The word STARK can mean one of several things.
This short review focuses on STARK $_2$.
At the heart of a STARK is the Algebraic Execution Trace (AET), which is a table of $\mathsf{w}$ columns and $T+1$ rows. Every column tracks the value of one register across time. Every row represents the state of the entire machine at one point in time. There are $T+1$ rows because the machine was initialized to the initial state before it was run for $T$ time steps.
The AET is integral if all Arithmetic Intermediate Representation (AIR) constraints are satisfied. These constraints come in several classes.
While the AIR constraints can be represented as multivariate polynomial equations, it helps to think of them as maps $\mathbb{F}^{\mathsf{w}} \rightarrow \mathbb{F}$ or $\mathbb{F}^{2\mathsf{w}} \rightarrow \mathbb{F}$ that must evaluate to zero on AETs that are integral and to something nonzero on AETs that are not. This perspective allows to generalize the argument from vectors of $\mathsf{w}$ or $2\mathsf{w}$ field elements to vectors of as many codewords or as many polynomials.
The prover runs a polynomial interpolation subprocedure to find, for every column, a low-degree polynomial that takes the value of the column $i$ rows down in the point $\omicron^i$, where $\omicron$ is a generator of a subgroup of order $T+1$. These polynomials are called the trace polynomials.
Evaluating the AIR constraints in the trace polynomials gives rise to boundary and transition polynomials. Moreover, every AIR constraint defines a support domain $D \subset \langle \omicron \rangle$ of points where it applies, and with it a zerofier $Z(X)$, which is the unique monic polynomial of degree $\vert D \vert$ that evaluates to zero on all of $D$ and no-where else. If the AET satisfies the AIR constraint then this zerofier divides the boundary or transition polynomial cleanly; if the AET does not satisfy the AIR constraint then this division has a nonzero remainder.
The prover continues with the quotient polynomials of the previous step. Specifically, he wishes to establish their bounded degree. If the AET satisfies the AIR constraints, then the quotient polynomials will be of low degree. If the AET does not satisfy the AIR constraints the malicious prover might be able to find impostor polynomials that agree with the division relation in enough points, but the point is that this impostor polynomial will necessarily have a high degree. And then the prover will fail to prove that its degree is low.
The prover commits to these polynomials as follows. First he evaluates them on a coset of the subgroup spanned by $\Omega$, whose order is $N > T+1$. This process of evaluation gives rise to Reed-Solomon codewords of length $N$. Next, a Merkle tree of this codeword is computed. The root is the commitment to the polynomial, and it is sent to the verifier.
One obvious optimization is available at this point. It is possible to zip the codewords before computing Merkle trees. In fact, after zipping, only one Merkle tree needs to be computed. The leafs of this Merkle tree correspond to tuples of field elements.
The next step is to combine the polynomials’ codewords into one codeword using random weights from the verifier. For every quotient polynomial $q_i(X)$, there is a degree bound $b_i$ originating from the known trace length $T$, the number of randomizers, and the AIR constraint degree. The prover combines the nonlinear combination \(\sum_{i=0} \alpha_i \cdot q_i(X) + \beta_i \cdot X^{\mathsf{d} - b_i} \cdot q_i(X) \enspace ,\) where the weights $\alpha_i$ and $\beta_i$ are provided by the verifier, and where $\mathsf{d}$ is the maximum degree bound provably be FRI. The codeword associated with this polynomial is the input to FRI.
In other words, the input to FRI is the sum of all codewords multiplied with weights $\alpha_i$ and again with weights $X^{\mathsf{d}-b_i} \beta_i$. The second part is necessary to adequately bound the individual degrees of $q_i(X)$.
FRI establishes that the input codeword has a low-degree defining polynomial. It does this by folding the working codeword in on itself using a random weight supplied by the verifier, over the course of several rounds. This folding procedure sends low-degree codewords to low-degree codewords, and high-degree codewords to high-degree codewords. The verifier checks the correct folding relation by inspecting the codewords in a series of indices.
