votes are generated only in one step, and so are the cert-votes. Thus they do not need to specify the corresponding step. The next-votes may be generated in multiple steps and need to specify the step numbers.

⁴¹The block structure in this manuscript is the same as in existing versions of Algorand, with the header of a block B, Header(B), containing everything in B except the actual payset. As in the original protocol, we can ask the potential leaders to propagate the headers of their proposed blocks together with the hashes, to allow users to start a round r + 1 before seeing the actual block B^r. In this case, u and T are of the same length as possible hash-header pairs. We will ignore the headers in the protocol description below. 9.4 The Actual Protocol Algorand2

In the protocol below, blocks and potential leaders’ credentials from different periods are selectively propagated as in Section 5.

Round r, Period 1

The following are period 1 instructions for a generic user i. If user i is not in the committee of a specific step, he still computes his vote in that step, but does not propagate it.

The moment user i starts his own round-r, he starts Period 1 and resets clock_s to 0.

STEP 1 : [Block Proposal] User i does the following when clocks = 0.

• If i is a potential leader, then he prepares his proposed block

which contains his signature SIGi{Q^{r l} , r) to define his proposed seed zE.⁴² He propagates H{H^rk¹) , and right after also propagates the block itself.

STEP 2: [The Filtering Step] User i does the following when clock_s = 2l.

• He identifies his leader ki ._{r p}. and soft-votes his leader’s proposed block hash.⁴³

STEP 3: [The Certifying Step] Let T = L + l. User i does the following when clock_s £ (2 X, T) .

• If i sees a valid block B, together with tn soft-votes for H(B), then i cert- votes H{B) . STEP 4: [The First Finishing Step] User i does the following when clock_s = T.

• If i has cert-voted some value v in Step 3, ⁴⁴ then he next- votes v

• Else he next- votes T.

STEP 5: [The Second Finishing Step] User i does the following when clock_s £ [T, T + L), where L equals, say, 50L.⁴⁵

• If i sees a valid block B, together with tn soft-votes for H(B) , then i next- votes H{B) . STEPS S > 6: [Consecutive Finishing Steps]

• If s is even then i follows the same voting instruction as in Step 4 when clocL· = T + (s— 4)L/2.

• Else (i.e., s is odd) , i follows the same voting instruction as in Step 5 when clock_s £ [T + (s - 5)L/2, T + (s - 3)L/2).

⁴² That is, Ql = H(SIGi(Q^{r 1}, r)).

⁴³It’s not required that i has seen the actual block before soft-voting.

⁴⁴If i is not in SV^r’^1,3, he may still have cert-voted some v“for himself”, without propagating his vote.

⁴⁵ If the Adversary cannot corrupt users dynamically, then it is enough to have Step 5 lasting till time oo and no further steps, as in the permissioned version of our protocol. Future steps are used to ensure that the protocol recovers after a partition is resolved, even if the Adversary has pocketed all Step 4 and Step 5 messages and corrupted all Step 4 and Step 5 committee members so that they cannot resend their messages after the partition. A smaller L helps speeding up partition resolution in the worst case, while a larger L reduces the amount of messages that need to be re-propagated once the partition is resolved. Round r, Period p > 2

The following are period-p instructions for a generic user i. Again, if user i is not in the committee of a specific step, he still computes his vote in the step but does not propagate it.

User i starts period p the moment he receives tn next-votes for some value v (which might be equal to _L) for the same step s of period p— 1, and only if he has not yet started a period p ' > p. User i sets his starting value for period p, st^p, to v. The moment i starts period p, he finishes all previous periods and resets clocki to 0.

STEP 1 : [Block Proposal] If user i is a potential leader, then he does the following when clocki =

0.

• If i has seen t}j next-votes for _L for the same step s of period p— l,⁴⁶ then i proposes a new block , which defines his proposed seed Q as H(Q^{r 1} , r) .⁴⁷ He propagates H(B^’^P) , and right after also propagates the block itself;

• Else (i.e., i only received tn next-votes for some v 1 for the same period s of period p— 1 , and st = v), i proposes st^p by propagating it;⁴⁸

STEP 2: [The Filtering Step] User i does the following when clocki = 2L.

