18. Bootstrapping & CKKS
Bootstrapping
Recall that BGV has a limit on the number of operations, so it cannot evaluate a circuit with a large depth. This was because of the growing noise, so we need a way to remove the noise.
An easy answer is decrypting the ciphertext and encrypting it again, but we want to do it without using the secret key.
Bootstrapping is a method to convert SHE into FHE.
Key Idea
The main idea is to homomorphically evaluate the decryption circuit over encrypted $\bf{s}$.
Let $\bf{c}$ be an encryption of $m \in \braces{0, 1}$, at the lowest level $0$. (Cannot perform multiplications anymore) The decryption algorithm, with a secret $\bf{s}$ fixed, is a function of $\bf{c}$.
Change the perspective and view it as a function of $\bf{s}$.
\[f(\bf{s}) = D(\bf{s}, \bf{c}) : \braces{0, 1}^n \ra \braces{0, 1}\]Then $f(\bf{s}) = m$.
Let $\bf{s}’ \in \braces{0, 1}^n$ be a new secret key. Generate the bootstrapping keys
\[BK = \braces{\bf{k}_i}_{i=1}^n, \qquad \bf{k}_i = E(\bf{s}', s_i).\]Then by the homomorphic property of $f$,
\[f(\bf{k_1}, \bf{k}_2, \dots, \bf{k}_n) = f\big( E(\bf{s}', s_1), \dots, E(\bf{s}', s_n) \big) = E\big( \bf{s}', f(s_1, \dots, s_n) \big) = E(\bf{s}', m).\]Example with BGV
Technically, the expression $f(\bf{k_1}, \bf{k}_2, \dots, \bf{k}_n)$ doesn’t make sense, but it works. Consider a message $m$ encrypted with secret $\bf{s}$ in the BGV scheme.
\[\bf{c} = (b, \bf{a}), \quad b = \span{\bf{a}, \bf{s}} + m + 2e \pmod q.\]The decryption is $r = b + \span{\bf{a}, \bf{s}} \pmod q$, and then taking the least significant bit. Consider it as a function
\[f(\bf{s}) = b + \span{\bf{a}, \bf{s}} = b + \sum_{i=1}^n a_is_i.\]For a new key $\bf{s}’ = (s_1’, \dots, s_n’)$, generate bootstrapping keys $\bf{k}_i = E(\bf{s}’, s_i)$ and plugging it in forcefully gives
\[\begin{aligned} f(\bf{k}_1, \dots, \bf{k}_n) &= b + \sum_{i=1}^n a_i E(\bf{s}', s_i) = b + \sum_{i=1}^n E(\bf{s}', a_is_i) \\ &=b + E\paren{\bf{s}', \sum_{i=1}^n a_is_i} = b + E\paren{\bf{s}', \span{\bf{a}, \bf{s}}}. \end{aligned}\]Since an encryption of $\span{\bf{a}, \bf{s}}$ with $\bf{s}’$ is $\span{\bf{a}’, \bf{s}’} + \span{\bf{a}, \bf{s}} + 2e’ \pmod q$, the above equation equals
\[\begin{aligned} b' &=b \span{\bf{a}', \bf{s}'} + \span{\bf{a}, \bf{s}} + 2e' \\ &= \span{\bf{a}', \bf{s}'} + m + 2(e + e') \pmod q. \end{aligned}\]Indeed, decrypting $b’$ will give $m$. So we have $E(\bf{s}’, m)$ from $f(\bf{k}_1, \dots, \bf{k}_n)$.^{1}
Bootstrapping Procedure
Given an encryption $\bf{c}$ of $m$ at level $0$, perform the following procedure.
Bootstrapping Key Generation
 Choose a new secret key $\bf{s}’ \in \braces{0, 1}^n$.
 Generate bootstrapping key ${} BK = \braces{\bf{k}i}{i=1}^n {}$ where $\bf{k}_i = E(\bf{s}’, s_i)$.
Bootstrapping
 Generate a circuit representation $f : \braces{0, 1}^n \ra \braces{0, 1}$ of the decryption function $D(\cdot, \bf{c})$.
 Compute and output $\bf{c}’ = f(\bf{k}_1, \dots, \bf{k}_n)$.
The bootstrapping procedure returns an encryption of $m$ under $\bf{s}’$, as shown above. The key idea here is that $\bf{k}_i$ are fresh ciphertexts at level $L$. Even though a few levels are consumed during the evaluation of $f$, the resulting ciphertext $\bf{c}’$ is not at level $0$ anymore, allowing us to do more computation.
