# Lazy Linear Knapsack

The continuous knapsack problem may be the simplest non-trivial linear programming problem:

$\max_{x \in [0, 1]^n} p’x$ subject to $w’x \leq b.$

It has a linear objective, one constraint, and each decision variable is bounded to ensure the optimum exists. Note the key difference from the binary knapsack problem: decision variables are allowed to take any value between 0 and 1. In other words, we can, e.g., stick half of a profitable but large item in the knapsack. That’s why this knapsack problem can be solved in linear time.

## Dual to primal is reasonable

Duality also lets us determine the shape of all optimal solutions to this problem. For each item $$i$$ with weight $$w_i$$ and profit $$p_i$$, let its profit ratio be $$r_i = p_i / w_i,$$ and let $$\lambda^\star$$ be the optimal dual (Lagrange or linear) multiplier associated with the capacity constraint $$w’x \leq b.$$ If $$\lambda^\star = 0,$$ we simply take all items with a positive profit ratio ($$r_i > 0$$) and a non-negative weight $$w_i \geq 0.$$ Otherwise, every item with a profit ratio $$r_i > \lambda^\star$$ will be at its weight upper bound (1 if $$w_i \geq 0$$, 0 otherwise), and items with $$r_i < \lambda^\star$$ will instead be at their lower bound (0 of $$w_i \leq 0$$, and 1 otherwise).

Critical items, items with $$r_i = \lambda^\star,$$ will take any value that results in $$w’x = b.$$ Given $$\lambda^\star,$$ we can derive the sum of weights for non-critical items; divide the remaining capacity for critical items by the total weight of critical items, and let that be the value for every critical item (with the appropriate sign for the weight).

For example, if we have capacity $$b = 10,$$ and the sum of weights for non-critical items in the knsapsack is $$8,$$ we’re left with another two units of capacity to distribute however we want among critical items (they all have the same profit ratio $$r_i = \lambda^\star,$$ so it doesn’t matter where that capacity goes). Say critical items with a positive weight have a collective weight of 4; we could then assign a value of $$2 / 4 = 0.5$$ to the corresponding decision variable (and 0 for critical items with a non-positive weight).

We could instead have $$b = 10,$$ and the sum of weights for non-critical items in the knapsack $$12$$: we must find two units of capacity among critical items (they all cost $$r_i = \lambda^\star$$ per unit, so it doesn’t matter which). If critical items with a negative weight have a collective weight of $$-3,$$ we could assign a value of $$-2 / -3 = 0.6\overline{6}$$ to the corresponding decision variables, and 0 for critical items with a non-negative weight.

The last case highlights something important about the knapsack: in general, we can’t assume that the weights or profits are positive. We could have an item with a non-positive weight and non-negative profit (that’s always worth taking), an item with positive weight and negative profit (never interesting), or weights and profits of the same sign. The last case is the only one that calls for actual decision making. Classically, items with negative weight and profit are rewritten away, by assuming they’re taken in the knapsack, and replacing them with a decision variable for the complementary decision of removing that item from the knapsack (i.e., removing the additional capacity in order to improve the profit). I’ll try to treat them directly as much as possible, because that reduction can be a significant fraction of solve times in practice.

The characterisation of optimal solutions above makes it easy to directly handle elements with a negative weight: just find the optimal multiplier, compute the contribution of non-critical elements (with decision variables at a bound) to the left-hand side of the capacity constraint, separately sums the negative and positive weights for critical elements, then do a final pass to distribute the remaining capacity to critical elements (and 0-weight / 0-value elements if one wishes).

## Solving the dual looks like selection

Finding the optimal multiplier $$\lambda^\star$$ is similar to a selection problem: the value is either 0 (the capacity constraint is redundant), or one of the profit ratios $$r_i,$$ and, given a multiplier value $$\lambda,$$ we can determine if it’s too high or too low in linear time. If the non-critical elements yield a left-hand side such that critical elements can’t add enough capacity (i.e., no solution with the optimal form can be feasible), $$\lambda$$ is too low. If the maximum weight of potentially optimal solutions is too low, $$\lambda$$ is too high.

We can thus sort the items by profit ratio $$r_i$$, compute the total weight corresponding to each ratio with a prefix sum (with a pre-pass to sum all negative weights), and perform a linear (or binary) search to find the critical profit ratio. Moreover, the status of non-critical items is monotonic as $$\lambda$$ grows: if an item with positive weight is taken at $$\lambda_0$$, it is also taken for every $$\lambda \leq \lambda_0$$, and a negative-weight item that’s taken at $$\lambda_0$$ is also taken for every $$\lambda \geq \lambda_0.$$ This means we can adapt selection algorithms like Quickselect to solve the continuous knapsack problem in linear time.