Rather than transmitting the codewords in the clear, the prover first compresses them using a Merkle tree, and then transmits only the root. After all the Merkle roots are in, the verifier announces the indices where he would like to inspect the committed codewords. The prover answers by sending the indicated leafs along with their authentication paths.
In the last round, the prover sends the codeword in the clear. The length of this codeword is what happens after $\mathsf{r}$ halvings – specifically, its length is $N/2^{\mathsf{r}}$, where $N$ as the length of the original codeword.
The STARK mechanics described above suffice for proving the integral evolution of a simple state, i.e., on that is fully determined by $\mathsf{w}$ registers. It suffices for a digital signature scheme based on a proof of knowledge of a preimage under a hash function, or a verifiable delay function. But there is a considerable gap between that and a general-purpose virtual machine.
For instance, a machine following the von Neumann architecture needs to
At best, the simple state evolution descibes the evolution of the machine’s register set, which takes place inside the processor. But how does the processor interact with external data sources and sinks, like reading from and writing to memory? More importantly, how to prove and verify the integrity of these interactions?
Let’s use as an illustrative example the case where the processor writes something to RAM and then later reads it. To define this interaction with the fewest superfluous components, we need three registers:
clk
whose value increases by one in every cycle.mp
whose value represents an address in RAM.mv
whose value corresponds to the value of the RAM at the address mp
.Note that this simple register set does not contain a register for the current instruction. In reality the interactions with the RAM are only necessary if the instruction pertains to it – specifically, if it is a read or write instruction. We will need to add a register to this effect in order for the instruction set to be capable of doing other things beyond reading from and writing to RAM. For the time being we can tolerate this omission by requiring that the value of mv
corresponds to the RAM at location mp
in every cycle, even cycles where the instruction has nothing to do with memory. Whenever mp
changes, mv
automatically changes also.
This execution trace illustrates two of the problems we want to address.
clk |
mp |
mv |
---|---|---|
0 | 0 | 0 |
1 | 0 | 5 |
2 | 1 | 0 |
3 | 0 | 5 |
The entire RAM is initialized to zero. At the start of the computation, mp
is set to zero as well. In cycle 1, the memory cell at location 0 (because mp
=0) is set to 5. In cycle 2, the memory pointer is set to 1, and the memory value register assumes the value of the RAM at this location, which happens to be zero. In cycle 3, the memory point is set back to 0, and the memory value register assumes the value of the RAM at this location, which now turns out to be 5 because it was set to this value in cycle 1.
The two problems we wish to address are these:
mp
is reset to a previous value, the memory value register mv
changes. Its new value now must be consistent with the value it had the last time the memory pointer pointed to it. This constraint pertains to two rows, but the problem is that these two rows are generally not consecutive and not even spaced apart by a static amount. Transition constraints can only apply to consecutive pairs of rows.Notice how both of these issues can be solved through traditional transition constraints if only the table’s rows were sorted not by clock cycle but by memory pointer first and then clock cycle. In that case we would have this table:
clk |
mp |
mv |
---|---|---|
0 | 0 | 0 |
1 | 0 | 5 |
3 | 0 | 5 |
2 | 1 | 0 |
The transition constraints that enforce the issues raised above can now be articulated.
mp
changes, the new memory value mv
is zero.mp
is unchanged and the clock cycle clk
jumps by more than one, the memory value mv
cannot change.The STARK prover obviously cannot commit to this table instead of the one sorted by cycle count because doing so would undermine the transition constraints that enforce the correct evolution of the register set through time. It seems as though we have to choose: we can either enforce the correct evolution of the register set through time, or the consistency of RAM accesses.
The solution is for the prover to commit to both tables. The verifier verifies two sets of AIR constraints, one for the processor table and one for the memory table. But one question remains: how does the verifier verify that these tables pertain to the same execution trace? A permutation argument is what establishes that the sets of rows of the two tables are identical.
Without loss of generality, we can assume that the prover has access to random scalars $a, b, c, …$ supplied by the verifier. The prover can use these scalars to compress multiple columns into one, simply by taking a weighted vectorial sum. The verifier uses the same scalars to compute the same weighted sum in select coordinates but obviously never accesses the entire vector. This construction of taking weighted sums allows us to reduce a claim about the identical sets of rows of two tables, to an equivalent but easier to prove claim about the identical sets of elements of two rows.
