Sunday, March 06, 2016

Solutions to Facebook Hacker Cup 2016

Since the official solutions aren't out yet, here are my solutions. Although I scored 0 in the contest, four of my solutions were basically correct, one I was able to solve fairly easily after another contestant pointed out that my approach was wrong, and I figured out Grundy Graphs on the way home.

Snake and Ladder

I'll describe two solutions to this. There are a few things one can start off with. Firstly, if there are rows at the top or bottom that are completely full of flowers, just cut them off — they make no difference. Then, handle N=1 specially: if K=1, the answer is 1, if K=0, it is 2. A number of people (including all the contest organisers!) messed up the first, and I messed up the second. Also, if one rung has two flowers on it, the answer is 0.

My contest solution used dynamic programming. If one cuts the ladder half-way between two rungs and considers the state of the bottom half, it can be divided into four cases:
  1. Only the left-hand side has the snake on it.
  2. Only the right-hand side has the snake on it.
  3. Both sides have the snake on it, and the two are connected somewhere below.
  4. Both sides have the snake on it, and the two do not connect.
It is reasonably straightforward (as long as one is careful) to determine the dynamic programming transitions to add a rung, for each possibility of a flower on the left, on the right, or no flower. Advancing one rung at a time will be too slow for the bound on N, but the transitions can be expressed as 4×4 matrices and large gaps between flowers can be crossed with fast matrix exponentiation. Finally, one must double the answer to account for the choice of which end is the head.

The alternative solution is to argue on a case-by-case basis. If there no flowers, then one can pick a row for the head and a row for the tail, and pick a column for the head (the column for the tail is then forced by parity). Once these choices are made, there is only one way to complete the snake. Also, the head and tail can only occupy the same row if it is the top or bottom row, giving 2N(N-1)+4 ways.

If there is at least one flower, then the head must be above the top flower and the tail below the bottom flower (or vice versa). Again, one can choose a row for each, with the column being forced by parity. One must also consider special cases for a flower in the top/bottom row.

Boomerang Crew

I think this was possibly the easiest of the problems, provided you thought clearly about it (which I failed to do). Clearly any opponents weaker than (or equal to) your champion can be defeated just be putting them up against your champion. Also, if you're going to defeat P players, they might as well be the weakest P players of the opposition. For a given P, can this be done? There will be a certain number of strong players (stronger than our champion), which we will have to sacrifice our weaker players against to wear them down. What's more, once we've worn one down and beaten him/her, we have to use the winner of that game to wear down the next strong opponent. Thus, we should match up the S strong players against our best S players in some order, and use our remaining players as sacrifices. Ideally we'd like to wear down each opponent to exactly the skill level of the player we will use to win; anything more is wasted effort. We can apply this idea greedily: for each opponent, find the remaining strong player of ours who can win by the smallest margin.

With that method to test a particular P, it is a standard binary search to determine the number of players we can beat.

Grundy Graph

I think this was the hardest problem: my submission wasn't a real solution and was just for a laugh; several other people I spoke to had submitted but didn't believe in their solution. During the contest I suspected that the solution would be related to 2-SAT, but it wasn't until the flight home that I could figure out how to incorporate turn ordering.

Let us construct an auxiliary directed graph A whose vertices correspond to those of the original. An edge u->v in A means that if u is black, then v must be black as well for Alice to win. All the edges in the original graph are clearly present in A. However, for each u->v in the original, we also add v'->u', where x' is the vertex assigned a colour at the same time as x. This is because if v' is black but u' is white, then it implies that u is black and v is white. This is the same as the contrapositive edge in a 2-SAT graph. Let x=>y indicate that y is reachable from x in A.