Suppose that the homomorphic evaluation of $f$ requires depth $d$, consuming $d$ levels. Then we say that the BGV scheme is bootstrappable if $d < L$. The output ciphertext $\bf{c}’$ will have level $l = L  d > 0$, which we call remaining level.
Thus, we need to set $L$ large enough in the BGV scheme so that it is bootstrappable. But larger $L$ results in larger $q$, reducing the security of the scheme. This is another reason we must use modulus switching. Without it, we wouldn’t have been able to set $L$ large enough for bootstrapping.
Fully Homomorphic Encryption
Thus, if BGV is bootstrappable, then we can apply bootstrapping on the ciphertext whenever its level reaches $0$. Now we can evaluate any circuit of finite depth.
There is a slight catch here. For every bootstrapping procedure, we need a bootstrapping key. This must be generated by the owner of the original secret. As a result, lots of secret keys are required to homomorphically evaluate a circuit.
\[\bf{s} \ra \bf{s}' \ra \bf{s}'' \ra \cdots\]Currently, we set $\bf{s}’ = \bf{s}$ and make the chain circular, so the bootstrapping keys are $E(\bf{s}, s_i)$. $\bf{s}$ is being encrypted by itself. We wonder if this is secure, but there is no known proof for this. This is used as an assumption called the circular security assumption.
Designing an FHE scheme without the circular security assumption is currently an open problem.
CKKS Scheme
The BGV scheme operates on $\Z_p$, so it doesn’t work on real numbers. CheonKimKimSong (CKKS) scheme works on real numbers using approximate computation.
Approximate Computation
Computers use floating point representations for real numbers.
\[2.9979 \times 10^8\]Here, $2.9979$ is the significand, $10$ is the base and $8$ is the exponent. We also call $10^8$ the scaling factor.
Floating point operations involve rounding, but rounding is not easy in homomorphic encryption. Using the BGV scheme on $\Z_p$, there are $2$ methods to do this.
 Bitwise Encryption
 $32$bit integer results in $32$ ciphertexts.
 Binary multiplier circuits can be used for multiplication.
 Rounding is easy if done this way.
 But this is extremely inefficient. Huge number of gates are required, consumes a lot of levels.
 Integer Encryption
 To encrypt the significant, use a modulus large enough, such as $p > 2^{32}$.
 For multiplication, use $p > 2^{64}$.
 But rounding is hard in $\Z_p$.
So our wish is to design an HE scheme that natively supports rounding operation!
CKKS Description
In the LWE problem, error was added for security. This can be exploited, since computing floating points allows some rounding errors.
Let $n, q, \sigma$ be parameters for LWE and set a scaling factor $\Delta > 0$.
Key Generation
 A secret key is chosen as $\bf{s} = (s_1, \dots, s_n) \in \braces{0, 1}^n$, with its linearization gadget.
Encryption: message $m \in \R$.
 Randomly sample $\bf{a} = (a_1, \dots, a_n) \la \Z_q^n$ and $e \la D_\sigma$.
 Compute $b = \span{\bf{a}, \bf{s}} + \round{\Delta \cdot m} + e \pmod q$.
 Output ciphertext $\bf{c} = (b, \bf{a}) \in \Z_q^{n+1}$.
Decryption
 Compute $\mu = b + \span{\bf{a}, \bf{s}} \pmod q$.
 Output $m’ = \Delta\inv \cdot \mu \in \R$.
Note that the decrypted output is $m’$, which is not equal to $m$. We have
\[\mu = \round{\Delta \cdot m} + e\]if $\mu$ is small. (ex. $\abs{\mu} < q/2$) But $m’ = \Delta\inv \cdot \mu \neq m$. The traditional correctness does not apply here, since $D(\bf{s}, \bf{c}) \neq m$.
Instead, CKKS is an approximate encryption. The exact $m$ is not recovered, but we get an approximation $m’$ with bounded error,
\[\abs{m  m'} \leq \frac{1}{\Delta} (0.5 + \abs{e}).\]This is okay, since small numerical errors are allowed in floatingpoint operations. Also, it can be seen from this inequality that $\Delta$ is sort of a precision.
Also, CKKS is secure under the LWE assumption.
Operations on Ciphertexts in CKKS
The overall process is similar to that of BGV, with some additional changes.
Remember that if $\bf{c} = (b, \bf{a})$ is an encryption of $m \in \R$, then
\[\mu = b + \span{\bf{a}, \bf{s}} \pmod q, \quad \mu \approx \Delta \cdot m.\]Addition in CKKS
Let $\bf{c} = (b, \bf{a})$ and $\bf{c}’ = (b’, \bf{a}’)$ be encryptions of $m, m’ \in \R$. Then, $\bf{c}_\rm{add} = \bf{c} + \bf{c}’$ is an encryption of $m + m’$.
