# Memory-Sample Lower Bounds for Learning Parity with Noise

In this work, we show, for the well-studied problem of learning parity under noise, where a learner tries to learn x=(x_1,…,x_n) ∈{0,1}^n from a stream of random linear equations over F_2 that are correct with probability 1/2+ε and flipped with probability 1/2-ε, that any learning algorithm requires either a memory of size Ω(n^2/ε) or an exponential number of samples. In fact, we study memory-sample lower bounds for a large class of learning problems, as characterized by [GRT'18], when the samples are noisy. A matrix M: A × X →{-1,1} corresponds to the following learning problem with error parameter ε: an unknown element x ∈ X is chosen uniformly at random. A learner tries to learn x from a stream of samples, (a_1, b_1), (a_2, b_2) …, where for every i, a_i ∈ A is chosen uniformly at random and b_i = M(a_i,x) with probability 1/2+ε and b_i = -M(a_i,x) with probability 1/2-ε (0<ε< 1/2). Assume that k,ℓ, r are such that any submatrix of M of at least 2^-k· |A| rows and at least 2^-ℓ· |X| columns, has a bias of at most 2^-r. We show that any learning algorithm for the learning problem corresponding to M, with error, requires either a memory of size at least Ω(k ·ℓ/ε), or at least 2^Ω(r) samples. In particular, this shows that for a large class of learning problems, same as those in [GRT'18], any learning algorithm requires either a memory of size at least Ω((log |X|) · (log |A|)/ε) or an exponential number of noisy samples. Our proof is based on adapting the arguments in [Raz'17,GRT'18] to the noisy case.

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## 1 Introduction

In this work, we study the number of samples needed for learning under noise and memory constraints. The study of the resources needed for learning, under memory constraints was initiated by Shamir [Sha14] and Steinhardt, Valiant and Wager [SVW16], and has been studied in the streaming setting. In addition to being a natural question in learning theory and complexity theory, lower bounds in this model also have direct applications to bounded storage cryptography [Raz16, VV16, KRT17, TT18, GZ19, JT19, DTZ20, GZ21]. [SVW16] conjectured that any algorithm for learning parities of size (that is, learning from a stream of random linear equations in ) requires either a memory of size or an exponential number of samples. This conjecture was proven in [Raz16] and in follow up works, this was generalized to learning sparse parities in [KRT17] and more general learning problems in [Raz17, MM17, MT17, GRT18, BGY18, DS18, MM18, SSV19, GRT19, DKS19, GRZ20].

In this work, we extend this line of work to noisy Boolean function learning problems. In particular, we consider the well-studied problem of learning parity under noise (LPN). In this problem, a learner wants to learn from independent and uniformly random linear equations in where the right hand sides are obtained by independently flipping the evaluation of an unknown parity function with probability . Learning Parity with Noise (LPN) is a central problem in Learning and Coding Theory (often referred to as decoding random linear codes) and has been extensively studied. Even without memory constraints, coming up with algorithms for the problem has proven to be challenging and the current state-of-the-art for solving the problem is still the celebrated work of Blum, Kalai and Wasserman [BKW03] that runs in time . Over time, the hardness of LPN (and its generalization to non-binary finite fields) has been used as a starting point in several hardness results [KKMS08, FGKP09] and constructing cryptographic primitives [Ale03]. On the other hand, lower-bounds for the problem are known only in restricted models such as Statistical Query Learning111The SQ model does not seem to distinguish between noisy and noiseless variants of parity learning and yields the same lower bound in both cases. [Kea98].

Learning under noise is at least as hard as learning without noise and thus, memory-sample lower bounds for parity learning [Raz16] holds for learning parity under noise too. It is natural to ask – can we get better space lower bounds for learning parities under noise? In this work, we are able to strengthen the memory lower bound to for parity learning with noise.

Our results actually extend to a broad class of learning problems under noise. As in [Raz17] and follow up works, we represent a learning problem using a matrix. Let , be two finite sets (where represents the concept-class that we are trying to learn and represents the set of possible samples). Let be a matrix. The matrix represents the following learning problem with error parameter (): An unknown element was chosen uniformly at random. A learner tries to learn from a stream of samples, , where for every , is chosen uniformly at random and with probability .