• If i has seen t}j next-votes for _L for the same step s of period p— l,⁴⁹ then i identifies his leader /y_p for period p and soft-votes the value v proposed by /y_p;

• Else (i.e., i only received tn next-votes for v _L for the same step s of period p— 1, and st^p = v), i soft- votes st^p.

STEP 3: [The Certifying Step] User i does the following when clocki £ (2L, T) , T = L + A.

• If i sees a valid block B, together with tn soft-votes for H(B), then i cert- votes H{B) . STEP 4: [The First Finishing Step] User i does the following when clocki = T.

• If i has cert- voted some value v in Step 3, then he next- votes v

• Else if i has seen tn next- votes for T for the same step s of period p— 1, he next-votes T.

• Else he next- votes st^p.

STEP 5.1 : [The Second Finishing Step] User i does the following when clocki £ [T, T + L) .

• If i sees a valid block B, together with tn soft-votes for H(B) , then i next- votes H{B) . STEP 5.2: [The Second Finishing Step] User i does the following when clocki £ [T, T + L) .⁵⁰

• If i has not cert-voted in Step 3 and he sees tn next- votes for T for the same step s of period p— 1, then i next- votes T.

⁴⁶In extreme cases, a user i may simultaneously see tn next-votes for _L and tn next-votes for some u F _L for period p 1. He is free to set his starting value to either _L or v in this case, but the Block Proposal step always gives priority to _L.

⁴⁷That is, a newly proposed block for a period p > 2 is as if there were no leader for it, and the corresponding seed Q^r is defined in the same way as in the original Algorand protocol when B^r is the empty block.

⁴⁸In this case, enough honest users have seen a valid block B such that v = H(B), and it is not required that i herself has seen B.

⁴⁹Note that these votes may reach i during his time (0, 2l] , thus i’s starting value is not necessarily _L.

⁵⁰Steps 5.1 and 5.2 are made into different steps so that they have different committees. STEPS S > 6: [Consecutive Finishing Steps]

• If s is even then i follows the same voting instruction as in Step 4 when clocki = T + (s— 4)L/2.

• Else (i.e., s is odd), i follows the same voting instructions as in Steps 5.1 and 5.2 in parallel when clocki £ [T + (s— 5)L/2, T + (s— 3)L/2) .

Changing Rounds

Instructions for every user i 6 PK:

• At any point while i is working on a round r, if i sees a string v 1 and a period r'.p with r' > r such that

i has received at least tu cert- votes for v for period r'.p,⁵¹ and

i has received the block header corresponding to v,⁵²

then i cert-votes v for period r' .p,⁵³ sets his own CERT^r ' to be the set of cert-votes for v and starts round r' + 1.

Referring to FIG. 6, a diagram 20’ shows a first plurality of computing workstations 22a- 22c connected to a first portion 24’ of a data network, such as the Internet, and a second plurality of computing workstations 22a’-22c’ connected to a second portion 24” of the data network. The workstations 22a-22c communicate with each other via the first portion 24’ of the network and the workstations 22a’-22c’ communicate with each other via the second portion 24” of the network. However, unlike the diagram 20 of FIG. 1, the workstations 22a-22c do not communicate with the workstations 22a’-22c’ because the network has been partitioned into the first part 24’ and the second part 24” by a partition mechanism 26. The partition mechanism 26 could be any mechanism that suppresses communication between the parts 24’, 24” of the network. Note that the partition mechanism 26 could be purposefully inserted into the network by an adversary or could occur due to unanticipated network interruption caused by natural or man-made phenomena, such as a power outage or a network switching error. In the system described herein, each of the workstations 22a-22c, 22a’-22c’ performs the steps set forth above, but the first plurality of workstations 22a-22c does not communicate with the second plurality of workstations 22a’-22c’. That is, each of the entities 22a-22c manages a transaction system locally within the part 24’ while each of the entities 22a’-22c’ manages a transaction system locally within the part 24”. Only one of the parts 24’, 24”, at most, has enough users (entities) to generate certified blocks. After the partition mechanism 26 is removed and the parts 24’, 24” converge to a single network (like the network 24, discussed above in connection with FIG. 1) , all of the workstations 22a-22c, 22a’-22c’ communicate with each other and are able to resume operation at the same point (block) of the chain. After the partition mechanism 26 is removed, each of the entities 22a-22c, 22a’-22c’ manages a single transaction system in the network where all of the entities 22a-22c, 22a’-22c’ communicate with each other.