I’m looking at large instances, so I would like to run these algorithms in parallel or even distributed on multiple machines, and ideally use GPUs or SIMD extensions. Unfortunately, selection doesn’t parallelise very well: we can run a distributed quickselect where every processor partitions the data in its local RAM, but that still requires a logarithmic number of iterations.

## Selection looks like quantile estimation; does the dual?

Lazy Select offers a completely different angle for the selection problem. Selecting the $$k$$th smallest element from a list of $$n$$ elements is the same as finding the $$k / n$$th quantile1 in that list of $$n$$ elements. We can use concentration bounds2 to estimate quantiles from a sample of, e.g., $$m = n^{3/4}$$ elements: the population quantile value is very probably between the $$qm - \frac{\log m}{\sqrt{m}}$$th and $$qm + \frac{\log m}{\sqrt{m}}$$th values of the sample. Moreover, this range very probably includes at most $$\mathcal{O}(n^{3/4})$$ elements3, so a second pass suffices to buffer all the elements around the quantile, and find the exact quantile. Even with a much smaller sample size $$m = \sqrt{n},$$ we would only need four passes.

Unfortunately, we can’t directly use that correspondence between selection and quantile estimation for the continuous knapsack.

I tried to apply a similar idea by sampling the knapsack elements equiprobably, and extrapolating from a solution to the sample. For every $$\lambda,$$ we can derive a selection function $$f_\lambda (i) = I[r_i \geq \lambda]w_i$$ (invert the condition if the weight is negative), and scale up $$\sum_i f(i)$$ from the sample to the population). As long as we sample independently of $$f$$, we can reuse the same sample for all $$f_\lambda.$$ The difficulty here is that, while the error for Lazy Select scales as a function of $$n,$$ the equivalent bounds with variable weights are a function of $$n(|\max_i w_i| + |\min_i w_i|)^2.$$ That doesn’t seem necessarily practical; scaling with $$\sum_i |w_i|$$ would be more reasonable.

Good news: we can hit that, thanks to linearity.

Let’s assume weights are all integers. Any item with weight $$w_i$$ is equivalent to $$w_i$$ subitems with unit weight (or $$-w_i$$ elements with negative unit weight), and the same profit ratio $$r_i$$, i.e., profit $$p_i / |w_i|$$. The range of subitem weights is now a constant.

We could sample uniformly from the subitems with a Bernoulli for each subitem, but that’s clearly linear time in the sum of weights, rather than the number of elements. If we wish to sample roughly $$m$$ elements from a total weight $$W = \sum_i |w_i|,$$ we can instead determine how many subitems (units of weight) to skip before sampling with a Geometric of success probability $$m / W.$$ This shows us how to lift the integrality constraint on weights: sample from an Exponential with the same parameter $$m / W!$$

That helps, but we could still end up spending much more than constant time on very heavy elements. The trick is to deterministically special-case these elements: stash any element with large weight $$w_i \geq W / m$$ to the side, exactly once. By Markov’s inequality,4 we know there aren’t too many heavy elements: at most $$m.$$

## Let’s test this out

The heart of the estimation problem can be formalised as follows: given a list of elements $$i \in [n]$$ with weight $$w_i \geq 0$$, generate a sample of $$m \leq n$$ elements ahead of time. After the sample has been generated, we want to accept an arbitrary predicate $$p \in \{0,1\}^n$$ and estimate $$\sum_{i\in [n]} p(i) w_i.$$

We just had a sketch of an algorithm for this problem. Let’s see what it looks like in Python. The initial sample logic has to determine the total weight, and sample items with probability proportional to their weight. Items heavier than the cutoff are not considered in the sample and instead saved to an auxiliary list.

We can assemble the resulting sample (and list of “large” elements) to compute a lower bound on the weight of items that satisfy any predicate that’s independent of the sampling decisions. The value for large elements is trivial: we have a list of all large elements. We can subtract the weight of all large elements from the total item weight, and determine how much we have to extrapolate up.

And finally, here’s how we can sample from an arbitrary list of items, compure a lower bound on the weight of items that satisfy a predicate, and compare that with the real lower bound.

How do we test that? Far too often, I see tests for randomised algorithms where the success rate is computed over randomly generated inputs. That’s too weak! For example, this approach could lead us to accept that the identity function is a randomised sort function, with success probability $$\frac{1}{n!}.$$

The property we’re looking for is that, for any input, the success rate (with the expectation over the pseudorandom sampling decisions) is as high as requested.