### Our Results

We use extractor-based characterization of the matrix to prove our lower bounds, as done in [GRT18]. Our main result can be stated as follows (Corollary 2): Assume that are such that any submatrix of of at least rows and at least columns, has a bias of at most . Then, any learning algorithm for the learning problem corresponding to with error parameter requires either a memory of size at least , or at least samples. Thus, we get an extra factor of in the space lower bound for all the bounds on learning problems that [GRT18] imply, some of which are as follows (see [GRT18] for details on why the corresponding matrices satisfy the extractor-based property):

1. Parities with noise: A learner tries to learn , from (a stream of) random linear equations over which are correct with probability and flipped with probability . Any learning algorithm requires either a memory of size or an exponential number of samples.

2. Sparse parities with noise: A learner tries to learn of sparsity , from (a stream of) random linear equations over which are correct with probability and flipped with probability . Any learning algorithm requires:

1. Assuming : either a memory of size or samples.

2. Assuming : either a memory of size or samples.

3. Learning from noisy sparse linear equations: A learner tries to learn , from (a stream of) random sparse linear equations, of sparsity , over , which are correct with probability and flipped with probability . Any learning algorithm requires:

1. Assuming : either a memory of size or samples.

2. Assuming : either a memory of size or samples.

4. Learning from noisy low-degree equations: A learner tries to learn , from (a stream of) random multilinear polynomial equations of degree at most , over , which are correct with probability and flipped with probability . We prove that if , any learning algorithm requires either a memory of size or samples (where ).

5. Low-degree polynomials with noise: A learner tries to learn an -variate multilinear polynomial of degree at most over , from (a stream of) random evaluations of over , which are correct with probability and flipped with probability . We prove that if , any learning algorithm requires either a memory of size or samples.

### Techniques

Our proof follows the proof of [Raz17, GRT18] very closely and builds on that proof. We extend the extractor-based result of  [GRT18] to the noisy case and a straightforward adaptation to its proof gives the stronger lower bound for the noisy case (which reflects on the strength of the current techniques). The main contribution of this paper is not a technical one but establishing stronger space lower bounds for a well-studied problem of learning parity with noise, using the current techniques.

### Discussion and Open Problem

Let’s look at a space upper bound for the problem of learning parity with noise, that is, a learner tries to learn from a stream of samples of the form , where is chosen uniformly at random and with probability and with probability (here, represents the inner product of and in , that is, ).

##### Upper Bound:

Consider the following algorithm : Store the first samples. Check for every , if for at least fraction of the samples , agrees with . Output the first that satisfies the check. In expectation, would agree with for fraction of the samples, and otherwise for , in expectation, would agree with for half the samples. Therefore, for large enough , using Chernoff bound and a union bound, with high probability () over the samples, satisfies the check if and only if , and outputs the correct answer under such an event. uses samples and bits of space.

In this paper, we prove that any algorithm that learns parity with noise from a stream of samples (as defined above) requires bits of space or exponential number of samples. Improving the lower bound to match the upper bound (or vice versa) is a fascinating open problem and we conjecture that the upper bound is tight. As each sample gives at most bits of information about , we can at least show that a learning algorithm requires samples to learn (which corresponds to using bits of space if each sample is stored).

###### Conjecture 1.1.

Any learner that tries to learn from a stream of samples of the form , where is chosen uniformly at random and with probability and with probability , requires either bits of memory or samples.

The proof of the conjecture, if true, would lead to new technical insights (beyond extractor-based techniques) into proving time-space (or memory-sample) lower bounds for learning problems.

### Outline of the Paper

In Section 2, we establish certain notations and definitions, which are borrowed from [Raz17, GRT18]. We give a proof overview in Section 3 and prove the main theorem in Section 4.

## 2 Preliminaries

Denote by

the uniform distribution over

. Denote by the logarithm to base

. For a random variable

and an event , we denote by the distribution of the random variables , and we denote by the distribution of the random variable conditioned on the event .

#### Viewing a Learning Problem, with error 12−ε, as a Matrix

Let , be two finite sets of size larger than 1. Let and .