⁵¹ A cert- vote from a player j is counted even if player i has also received a message from j cert- voting for a different value for the same period. As shown in the analysis, this is to ensure that all honest users know CERT within time l from each other.

⁵²The header defines the seed Q^r and, as mentioned earlier, strictly speaking the header is part of v itself.

⁵³For the theoretical analysis it is not necessary for i to propagate his vote, because he already helped propagating the tn cert-votes he received. However, having i propagate his vote when i G SV^r’^p’ⁱ helps other users see a certificate faster in reality. Analysis. Even when the network is partitioned, the new embodiment remains secure that is, at most one block is certified for each round r. At a high level, this is because honest users cert-vote at most once in each period of round r, and the choices of the committee size n and the threshold tn guarantee that in each period r.p, at most one hash value H{B ) of a valid block B can get a certificate. To see it from a different direction, assume a period r.p has generated t cert-votes for one valid block B and tn cert-votes for another valid block B' relative to B°, . . . , B^{1 1}, then the conditions for how n and tu are chosen imply that at least one honest user has cert- voted for the hash values of both B and B' , which contradicts the fact that an honest user cert- votes at most once in period r.p.

If in some period r.p, a valid block B has gotten a certificate that is, at least tn cert-votes for its hash value H{B) , then in all future periods p' > p (if they are ever reached) , B will be the only block that may get a certificate in period r.p' . Indeed, an honest user does not next-vote for _L in period r.p if he has cert-voted for H{B) in r.p. Thus by the same choices of n and tjj, in no step s > 4 of period r.p will there be tjj next-votes for _L or for any other value v H{B). So an honest user moves to period r.(p + 1) only if he has received at least tjj next-votes for H{B) . Accordingly, H{B) will be the only value that is (re-)proposed in step 1 of period r.(p + 1), the only value that honest users will soft- vote for in step 2 of period r.(p + 1) , and consequently the only value they will cert-vote in step 3 and next-vote in step s > 4 of period r.(p + 1) . By an inductive argument, the same is true for all consecutive periods.

Notice that, when the network is partitioned, B having gotten a certificate in a period r.p does not mean that the honest users will receive the certificate. Indeed, during a network partition an adversary controls how messages are delivered in the system. He may, for example, allow all messages to be delivered properly except the cert- votes, where he does not allow cert- votes from one group of users to be delivered to other groups. Nevertheless, B having gotten a certificate implies that enough honest users have cert-voted for it and will not next-vote for anything else, which prevents any other block from being certified in period r.p and any future period.

The efficiency of the new embodiment comes from two parts. First, when the network is not partitioned, consensus about the rth block is reached quickly. Indeed, if the leader of step 1 of period r.l is honest, then all honest users immediately cert-vote for his proposed block B, B has gotten a certificate after step 3 of period r.p, and all honest users finish round r afterward.

Similarly, if round r has reached a period p > 2 and the leader ί for period p is honest, then the block newly proposed or re-proposed by ί is certified in step 3 and all honest users finish round r afterward. This is so because, if ί has seen tn next- votes for _L from period p— 1 and proposed a new block, then those next-votes will reach all honest users within time l and they will all soft-vote for Fs proposal. Otherwise, ί has only seen tn next- votes for the hash H{B) of a valid block B from period p— 1 and has re-proposed H{B) . All honest users’ starting values of period p are either _L or H{B) , and no matter which is the case, they soft- vote for H{B ) because they are following the leader’s proposal in the former case, and because they are voting for their own starting values in the latter case.