Both tables are now reduced to a single column each. Let the elements in these columns be $(c_ i)_ i$ and $(k_ i)_ i$. The claim is that the only difference is the order, and that as sets $\lbrace c_ i \rbrace_ i = \lbrace k_ i\rbrace_ i$. Let $\alpha$ be another random scalar supplied by the verifier. Both tables, which are now reduced to a single column each, are extended with a new column that computes the products $\prod_i (\alpha - c_ i)$ and $\prod_i (\alpha - k_ i)$, respectively, by integrating a new factor $\alpha - c_i$ or $\alpha - k_i$ into a running product in each row. Specifically, the transition constraint for this extension column is $\forall i > 0: e_{i-1} \cdot (\alpha - c_ i) = e_ i$ (and analogously for the other table), where $e_i$ represents the value from the extension column.
$c_i$ | $e$ |
---|---|
$w$ | $\alpha - w$ |
$x$ | $(\alpha - w)(\alpha - x)$ |
$y$ | $(\alpha - w)(\alpha - x)(\alpha - y)$ |
$z$ | $(\alpha - w)(\alpha - x)(\alpha - y)(\alpha - z)$ |
$k_i$ | $e$ |
---|---|
$x$ | $\alpha - x$ |
$y$ | $(\alpha - x)(\alpha - y)$ |
$w$ | $(\alpha - x)(\alpha - y)(\alpha - w)$ |
$z$ | $(\alpha - x)(\alpha - y)(\alpha - w)(\alpha - z)$ |
If the two columns really do contain the same values but in a different order, then the products should be the same: \(\prod_i (\alpha - c_i) = \prod_i (\alpha - k_i) \enspace .\) To see this, just move the factors around until you get identically the same expression.
It is clear upon inspection that whenever we start from two tables whose sets of rows are identical, this technique will suffice to convince the verifier. That’s the completeness property. What about soundness? If we start from a pair of tables whose sets of rows are distinct, is it possible to trick the verifier into accepting? It turns out that this technique does induce a soundness degradation originating from two possible events.
In summary, as long as the cardinality of $\mathbb{F}$ is on the order of $2^\lambda$ where $\lambda$ is the security parameter, the soundness degradation is negligible.
The permutation argument establishes that two tables with the same height have the same rows up to order. In some cases we are interested in an alternative but linked relation: when one table’s list of rows appears in order as a sublist of another table’s list of rows. This relation occurs, for instance, when the processor reads input from a designated input tape. In between reading input symbols, the processor runs for an arbitrary number of cycles. The evaluation argument is what establishes relations such as these.
Without loss of generality, both tables columns are compressed into one column with the same technique to produce a linear combination with verifier-supplied weights $a, b, c, …$. Additionally, the larger table has to have a virtual or explicit indicator column, with values denoted by $\jmath_i$, that takes the value 1 if row $i$ is part of the sublist relation and 0 if it is not. Reusing the same notation as the previous section, the claimed relation is $(c_ i)_ {i \, \vert \, \jmath_ i = 1} = (k_ j)_ j$.
Like with the permutation argument both tables will be extended with a new column. Unlike the permutation argument, the evaluation argument interprets the row elements as the coefficients in reverse order, rather than the roots, of a polynomial whose value in $\alpha$ is computed step by step. Specifically, the transition constraints are given by
Note that the factors $\jmath_ {i-1}$ and $1-\jmath_ {i-1}$ enforce the conditional accumulation of a new term. Specifically, in un-indicated rows the running sum does not change whereas in indicated rows it changes in the same way that it changes in the smaller table.