Let be a matrix. The matrix corresponds to the following learning problem with error parameter (). There is an unknown element that was chosen uniformly at random. A learner tries to learn from samples , where is chosen uniformly at random, and with probability and with probability . That is, the learning algorithm is given a stream of samples, , where each  is uniformly distributed, and with probability and with probability .

#### Norms and Inner Products

Let . For a function , denote by the norm of , with respect to the uniform distribution over , that is:

 ∥f∥p=(E\/x∈RX[|f(x)|p])1/p.

For two functions , define their inner product with respect to the uniform distribution over as

 ⟨f,g⟩=E\/x∈RX[f(x)⋅g(x)].

For a matrix and a row , we denote by the function corresponding to the -th row of . Note that for a function , we have . Here, represents the matrix multiplication of with .

#### L2-Extractors and L∞-Extractors

###### Definition 2.1.

-Extractor: Let be two finite sets. A matrix is a --Extractor with error , if for every non-negative with there are at most rows in with

 |⟨Ma,f⟩|∥f∥1≥2−r.

Let be a finite set. We denote a distribution over as a function such that . We say that a distribution has min-entropy if for all , we have .

###### Definition 2.2.

Extractor: Let be two finite sets. A matrix is a --Extractor if for every distribution with min-entropy at least and every distribution with min-entropy at least ,

 ∣∣∣∑a′∈A∑x′∈Xpa(a′)⋅px(x′)⋅M(a′,x′)∣∣∣≤2−r.

#### Branching Program for a Learning Problem

In the following definition, we model the learner for the learning problem that corresponds to the matrix , by a branching program, as done by previous papers starting with [Raz16].

###### Definition 2.3.

Branching Program for a Learning Problem: A branching program of length and width , for learning, is a directed (multi) graph with vertices arranged in layers containing at most vertices each. In the first layer, that we think of as layer 0, there is only one vertex, called the start vertex. A vertex of outdegree 0 is called a leaf. All vertices in the last layer are leaves (but there may be additional leaves). Every non-leaf vertex in the program has outgoing edges, labeled by elements , with exactly one edge labeled by each such , and all these edges going into vertices in the next layer. Each leaf in the program is labeled by an element , that we think of as the output of the program on that leaf.

Computation-Path: The samples that are given as input, define a computation-path in the branching program, by starting from the start vertex and following at step  the edge labeled by , until reaching a leaf. The program outputs the label of the leaf reached by the computation-path.

Success Probability: The success probability of the program is the probability that , where is the element that the program outputs, and the probability is over (where is uniformly distributed over and are uniformly distributed over , and for every , with probability and with probability ).

A learning algorithm, using samples and a memory of bits, can be modeled as a branching program222The lower bound holds for randomized learning algorithms because a branching program is a non-uniform model of computation, and we can fix a good randomization for the computation without affecting the width. of length and width . Thus, we will focus on proving width-length tradeoffs for any branching program that learns an extractor-based learning problem with noise, and such tradeoffs would translate into memory-sample tradeoffs for the learning algorithms.

## 3 Overview of the Proof

The proof adapts the extractor-based time-space lower bound of [GRT18] to the noisy case, which in turn built on [Raz17] that gave a general technique for proving memory-samples lower bounds. We recall the arguments in [Raz17, GRT18] for convenience.

Assume that is a --extractor with error , and let . Let be a branching program for the noisy learning problem that corresponds to the matrix . We want to prove that has at least length or requires at least width (that is, any learning algorithm solving the learning problem corresponding to the matrix with error parameter , requires either memory or exponential number of samples). Assume for a contradiction that is of length and width , where is a small constant.

We define the truncated-path, , to be the same as the computation-path of , except that it sometimes stops before reaching a leaf. Roughly speaking, stops before reaching a leaf if certain “bad” events occur. Nevertheless, we show that the probability that stops before reaching a leaf is negligible, so we can think of as almost identical to the computation-path.

For a vertex of , we denote by the event that reaches the vertex . We denote by the probability for (where the probability is over ), and we denote by the distribution of the random variable conditioned on the event . Similarly, for an edge  of the branching program , let be the event that traverses the edge . Denote, , and .