Moreover, if a certificate for some valid block B is generated in a period r.p, then all honest users finish round r soon after that. This is so because, if at least tn cert-votes for B come from honest users, then all honest users will receive them within time l and will finish round r with B being the rth block. If any set of tn cert-votes for B contains at least one malicious user, the malicious users may choose to not send their cert-votes and the honest users do not receive a certificate for B right away. However, the choices of n and tn ensure that a certificate for B contain at least one honest user i, and i has received tn soft-votes for B before he cert-voted it. Since i propagated these soft- votes and the network is not partitioned, all honest users will receive them within time l and will next-vote for H{B) in step 5.1. Accordingly, all honest users start period r.(p + 1) with starting value H{B ) and will soft-vote for it in step 2 of period r.(p + 1) , regardless whether the leader of period r.(p + 1) is honest or not. As a result, honest users will cert- vote for H{B ) in step 3, B now has a certificate from honest users, and all honest users will receive them and finish round r within time l.

The choices of the seeds Q^r and the cryptographic sortition used in the new embodiment ensure that each period p of round r has an honest leader with high probability, as in the original Algorand protocol, and the missing detailed analysis about the new embodiment’s efficiency when there is no network partition follow from there.

Second, after the network partition is resolved, the protocol will recover and reach consensus quickly. Indeed, if some honest users have received a certificate for a block B in round r during the partition and moved to round r + 1, then once the partition is resolved all honest users will receive such a certificate for B and move to round r + 1. Moreover, let p be the furthest period in round r + 1 where an honest user i has reached during the partition. Then all the next-votes that allowed i to move to period (r + l) .p will reach other honest users after the partition is resolved, and they will also move to period {r + l) .p. The protocol will then continue from there as usual, following the same analysis as when there is not network partition. If no honest user moved from round r to round r + 1 during the partition, then all honest users are in the same round but perhaps different periods. In this case, let p be the furthest period in round r where an honest user i has reached during the partition. Similarly, after the partition is resolved, all the next-votes that allowed i to move to period r.p will reach other honest users and they will also move to period r.p. Again the protocol will continue from there as usual.

To summarize, the new embodiment is secure and does not soft-fork even when the network is partitioned. It generates block efficiently when the network is not partitioned, and recovers quickly after a network partition is resolved.

10 Scope

Note that the mechanism described herein is applicable to other blockchain systems where it is desirable to prevent more than one blocks are certified during a network partition and to restore liveness quickly after a partition is resolved. Thus, the system described herein may be adapted to other blockchain schemes, even schemes that do not relate directly to currency.

The system described herein may be adapted to be applied to and combined with mechanisms set forth in any or all of PCT/US2017/031037, filed on May 4, 2017, 15/551,678 filed August 17, 2017, 16/096,107 filed on October 24, 2018, PCT/US2018/053360 filed on September 28, 2018, PCT/US2018/054311 filed on October 4, 2018, 62/632,944 filed on February 20, 2018, 62/643,331 filed on March 15, 2018, 62/777,410 filed on December 10, 2018, and 62/778,482 filed on December 12, 2018, all of which are incorporated by reference herein.

Software implementations of the system described herein may include executable code that is stored in a computer readable medium and executed by one or more processors. The computer readable medium may be non-transitory and include a computer hard drive, ROM, RAM, flash memory, portable computer storage media such as a CD-ROM, a DVD-ROM, a flash drive, an SD card and/or other drive with, for example, a universal serial bus (USB) interface, and/or any other appropriate tangible or non-transitory computer readable medium or computer memory on which executable code may be stored and executed by a processor. The system described herein may be used in connection with any appropriate operating system.

Other embodiments of the invention will be apparent to those skilled in the art from a consideration of the specification or practice of the invention disclosed herein. It is intended that the specification and examples be considered as exemplary only, with the true scope and spirit of the invention being indicated by the following claims.

Claims

What is claimed is:

1. A method for an entity to manage a transaction system in which transactions are organized in a sequence of blocks that are certified by digital signatures of a sufficient number of verifiers, the method comprising:

the entity proposing a hash of a block B’ that includes new valid transactions relative to a sequence of certified blocks B°, . . . , B ^r_1 if no rth block B^r has been certified; and

the entity proposing a hash of the block B^r if the rth block B^r has been verified by a sufficient number of other entities.

2. A method, according to claim 1, wherein a block is certified by the entity only in response to confirming transactions for the block and confirming that the block was constructed and propagated by an entity entitled to construct and propagate the block.