$\jmath_ i$ | $c_ i$ | $e_ i$ |
---|---|---|
0 | $u$ | 0 |
1 | $x$ | $x$ |
0 | $v$ | $x$ |
1 | $y$ | $\alpha x + y$ |
0 | $w$ | $\alpha x + y$ |
1 | $z$ | $\alpha^2 x + \alpha y + z$ |
$c_ j$ | $e_ j$ |
---|---|
$x$ | $x$ |
$y$ | $\alpha x + y$ |
$z$ | $\alpha^2 x + \alpha y + z$ |
Like before, the honest prover is guaranteed to succeed and therefore there is no completeness degradation. The soundness degradation is bounded by an analogous argument to the one in the case of the permutation argument.
Conditioning – the function achieved by the variable $\jmath_ i$ in the previous section – applies to permutation arguments as well as to evaluation arguments. In this case what is being proved is that a subset of the rows of one table is equal to the set of rows of another table. Or indeed, it can be shown that two tables have a common subset or sublist. The combination with a univariate sumcheck argument can establish that the sublist or subset in question has a given cardinality.
It should be noted that not all columns need to be included in the random linear combination. Extending the previous examples, it makes no sense to include the input symbols in the permutation check to show the consistency of memory accesses. Not all columns must be present in all tables.
The relation that is proved by a permutation argument is actually multiset equality rather than set equality – multiplicity matters. When including the clk
column, the two notions are one and the same because clk
guarantees that every row in a table is unique.
This observation raises the question, is it possible to prove set equality when there is no column like clk
to make each row unique? The answer is yes! The following strategy accomplishes this task.
This deduplication technique gives rise to a table-lookup argument ^{1}. The lookup table consists of two columns, input and output. The processor (or whatever table is using lookups) has two similar columns along with an explicit indicator. The argument shows that the set of all (input, output) pairs from the indicated rows of the processor table is a subset of the (input, output) pairs defined by the lookup table. Note that the lookup table can be precomputed.
It should be noted that every permutation or evaluation argument leaks one linear relation worth of information about the contents of the columns. This information leakage comes from the terminal, which is the last value of the running product or running sum. The prover needs to send this value to the verifier explicitly.
In order to guarantee that this terminal value is independent of the cells of the column, it is important to start the running sum or product with a uniformly random initial value. The prover does not send this initial value to the verifier. Instead, the verifier verifies that the two extension columns start with the same initial value by subtracting their interpolating polynomials. This difference polynomial evaluates to zero in $X=1$ iff the extension columns start with the same initial value. So it is divisible by $X-1$ and this division gives rise to a difference quotient, which is to be included in the nonlinear combination.
In the diagram representation of an automaton (far) above every modular component maps onto a table. Every table comes with AIR constraints that establish the correct evolution of the information contained inside. Moreover, table relations assert the correctness of interactions between them.
Translating this notion to source code, it makes sense to create an abstract Table
class or module that captures important data and constraints interactions to it to a specific interface. So what is the correct interface for tables?
In terms of fields, which fields should and should not be part of a Table
is largely a matter of taste. However it cannot hurt to list the various data objects that a Table
pertains to, and suggest names for them.
matrix
is a two-dimensional array of field elements that represents the part of the algebraic execution trace that this table captures. This matrix is smaller before it is extended, and larger afterwards.base_width
is the width of the matrix before extension.full_width
is the width of the matrix after extension.height
is the height of the matrix. This dimension is independent of table extensions. Note that height
is always a power of 2 because the table is padded to have this height. The virtual machine can run any number of steps. The term length
denotes the length of the execution trace, but when transforming the execution trace into tables the tables are padded in a way that is consistent with the AIR.generator
is the generator of the multiplicative group of the finite field, common to all tables, used to derive the appropriate cosets over which to interpolate the columns or evaluate the polynomials resulting from that operation.order
is the multiplicative order of generator
.omicron
is the generator of the subgroup of order height
over which the columns are interpolated.num_randomizers
is the number of randomizers used when interpolating.Speaking of interpolating – one essential functionality associated with Table
s is the interpolation of its columns. This method, which might be called interpolate_columns
outputs one polynomial for each column such that the polynomial passes through all points $(x,y)$ defined by the powers of omicron
($x$-coordinates) and the value of the given column ($y$-coordinate). Additionally, the polynomial passes through num_randomizers
-many points $(x,y)$ defined by fixed $x$-coordinates (that do not coincide with powers of omicron
) and uniformly random $y$-coordinates. This process of interpolation produces polynomials of degree height
+ num_randomizers
- 1.