A vertex of is called significant if

 ∥∥Px|v∥∥2>2ℓ′⋅2−n.

Roughly speaking, this means that conditioning on the event that reaches the vertex , a non-negligible amount of information is known about . In order to guess with a non-negligible success probability, must reach a significant vertex. Lemma 4.1 shows that the probability that reaches any significant vertex is negligible, and thus the main result follows.

To prove Lemma 4.1, we show that for every fixed significant vertex , the probability that reaches is at most (which is smaller than one over the number of vertices in ). Hence, we can use a union bound to prove the lemma.

The proof that the probability that reaches is extremely small is the main part of the proof. To that end, we use the following functions to measure the progress made by the branching program towards reaching .

Let be the set of vertices in layer- of , such that . Let be the set of edges from layer- of to layer- of , such that . Let

 Zi=∑v∈LiPr(v)⋅⟨Px|v,Px|s⟩k′/2ε,
 Z′i=∑e∈ΓiPr(e)⋅⟨Px|e,Px|s⟩k′/2ε.

We think of as measuring the progress made by the branching program, towards reaching a state with distribution similar to .

We show that each may only be negligibly larger than . Hence, since it’s easy to calculate that , it follows that is close to , for every . On the other hand, if is in layer- then is at least . Thus, cannot be much larger than . Since is significant, and hence is at most .

The proof that may only be negligibly larger than is done in two steps: Claim 4.12 shows by a simple convexity argument that . The hard part, that is done in Claim 4.10 and Claim 4.11, is to prove that may only be negligibly larger than .

For this proof, we define for every vertex , the set of edges that are going out of , such that . Claim 4.10 shows that for every vertex ,

 ∑e∈Γout(v)Pr(e)⋅⟨Px|e,Px|s⟩k′/2ε

may only be negligibly higher than

 Pr(v)⋅⟨Px|v,Px|s⟩k′/2ε.

For the proof of Claim 4.10, which is the hardest proof in the paper, we follow [Raz17, GRT18] and consider the function . We first show how to bound . We then consider two cases: If is negligible, then is negligible and doesn’t contribute much, and we show that for every , is also negligible and doesn’t contribute much. If is non-negligible, we use the bound on and the assumption that is a --extractor to show that for almost all edges , we have that is very close to . Only an exponentially small () fraction of edges are “bad” and give a significantly larger . In the noiseless case, any “bad” edge can increase by a factor of 2 in the worst case, and hence [GRT18] raised and to the power of , as it is the largest power for which the contribution of the “bad” edges is still small (as their fraction is ). But in the noisy case, any “bad” edge can increase by a factor of at most in the worst case, and thus, we can afford to raise and to the power of . This is where our proof differs from that of [GRT18].

This outline oversimplifies many details. To make the argument work, we force to stop at significant vertices and whenever is large, that is, at significant values, as done in previous papers. And we force to stop before traversing some edges, that are so “bad” that their contribution to is huge and they cannot be ignored. We show that the total probability that stops before reaching a leaf is negligible.

## 4 Main Result

###### Theorem 1.

Let . Fix to be such that . Let , be two finite sets. Let . Let be a matrix which is a --extractor with error , for sufficiently large333By “sufficiently large” we mean that are larger than some constant that depends on . and , where . Let

 r:=min{r′2,(1−γ)k′2,(1−γ)ℓ′2−1}. (1)

Let be a branching program, of length at most and width at most , for the learning problem that corresponds to the matrix with error parameter . Then, the success probability of is at most .

###### Proof.

We recall the proof in [GRT18, Raz17] and adapt it to the noisy case. Let

 k:=γln22εk′andℓ:=γℓ′/3. (2)

Our proof differs from [GRT18] starting with Claim 4.5, which allows us to set to a larger value of instead of as set in [GRT18]. Note that by the assumption that and are sufficiently large, we get that and  are also sufficiently large. Since , we have . Thus,

 r

Let be a branching program of length and width444width lower bound is vacuous for as regardless of the width, samples are needed to learn. for the learning problem that corresponds to the matrix with error parameter . We will show that the success probability of is at most .