3. A method, according to claim 1, wherein the entity proposes a hash value by digitally signing the hash value to provide a digitally-signed version of the hash value and wherein the entity propagates the digitally-signed version of the hash value to a network that includes other entities.

4. A method, according to claim 3, wherein if no rth block B^r has been certified, the entity also digitally signs and propagates the block B’.

5. A method, according to claim 1, wherein the entity determines a quantity Q from the prior blocks and uses a secret key in order to compute a string S uniquely associated with Q and computes from S a quantity T that is at least one of: S itself, a function of S, and hash value of S and wherein the entity determines whether to propose a hash value by determining whether T possesses a given property.

6. A method, according to claim 5, wherein S is a signature of Q under a secret key of the entity, T is a hash of S and T possesses the given property if T is less than a given threshold.

7. A method, according to claim 1, wherein the entity is part of a network of entities and wherein a particular one of the entities constructs and propagates the block B^r .

8. A method, according to claim 7, wherein the rth block B^r is determined to be certified by the entity if the entity receives an indication that at least a predetermined number of the entities individually certify a hash value corresponding to the rth block B^r .

9. A method, according to claim 8, wherein, in response to the entity receiving the indication that a predetermined number of the entities individually certified the rth block B^r , the entity increments r to begin adding additional blocks to the sequence of blocks.

10. A method, according to claim 7, wherein the particular one of the entities is individually chosen by a predetermined number of the entities to be a leader.

11. A method, according to claim 10, wherein the rth block B^r is determined to be certifiable by the entity if the entity receives an indication that at least a predetermined number of the entities individually verify receiving an indication that the particular one of the entities has provided a hash value corresponding to the rth block B^r to each of the predetermined number of the entities.

12. A method for an entity to manage a transaction system in which transactions are organized in a sequence of certified blocks, the method comprising:

the entity receiving a hash value of a block B^r from an other entity that generated the block based on new valid transactions relative to a sequence of certified blocks B°, . . . , B^{1 1}

the entity certifying the block B^r in response to a sufficient number of other entities having indicated receipt of the hash value of the block B^r from the other entity and the hash value being valid for the block B^r

the entity generating a new block B’ based on new valid transactions relative to a sequence of certified blocks B°, . . . , B ^{r_ 1} in response to an insufficient number of the other entities indicating receipt of the hash value of the block B^r from the other entity, wherein B’ is different from B^r and

the entity incrementing r to begin adding additional blocks to the sequence of blocks in response to the entity receiving the indication that a predetermined number of the entities individually certified the rth block B^r or a predetermined number of the entities individually certified the new block B’.

13. A method, according to claim 12, wherein the blocks are certified by digital signatures.

14. A method, according to claim 12, wherein new blocks are proposed by different ones of the entities until receiving the indication that a predetermined number of the entities individually certified a previously proposed block.

15. A method, according to claim 12, wherein the entity provides an indication that a new block should be generated in response to the hash value not being valid for the block B^r .

16. A method, according to claim 15, wherein the entity generates a new block B’ based on new valid transactions relative to a sequence of certified blocks B°, . . . , B ^r_1 in response to a sufficient number of the other entities providing an indication that a new block should be generated.

17. A method, according to claim 12, wherein the entity provides an indication that the hash value of the block B^r should be propagated in response to a sufficient number of the other entities having indicated receipt of the hash value of the block B^r from the other entity and the hash value being valid for the block B^r .

18. A method for an entity to verify a proposed hash value of a new block B^r of transactions relative to a given a sequence of blocks, B°, . . . , without access to the new block B^r in a transaction system in which transactions are organized in blocks and blocks are certified by a set of digital signatures, the method comprising:

having the entity determine a quantity Q from the prior blocks;

having the entity compute a digital signature S of Q;

having the entity compute from S a quantity T that is at least one of: S itself, a function of S, and hash value of S;

having the entity determine whether T possesses a given property; and

if T possesses the given property, having the entity verify the proposed hash value of the new block B^r independent of confirming whether the proposed hash value corresponds to the new block B^r .

19. A method, according to claim 18, wherein the entity propagates the proposed hash value of the new block B^r prior to receiving the new block B^r .

20. Computer software, provided in a non-transitory computer-readable medium, comprising: executable code that implements the method of one of the preceding claims 1-19.