For a given input (list of items and predicate), we can use the Confidence sequence method (CSM) to confirm that the lower bound is valid at least $$1 - \alpha$$ of the time.

With a false positive rate of at most one in a million,5 we can run automated tests against check_bounds. I’ll use Hypothesis to generate list of pairs of weight and predicate value:

Bimodal inputs tend to be harder, so we can add a specialised test generator.

Again, we use Hypothesis to generate inputs, and the Confidence sequence method (available in C, Common Lisp, and Python) to check that the lower bound is valid with probability at least $$1 - \alpha$$. The CSM tests for this statistical property with power 1 and adjustable error rate (in our case, one in a million): we only provide a generator for success values, and the driver adaptively determines when it makes sense to make a call and stop generating more data, while accounting for multiple hypothesis testing.

TL;DR: the estimation algorithm for individual sampling passes works, and the combination of Hypothesis and Confidence Sequence Method lets us painlessly test for a statistical property.

We can iteratively use this sampling procedure to derive lower and (symmetrically) upper bounds for the optimal Lagrange multiplier $$\lambda^\star,$$ and Hoeffding’s inequality lets us control the probability that the lower and upper bounds are valid. Typically, we’d use a tolerance of $$\sqrt{\log(n) / n},$$ for an error rate of $$1 / n^2.$$ I prefer to simply use something like $$7 / \sqrt{n}:$$ the error rate is then less than $$10^{-42},$$ orders of manitude smaller than the probability of hardware failure in any given nanosecond.6 We can still check for failure of our Las Vegas algorithm, but if something went wrong, it’s much more likely that we detected a hardware failure than anything else. It’s like running SuperPi to stress test a computer, except the work is useful. 😉

## Repeat as necessary to solve a knapsack

How many sampling passes do we need? Our bounds are in terms of the sum of item weight: if we let our sample size be in $$\Theta(\sqrt{n}),$$ the sum of weights $$\sum_i |w_i|$$ for unfathomed items (that may or may not be chosen depending on the exact optimal multiplier $$\lambda^\star$$ in the current range) will very probably shrink by a factor of $$\Omega(n^{1/4}).$$ The initial sum can, in the worst case, be exponentially larger than the bitlength of the input, so even a division by $$n^{1/4}$$ isn’t necessarily that great.

I intend to apply this Lazy Linear Knapsack algorithm on subproblems in a more interesting solver, and I know that the sum of weights is bounded by the size of the initial problem, so that’s good enough for me! After a constant ($$\approx 4$$) number of passes, the difference in item weight between the lower and upper bound on $$\lambda^\star$$ should also be at most 1. One or two additional passes will get me near optimality (e.g., within $$10^{-4}$$), and the lower bound on $$\lambda^\star$$ should thus yield a super-optimal solution that’s infeasible by at most $$10^{-4},$$ which is, for my intended usage (again), good enough.

Given an optimal enough $$\lambda^\star,$$ we can construct an explicit solution in one pass, plus a simple fixup for critical items. This Lazy Knapsack seems pretty reasonable for parallel or GPU computing: each sampling pass only needs to read the items (i.e., no partitioning-like shuffling) before writing a fraction of the data to a sample buffer, and we only need a constant number of passes (around 6 or 7) in the worst case.

1. It’s more like a fractional percentile, but you know what I mean: the value such that the distribution function at that point equals $$k / n$$.

2. Binomial bounds offer even stronger confidence intervals when the estimate is close to 0 or 1 (where Hoeffding’s bound would yield a confidence interval that juts outside $$[0, 1]$$), but don’t impact worst-case performance.

3. Thanks to Hoeffding’s inequality, again.

4. That’s a troll. I think any self-respecting computer person would rather see it as a sort of pigeonhole argument.

5. We’re juggling a handful of error rates here. We’re checking whether the success rate for the Lazy Knapsack sampling subroutine is at least as high as $$1 - \alpha,$$ as requested in the test parameters, and we’re doing so with another randomised procedure that will give an incorrect conclusion at most once every one million invocation.

6. This classic Google study found 8% of DIMMs hit at least one error per year; that’s more than one single-bit error every $$10^9$$ DIMM-second, and they’re mostly hard errors. More recently, Facebook reported that uncorrectable errors affect 0.03% of servers each month; that’s more than one uncorrectable error every $$10^{10}$$ server-second. If we performed one statistical test every nanosecond, the probability of memory failure alone would still dominate statistical errors by $$10^{20}!$$

Jan 20th, 2020