The number of randomizers has to be set in accordance with the number of nonlinear combination checks for verifying the correct nonlinear combination of codewords. Specifically, every nonlinear combination check generates up to two accesses in each of the trace polynomials, because of the transition constraints. We want the value of the trace polynomials at these locations to be uniformly distributed. So we need two times the number of nonlinear combination checks – that number of randomizers.
That being said, for reasons of performance the immediate next step following the polynomial interpolation is polynomial evaluation but on a different domain, resultin in Reed-Solomon codewords. The reason for the performance benefit is that arithmetic on polynomials is faster when they are represented as codewords rather than vectors of coefficients. The process of interpolation followed by evaluation on a larger domain is called low-degree extension and abbreviated lde
.
A Table
’s AIR constraints come in three varieties: boundary, transition, and terminal constraints. All types are represented by multivariate polynomials; the difference is where they apply.
boundary_constraints
returns these.transition_constraints
returns these.terminal_constraints
returns these.Moreover, all terminal constraints and some transition constraints pertain to the extension columns and can only be articulated relative to the verifier’s random scalars $a,b,c, \alpha…$. The suffix _ext
denotes the inclusion of extension AIR constraints. In general we always want the whole AIR, i.e., including the part covering the extension columns; nevertheless, it is worthwhile having having access to the “base” AIR (i.e., the part covering just the base table) for testing purposes.
Given knowledge of the polynomials (in either representation) and knowledge of the AIR constraints, it is possible to compute the quotients. These are found by evaluating them and dividing out the matching zerofiers. The methods {evaluate_
}+{boundary
,transition
,terminal
}+_constraints
perform this evaluation; {boundary
, transition
, terminal
}+_quotients
produce the quotients. Since we are interested in all of them, all_quotients
returns the concatenation of these lists.
Every quotient polynomial necessarily comes with a degree bound that depends on a) the degree of the polynomial that interpolates the column; and b) the concrete AIR that generated the quotient. These degree bounds are used to determine the appropriate shifts in the nonlinear combination. The methods {boundary_
, transition_
, terminal_
}+quotient_degree_bounds
compute the lists of matching degree bounds. Likewise, we are interested in all of them, so all_quotient_degree_bounds
returns the concatenation of these lists.
There are a number of notable differences between this tutorial and the Anatomy of a STARK. For convenience this predecessor tutorial is referred to as the Anatomy.
In the present tutorial the field is chosen as $\mathbb{F}_p$ where $p$ is the amazing prime $2^{64}-2^{32}+1$. This prime is mistakenly ^{2} called the Goldilocks prime by some people.
One of the drawbacks of using this field is that its cardinality is smaller than any reasonable security level. As a result, whenever random scalars from the verifier are needed, they should be sampled from an extension of this field. To this end $\mathbb{F}_{p^3} \cong \frac{\mathbb{F}[X]}{\langle X^3 -X + 1\rangle}$ is used. Moreover, the prover’s randomizer polynomial must have coefficients drawn uniformly at random from this field as well, not to mention the random initial values for permutation or evaluation arguments.
When zero-knowledge is important, the authentication paths in the Merkle tree of the zipped codeword leak small amounts of information. This step involves appending raw randomness to every leaf before computing the Merkle tree. With this option enabled, an authentication path for one leaf leaks almost no information about the leaf’s sibling – and exactly zero bits of information in the random oracle model, which is when an idealized hash function is used for the security proof.
The Anatomy computes a separate Merkle tree for each codeword. This tutorial zips the codewords together before computing the Merkle tree, so that each leaf contains a tuple of field elements. This change saves time for both prover and verifier and moreover generates smaller proofs.
More than one Merkle tree is being computed outside of the FRI subprotocol. The first Merkle tree contains the randomizer codeword and the codewords of all base columns. This Merkle root is sent to the verifier who responds with random scalars $a, b, c, \alpha, \ldots$ These random scalars are needed to compute the table extensions, and so the second Merkle tree contains the codewords of the extension columns.