Proof. Decrypt $\bf{c}_\rm{add} = (b + b’, \bf{a} + \bf{a}’)$.
\[\mu_\rm{add} = \mu + \mu' = (b + b') + \span{\bf{a} + \bf{a}', \bf{s}} \pmod q.\]If $\abs{\mu + \mu’} < q/2$, then
\[\mu_\rm{add} = \mu + \mu' = \Delta \cdot (m + m'),\]so the decryption results in $\Delta\inv \cdot (\mu + \mu’) \approx m + m’$.
Multiplication in CKKS
We also use tensor products, and their properties.
Let $\bf{c} = (b, \bf{a})$ and $\bf{c}’ = (b’, \bf{a}’)$ be encryptions of $m, m’ \in \R$. Then,
\[\bf{c}_\rm{mul} = \bf{c} \otimes \bf{c}' = (bb', b\bf{a}' + b' \bf{a}, \bf{a} \otimes \bf{a}')\]is an encryption of $mm’$ with $(1, \bf{s}, \bf{s} \otimes \bf{s})$.
Proof. It can be checked that
\[\begin{aligned} \mu_\rm{mul} &= \mu\mu' = (b + \span{\bf{a}, \bf{s}})(b' + \span{\bf{a}', \bf{s}}) \\ &= bb' + \span{b\bf{a}' + b' \bf{a}, \bf{s}} + \span{\bf{a} \otimes \bf{a}', \bf{s} \otimes \bf{s}'} \pmod q \end{aligned}\]if $\abs{\mu\mu’} < q/2$. Then
\[\mu_\rm{mul} = \mu\mu' \approx (\Delta \cdot m) \cdot (\Delta \cdot m') = \Delta^2 \cdot mm'.\]So $mm’ \approx \Delta^{2} \cdot \mu_\rm{mul}$.
We have issues with multiplication, as we did in BGV.
 The dimension of the ciphertext has increased to $n^2$.
 The scaling factor has increased to $\Delta^2$.
 A larger scaling factor means more significant digits to calculate.
Dimension Reduction
The relinearization procedure is almost the same as in BGV relinearization.
For convenience, let $a_{i, j} = a_i a_j’$.
Relinearization Keys: for $1 \leq i, j \leq n$ and $0 \leq k < \ceil{\log q}$, perform the following.
 Sample $\bf{u}{i, j, k} \la \Z_q^{n}$ and ${} e{i, j, k} \la D_\sigma {}$.
 Compute ${} v_{i, j, k} = \span{\bf{u}{i, j, k}, \bf{s}} + 2^k \cdot s_i s_j + e{i, j, k} \pmod q {}$.
 Output ${} \bf{w}{i, j, k} = (v{i, j, k}, \bf{u}_{i, j, k}) {}$.
Linearization: given $\bf{c}\rm{mul} = (bb’, b\bf{a}’ + b’ \bf{a}, \bf{a} \otimes \bf{a}’)$, $\bf{w}{i, j, k}$ for $1 \leq i, j \leq n$ and $0 \leq k < \ceil{\log q}$, output the following.
\[\bf{c}_\rm{mul}^\ast = (b_\rm{mul}^\ast, \bf{a}_\rm{mul}^\ast) = (bb', b\bf{a}' + b'\bf{a}) + \sum_{i=1}^n \sum_{j=1}^n \sum_{k=0}^{\ceil{\log q}} a_{i, j}[k] \bf{w}_{i, j, k} \pmod q.\]
Correctness can be checked. The bounds for summations are omitted for brevity. They range from $1 \leq i, j \leq n$ and $0 \leq k < \ceil{\log q}$.