### 4.1 The Truncated-Path and Additional Definitions and Notation

We will define the truncated-path, , to be the same as the computation-path of , except that it sometimes stops before reaching a leaf. Formally, we define , together with several other definitions and notations, by induction on the layers of the branching program .

Assume that we already defined the truncated-path , until it reaches layer- of . For a vertex in layer- of , let be the event that reaches the vertex . For simplicity, we denote by the probability for (where the probability is over ), and we denote by the distribution of the random variable conditioned on the event .

There will be three cases in which the truncated-path stops on a non-leaf :

1. If is a, so called, significant vertex, where the norm of is non-negligible. (Intuitively, this means that conditioned on the event that reaches , a non-negligible amount of information is known about ).

2. If is non-negligible. (Intuitively, this means that conditioned on the event that reaches , the correct element could have been guessed with a non-negligible probability).

3. If is non-negligible. (Intuitively, this means that is about to traverse a “bad” edge, which is traversed with a non-negligibly higher or lower probability than probability of traversal under uniform distribution on ).

Next, we describe these three cases more formally.

#### Significant Vertices

We say that a vertex in layer- of is significant if

 ∥∥Px|v∥∥2>2ℓ⋅2−n.

#### Significant Values

Even if is not significant, may have relatively large values. For a vertex in layer- of , denote by the set of all , such that,

 Px|v(x′)>22ℓ+2r⋅2−n.

#### Bad Edges

For a vertex in layer- of , denote by the set of all , such that,

 ∣∣(M⋅Px|v)(α)∣∣≥2−r′.

#### The Truncated-Path T

We define by induction on the layers of the branching program . Assume that we already defined until it reaches a vertex in layer- of . The path stops on if (at least) one of the following occurs:

1. is significant.

2. .

3. .

4. is a leaf.

Otherwise, proceeds by following the edge labeled by  (same as the computational-path).

### 4.2 Proof of Theorem 1

Since follows the computation-path of , except that it sometimes stops before reaching a leaf, the success probability of is bounded (from above) by the probability that stops before reaching a leaf, plus the probability that reaches a leaf and .

The main lemma needed for the proof of Theorem 1 is Lemma 4.1 that shows that the probability that reaches a significant vertex is at most .

###### Lemma 4.1.

The probability that reaches a significant vertex is at most .

Lemma 4.1 is proved in Section 4.3. We will now show how the proof of Theorem 1 follows from that lemma.

Lemma 4.1 shows that the probability that stops on a non-leaf vertex, because of the first reason (i.e., that the vertex is significant), is small. The next two lemmas imply that the probabilities that stops on a non-leaf vertex, because of the second and third reasons, are also small.

###### Claim 4.2.

If is a non-significant vertex of then

 Prx[x∈Sig(v)|Ev]≤2−2r.
###### Proof.

Since is not significant,

 E\/x′∼Px|v[Px|v(x′)]=∑x′∈X[Px|v(x′)2]=2n⋅E\/x′∈RX[Px|v(x′)2]≤22ℓ⋅2−n.

Hence, by Markov’s inequality,

 Prx′∼Px|v[Px|v(x′)>22r⋅22ℓ⋅2−n]≤2−2r.

Since conditioned on , the distribution of is , we obtain

 Prx[x∈Sig(v)∣∣Ev]=Prx[(Px|v(x)>22r⋅22ℓ⋅2−n)∣∣Ev]≤2−2r.\qed
###### Claim 4.3.

If is a non-significant vertex of then

 Prai+1[ai+1∈Bad(v)]≤2−2r.
###### Proof.

Since is not significant, . Since is a distribution, . Thus,

 ∥∥Px|v∥∥2∥∥Px|v∥∥1≤2ℓ≤2ℓ′.