This two-stage approach stands in contrast to the one-stage approach used in the Anatomy. The extra step of interaction makes the present qualify as a Randomized AIR with Preprocessing (RAP).
The third Merkle tree coincides with the codeword that goes into FRI. However, since the interface to FRI is being changed (see the next section), this Merkle tree is actually being computed prior to FRI.
In the Anatomy, FRI returns a list of indices where the top level codeword is probed to check the correct folding with respect to the next codeword. The key point is that the same indices are used to verify the correct nonlinear combination of the codewords that were committed to prior to FRI.
In contrast, the present tutorial uses a different set of indices. Now FRI is a completely standalone protocol that establishes that a given Merkle root decommits to a codeword corresponding to a low-degree polynomial. Likewise, the AET/AIR part of STARK outputs a Merkle root that decommits to a nonlinear combination of codewords committed to earlier. This nonlinear combination is still checked, but in a list of uniformly random indices sampled independently from FRI. While this change does increase the proof size as well as prover and verifier work, the interface between the FRI and AET/AIR parts is cleaner. This change is conducive to Halo-style recursion, whereby the recursive verifier verifies the nonlinear combination only, and the FRI part is postponed until the last possible step.
This cleaner interface to FRI raises an important question: how many incides do we sample to verify the nonlinear combination in order to target a security of $\lambda$ bits?
Let $s$ be the number of colinearity checks used in FRI, $t$ be the number of nonlinear combination checks in the AET/AIR part, $\varrho$ be the proportion of points where the codeword in question agrees with the nonlinear combination, and $N$ the length of the codeword. We want to upper bound the malicious prover’s success probability, so we are working under the assumption that this nonlinear combination does not correspond to a low-degree polynomial (because otherwise the prover is not being malicous).
Assume that the nonlinear combination codeword has maximal distance from the nearest low-degree codewords. While this assumption is rather heuristic, it is justified by condering that the cheating prover must include at least one unclean division in the nonlinear combination; the probability that this unclean codeword does not have maximal distance is vanishingly small. Under this assumption, the Hamming distance to the nearest low-degree codeword is $(1-\rho) N$, where $\rho$ is the expansion factor of FRI.
We distinguish two cases:
Case 1: $\varrho N \leq \rho N$. In this case there is a low-degree codeword that is consistent with the $\varrho N$ nonlinear combination constraints and by selecting this codeword the malicious prover guarantees that the FRI step succeeds. So the probability of success is the probability of sampling $t$ consistent points when only $\varrho N$ out of $N$ points are consistent, or $\frac{\binom{\varrho N}{t}}{\binom{N}{t}} = \frac{(\varrho N)!t!(N-t)!}{(\varrho N - t)! t! N!} = \prod_{i=0}^{t-1} \frac{\varrho N - i}{N - i} < \varrho^t$.
Case 2: $\varrho N > \rho N$. In this case there is no low-degree codeword that is consistent with all $\varrho N$ nonlinear combination constraints. However, there are low-degree codewords at Hamming distance $(\varrho - \rho)N$ from from this nonlinear combination codeword. Assume without loss of generality that the erroneous coordinates come in pairs $(i, N/2+i)$. The probability of malicious prover success in this case is $\frac{\binom{N/2 - (\varrho - \rho)N/2}{s}}{\binom{N/2}{s}}$ to pass FRI and $\frac{\binom{\varrho N}{t}}{\binom{N}{t}}$ to pass the nonlinear combination checks. A simplified upper bound on the probability of passing both is $(1 - (\varrho - \rho))^s \cdot \varrho^t$.
To find the optimal value of $\varrho$, derive this expression and equate it to zero. This operation gives \(0 = \frac{\mathsf{d} (1 - (\varrho - \rho))^s \cdot \varrho^t}{\mathsf{d}\varrho}\) \(= - s (1 - (\varrho - \rho))^{s-1} \varrho^t + t (1 - (\varrho - \rho))^s \varrho^{t-1}\) \(= -s \varrho + t (1 - (\varrho - \rho)) = t - (s + t)\varrho + t \rho\) and so the optimal value is $\varrho = \frac{t(\rho + 1)}{s + t}$.