\[\begin{aligned} b_\rm{mul}^\ast + \span{\bf{a}_\rm{mul}^\ast, \bf{s}} &= bb' + \sum_{i, j, k} a_{i, j}[k] \cdot v_{i, j, k} + \span{b\bf{a}' + b'\bf{a}, \bf{s}} + \sum_{i, j, k} a_{i, j}[k] \cdot \span{\bf{u}_{i, j, k}, \bf{s}} \\ &= bb' + \span{b\bf{a}' + b'\bf{a}, \bf{s}} + \sum_{i, j, k} a_{i, j}[k] \cdot \paren{v_{i, j, k} + \span{\bf{u}_{i, j, k}, \bf{s}}} \\ &= bb' + \span{b\bf{a}' + b'\bf{a}, \bf{s}} + \sum_{i, j, k} a_{i, j}[k] \paren{2^k \cdot s_is_j + e_{i, j, k}} \\ &= bb' + \span{b\bf{a}' + b'\bf{a}, \bf{s}} + \sum_{i, j} a_{i, j}s_i s_j + \sum_{i, j, k} a_{i, j}[k] \cdot e_{i, j, k} \\ &= bb' + \span{b\bf{a}' + b'\bf{a}, \bf{s}} + \span{\bf{a} \otimes \bf{a}', \bf{s} \otimes \bf{s}} + e\conj \\ &= \mu_\rm{mul} + e\conj\pmod q. \end{aligned}\]Since
\[e\conj = \sum_{i, j, k} a_{i, j}[k] \cdot e_{i, j, k} \ll q,\]we have
\[\mu_\rm{mul}^\ast = \mu_\rm{mul} + e\conj \approx \mu\mu' \approx \Delta^2 \cdot mm'.\]Note that the proof is identical to that of BGV linearization, except for missing constant factor $2$ in the error.
Scaling Factor Reduction
In BGV, we used modulus switching for noise reduction. It was for reducing the error and preserving the message. We also use modulus switching here, but for a different purpose. The message can have small numerical errors, we just want to reduce the scaling factor. This operation is called rescaling.
Given $\bf{c} = (b, \bf{a}) \in \Z_q^{n+1}$ such that $b + \span{\bf{a}, \bf{s}} = \mu \pmod q$ and $\mu \approx \Delta^2 \cdot m$, we want to generate a new ciphertext of $m’ \approx m$ that has a scaling factor reduced to $\Delta$. This can be done by dividing the ciphertext by $\Delta$ and then rounding it appropriately.
Modulus Switching: let $\bf{c} = (b, \bf{a}) \in \Z_q^{n+1}$ be given.
 Let $q’ = \Delta \inv \cdot q$.^{2}
 Output $\bf{c}’ = \round{\Delta\inv \cdot \bf{c}} \in \Z_{q’}^{n+1}$.
Note that the modulus has been switched to $q’$. Constant multiplication and rounding is done componentwise on $\bf{c}$.
We check that $\bf{c}’$ has scaling factor $\Delta$. We know that $\mu’ = b’ + \span{\bf{a}’, \bf{s}} \pmod{q’}$.
Let $k \in \Z$ such that $b + \span{\bf{a}, \bf{s}} = \mu + kq$. By the choice of $b’$ and $\bf{a}’$, we have
\[b' = \Delta\inv \cdot b + \epsilon_0, \quad a_i' = \Delta\inv \cdot a_i + \epsilon_i\]for some $\epsilon_i$ such that $\abs{\epsilon_i} \leq 0.5$. So we have
\[\begin{aligned} \mu' &= \Delta\inv \cdot \paren{b + \sum_{i=1}^n a_i s_i} + \epsilon_0 + \sum_{i=1}^n \epsilon_i s_i \\ &= \Delta\inv \cdot (\mu + kq) + \epsilon \approx \Delta \inv \cdot (\Delta^2 \cdot m) + kq' = \Delta \cdot m \pmod{q'}, \end{aligned}\]since $\epsilon = \epsilon_0 + \sum_{i=1}^n \epsilon_i s_i$ is small.
Modulus Chain
Using modulus switching, we can set ${} q_L = \Delta^{L+1} {}$ where $L$ is the maximal level for multiplication. After each multiplication, the modulus is switched to $q_{k1} = q_k / \Delta$.
Multiplication increases the scaling factor to $\Delta^2$, and then rescaling operation reduces the scaling factor back to $\Delta$.
So we have a modulus chain,
\[\Delta^{L+1} \ra \Delta^L \ra \cdots \ra \Delta.\]When we reach $q_0 = \Delta$, we cannot perform any multiplications, so we apply bootstrapping here.
Multiplication in CKKS (Summary)
 Set up a modulus chain ${} q_k = \Delta^{k+1} {}$ for $k = 0, \dots, L$.

Given two ciphertexts $\bf{c} = (b, \bf{a}) \in \Z_{q_k}^{n+1}$ and $\bf{c}’ = (b’, \bf{a}’) \in \Z_{q_k}^{n+1}$ with modulus $q_k$ and scaling factor $\Delta$.
 (Tensor Product) $\bf{c}_\rm{mul} = \bf{c} \otimes \bf{c}’ \pmod{q_k}$.
 Now we have $n^2$ dimensions and scaling factor $\Delta^2$.
 (Relinearization)
 Back to $n$ dimensions and scaling factor $\Delta^2$.
 (Modulus Switching; Rescaling)
 Modulus is switched to $q_{k1}$ and scaling factor is back to $\Delta$.