There are a number of situations in which we can see that Bob can win:
  • If any strongly connected component contains both x and x', then Bob wins regardless of what Alice and Bob play.
  • If Bob controls vertices x and y and x=>y, then Bob wins by assigning x black and y white, regardless of Alice's play.
  • If Bob controls x, Alice controls y, x=>y, x' => y', and y is played before x, then Bob wins by assigning y the opposite colour to x.
We claim that if none of the above apply, then Alice wins. Her strategy is that on her turn, she must colour x black if there is any u such that u => x and either u has already been coloured black, u could later be coloured black by Bob, or u = x'. This ensures that it does not immediately allow Bob to win, nor allows Bob to use it to win later. The only way this could fail is if both x and x' are forced to be black (if neither is forced, Alice can pick freely).

Suppose u => x and v => x', where u and v are as described above. Then x => v', so u => v'. Firstly consider u = x'. We cannot have v = x (because otherwise x, x' are in the same SCC). Now x' => v', so v => v'. If Bob controls v, then we hit the second case above; if Alice controls v, then v => v' means she would never have chosen v. So we have a contradiction. Now, if u ≠ x' (and by symmetry, v ≠ x), then consider u => v' again. Alice cannot control v/v', since she would then have chosen v' rather than v. But similarly, v => u', so Alice couldn't have chosen u. It follows that Bob controls u and v, which can only be possible if u = v'. But u and u' can only both be relevant if Bob hasn't decided their colour yet, which is the third case above.

The implementation details are largely left as an exercise. Hint: construct the SCCs of A, and the corresponding DAG of the SCCs, then walk it in topological order to check the conditions. The whole thing takes linear time.


The concept here is a reasonably straightforward exponential DP, but needs a careful implementation to get all the formulae right, particularly in the face of potential numeric instability. The graph itself isn't particularly relevant; the only "interesting" vertices are the start and the gifts, and the only useful information is the distance from each interesting vertex to each other one (if reachable) and the distance from each gift to the nearest dead-end. The DP state is the set of gifts collected and the current position (one of the interesting vertices). When one is at a gift, there are two possible options:
  1. Pick another gift, and make towards it along the shortest path. Based on the path length L, there is a probability P^L of reaching it in time DL, and a probability 1 - P^L of respawning, with an expected time computable from a formula.
  2. Aim to respawn by heading for the nearest dead-end. Again, one can work out a formula for the expected time to respawn.
When one is at the start, one should of course pick a gift and head for it. Again, one can compute an expected time to reach it (one obtains a formula that includes its own result, but a little algebra sorts that out).

Maximinimax flow

I think I liked this problem the best, even though I screwed up the limits on the binary search. Firstly, what is the minimax flow? For a general graph it would be nasty to compute, but this graph has the same number of edges as vertices. This means that it has exactly one cycle with trees hanging off of it. It shouldn't be too hard to convince yourself that the minimax flow is the smaller of the smallest edge outside the cycle and the sum of the two smallest edges inside the cycle.

Veterans will immediately reach for a binary search, testing whether it is possible to raise the minimax flow to a given level. For the edges outside the ring, this is a somewhat standard problem that can be solved with query structures e.g. Fenwick trees, indexed by edge value. To raise the minimum to a given level, one needs to know the number of edges below that level (one Fenwick tree, storing a 1 for each edge), and their current sum (a second tree, storing the edge weight for each edge).

For the vertices inside the cycle it is only slightly more complex. If the required level is less than double the second-smallest edge, then one need only augment the smallest edge. Otherwise one raises all edges to half the target, with a bit of special-casing when the target is odd. The Fenwick tree can also be queried to find the two smallest edges, but I just added a std::multiset to look this up.

Rainbow strings

This seemed to be one of the most-solved problems, and is possibly the easiest if you have a suffix tree/array routine in your toolbox. A right-to-left sweep will yield the shortest and longest possible substring for each starting position (next green for the shortest, next red for the longest). The trick is to process queries in order of length. Each entry in the suffix array will become usable at some length, and unusable again at another length, and these events, along with the queries, can be sorted by length for processing. The only remaining work is to determine the Kth of the active suffixes, which can be done with a Fenwick tree or other query structure.

Finally, one needs the frequency table for each prefix to be able to quickly count the frequency in each substring.