At this point the malicious prover’s success probability is bounded by $\left( \frac{s(1 + \rho)}{s + t} \right)^s \left( \frac{t(1 + \rho)}{s+t}\right)^t$ and for $\lambda$ bits of security this value must be less than $2^{-\lambda}$. This expression implies that one can minimize $s$ or $t$ but not both simultaneously. Let’s consider some special cases.
The Anatomy only includes the boundary quotient polynomials in the nonlinear combination, and not the trace polynomials from which they are derived, at least when these boundary quotients exist. As a result, the trace polynomials are not proved to be low degree. The omission there does not degrade security because its low degree is implied by the low degree of the transition quotient polynomials and that of the boundary quotient polynomials.
In the present tutorial, the concrete boundary constraints are are intentionally abstract. While it may be possible to omit some trace polynomials from the nonlinear combination, whether this is secure depends on the concrete arithmetization of the virtual machine. This tutorial’s inclusion in the nonlinear combination supports making abstraction of the underlying VM.
Whenever a polynomial is computed, it is instantly transformed into a Reed-Solomon codeword of sufficient length. As a result, multiplication and division operations can be performed element-wise, resulting in a non-trivial speedup.
In principle it could be worthwhile to have two instances of low-degree extension (LDE). In the first, the interpolated columns are evaluated on a slightly larger intermediate domain to find codeword representatives of the interpolants. These codewords are used for polynomial arithmetic. In the second LDE step these codewords are extended to match with the FRI domain. This tutorial decided against using two separate domains for the sake of simplicity.
The verifier needs to evaluate zerofiers, and there are two options to make this process fast. The Anatomy uses preprocessed zerofiers because the number of cycles there is far from a power of two. In that context this choice makes sense because the number of cycles is fixed and is not close to a power of two. For the present tutorial, the number of cycles is not fixed. To deal with this problem, the tables are padded until the next power of two, making them compatible with group-theoretical zerofiers, which can be evaluated in $O(\log N)$ time.
If you are a person who finds it easier to read code than (maths-heavy) text, then you can check out the repository for a fully functional python implementation. Developed prior to writing the present explanation. However, the text makes abstraction of the particular implementation and in fact, there will be no code snippets interspersed with the text.
The purpose of this tutorial is didactical, not performance. That was one of the reasons for selecting python. It is also the reason why we are not reducing the AIR degree as far as we possibly can. And it’s also the reason why we chose Brainfuck as the target instruction set architecture. However, all of these choices do make for a very poorly performing prover. To keep running times manageable, the target security is set to impossibly small – 2 bits. The good news is that it’s a variable parameter; you can achieve any target security level just by modifying the right line in brainfuck_stark.py
.
A STARK engine consists of three parts: a virtual machine, a prover, and a verifier. While an execution of the virtual machine does not need be followed up with a prover, the fact that this option exists implies that the instruction set architecture had better be conducive to proving and verifying. Specifically, it should be easily arithmetizable – describable as the zero set of a small number of low-degree multivariate polynomials.
The virtual machine has two modes of operation:
run
takes the program and the input, executes the program, and returns the output.simulate
does the same but additionally outputs the execution trace.The tuple (input, program, output) is the computational integrity claim, and both the prover and verifier receive it as input. The execution trace is the secret additional input only for the prover.
The prover follows the workflow sketched below. This workflow implicitly defines the matching operations of the verifier. It is presented in an interactive language, but this protocol can obviously be made non-interactive with the Fiat-Shamir transform.
Next up: Part II: Brainfuck |
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0 - 1 - 2 - 3 - 4 - 5 |
This table-lookup argument is similar to Plookup except that it uses the element-wise inverse column along with an evaluation argument, whereas Plookup uses a custom argument to establish the correct order of the nonzero consecutive differences. ↩
Let’s set the record straight: Mike Hamburg coined the term “the Goldilocks prime” to refer specifically to $2^{448} - 2^{224} - 1$. ↩