Since is a --extractor with error , there are at most elements with

 ∣∣⟨Mα,Px|v⟩∣∣≥2−r′⋅∥∥Px|v∥∥1=2−r′⋅2−n

The claim follows since is uniformly distributed over and since (Equation (1)). ∎

We can now use Lemma 4.1, Claim 4.2 and Claim 4.3 to prove that the probability that  stops before reaching a leaf is at most . Lemma 4.1 shows that the probability that  reaches a significant vertex and hence stops because of the first reason, is at most . Assuming that doesn’t reach any significant vertex (in which case it would have stopped because of the first reason), Claim 4.2 shows that in each step, the probability that stops because of the second reason, is at most . Taking a union bound over the steps, the total probability that stops because of the second reason, is at most . In the same way, assuming that doesn’t reach any significant vertex (in which case it would have stopped because of the first reason), Claim 4.3 shows that in each step, the probability that stops because of the third reason, is at most . Again, taking a union bound over the steps, the total probability that stops because of the third reason, is at most . Thus, the total probability that  stops (for any reason) before reaching a leaf is at most .

Recall that if doesn’t stop before reaching a leaf, it just follows the computation-path of . Recall also that by Lemma 4.1, the probability that reaches a significant leaf is at most . Thus, to bound (from above) the success probability of by , it remains to bound the probability that reaches a non-significant leaf and . Claim 4.4 shows that for any non-significant leaf , conditioned on the event that reaches , the probability for is at most , which completes the proof of Theorem 1.

###### Claim 4.4.

If is a non-significant leaf of then

 Pr[~x(v)=x|Ev]≤2−r.
###### Proof.

Since is not significant,

 E\/x′∈RX[Px|v(x′)2]≤22ℓ⋅2−2n.

Hence, for every ,

 Pr[x=x′|Ev]=Px|v(x′)≤2ℓ⋅2−n/2≤2−r

since (Equation (3)). In particular,

 Pr[~x(v)=x|Ev]≤2−r.\qed

This completes the proof of Theorem 1. ∎

### 4.3 Proof of Lemma 4.1

###### Proof.

We need to prove that the probability that reaches any significant vertex is at most . Let be a significant vertex of . We will bound from above the probability that  reaches , and then use a union bound over all significant vertices of . Interestingly, the upper bound on the width of is used only in the union bound.

#### The Distributions Px|v and Px|e

Recall that for a vertex of , we denote by the event that reaches the vertex . For simplicity, we denote by the probability for (where the probability is over ), and we denote by the distribution of the random variable conditioned on the event .

Similarly, for an edge  of the branching program , let be the event that traverses the edge . Denote, (where the probability is over ), and .

###### Claim 4.5.

For any edge  of , labeled by , such that , for any ,

where is a normalization factor that satisfies,

 ce≥1−4⋅2−2r.
###### Proof.

Let be an edge of , labeled by , and such that . Since , the vertex is not significant (as otherwise always stops on and hence ). Also, since , we know that (as otherwise never traverses  and hence ).

If reaches , it traverses the edge if and only if: (as otherwise stops on ) and , . Therefore, by Bayes’ rule, for any ,

where is a normalization factor, given by

 ce =∑{x′:x′∉Sig(v)∧M(a,x′)=b}Px|v(x′)(1+2ε)+∑{x′:x′∉Sig(v)∧M(a,x′)≠b}Px|v(x′)(1−2ε) =(1+2ε)⋅Prx[(x∉Sig(v))∧(M(a,x)=b)|Ev]+(1−2ε)⋅Prx[(x∉Sig(v))∧(M(a,x)≠b)|Ev].

Since is not significant, by Claim 4.2,

 Prx[x∈Sig(v)|Ev]≤2−2r.

Since ,

 ∣∣Prx[M(a,x)=1|Ev]−Prx[M(a,x)=−1|Ev]∣∣=∣∣(M⋅Px|v)(a)∣∣≤2−r′,

and hence for every ,

 Prx[M(a,x)=b′|Ev]≥12−2−r′.

Hence, by the union bound,

 ce≥(1+2ε)⋅(12−2−r′−2−2r)+(1−2ε)⋅(12−2−r′−2−2r)≥1−4⋅2−2r

(where the last inequality follows since , by Equation (1)). ∎

#### Bounding the Norm of Px|s

We will show that cannot be too large. Towards this, we will first prove that for every edge of that is traversed by with probability larger than zero, cannot be too large.

###### Claim 4.6.

For any edge  of , such that ,

 ∥∥Px|e∥∥2≤4⋅2ℓ⋅2−n.
###### Proof.

Let be an edge of , labeled by , and such that . Since , the vertex is not significant (as otherwise always stops on and hence